linux-stable/fs/btrfs/ctree.h

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/* SPDX-License-Identifier: GPL-2.0 */
/*
* Copyright (C) 2007 Oracle. All rights reserved.
*/
#ifndef BTRFS_CTREE_H
#define BTRFS_CTREE_H
#include <linux/mm.h>
#include <linux/sched/signal.h>
#include <linux/highmem.h>
#include <linux/fs.h>
#include <linux/rwsem.h>
#include <linux/semaphore.h>
#include <linux/completion.h>
#include <linux/backing-dev.h>
#include <linux/wait.h>
include cleanup: Update gfp.h and slab.h includes to prepare for breaking implicit slab.h inclusion from percpu.h percpu.h is included by sched.h and module.h and thus ends up being included when building most .c files. percpu.h includes slab.h which in turn includes gfp.h making everything defined by the two files universally available and complicating inclusion dependencies. percpu.h -> slab.h dependency is about to be removed. Prepare for this change by updating users of gfp and slab facilities include those headers directly instead of assuming availability. As this conversion needs to touch large number of source files, the following script is used as the basis of conversion. http://userweb.kernel.org/~tj/misc/slabh-sweep.py The script does the followings. * Scan files for gfp and slab usages and update includes such that only the necessary includes are there. ie. if only gfp is used, gfp.h, if slab is used, slab.h. * When the script inserts a new include, it looks at the include blocks and try to put the new include such that its order conforms to its surrounding. It's put in the include block which contains core kernel includes, in the same order that the rest are ordered - alphabetical, Christmas tree, rev-Xmas-tree or at the end if there doesn't seem to be any matching order. * If the script can't find a place to put a new include (mostly because the file doesn't have fitting include block), it prints out an error message indicating which .h file needs to be added to the file. The conversion was done in the following steps. 1. The initial automatic conversion of all .c files updated slightly over 4000 files, deleting around 700 includes and adding ~480 gfp.h and ~3000 slab.h inclusions. The script emitted errors for ~400 files. 2. Each error was manually checked. Some didn't need the inclusion, some needed manual addition while adding it to implementation .h or embedding .c file was more appropriate for others. This step added inclusions to around 150 files. 3. The script was run again and the output was compared to the edits from #2 to make sure no file was left behind. 4. Several build tests were done and a couple of problems were fixed. e.g. lib/decompress_*.c used malloc/free() wrappers around slab APIs requiring slab.h to be added manually. 5. The script was run on all .h files but without automatically editing them as sprinkling gfp.h and slab.h inclusions around .h files could easily lead to inclusion dependency hell. Most gfp.h inclusion directives were ignored as stuff from gfp.h was usually wildly available and often used in preprocessor macros. Each slab.h inclusion directive was examined and added manually as necessary. 6. percpu.h was updated not to include slab.h. 7. Build test were done on the following configurations and failures were fixed. CONFIG_GCOV_KERNEL was turned off for all tests (as my distributed build env didn't work with gcov compiles) and a few more options had to be turned off depending on archs to make things build (like ipr on powerpc/64 which failed due to missing writeq). * x86 and x86_64 UP and SMP allmodconfig and a custom test config. * powerpc and powerpc64 SMP allmodconfig * sparc and sparc64 SMP allmodconfig * ia64 SMP allmodconfig * s390 SMP allmodconfig * alpha SMP allmodconfig * um on x86_64 SMP allmodconfig 8. percpu.h modifications were reverted so that it could be applied as a separate patch and serve as bisection point. Given the fact that I had only a couple of failures from tests on step 6, I'm fairly confident about the coverage of this conversion patch. If there is a breakage, it's likely to be something in one of the arch headers which should be easily discoverable easily on most builds of the specific arch. Signed-off-by: Tejun Heo <tj@kernel.org> Guess-its-ok-by: Christoph Lameter <cl@linux-foundation.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: Lee Schermerhorn <Lee.Schermerhorn@hp.com>
2010-03-24 08:04:11 +00:00
#include <linux/slab.h>
#include <linux/kobject.h>
Btrfs: add initial tracepoint support for btrfs Tracepoints can provide insight into why btrfs hits bugs and be greatly helpful for debugging, e.g dd-7822 [000] 2121.641088: btrfs_inode_request: root = 5(FS_TREE), gen = 4, ino = 256, blocks = 8, disk_i_size = 0, last_trans = 8, logged_trans = 0 dd-7822 [000] 2121.641100: btrfs_inode_new: root = 5(FS_TREE), gen = 8, ino = 257, blocks = 0, disk_i_size = 0, last_trans = 0, logged_trans = 0 btrfs-transacti-7804 [001] 2146.935420: btrfs_cow_block: root = 2(EXTENT_TREE), refs = 2, orig_buf = 29368320 (orig_level = 0), cow_buf = 29388800 (cow_level = 0) btrfs-transacti-7804 [001] 2146.935473: btrfs_cow_block: root = 1(ROOT_TREE), refs = 2, orig_buf = 29364224 (orig_level = 0), cow_buf = 29392896 (cow_level = 0) btrfs-transacti-7804 [001] 2146.972221: btrfs_transaction_commit: root = 1(ROOT_TREE), gen = 8 flush-btrfs-2-7821 [001] 2155.824210: btrfs_chunk_alloc: root = 3(CHUNK_TREE), offset = 1103101952, size = 1073741824, num_stripes = 1, sub_stripes = 0, type = DATA flush-btrfs-2-7821 [001] 2155.824241: btrfs_cow_block: root = 2(EXTENT_TREE), refs = 2, orig_buf = 29388800 (orig_level = 0), cow_buf = 29396992 (cow_level = 0) flush-btrfs-2-7821 [001] 2155.824255: btrfs_cow_block: root = 4(DEV_TREE), refs = 2, orig_buf = 29372416 (orig_level = 0), cow_buf = 29401088 (cow_level = 0) flush-btrfs-2-7821 [000] 2155.824329: btrfs_cow_block: root = 3(CHUNK_TREE), refs = 2, orig_buf = 20971520 (orig_level = 0), cow_buf = 20975616 (cow_level = 0) btrfs-endio-wri-7800 [001] 2155.898019: btrfs_cow_block: root = 5(FS_TREE), refs = 2, orig_buf = 29384704 (orig_level = 0), cow_buf = 29405184 (cow_level = 0) btrfs-endio-wri-7800 [001] 2155.898043: btrfs_cow_block: root = 7(CSUM_TREE), refs = 2, orig_buf = 29376512 (orig_level = 0), cow_buf = 29409280 (cow_level = 0) Here is what I have added: 1) ordere_extent: btrfs_ordered_extent_add btrfs_ordered_extent_remove btrfs_ordered_extent_start btrfs_ordered_extent_put These provide critical information to understand how ordered_extents are updated. 2) extent_map: btrfs_get_extent extent_map is used in both read and write cases, and it is useful for tracking how btrfs specific IO is running. 3) writepage: __extent_writepage btrfs_writepage_end_io_hook Pages are cirtical resourses and produce a lot of corner cases during writeback, so it is valuable to know how page is written to disk. 4) inode: btrfs_inode_new btrfs_inode_request btrfs_inode_evict These can show where and when a inode is created, when a inode is evicted. 5) sync: btrfs_sync_file btrfs_sync_fs These show sync arguments. 6) transaction: btrfs_transaction_commit In transaction based filesystem, it will be useful to know the generation and who does commit. 7) back reference and cow: btrfs_delayed_tree_ref btrfs_delayed_data_ref btrfs_delayed_ref_head btrfs_cow_block Btrfs natively supports back references, these tracepoints are helpful on understanding btrfs's COW mechanism. 8) chunk: btrfs_chunk_alloc btrfs_chunk_free Chunk is a link between physical offset and logical offset, and stands for space infomation in btrfs, and these are helpful on tracing space things. 9) reserved_extent: btrfs_reserved_extent_alloc btrfs_reserved_extent_free These can show how btrfs uses its space. Signed-off-by: Liu Bo <liubo2009@cn.fujitsu.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2011-03-24 11:18:59 +00:00
#include <trace/events/btrfs.h>
#include <asm/kmap_types.h>
#include <linux/pagemap.h>
#include <linux/btrfs.h>
#include <linux/btrfs_tree.h>
Btrfs: reclaim the reserved metadata space at background Before applying this patch, the task had to reclaim the metadata space by itself if the metadata space was not enough. And When the task started the space reclamation, all the other tasks which wanted to reserve the metadata space were blocked. At some cases, they would be blocked for a long time, it made the performance fluctuate wildly. So we introduce the background metadata space reclamation, when the space is about to be exhausted, we insert a reclaim work into the workqueue, the worker of the workqueue helps us to reclaim the reserved space at the background. By this way, the tasks needn't reclaim the space by themselves at most cases, and even if the tasks have to reclaim the space or are blocked for the space reclamation, they will get enough space more quickly. Here is my test result(Tested by compilebench): Memory: 2GB CPU: 2Cores * 1CPU Partition: 40GB(SSD) Test command: # compilebench -D <mnt> -m Without this patch: intial create total runs 30 avg 54.36 MB/s (user 0.52s sys 2.44s) compile total runs 30 avg 123.72 MB/s (user 0.13s sys 1.17s) read compiled tree total runs 3 avg 81.15 MB/s (user 0.74s sys 4.89s) delete compiled tree total runs 30 avg 5.32 seconds (user 0.35s sys 4.37s) With this patch: intial create total runs 30 avg 59.80 MB/s (user 0.52s sys 2.53s) compile total runs 30 avg 151.44 MB/s (user 0.13s sys 1.11s) read compiled tree total runs 3 avg 83.25 MB/s (user 0.76s sys 4.91s) delete compiled tree total runs 30 avg 5.29 seconds (user 0.34s sys 4.34s) Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-05-14 00:29:04 +00:00
#include <linux/workqueue.h>
#include <linux/security.h>
#include <linux/sizes.h>
#include <linux/dynamic_debug.h>
#include <linux/refcount.h>
btrfs: Remove custom crc32c init code The custom crc32 init code was introduced in 14a958e678cd ("Btrfs: fix btrfs boot when compiled as built-in") to enable using btrfs as a built-in. However, later as pointed out by 60efa5eb2e88 ("Btrfs: use late_initcall instead of module_init") this wasn't enough and finally btrfs was switched to late_initcall which comes after the generic crc32c implementation is initiliased. The latter commit superseeded the former. Now that we don't have to maintain our own code let's just remove it and switch to using the generic implementation. Despite touching a lot of files the patch is really simple. Here is the gist of the changes: 1. Select LIBCRC32C rather than the low-level modules. 2. s/btrfs_crc32c/crc32c/g 3. replace hash.h with linux/crc32c.h 4. Move the btrfs namehash funcs to ctree.h and change the tree accordingly. I've tested this with btrfs being both a module and a built-in and xfstest doesn't complain. Does seem to fix the longstanding problem of not automatically selectiong the crc32c module when btrfs is used. Possibly there is a workaround in dracut. The modinfo confirms that now all the module dependencies are there: before: depends: zstd_compress,zstd_decompress,raid6_pq,xor,zlib_deflate after: depends: libcrc32c,zstd_compress,zstd_decompress,raid6_pq,xor,zlib_deflate Signed-off-by: Nikolay Borisov <nborisov@suse.com> Reviewed-by: David Sterba <dsterba@suse.com> [ add more info to changelog from mails ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-01-08 09:45:05 +00:00
#include <linux/crc32c.h>
#include "extent_io.h"
#include "extent_map.h"
#include "async-thread.h"
struct btrfs_trans_handle;
struct btrfs_transaction;
struct btrfs_pending_snapshot;
btrfs: handle delayed ref head accounting cleanup in abort We weren't doing any of the accounting cleanup when we aborted transactions. Fix this by making cleanup_ref_head_accounting global and calling it from the abort code, this fixes the issue where our accounting was all wrong after the fs aborts. The test generic/475 on a 2G VM can trigger the problems eg.: [ 8502.136957] WARNING: CPU: 0 PID: 11064 at fs/btrfs/extent-tree.c:5986 btrfs_free_block_grou +ps+0x3dc/0x410 [btrfs] [ 8502.148372] CPU: 0 PID: 11064 Comm: umount Not tainted 5.0.0-rc1-default+ #394 [ 8502.150807] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS rel-1.11.2-0-gf9626 +cc-prebuilt.qemu-project.org 04/01/2014 [ 8502.154317] RIP: 0010:btrfs_free_block_groups+0x3dc/0x410 [btrfs] [ 8502.160623] RSP: 0018:ffffb1ab84b93de8 EFLAGS: 00010206 [ 8502.161906] RAX: 0000000001000000 RBX: ffff9f34b1756400 RCX: 0000000000000000 [ 8502.163448] RDX: 0000000000000002 RSI: 0000000000000001 RDI: ffff9f34b1755400 [ 8502.164906] RBP: ffff9f34b7e8c000 R08: 0000000000000001 R09: 0000000000000000 [ 8502.166716] R10: 0000000000000000 R11: 0000000000000001 R12: ffff9f34b7e8c108 [ 8502.168498] R13: ffff9f34b7e8c158 R14: 0000000000000000 R15: dead000000000100 [ 8502.170296] FS: 00007fb1cf15ffc0(0000) GS:ffff9f34bd400000(0000) knlGS:0000000000000000 [ 8502.172439] CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033 [ 8502.173669] CR2: 00007fb1ced507b0 CR3: 000000002f7a6000 CR4: 00000000000006f0 [ 8502.175094] Call Trace: [ 8502.175759] close_ctree+0x17f/0x350 [btrfs] [ 8502.176721] generic_shutdown_super+0x64/0x100 [ 8502.177702] kill_anon_super+0x14/0x30 [ 8502.178607] btrfs_kill_super+0x12/0xa0 [btrfs] [ 8502.179602] deactivate_locked_super+0x29/0x60 [ 8502.180595] cleanup_mnt+0x3b/0x70 [ 8502.181406] task_work_run+0x98/0xc0 [ 8502.182255] exit_to_usermode_loop+0x83/0x90 [ 8502.183113] do_syscall_64+0x15b/0x180 [ 8502.183919] entry_SYSCALL_64_after_hwframe+0x49/0xbe Corresponding to release_global_block_rsv() { ... WARN_ON(fs_info->delayed_refs_rsv.reserved > 0); CC: stable@vger.kernel.org Signed-off-by: Josef Bacik <josef@toxicpanda.com> [ add log dump ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-11-21 19:05:41 +00:00
struct btrfs_delayed_ref_root;
extern struct kmem_cache *btrfs_trans_handle_cachep;
extern struct kmem_cache *btrfs_bit_radix_cachep;
extern struct kmem_cache *btrfs_path_cachep;
extern struct kmem_cache *btrfs_free_space_cachep;
struct btrfs_ordered_sum;
#define BTRFS_MAGIC 0x4D5F53665248425FULL /* ascii _BHRfS_M, no null */
#define BTRFS_MAX_MIRRORS 3
#define BTRFS_MAX_LEVEL 8
#define BTRFS_OLDEST_GENERATION 0ULL
/*
* the max metadata block size. This limit is somewhat artificial,
* but the memmove costs go through the roof for larger blocks.
*/
#define BTRFS_MAX_METADATA_BLOCKSIZE 65536
/*
* we can actually store much bigger names, but lets not confuse the rest
* of linux
*/
#define BTRFS_NAME_LEN 255
/*
* Theoretical limit is larger, but we keep this down to a sane
* value. That should limit greatly the possibility of collisions on
* inode ref items.
*/
#define BTRFS_LINK_MAX 65535U
/* four bytes for CRC32 */
static const int btrfs_csum_sizes[] = { 4 };
#define BTRFS_EMPTY_DIR_SIZE 0
/* ioprio of readahead is set to idle */
#define BTRFS_IOPRIO_READA (IOPRIO_PRIO_VALUE(IOPRIO_CLASS_IDLE, 0))
#define BTRFS_DIRTY_METADATA_THRESH SZ_32M
btrfs: use customized batch size for total_bytes_pinned In commit b150a4f10d878 ("Btrfs: use a percpu to keep track of possibly pinned bytes") we use total_bytes_pinned to track how many bytes we are going to free in this transaction. When we are close to ENOSPC, we check it and know if we can make the allocation by commit the current transaction. For every data/metadata extent we are going to free, we add total_bytes_pinned in btrfs_free_extent() and btrfs_free_tree_block(), and release it in unpin_extent_range() when we finish the transaction. So this is a variable we frequently update but rarely read - just the suitable use of percpu_counter. But in previous commit we update total_bytes_pinned by default 32 batch size, making every update essentially a spin lock protected update. Since every spin lock/unlock operation involves syncing a globally used variable and some kind of barrier in a SMP system, this is more expensive than using total_bytes_pinned as a simple atomic64_t. So fix this by using a customized batch size. Since we only read total_bytes_pinned when we are close to ENOSPC and fail to allocate new chunk, we can use a really large batch size and have nearly no penalty in most cases. [Test] We tested the patch on a 4-cores x86 machine: 1. fallocate a 16GiB size test file 2. take snapshot (so all following writes will be COW) 3. run a 180 sec, 4 jobs, 4K random write fio on test file We also added a temporary lockdep class on percpu_counter's spin lock used by total_bytes_pinned to track it by lock_stat. [Results] unpatched: lock_stat version 0.4 ----------------------------------------------------------------------- class name con-bounces contentions waittime-min waittime-max waittime-total waittime-avg acq-bounces acquisitions holdtime-min holdtime-max holdtime-total holdtime-avg total_bytes_pinned_percpu: 82 82 0.21 0.61 29.46 0.36 298340 635973 0.09 11.01 173476.25 0.27 patched: lock_stat version 0.4 ----------------------------------------------------------------------- class name con-bounces contentions waittime-min waittime-max waittime-total waittime-avg acq-bounces acquisitions holdtime-min holdtime-max holdtime-total holdtime-avg total_bytes_pinned_percpu: 1 1 0.62 0.62 0.62 0.62 13601 31542 0.14 9.61 11016.90 0.35 [Analysis] Since the spin lock only protects a single in-memory variable, the contentions (number of lock acquisitions that had to wait) in both unpatched and patched version are low. But when we see acquisitions and acq-bounces, we get much lower counts in patched version. Here the most important metric is acq-bounces. It means how many times the lock gets transferred between different cpus, so the patch can really reduce cacheline bouncing of spin lock (also the global counter of percpu_counter) in a SMP system. Fixes: b150a4f10d878 ("Btrfs: use a percpu to keep track of possibly pinned bytes") Signed-off-by: Ethan Lien <ethanlien@synology.com> Reviewed-by: Nikolay Borisov <nborisov@suse.com> Reviewed-by: David Sterba <dsterba@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2018-07-13 08:50:42 +00:00
/*
* Use large batch size to reduce overhead of metadata updates. On the reader
* side, we only read it when we are close to ENOSPC and the read overhead is
* mostly related to the number of CPUs, so it is OK to use arbitrary large
* value here.
*/
#define BTRFS_TOTAL_BYTES_PINNED_BATCH SZ_128M
#define BTRFS_MAX_EXTENT_SIZE SZ_128M
btrfs: Remove custom crc32c init code The custom crc32 init code was introduced in 14a958e678cd ("Btrfs: fix btrfs boot when compiled as built-in") to enable using btrfs as a built-in. However, later as pointed out by 60efa5eb2e88 ("Btrfs: use late_initcall instead of module_init") this wasn't enough and finally btrfs was switched to late_initcall which comes after the generic crc32c implementation is initiliased. The latter commit superseeded the former. Now that we don't have to maintain our own code let's just remove it and switch to using the generic implementation. Despite touching a lot of files the patch is really simple. Here is the gist of the changes: 1. Select LIBCRC32C rather than the low-level modules. 2. s/btrfs_crc32c/crc32c/g 3. replace hash.h with linux/crc32c.h 4. Move the btrfs namehash funcs to ctree.h and change the tree accordingly. I've tested this with btrfs being both a module and a built-in and xfstest doesn't complain. Does seem to fix the longstanding problem of not automatically selectiong the crc32c module when btrfs is used. Possibly there is a workaround in dracut. The modinfo confirms that now all the module dependencies are there: before: depends: zstd_compress,zstd_decompress,raid6_pq,xor,zlib_deflate after: depends: libcrc32c,zstd_compress,zstd_decompress,raid6_pq,xor,zlib_deflate Signed-off-by: Nikolay Borisov <nborisov@suse.com> Reviewed-by: David Sterba <dsterba@suse.com> [ add more info to changelog from mails ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-01-08 09:45:05 +00:00
/*
* Count how many BTRFS_MAX_EXTENT_SIZE cover the @size
*/
static inline u32 count_max_extents(u64 size)
{
return div_u64(size + BTRFS_MAX_EXTENT_SIZE - 1, BTRFS_MAX_EXTENT_SIZE);
}
struct btrfs_mapping_tree {
struct extent_map_tree map_tree;
};
static inline unsigned long btrfs_chunk_item_size(int num_stripes)
{
BUG_ON(num_stripes == 0);
return sizeof(struct btrfs_chunk) +
sizeof(struct btrfs_stripe) * (num_stripes - 1);
}
/*
* Runtime (in-memory) states of filesystem
*/
enum {
/* Global indicator of serious filesystem errors */
BTRFS_FS_STATE_ERROR,
/*
* Filesystem is being remounted, allow to skip some operations, like
* defrag
*/
BTRFS_FS_STATE_REMOUNTING,
/* Track if a transaction abort has been reported on this filesystem */
BTRFS_FS_STATE_TRANS_ABORTED,
/*
* Bio operations should be blocked on this filesystem because a source
* or target device is being destroyed as part of a device replace
*/
BTRFS_FS_STATE_DEV_REPLACING,
/* The btrfs_fs_info created for self-tests */
BTRFS_FS_STATE_DUMMY_FS_INFO,
};
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
#define BTRFS_BACKREF_REV_MAX 256
#define BTRFS_BACKREF_REV_SHIFT 56
#define BTRFS_BACKREF_REV_MASK (((u64)BTRFS_BACKREF_REV_MAX - 1) << \
BTRFS_BACKREF_REV_SHIFT)
#define BTRFS_OLD_BACKREF_REV 0
#define BTRFS_MIXED_BACKREF_REV 1
/*
* every tree block (leaf or node) starts with this header.
*/
struct btrfs_header {
/* these first four must match the super block */
u8 csum[BTRFS_CSUM_SIZE];
u8 fsid[BTRFS_FSID_SIZE]; /* FS specific uuid */
__le64 bytenr; /* which block this node is supposed to live in */
__le64 flags;
/* allowed to be different from the super from here on down */
u8 chunk_tree_uuid[BTRFS_UUID_SIZE];
__le64 generation;
__le64 owner;
__le32 nritems;
u8 level;
} __attribute__ ((__packed__));
/*
* this is a very generous portion of the super block, giving us
* room to translate 14 chunks with 3 stripes each.
*/
#define BTRFS_SYSTEM_CHUNK_ARRAY_SIZE 2048
/*
* just in case we somehow lose the roots and are not able to mount,
* we store an array of the roots from previous transactions
* in the super.
*/
#define BTRFS_NUM_BACKUP_ROOTS 4
struct btrfs_root_backup {
__le64 tree_root;
__le64 tree_root_gen;
__le64 chunk_root;
__le64 chunk_root_gen;
__le64 extent_root;
__le64 extent_root_gen;
__le64 fs_root;
__le64 fs_root_gen;
__le64 dev_root;
__le64 dev_root_gen;
__le64 csum_root;
__le64 csum_root_gen;
__le64 total_bytes;
__le64 bytes_used;
__le64 num_devices;
/* future */
__le64 unused_64[4];
u8 tree_root_level;
u8 chunk_root_level;
u8 extent_root_level;
u8 fs_root_level;
u8 dev_root_level;
u8 csum_root_level;
/* future and to align */
u8 unused_8[10];
} __attribute__ ((__packed__));
/*
* the super block basically lists the main trees of the FS
* it currently lacks any block count etc etc
*/
struct btrfs_super_block {
/* the first 4 fields must match struct btrfs_header */
btrfs: Introduce support for FSID change without metadata rewrite This field is going to be used when the user wants to change the UUID of the filesystem without having to rewrite all metadata blocks. This field adds another level of indirection such that when the FSID is changed what really happens is the current UUID (the one with which the fs was created) is copied to the 'metadata_uuid' field in the superblock as well as a new incompat flag is set METADATA_UUID. When the kernel detects this flag is set it knows that the superblock in fact has 2 UUIDs: 1. Is the UUID which is user-visible, currently known as FSID. 2. Metadata UUID - this is the UUID which is stamped into all on-disk datastructures belonging to this file system. When the new incompat flag is present device scanning checks whether both fsid/metadata_uuid of the scanned device match any of the registered filesystems. When the flag is not set then both UUIDs are equal and only the FSID is retained on disk, metadata_uuid is set only in-memory during mount. Additionally a new metadata_uuid field is also added to the fs_info struct. It's initialised either with the FSID in case METADATA_UUID incompat flag is not set or with the metdata_uuid of the superblock otherwise. This commit introduces the new fields as well as the new incompat flag and switches all users of the fsid to the new logic. Signed-off-by: Nikolay Borisov <nborisov@suse.com> Reviewed-by: David Sterba <dsterba@suse.com> [ minor updates in comments ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-10-30 14:43:23 +00:00
u8 csum[BTRFS_CSUM_SIZE];
/* FS specific UUID, visible to user */
u8 fsid[BTRFS_FSID_SIZE];
__le64 bytenr; /* this block number */
__le64 flags;
/* allowed to be different from the btrfs_header from here own down */
__le64 magic;
__le64 generation;
__le64 root;
__le64 chunk_root;
__le64 log_root;
/* this will help find the new super based on the log root */
__le64 log_root_transid;
__le64 total_bytes;
__le64 bytes_used;
__le64 root_dir_objectid;
__le64 num_devices;
__le32 sectorsize;
__le32 nodesize;
__le32 __unused_leafsize;
__le32 stripesize;
__le32 sys_chunk_array_size;
__le64 chunk_root_generation;
__le64 compat_flags;
__le64 compat_ro_flags;
__le64 incompat_flags;
__le16 csum_type;
u8 root_level;
u8 chunk_root_level;
u8 log_root_level;
struct btrfs_dev_item dev_item;
char label[BTRFS_LABEL_SIZE];
__le64 cache_generation;
__le64 uuid_tree_generation;
btrfs: Introduce support for FSID change without metadata rewrite This field is going to be used when the user wants to change the UUID of the filesystem without having to rewrite all metadata blocks. This field adds another level of indirection such that when the FSID is changed what really happens is the current UUID (the one with which the fs was created) is copied to the 'metadata_uuid' field in the superblock as well as a new incompat flag is set METADATA_UUID. When the kernel detects this flag is set it knows that the superblock in fact has 2 UUIDs: 1. Is the UUID which is user-visible, currently known as FSID. 2. Metadata UUID - this is the UUID which is stamped into all on-disk datastructures belonging to this file system. When the new incompat flag is present device scanning checks whether both fsid/metadata_uuid of the scanned device match any of the registered filesystems. When the flag is not set then both UUIDs are equal and only the FSID is retained on disk, metadata_uuid is set only in-memory during mount. Additionally a new metadata_uuid field is also added to the fs_info struct. It's initialised either with the FSID in case METADATA_UUID incompat flag is not set or with the metdata_uuid of the superblock otherwise. This commit introduces the new fields as well as the new incompat flag and switches all users of the fsid to the new logic. Signed-off-by: Nikolay Borisov <nborisov@suse.com> Reviewed-by: David Sterba <dsterba@suse.com> [ minor updates in comments ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-10-30 14:43:23 +00:00
/* the UUID written into btree blocks */
u8 metadata_uuid[BTRFS_FSID_SIZE];
/* future expansion */
btrfs: Introduce support for FSID change without metadata rewrite This field is going to be used when the user wants to change the UUID of the filesystem without having to rewrite all metadata blocks. This field adds another level of indirection such that when the FSID is changed what really happens is the current UUID (the one with which the fs was created) is copied to the 'metadata_uuid' field in the superblock as well as a new incompat flag is set METADATA_UUID. When the kernel detects this flag is set it knows that the superblock in fact has 2 UUIDs: 1. Is the UUID which is user-visible, currently known as FSID. 2. Metadata UUID - this is the UUID which is stamped into all on-disk datastructures belonging to this file system. When the new incompat flag is present device scanning checks whether both fsid/metadata_uuid of the scanned device match any of the registered filesystems. When the flag is not set then both UUIDs are equal and only the FSID is retained on disk, metadata_uuid is set only in-memory during mount. Additionally a new metadata_uuid field is also added to the fs_info struct. It's initialised either with the FSID in case METADATA_UUID incompat flag is not set or with the metdata_uuid of the superblock otherwise. This commit introduces the new fields as well as the new incompat flag and switches all users of the fsid to the new logic. Signed-off-by: Nikolay Borisov <nborisov@suse.com> Reviewed-by: David Sterba <dsterba@suse.com> [ minor updates in comments ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-10-30 14:43:23 +00:00
__le64 reserved[28];
u8 sys_chunk_array[BTRFS_SYSTEM_CHUNK_ARRAY_SIZE];
struct btrfs_root_backup super_roots[BTRFS_NUM_BACKUP_ROOTS];
} __attribute__ ((__packed__));
/*
* Compat flags that we support. If any incompat flags are set other than the
* ones specified below then we will fail to mount
*/
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
#define BTRFS_FEATURE_COMPAT_SUPP 0ULL
#define BTRFS_FEATURE_COMPAT_SAFE_SET 0ULL
#define BTRFS_FEATURE_COMPAT_SAFE_CLEAR 0ULL
#define BTRFS_FEATURE_COMPAT_RO_SUPP \
(BTRFS_FEATURE_COMPAT_RO_FREE_SPACE_TREE | \
BTRFS_FEATURE_COMPAT_RO_FREE_SPACE_TREE_VALID)
#define BTRFS_FEATURE_COMPAT_RO_SAFE_SET 0ULL
#define BTRFS_FEATURE_COMPAT_RO_SAFE_CLEAR 0ULL
#define BTRFS_FEATURE_INCOMPAT_SUPP \
(BTRFS_FEATURE_INCOMPAT_MIXED_BACKREF | \
BTRFS_FEATURE_INCOMPAT_DEFAULT_SUBVOL | \
BTRFS_FEATURE_INCOMPAT_MIXED_GROUPS | \
BTRFS_FEATURE_INCOMPAT_BIG_METADATA | \
BTRFS_FEATURE_INCOMPAT_COMPRESS_LZO | \
btrfs: Add zstd support Add zstd compression and decompression support to BtrFS. zstd at its fastest level compresses almost as well as zlib, while offering much faster compression and decompression, approaching lzo speeds. I benchmarked btrfs with zstd compression against no compression, lzo compression, and zlib compression. I benchmarked two scenarios. Copying a set of files to btrfs, and then reading the files. Copying a tarball to btrfs, extracting it to btrfs, and then reading the extracted files. After every operation, I call `sync` and include the sync time. Between every pair of operations I unmount and remount the filesystem to avoid caching. The benchmark files can be found in the upstream zstd source repository under `contrib/linux-kernel/{btrfs-benchmark.sh,btrfs-extract-benchmark.sh}` [1] [2]. I ran the benchmarks on a Ubuntu 14.04 VM with 2 cores and 4 GiB of RAM. The VM is running on a MacBook Pro with a 3.1 GHz Intel Core i7 processor, 16 GB of RAM, and a SSD. The first compression benchmark is copying 10 copies of the unzipped Silesia corpus [3] into a BtrFS filesystem mounted with `-o compress-force=Method`. The decompression benchmark times how long it takes to `tar` all 10 copies into `/dev/null`. The compression ratio is measured by comparing the output of `df` and `du`. See the benchmark file [1] for details. I benchmarked multiple zstd compression levels, although the patch uses zstd level 1. | Method | Ratio | Compression MB/s | Decompression speed | |---------|-------|------------------|---------------------| | None | 0.99 | 504 | 686 | | lzo | 1.66 | 398 | 442 | | zlib | 2.58 | 65 | 241 | | zstd 1 | 2.57 | 260 | 383 | | zstd 3 | 2.71 | 174 | 408 | | zstd 6 | 2.87 | 70 | 398 | | zstd 9 | 2.92 | 43 | 406 | | zstd 12 | 2.93 | 21 | 408 | | zstd 15 | 3.01 | 11 | 354 | The next benchmark first copies `linux-4.11.6.tar` [4] to btrfs. Then it measures the compression ratio, extracts the tar, and deletes the tar. Then it measures the compression ratio again, and `tar`s the extracted files into `/dev/null`. See the benchmark file [2] for details. | Method | Tar Ratio | Extract Ratio | Copy (s) | Extract (s)| Read (s) | |--------|-----------|---------------|----------|------------|----------| | None | 0.97 | 0.78 | 0.981 | 5.501 | 8.807 | | lzo | 2.06 | 1.38 | 1.631 | 8.458 | 8.585 | | zlib | 3.40 | 1.86 | 7.750 | 21.544 | 11.744 | | zstd 1 | 3.57 | 1.85 | 2.579 | 11.479 | 9.389 | [1] https://github.com/facebook/zstd/blob/dev/contrib/linux-kernel/btrfs-benchmark.sh [2] https://github.com/facebook/zstd/blob/dev/contrib/linux-kernel/btrfs-extract-benchmark.sh [3] http://sun.aei.polsl.pl/~sdeor/index.php?page=silesia [4] https://cdn.kernel.org/pub/linux/kernel/v4.x/linux-4.11.6.tar.xz zstd source repository: https://github.com/facebook/zstd Signed-off-by: Nick Terrell <terrelln@fb.com> Signed-off-by: Chris Mason <clm@fb.com>
2017-08-10 02:39:02 +00:00
BTRFS_FEATURE_INCOMPAT_COMPRESS_ZSTD | \
BTRFS_FEATURE_INCOMPAT_RAID56 | \
BTRFS_FEATURE_INCOMPAT_EXTENDED_IREF | \
BTRFS_FEATURE_INCOMPAT_SKINNY_METADATA | \
btrfs: Introduce support for FSID change without metadata rewrite This field is going to be used when the user wants to change the UUID of the filesystem without having to rewrite all metadata blocks. This field adds another level of indirection such that when the FSID is changed what really happens is the current UUID (the one with which the fs was created) is copied to the 'metadata_uuid' field in the superblock as well as a new incompat flag is set METADATA_UUID. When the kernel detects this flag is set it knows that the superblock in fact has 2 UUIDs: 1. Is the UUID which is user-visible, currently known as FSID. 2. Metadata UUID - this is the UUID which is stamped into all on-disk datastructures belonging to this file system. When the new incompat flag is present device scanning checks whether both fsid/metadata_uuid of the scanned device match any of the registered filesystems. When the flag is not set then both UUIDs are equal and only the FSID is retained on disk, metadata_uuid is set only in-memory during mount. Additionally a new metadata_uuid field is also added to the fs_info struct. It's initialised either with the FSID in case METADATA_UUID incompat flag is not set or with the metdata_uuid of the superblock otherwise. This commit introduces the new fields as well as the new incompat flag and switches all users of the fsid to the new logic. Signed-off-by: Nikolay Borisov <nborisov@suse.com> Reviewed-by: David Sterba <dsterba@suse.com> [ minor updates in comments ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-10-30 14:43:23 +00:00
BTRFS_FEATURE_INCOMPAT_NO_HOLES | \
BTRFS_FEATURE_INCOMPAT_METADATA_UUID)
#define BTRFS_FEATURE_INCOMPAT_SAFE_SET \
(BTRFS_FEATURE_INCOMPAT_EXTENDED_IREF)
#define BTRFS_FEATURE_INCOMPAT_SAFE_CLEAR 0ULL
/*
* A leaf is full of items. offset and size tell us where to find
* the item in the leaf (relative to the start of the data area)
*/
struct btrfs_item {
struct btrfs_disk_key key;
__le32 offset;
__le32 size;
} __attribute__ ((__packed__));
/*
* leaves have an item area and a data area:
* [item0, item1....itemN] [free space] [dataN...data1, data0]
*
* The data is separate from the items to get the keys closer together
* during searches.
*/
struct btrfs_leaf {
struct btrfs_header header;
struct btrfs_item items[];
} __attribute__ ((__packed__));
/*
* all non-leaf blocks are nodes, they hold only keys and pointers to
* other blocks
*/
struct btrfs_key_ptr {
struct btrfs_disk_key key;
__le64 blockptr;
__le64 generation;
} __attribute__ ((__packed__));
struct btrfs_node {
struct btrfs_header header;
struct btrfs_key_ptr ptrs[];
} __attribute__ ((__packed__));
/*
* btrfs_paths remember the path taken from the root down to the leaf.
* level 0 is always the leaf, and nodes[1...BTRFS_MAX_LEVEL] will point
* to any other levels that are present.
*
* The slots array records the index of the item or block pointer
* used while walking the tree.
*/
enum { READA_NONE, READA_BACK, READA_FORWARD };
struct btrfs_path {
struct extent_buffer *nodes[BTRFS_MAX_LEVEL];
int slots[BTRFS_MAX_LEVEL];
/* if there is real range locking, this locks field will change */
u8 locks[BTRFS_MAX_LEVEL];
u8 reada;
/* keep some upper locks as we walk down */
u8 lowest_level;
/*
* set by btrfs_split_item, tells search_slot to keep all locks
* and to force calls to keep space in the nodes
*/
unsigned int search_for_split:1;
unsigned int keep_locks:1;
unsigned int skip_locking:1;
unsigned int leave_spinning:1;
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
unsigned int search_commit_root:1;
unsigned int need_commit_sem:1;
unsigned int skip_release_on_error:1;
};
#define BTRFS_MAX_EXTENT_ITEM_SIZE(r) ((BTRFS_LEAF_DATA_SIZE(r->fs_info) >> 4) - \
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
sizeof(struct btrfs_item))
struct btrfs_dev_replace {
u64 replace_state; /* see #define above */
time64_t time_started; /* seconds since 1-Jan-1970 */
time64_t time_stopped; /* seconds since 1-Jan-1970 */
atomic64_t num_write_errors;
atomic64_t num_uncorrectable_read_errors;
u64 cursor_left;
u64 committed_cursor_left;
u64 cursor_left_last_write_of_item;
u64 cursor_right;
u64 cont_reading_from_srcdev_mode; /* see #define above */
int is_valid;
int item_needs_writeback;
struct btrfs_device *srcdev;
struct btrfs_device *tgtdev;
struct mutex lock_finishing_cancel_unmount;
struct rw_semaphore rwsem;
struct btrfs_scrub_progress scrub_progress;
struct percpu_counter bio_counter;
wait_queue_head_t replace_wait;
};
/* For raid type sysfs entries */
struct raid_kobject {
u64 flags;
struct kobject kobj;
struct list_head list;
};
struct btrfs_space_info {
spinlock_t lock;
u64 total_bytes; /* total bytes in the space,
this doesn't take mirrors into account */
u64 bytes_used; /* total bytes used,
this doesn't take mirrors into account */
u64 bytes_pinned; /* total bytes pinned, will be freed when the
transaction finishes */
u64 bytes_reserved; /* total bytes the allocator has reserved for
current allocations */
u64 bytes_may_use; /* number of bytes that may be used for
Btrfs: proper -ENOSPC handling At the start of a transaction we do a btrfs_reserve_metadata_space() and specify how many items we plan on modifying. Then once we've done our modifications and such, just call btrfs_unreserve_metadata_space() for the same number of items we reserved. For keeping track of metadata needed for data I've had to add an extent_io op for when we merge extents. This lets us track space properly when we are doing sequential writes, so we don't end up reserving way more metadata space than what we need. The only place where the metadata space accounting is not done is in the relocation code. This is because Yan is going to be reworking that code in the near future, so running btrfs-vol -b could still possibly result in a ENOSPC related panic. This patch also turns off the metadata_ratio stuff in order to allow users to more efficiently use their disk space. This patch makes it so we track how much metadata we need for an inode's delayed allocation extents by tracking how many extents are currently waiting for allocation. It introduces two new callbacks for the extent_io tree's, merge_extent_hook and split_extent_hook. These help us keep track of when we merge delalloc extents together and split them up. Reservations are handled prior to any actually dirty'ing occurs, and then we unreserve after we dirty. btrfs_unreserve_metadata_for_delalloc() will make the appropriate unreservations as needed based on the number of reservations we currently have and the number of extents we currently have. Doing the reservation outside of doing any of the actual dirty'ing lets us do things like filemap_flush() the inode to try and force delalloc to happen, or as a last resort actually start allocation on all delalloc inodes in the fs. This has survived dbench, fs_mark and an fsx torture test. Signed-off-by: Josef Bacik <jbacik@redhat.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-09-11 20:12:44 +00:00
delalloc/allocations */
u64 bytes_readonly; /* total bytes that are read only */
u64 max_extent_size; /* This will hold the maximum extent size of
the space info if we had an ENOSPC in the
allocator. */
unsigned int full:1; /* indicates that we cannot allocate any more
chunks for this space */
unsigned int chunk_alloc:1; /* set if we are allocating a chunk */
unsigned int flush:1; /* set if we are trying to make space */
unsigned int force_alloc; /* set if we need to force a chunk
alloc for this space */
u64 disk_used; /* total bytes used on disk */
u64 disk_total; /* total bytes on disk, takes mirrors into
account */
u64 flags;
/*
* bytes_pinned is kept in line with what is actually pinned, as in
* we've called update_block_group and dropped the bytes_used counter
* and increased the bytes_pinned counter. However this means that
* bytes_pinned does not reflect the bytes that will be pinned once the
* delayed refs are flushed, so this counter is inc'ed every time we
* call btrfs_free_extent so it is a realtime count of what will be
* freed once the transaction is committed. It will be zeroed every
* time the transaction commits.
*/
struct percpu_counter total_bytes_pinned;
struct list_head list;
/* Protected by the spinlock 'lock'. */
struct list_head ro_bgs;
struct list_head priority_tickets;
struct list_head tickets;
/*
* tickets_id just indicates the next ticket will be handled, so note
* it's not stored per ticket.
*/
u64 tickets_id;
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
struct rw_semaphore groups_sem;
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
/* for block groups in our same type */
struct list_head block_groups[BTRFS_NR_RAID_TYPES];
wait_queue_head_t wait;
struct kobject kobj;
struct kobject *block_group_kobjs[BTRFS_NR_RAID_TYPES];
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
};
/*
* Types of block reserves
*/
enum {
BTRFS_BLOCK_RSV_GLOBAL,
BTRFS_BLOCK_RSV_DELALLOC,
BTRFS_BLOCK_RSV_TRANS,
BTRFS_BLOCK_RSV_CHUNK,
BTRFS_BLOCK_RSV_DELOPS,
btrfs: introduce delayed_refs_rsv Traditionally we've had voodoo in btrfs to account for the space that delayed refs may take up by having a global_block_rsv. This works most of the time, except when it doesn't. We've had issues reported and seen in production where sometimes the global reserve is exhausted during transaction commit before we can run all of our delayed refs, resulting in an aborted transaction. Because of this voodoo we have equally dubious flushing semantics around throttling delayed refs which we often get wrong. So instead give them their own block_rsv. This way we can always know exactly how much outstanding space we need for delayed refs. This allows us to make sure we are constantly filling that reservation up with space, and allows us to put more precise pressure on the enospc system. Instead of doing math to see if its a good time to throttle, the normal enospc code will be invoked if we have a lot of delayed refs pending, and they will be run via the normal flushing mechanism. For now the delayed_refs_rsv will hold the reservations for the delayed refs, the block group updates, and deleting csums. We could have a separate rsv for the block group updates, but the csum deletion stuff is still handled via the delayed_refs so that will stay there. Historical background: The global reserve has grown to cover everything we don't reserve space explicitly for, and we've grown a lot of weird ad-hoc heuristics to know if we're running short on space and when it's time to force a commit. A failure rate of 20-40 file systems when we run hundreds of thousands of them isn't super high, but cleaning up this code will make things less ugly and more predictible. Thus the delayed refs rsv. We always know how many delayed refs we have outstanding, and although running them generates more we can use the global reserve for that spill over, which fits better into it's desired use than a full blown reservation. This first approach is to simply take how many times we're reserving space for and multiply that by 2 in order to save enough space for the delayed refs that could be generated. This is a niave approach and will probably evolve, but for now it works. Signed-off-by: Josef Bacik <jbacik@fb.com> Reviewed-by: David Sterba <dsterba@suse.com> # high-level review [ added background notes from the cover letter ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-12-03 15:20:33 +00:00
BTRFS_BLOCK_RSV_DELREFS,
BTRFS_BLOCK_RSV_EMPTY,
BTRFS_BLOCK_RSV_TEMP,
};
struct btrfs_block_rsv {
u64 size;
u64 reserved;
struct btrfs_space_info *space_info;
spinlock_t lock;
unsigned short full;
unsigned short type;
unsigned short failfast;
/*
* Qgroup equivalent for @size @reserved
*
* Unlike normal @size/@reserved for inode rsv, qgroup doesn't care
* about things like csum size nor how many tree blocks it will need to
* reserve.
*
* Qgroup cares more about net change of the extent usage.
*
* So for one newly inserted file extent, in worst case it will cause
* leaf split and level increase, nodesize for each file extent is
* already too much.
*
* In short, qgroup_size/reserved is the upper limit of possible needed
* qgroup metadata reservation.
*/
u64 qgroup_rsv_size;
u64 qgroup_rsv_reserved;
};
/*
* free clusters are used to claim free space in relatively large chunks,
btrfs: Do not use data_alloc_cluster in ssd mode This patch provides a band aid to improve the 'out of the box' behaviour of btrfs for disks that are detected as being an ssd. In a general purpose mixed workload scenario, the current ssd mode causes overallocation of available raw disk space for data, while leaving behind increasing amounts of unused fragmented free space. This situation leads to early ENOSPC problems which are harming user experience and adoption of btrfs as a general purpose filesystem. This patch modifies the data extent allocation behaviour of the ssd mode to make it behave identical to nossd mode. The metadata behaviour and additional ssd_spread option stay untouched so far. Recommendations for future development are to reconsider the current oversimplified nossd / ssd distinction and the broken detection mechanism based on the rotational attribute in sysfs and provide experienced users with a more flexible way to choose allocator behaviour for data and metadata, optimized for certain use cases, while keeping sane 'out of the box' default settings. The internals of the current btrfs code have more potential than what currently gets exposed to the user to choose from. The SSD story... In the first year of btrfs development, around early 2008, btrfs gained a mount option which enables specific functionality for filesystems on solid state devices. The first occurance of this functionality is in commit e18e4809, labeled "Add mount -o ssd, which includes optimizations for seek free storage". The effect on allocating free space for doing (data) writes is to 'cluster' writes together, writing them out in contiguous space, as opposed to a 'tetris' way of putting all separate writes into any free space fragment that fits (which is what the -o nossd behaviour does). A somewhat simplified explanation of what happens is that, when for example, the 'cluster' size is set to 2MiB, when we do some writes, the data allocator will search for a free space block that is 2MiB big, and put the writes in there. The ssd mode itself might allow a 2MiB cluster to be composed of multiple free space extents with some existing data in between, while the additional ssd_spread mount option kills off this option and requires fully free space. The idea behind this is (commit 536ac8ae): "The [...] clusters make it more likely a given IO will completely overwrite the ssd block, so it doesn't have to do an internal rwm cycle."; ssd block meaning nand erase block. So, effectively this means applying a "locality based algorithm" and trying to outsmart the actual ssd. Since then, various changes have been made to the involved code, but the basic idea is still present, and gets activated whenever the ssd mount option is active. This also happens by default, when the rotational flag as seen at /sys/block/<device>/queue/rotational is set to 0. However, there's a number of problems with this approach. First, what the optimization is trying to do is outsmart the ssd by assuming there is a relation between the physical address space of the block device as seen by btrfs and the actual physical storage of the ssd, and then adjusting data placement. However, since the introduction of the Flash Translation Layer (FTL) which is a part of the internal controller of an ssd, these attempts are futile. The use of good quality FTL in consumer ssd products might have been limited in 2008, but this situation has changed drastically soon after that time. Today, even the flash memory in your automatic cat feeding machine or your grandma's wheelchair has a full featured one. Second, the behaviour as described above results in the filesystem being filled up with badly fragmented free space extents because of relatively small pieces of space that are freed up by deletes, but not selected again as part of a 'cluster'. Since the algorithm prefers allocating a new chunk over going back to tetris mode, the end result is a filesystem in which all raw space is allocated, but which is composed of underutilized chunks with a 'shotgun blast' pattern of fragmented free space. Usually, the next problematic thing that happens is the filesystem wanting to allocate new space for metadata, which causes the filesystem to fail in spectacular ways. Third, the default mount options you get for an ssd ('ssd' mode enabled, 'discard' not enabled), in combination with spreading out writes over the full address space and ignoring freed up space leads to worst case behaviour in providing information to the ssd itself, since it will never learn that all the free space left behind is actually free. There are two ways to let an ssd know previously written data does not have to be preserved, which are sending explicit signals using discard or fstrim, or by simply overwriting the space with new data. The worst case behaviour is the btrfs ssd_spread mount option in combination with not having discard enabled. It has a side effect of minimizing the reuse of free space previously written in. Fourth, the rotational flag in /sys/ does not reliably indicate if the device is a locally attached ssd. For example, iSCSI or NBD displays as non-rotational, while a loop device on an ssd shows up as rotational. The combination of the second and third problem effectively means that despite all the good intentions, the btrfs ssd mode reliably causes the ssd hardware and the filesystem structures and performance to be choked to death. The clickbait version of the title of this story would have been "Btrfs ssd optimizations considered harmful for ssds". The current nossd 'tetris' mode (even still without discard) allows a pattern of overwriting much more previously used space, causing many more implicit discards to happen because of the overwrite information the ssd gets. The actual location in the physical address space, as seen from the point of view of btrfs is irrelevant, because the actual writes to the low level flash are reordered anyway thanks to the FTL. Changes made in the code 1. Make ssd mode data allocation identical to tetris mode, like nossd. 2. Adjust and clean up filesystem mount messages so that we can easily identify if a kernel has this patch applied or not, when providing support to end users. Also, make better use of the *_and_info helpers to only trigger messages on actual state changes. Backporting notes Notes for whoever wants to backport this patch to their 4.9 LTS kernel: * First apply commit 951e7966 "btrfs: drop the nossd flag when remounting with -o ssd", or fixup the differences manually. * The rest of the conflicts are because of the fs_info refactoring. So, for example, instead of using fs_info, it's root->fs_info in extent-tree.c Signed-off-by: Hans van Kranenburg <hans.van.kranenburg@mendix.com> Signed-off-by: David Sterba <dsterba@suse.com>
2017-07-28 06:31:28 +00:00
* allowing us to do less seeky writes. They are used for all metadata
* allocations. In ssd_spread mode they are also used for data allocations.
*/
struct btrfs_free_cluster {
spinlock_t lock;
spinlock_t refill_lock;
struct rb_root root;
/* largest extent in this cluster */
u64 max_size;
/* first extent starting offset */
u64 window_start;
/* We did a full search and couldn't create a cluster */
bool fragmented;
struct btrfs_block_group_cache *block_group;
/*
* when a cluster is allocated from a block group, we put the
* cluster onto a list in the block group so that it can
* be freed before the block group is freed.
*/
struct list_head block_group_list;
};
Btrfs: async block group caching This patch moves the caching of the block group off to a kthread in order to allow people to allocate sooner. Instead of blocking up behind the caching mutex, we instead kick of the caching kthread, and then attempt to make an allocation. If we cannot, we wait on the block groups caching waitqueue, which the caching kthread will wake the waiting threads up everytime it finds 2 meg worth of space, and then again when its finished caching. This is how I tested the speedup from this mkfs the disk mount the disk fill the disk up with fs_mark unmount the disk mount the disk time touch /mnt/foo Without my changes this took 11 seconds on my box, with these changes it now takes 1 second. Another change thats been put in place is we lock the super mirror's in the pinned extent map in order to keep us from adding that stuff as free space when caching the block group. This doesn't really change anything else as far as the pinned extent map is concerned, since for actual pinned extents we use EXTENT_DIRTY, but it does mean that when we unmount we have to go in and unlock those extents to keep from leaking memory. I've also added a check where when we are reading block groups from disk, if the amount of space used == the size of the block group, we go ahead and mark the block group as cached. This drastically reduces the amount of time it takes to cache the block groups. Using the same test as above, except doing a dd to a file and then unmounting, it used to take 33 seconds to umount, now it takes 3 seconds. This version uses the commit_root in the caching kthread, and then keeps track of how many async caching threads are running at any given time so if one of the async threads is still running as we cross transactions we can wait until its finished before handling the pinned extents. Thank you, Signed-off-by: Josef Bacik <jbacik@redhat.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-07-14 01:29:25 +00:00
enum btrfs_caching_type {
BTRFS_CACHE_NO,
BTRFS_CACHE_STARTED,
BTRFS_CACHE_FAST,
BTRFS_CACHE_FINISHED,
BTRFS_CACHE_ERROR,
Btrfs: async block group caching This patch moves the caching of the block group off to a kthread in order to allow people to allocate sooner. Instead of blocking up behind the caching mutex, we instead kick of the caching kthread, and then attempt to make an allocation. If we cannot, we wait on the block groups caching waitqueue, which the caching kthread will wake the waiting threads up everytime it finds 2 meg worth of space, and then again when its finished caching. This is how I tested the speedup from this mkfs the disk mount the disk fill the disk up with fs_mark unmount the disk mount the disk time touch /mnt/foo Without my changes this took 11 seconds on my box, with these changes it now takes 1 second. Another change thats been put in place is we lock the super mirror's in the pinned extent map in order to keep us from adding that stuff as free space when caching the block group. This doesn't really change anything else as far as the pinned extent map is concerned, since for actual pinned extents we use EXTENT_DIRTY, but it does mean that when we unmount we have to go in and unlock those extents to keep from leaking memory. I've also added a check where when we are reading block groups from disk, if the amount of space used == the size of the block group, we go ahead and mark the block group as cached. This drastically reduces the amount of time it takes to cache the block groups. Using the same test as above, except doing a dd to a file and then unmounting, it used to take 33 seconds to umount, now it takes 3 seconds. This version uses the commit_root in the caching kthread, and then keeps track of how many async caching threads are running at any given time so if one of the async threads is still running as we cross transactions we can wait until its finished before handling the pinned extents. Thank you, Signed-off-by: Josef Bacik <jbacik@redhat.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-07-14 01:29:25 +00:00
};
enum btrfs_disk_cache_state {
BTRFS_DC_WRITTEN,
BTRFS_DC_ERROR,
BTRFS_DC_CLEAR,
BTRFS_DC_SETUP,
};
struct btrfs_caching_control {
struct list_head list;
struct mutex mutex;
wait_queue_head_t wait;
struct btrfs_work work;
struct btrfs_block_group_cache *block_group;
u64 progress;
refcount_t count;
};
/* Once caching_thread() finds this much free space, it will wake up waiters. */
#define CACHING_CTL_WAKE_UP SZ_2M
struct btrfs_io_ctl {
void *cur, *orig;
struct page *page;
struct page **pages;
struct btrfs_fs_info *fs_info;
struct inode *inode;
unsigned long size;
int index;
int num_pages;
int entries;
int bitmaps;
unsigned check_crcs:1;
};
/*
* Tree to record all locked full stripes of a RAID5/6 block group
*/
struct btrfs_full_stripe_locks_tree {
struct rb_root root;
struct mutex lock;
};
struct btrfs_block_group_cache {
struct btrfs_key key;
struct btrfs_block_group_item item;
Btrfs: async block group caching This patch moves the caching of the block group off to a kthread in order to allow people to allocate sooner. Instead of blocking up behind the caching mutex, we instead kick of the caching kthread, and then attempt to make an allocation. If we cannot, we wait on the block groups caching waitqueue, which the caching kthread will wake the waiting threads up everytime it finds 2 meg worth of space, and then again when its finished caching. This is how I tested the speedup from this mkfs the disk mount the disk fill the disk up with fs_mark unmount the disk mount the disk time touch /mnt/foo Without my changes this took 11 seconds on my box, with these changes it now takes 1 second. Another change thats been put in place is we lock the super mirror's in the pinned extent map in order to keep us from adding that stuff as free space when caching the block group. This doesn't really change anything else as far as the pinned extent map is concerned, since for actual pinned extents we use EXTENT_DIRTY, but it does mean that when we unmount we have to go in and unlock those extents to keep from leaking memory. I've also added a check where when we are reading block groups from disk, if the amount of space used == the size of the block group, we go ahead and mark the block group as cached. This drastically reduces the amount of time it takes to cache the block groups. Using the same test as above, except doing a dd to a file and then unmounting, it used to take 33 seconds to umount, now it takes 3 seconds. This version uses the commit_root in the caching kthread, and then keeps track of how many async caching threads are running at any given time so if one of the async threads is still running as we cross transactions we can wait until its finished before handling the pinned extents. Thank you, Signed-off-by: Josef Bacik <jbacik@redhat.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-07-14 01:29:25 +00:00
struct btrfs_fs_info *fs_info;
struct inode *inode;
spinlock_t lock;
u64 pinned;
u64 reserved;
Btrfs: fix broken free space cache after the system crashed When we mounted the filesystem after the crash, we got the following message: BTRFS error (device xxx): block group xxxx has wrong amount of free space BTRFS error (device xxx): failed to load free space cache for block group xxx It is because we didn't update the metadata of the allocated space (in extent tree) until the file data was written into the disk. During this time, there was no information about the allocated spaces in either the extent tree nor the free space cache. when we wrote out the free space cache at this time (commit transaction), those spaces were lost. In fact, only the free space that is used to store the file data had this problem, the others didn't because the metadata of them is updated in the same transaction context. There are many methods which can fix the above problem - track the allocated space, and write it out when we write out the free space cache - account the size of the allocated space that is used to store the file data, if the size is not zero, don't write out the free space cache. The first one is complex and may make the performance drop down. This patch chose the second method, we use a per-block-group variant to account the size of that allocated space. Besides that, we also introduce a per-block-group read-write semaphore to avoid the race between the allocation and the free space cache write out. Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-06-19 02:42:50 +00:00
u64 delalloc_bytes;
u64 bytes_super;
u64 flags;
Btrfs: inline checksums into the disk free space cache Yeah yeah I know this is how we used to do it and then I changed it, but damnit I'm changing it back. The fact is that writing out checksums will modify metadata, which could cause us to dirty a block group we've already written out, so we have to truncate it and all of it's checksums and re-write it which will write new checksums which could dirty a blockg roup that has already been written and you see where I'm going with this? This can cause unmount or really anything that depends on a transaction to commit to take it's sweet damned time to happen. So go back to the way it was, only this time we're specifically setting NODATACOW because we can't go through the COW pathway anyway and we're doing our own built-in cow'ing by truncating the free space cache. The other new thing is once we truncate the old cache and preallocate the new space, we don't need to do that song and dance at all for the rest of the transaction, we can just overwrite the existing space with the new cache if the block group changes for whatever reason, and the NODATACOW will let us do this fine. So keep track of which transaction we last cleared our cache in and if we cleared it in this transaction just say we're all setup and carry on. This survives xfstests and stress.sh. The inode cache will continue to use the normal csum infrastructure since it only gets written once and there will be no more modifications to the fs tree in a transaction commit. Signed-off-by: Josef Bacik <josef@redhat.com>
2011-10-06 12:58:24 +00:00
u64 cache_generation;
/*
* If the free space extent count exceeds this number, convert the block
* group to bitmaps.
*/
u32 bitmap_high_thresh;
/*
* If the free space extent count drops below this number, convert the
* block group back to extents.
*/
u32 bitmap_low_thresh;
Btrfs: fix broken free space cache after the system crashed When we mounted the filesystem after the crash, we got the following message: BTRFS error (device xxx): block group xxxx has wrong amount of free space BTRFS error (device xxx): failed to load free space cache for block group xxx It is because we didn't update the metadata of the allocated space (in extent tree) until the file data was written into the disk. During this time, there was no information about the allocated spaces in either the extent tree nor the free space cache. when we wrote out the free space cache at this time (commit transaction), those spaces were lost. In fact, only the free space that is used to store the file data had this problem, the others didn't because the metadata of them is updated in the same transaction context. There are many methods which can fix the above problem - track the allocated space, and write it out when we write out the free space cache - account the size of the allocated space that is used to store the file data, if the size is not zero, don't write out the free space cache. The first one is complex and may make the performance drop down. This patch chose the second method, we use a per-block-group variant to account the size of that allocated space. Besides that, we also introduce a per-block-group read-write semaphore to avoid the race between the allocation and the free space cache write out. Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-06-19 02:42:50 +00:00
/*
* It is just used for the delayed data space allocation because
* only the data space allocation and the relative metadata update
* can be done cross the transaction.
*/
struct rw_semaphore data_rwsem;
/* for raid56, this is a full stripe, without parity */
unsigned long full_stripe_len;
unsigned int ro;
unsigned int iref:1;
Btrfs: fix crash caused by block group removal If we remove a block group (because it became empty), we might have left a caching_ctl structure in fs_info->caching_block_groups that points to the block group and is accessed at transaction commit time. This results in accessing an invalid or incorrect block group. This issue became visible after Josef's patch "Btrfs: remove empty block groups automatically". So if the block group is removed make sure we don't leave a dangling caching_ctl in caching_block_groups. Sample crash trace: [58380.439449] BUG: unable to handle kernel paging request at ffff8801446eaeb8 [58380.439707] IP: [<ffffffffa03f6d05>] block_group_cache_done.isra.21+0xc/0x1c [btrfs] [58380.440879] PGD 1acb067 PUD 23f5ff067 PMD 23f5db067 PTE 80000001446ea060 [58380.441220] Oops: 0000 [#1] SMP DEBUG_PAGEALLOC [58380.441486] Modules linked in: btrfs crc32c_generic xor raid6_pq nfsd auth_rpcgss oid_registry nfs_acl nfs lockd fscache sunrpc loop psmouse processor i2c_piix4 parport_pc parport pcspkr serio_raw evdev i2ccore thermal_sys microcode button ext4 crc16 jbd2 mbcache sr_mod cdrom ata_generic sg sd_mod crc_t10dif crct10dif_generic crct10dif_common virtio_scsi floppy ata_piix e1000 libata virtio_pci scsi_mod virtio_ring virtio [last unloaded: btrfs] [58380.443238] CPU: 3 PID: 25728 Comm: btrfs-transacti Tainted: G W 3.17.0-rc5-btrfs-next-1+ #1 [58380.443238] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS rel-1.7.5-0-ge51488c-20140602_164612-nilsson.home.kraxel.org 04/01/2014 [58380.443238] task: ffff88013ac82090 ti: ffff88013896c000 task.ti: ffff88013896c000 [58380.443238] RIP: 0010:[<ffffffffa03f6d05>] [<ffffffffa03f6d05>] block_group_cache_done.isra.21+0xc/0x1c [btrfs] [58380.443238] RSP: 0018:ffff88013896fdd8 EFLAGS: 00010283 [58380.443238] RAX: ffff880222cae850 RBX: ffff880119ba74c0 RCX: 0000000000000000 [58380.443238] RDX: 0000000000000000 RSI: ffff880185e16800 RDI: ffff8801446eaeb8 [58380.443238] RBP: ffff88013896fdd8 R08: ffff8801a9ca9fa8 R09: ffff88013896fc60 [58380.443238] R10: ffff88013896fd28 R11: 0000000000000000 R12: ffff880222cae000 [58380.443238] R13: ffff880222cae850 R14: ffff880222cae6b0 R15: ffff8801446eae00 [58380.443238] FS: 0000000000000000(0000) GS:ffff88023ed80000(0000) knlGS:0000000000000000 [58380.443238] CS: 0010 DS: 0000 ES: 0000 CR0: 000000008005003b [58380.443238] CR2: ffff8801446eaeb8 CR3: 0000000001811000 CR4: 00000000000006e0 [58380.443238] Stack: [58380.443238] ffff88013896fe18 ffffffffa03fe2d5 ffff880222cae850 ffff880185e16800 [58380.443238] ffff88000dc41c20 0000000000000000 ffff8801a9ca9f00 0000000000000000 [58380.443238] ffff88013896fe80 ffffffffa040fbcf ffff88018b0dcdb0 ffff88013ac82090 [58380.443238] Call Trace: [58380.443238] [<ffffffffa03fe2d5>] btrfs_prepare_extent_commit+0x5a/0xd7 [btrfs] [58380.443238] [<ffffffffa040fbcf>] btrfs_commit_transaction+0x45c/0x882 [btrfs] [58380.443238] [<ffffffffa040c058>] transaction_kthread+0xf2/0x1a4 [btrfs] [58380.443238] [<ffffffffa040bf66>] ? btrfs_cleanup_transaction+0x3d8/0x3d8 [btrfs] [58380.443238] [<ffffffff8105966b>] kthread+0xb7/0xbf [58380.443238] [<ffffffff810595b4>] ? __kthread_parkme+0x67/0x67 [58380.443238] [<ffffffff813ebeac>] ret_from_fork+0x7c/0xb0 [58380.443238] [<ffffffff810595b4>] ? __kthread_parkme+0x67/0x67 Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-11-26 15:28:51 +00:00
unsigned int has_caching_ctl:1;
Btrfs: fix race between fs trimming and block group remove/allocation Our fs trim operation, which is completely transactionless (doesn't start or joins an existing transaction) consists of visiting all block groups and then for each one to iterate its free space entries and perform a discard operation against the space range represented by the free space entries. However before performing a discard, the corresponding free space entry is removed from the free space rbtree, and when the discard completes it is added back to the free space rbtree. If a block group remove operation happens while the discard is ongoing (or before it starts and after a free space entry is hidden), we end up not waiting for the discard to complete, remove the extent map that maps logical address to physical addresses and the corresponding chunk metadata from the the chunk and device trees. After that and before the discard completes, the current running transaction can finish and a new one start, allowing for new block groups that map to the same physical addresses to be allocated and written to. So fix this by keeping the extent map in memory until the discard completes so that the same physical addresses aren't reused before it completes. If the physical locations that are under a discard operation end up being used for a new metadata block group for example, and dirty metadata extents are written before the discard finishes (the VM might call writepages() of our btree inode's i_mapping for example, or an fsync log commit happens) we end up overwriting metadata with zeroes, which leads to errors from fsck like the following: checking extents Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 read block failed check_tree_block owner ref check failed [833912832 16384] Errors found in extent allocation tree or chunk allocation checking free space cache checking fs roots Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 read block failed check_tree_block root 5 root dir 256 error root 5 inode 260 errors 2001, no inode item, link count wrong unresolved ref dir 256 index 0 namelen 8 name foobar_3 filetype 1 errors 6, no dir index, no inode ref root 5 inode 262 errors 2001, no inode item, link count wrong unresolved ref dir 256 index 0 namelen 8 name foobar_5 filetype 1 errors 6, no dir index, no inode ref root 5 inode 263 errors 2001, no inode item, link count wrong (...) Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-11-27 21:14:15 +00:00
unsigned int removed:1;
int disk_cache_state;
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
Btrfs: async block group caching This patch moves the caching of the block group off to a kthread in order to allow people to allocate sooner. Instead of blocking up behind the caching mutex, we instead kick of the caching kthread, and then attempt to make an allocation. If we cannot, we wait on the block groups caching waitqueue, which the caching kthread will wake the waiting threads up everytime it finds 2 meg worth of space, and then again when its finished caching. This is how I tested the speedup from this mkfs the disk mount the disk fill the disk up with fs_mark unmount the disk mount the disk time touch /mnt/foo Without my changes this took 11 seconds on my box, with these changes it now takes 1 second. Another change thats been put in place is we lock the super mirror's in the pinned extent map in order to keep us from adding that stuff as free space when caching the block group. This doesn't really change anything else as far as the pinned extent map is concerned, since for actual pinned extents we use EXTENT_DIRTY, but it does mean that when we unmount we have to go in and unlock those extents to keep from leaking memory. I've also added a check where when we are reading block groups from disk, if the amount of space used == the size of the block group, we go ahead and mark the block group as cached. This drastically reduces the amount of time it takes to cache the block groups. Using the same test as above, except doing a dd to a file and then unmounting, it used to take 33 seconds to umount, now it takes 3 seconds. This version uses the commit_root in the caching kthread, and then keeps track of how many async caching threads are running at any given time so if one of the async threads is still running as we cross transactions we can wait until its finished before handling the pinned extents. Thank you, Signed-off-by: Josef Bacik <jbacik@redhat.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-07-14 01:29:25 +00:00
/* cache tracking stuff */
int cached;
struct btrfs_caching_control *caching_ctl;
u64 last_byte_to_unpin;
Btrfs: async block group caching This patch moves the caching of the block group off to a kthread in order to allow people to allocate sooner. Instead of blocking up behind the caching mutex, we instead kick of the caching kthread, and then attempt to make an allocation. If we cannot, we wait on the block groups caching waitqueue, which the caching kthread will wake the waiting threads up everytime it finds 2 meg worth of space, and then again when its finished caching. This is how I tested the speedup from this mkfs the disk mount the disk fill the disk up with fs_mark unmount the disk mount the disk time touch /mnt/foo Without my changes this took 11 seconds on my box, with these changes it now takes 1 second. Another change thats been put in place is we lock the super mirror's in the pinned extent map in order to keep us from adding that stuff as free space when caching the block group. This doesn't really change anything else as far as the pinned extent map is concerned, since for actual pinned extents we use EXTENT_DIRTY, but it does mean that when we unmount we have to go in and unlock those extents to keep from leaking memory. I've also added a check where when we are reading block groups from disk, if the amount of space used == the size of the block group, we go ahead and mark the block group as cached. This drastically reduces the amount of time it takes to cache the block groups. Using the same test as above, except doing a dd to a file and then unmounting, it used to take 33 seconds to umount, now it takes 3 seconds. This version uses the commit_root in the caching kthread, and then keeps track of how many async caching threads are running at any given time so if one of the async threads is still running as we cross transactions we can wait until its finished before handling the pinned extents. Thank you, Signed-off-by: Josef Bacik <jbacik@redhat.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-07-14 01:29:25 +00:00
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
struct btrfs_space_info *space_info;
/* free space cache stuff */
struct btrfs_free_space_ctl *free_space_ctl;
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
/* block group cache stuff */
struct rb_node cache_node;
/* for block groups in the same raid type */
struct list_head list;
/* usage count */
atomic_t count;
/* List of struct btrfs_free_clusters for this block group.
* Today it will only have one thing on it, but that may change
*/
struct list_head cluster_list;
/* For delayed block group creation or deletion of empty block groups */
struct list_head bg_list;
/* For read-only block groups */
struct list_head ro_list;
Btrfs: fix race between fs trimming and block group remove/allocation Our fs trim operation, which is completely transactionless (doesn't start or joins an existing transaction) consists of visiting all block groups and then for each one to iterate its free space entries and perform a discard operation against the space range represented by the free space entries. However before performing a discard, the corresponding free space entry is removed from the free space rbtree, and when the discard completes it is added back to the free space rbtree. If a block group remove operation happens while the discard is ongoing (or before it starts and after a free space entry is hidden), we end up not waiting for the discard to complete, remove the extent map that maps logical address to physical addresses and the corresponding chunk metadata from the the chunk and device trees. After that and before the discard completes, the current running transaction can finish and a new one start, allowing for new block groups that map to the same physical addresses to be allocated and written to. So fix this by keeping the extent map in memory until the discard completes so that the same physical addresses aren't reused before it completes. If the physical locations that are under a discard operation end up being used for a new metadata block group for example, and dirty metadata extents are written before the discard finishes (the VM might call writepages() of our btree inode's i_mapping for example, or an fsync log commit happens) we end up overwriting metadata with zeroes, which leads to errors from fsck like the following: checking extents Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 read block failed check_tree_block owner ref check failed [833912832 16384] Errors found in extent allocation tree or chunk allocation checking free space cache checking fs roots Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 Check tree block failed, want=833912832, have=0 read block failed check_tree_block root 5 root dir 256 error root 5 inode 260 errors 2001, no inode item, link count wrong unresolved ref dir 256 index 0 namelen 8 name foobar_3 filetype 1 errors 6, no dir index, no inode ref root 5 inode 262 errors 2001, no inode item, link count wrong unresolved ref dir 256 index 0 namelen 8 name foobar_5 filetype 1 errors 6, no dir index, no inode ref root 5 inode 263 errors 2001, no inode item, link count wrong (...) Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-11-27 21:14:15 +00:00
atomic_t trimming;
/* For dirty block groups */
struct list_head dirty_list;
struct list_head io_list;
struct btrfs_io_ctl io_ctl;
Btrfs: don't do unnecessary delalloc flushes when relocating Before we start the actual relocation process of a block group, we do calls to flush delalloc of all inodes and then wait for ordered extents to complete. However we do these flush calls just to make sure we don't race with concurrent tasks that have actually already started to run delalloc and have allocated an extent from the block group we want to relocate, right before we set it to readonly mode, but have not yet created the respective ordered extents. The flush calls make us wait for such concurrent tasks because they end up calling filemap_fdatawrite_range() (through btrfs_start_delalloc_roots() -> __start_delalloc_inodes() -> btrfs_alloc_delalloc_work() -> btrfs_run_delalloc_work()) which ends up serializing us with those tasks due to attempts to lock the same pages (and the delalloc flush procedure calls the allocator and creates the ordered extents before unlocking the pages). These flushing calls not only make us waste time (cpu, IO) but also reduce the chances of writing larger extents (applications might be writing to contiguous ranges and we flush before they finish dirtying the whole ranges). So make sure we don't flush delalloc and just wait for concurrent tasks that have already started flushing delalloc and have allocated an extent from the block group we are about to relocate. This change also ends up fixing a race with direct IO writes that makes relocation not wait for direct IO ordered extents. This race is illustrated by the following diagram: CPU 1 CPU 2 btrfs_relocate_block_group(bg X) starts direct IO write, target inode currently has no ordered extents ongoing nor dirty pages (delalloc regions), therefore the root for our inode is not in the list fs_info->ordered_roots btrfs_direct_IO() __blockdev_direct_IO() btrfs_get_blocks_direct() btrfs_lock_extent_direct() locks range in the io tree btrfs_new_extent_direct() btrfs_reserve_extent() --> extent allocated from bg X btrfs_inc_block_group_ro(bg X) btrfs_start_delalloc_roots() __start_delalloc_inodes() --> does nothing, no dealloc ranges in the inode's io tree so the inode's root is not in the list fs_info->delalloc_roots btrfs_wait_ordered_roots() --> does not find the inode's root in the list fs_info->ordered_roots --> ends up not waiting for the direct IO write started by the task at CPU 2 relocate_block_group(rc->stage == MOVE_DATA_EXTENTS) prepare_to_relocate() btrfs_commit_transaction() iterates the extent tree, using its commit root and moves extents into new locations btrfs_add_ordered_extent_dio() --> now a ordered extent is created and added to the list root->ordered_extents and the root added to the list fs_info->ordered_roots --> this is too late and the task at CPU 1 already started the relocation btrfs_commit_transaction() btrfs_finish_ordered_io() btrfs_alloc_reserved_file_extent() --> adds delayed data reference for the extent allocated from bg X relocate_block_group(rc->stage == UPDATE_DATA_PTRS) prepare_to_relocate() btrfs_commit_transaction() --> delayed refs are run, so an extent item for the allocated extent from bg X is added to extent tree --> commit roots are switched, so the next scan in the extent tree will see the extent item sees the extent in the extent tree When this happens the relocation produces the following warning when it finishes: [ 7260.832836] ------------[ cut here ]------------ [ 7260.834653] WARNING: CPU: 5 PID: 6765 at fs/btrfs/relocation.c:4318 btrfs_relocate_block_group+0x245/0x2a1 [btrfs]() [ 7260.838268] Modules linked in: btrfs crc32c_generic xor ppdev raid6_pq psmouse sg acpi_cpufreq evdev i2c_piix4 tpm_tis serio_raw tpm i2c_core pcspkr parport_pc [ 7260.850935] CPU: 5 PID: 6765 Comm: btrfs Not tainted 4.5.0-rc6-btrfs-next-28+ #1 [ 7260.852998] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS by qemu-project.org 04/01/2014 [ 7260.852998] 0000000000000000 ffff88020bf57bc0 ffffffff812648b3 0000000000000000 [ 7260.852998] 0000000000000009 ffff88020bf57bf8 ffffffff81051608 ffffffffa03c1b2d [ 7260.852998] ffff8800b2bbb800 0000000000000000 ffff8800b17bcc58 ffff8800399dd000 [ 7260.852998] Call Trace: [ 7260.852998] [<ffffffff812648b3>] dump_stack+0x67/0x90 [ 7260.852998] [<ffffffff81051608>] warn_slowpath_common+0x99/0xb2 [ 7260.852998] [<ffffffffa03c1b2d>] ? btrfs_relocate_block_group+0x245/0x2a1 [btrfs] [ 7260.852998] [<ffffffff810516d4>] warn_slowpath_null+0x1a/0x1c [ 7260.852998] [<ffffffffa03c1b2d>] btrfs_relocate_block_group+0x245/0x2a1 [btrfs] [ 7260.852998] [<ffffffffa039d9de>] btrfs_relocate_chunk.isra.29+0x66/0xdb [btrfs] [ 7260.852998] [<ffffffffa039f314>] btrfs_balance+0xde1/0xe4e [btrfs] [ 7260.852998] [<ffffffff8127d671>] ? debug_smp_processor_id+0x17/0x19 [ 7260.852998] [<ffffffffa03a9583>] btrfs_ioctl_balance+0x255/0x2d3 [btrfs] [ 7260.852998] [<ffffffffa03ac96a>] btrfs_ioctl+0x11e0/0x1dff [btrfs] [ 7260.852998] [<ffffffff811451df>] ? handle_mm_fault+0x443/0xd63 [ 7260.852998] [<ffffffff81491817>] ? _raw_spin_unlock+0x31/0x44 [ 7260.852998] [<ffffffff8108b36a>] ? arch_local_irq_save+0x9/0xc [ 7260.852998] [<ffffffff811876ab>] vfs_ioctl+0x18/0x34 [ 7260.852998] [<ffffffff81187cb2>] do_vfs_ioctl+0x550/0x5be [ 7260.852998] [<ffffffff81190c30>] ? __fget_light+0x4d/0x71 [ 7260.852998] [<ffffffff81187d77>] SyS_ioctl+0x57/0x79 [ 7260.852998] [<ffffffff81492017>] entry_SYSCALL_64_fastpath+0x12/0x6b [ 7260.893268] ---[ end trace eb7803b24ebab8ad ]--- This is because at the end of the first stage, in relocate_block_group(), we commit the current transaction, which makes delayed refs run, the commit roots are switched and so the second stage will find the extent item that the ordered extent added to the delayed refs. But this extent was not moved (ordered extent completed after first stage finished), so at the end of the relocation our block group item still has a positive used bytes counter, triggering a warning at the end of btrfs_relocate_block_group(). Later on when trying to read the extent contents from disk we hit a BUG_ON() due to the inability to map a block with a logical address that belongs to the block group we relocated and is no longer valid, resulting in the following trace: [ 7344.885290] BTRFS critical (device sdi): unable to find logical 12845056 len 4096 [ 7344.887518] ------------[ cut here ]------------ [ 7344.888431] kernel BUG at fs/btrfs/inode.c:1833! [ 7344.888431] invalid opcode: 0000 [#1] PREEMPT SMP DEBUG_PAGEALLOC [ 7344.888431] Modules linked in: btrfs crc32c_generic xor ppdev raid6_pq psmouse sg acpi_cpufreq evdev i2c_piix4 tpm_tis serio_raw tpm i2c_core pcspkr parport_pc [ 7344.888431] CPU: 0 PID: 6831 Comm: od Tainted: G W 4.5.0-rc6-btrfs-next-28+ #1 [ 7344.888431] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS by qemu-project.org 04/01/2014 [ 7344.888431] task: ffff880215818600 ti: ffff880204684000 task.ti: ffff880204684000 [ 7344.888431] RIP: 0010:[<ffffffffa037c88c>] [<ffffffffa037c88c>] btrfs_merge_bio_hook+0x54/0x6b [btrfs] [ 7344.888431] RSP: 0018:ffff8802046878f0 EFLAGS: 00010282 [ 7344.888431] RAX: 00000000ffffffea RBX: 0000000000001000 RCX: 0000000000000001 [ 7344.888431] RDX: ffff88023ec0f950 RSI: ffffffff8183b638 RDI: 00000000ffffffff [ 7344.888431] RBP: ffff880204687908 R08: 0000000000000001 R09: 0000000000000000 [ 7344.888431] R10: ffff880204687770 R11: ffffffff82f2d52d R12: 0000000000001000 [ 7344.888431] R13: ffff88021afbfee8 R14: 0000000000006208 R15: ffff88006cd199b0 [ 7344.888431] FS: 00007f1f9e1d6700(0000) GS:ffff88023ec00000(0000) knlGS:0000000000000000 [ 7344.888431] CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033 [ 7344.888431] CR2: 00007f1f9dc8cb60 CR3: 000000023e3b6000 CR4: 00000000000006f0 [ 7344.888431] Stack: [ 7344.888431] 0000000000001000 0000000000001000 ffff880204687b98 ffff880204687950 [ 7344.888431] ffffffffa0395c8f ffffea0004d64d48 0000000000000000 0000000000001000 [ 7344.888431] ffffea0004d64d48 0000000000001000 0000000000000000 0000000000000000 [ 7344.888431] Call Trace: [ 7344.888431] [<ffffffffa0395c8f>] submit_extent_page+0xf5/0x16f [btrfs] [ 7344.888431] [<ffffffffa03970ac>] __do_readpage+0x4a0/0x4f1 [btrfs] [ 7344.888431] [<ffffffffa039680d>] ? btrfs_create_repair_bio+0xcb/0xcb [btrfs] [ 7344.888431] [<ffffffffa037eeb4>] ? btrfs_writepage_start_hook+0xbc/0xbc [btrfs] [ 7344.888431] [<ffffffff8108df55>] ? trace_hardirqs_on+0xd/0xf [ 7344.888431] [<ffffffffa039728c>] __do_contiguous_readpages.constprop.26+0xc2/0xe4 [btrfs] [ 7344.888431] [<ffffffffa037eeb4>] ? btrfs_writepage_start_hook+0xbc/0xbc [btrfs] [ 7344.888431] [<ffffffffa039739b>] __extent_readpages.constprop.25+0xed/0x100 [btrfs] [ 7344.888431] [<ffffffff81129d24>] ? lru_cache_add+0xe/0x10 [ 7344.888431] [<ffffffffa0397ea8>] extent_readpages+0x160/0x1aa [btrfs] [ 7344.888431] [<ffffffffa037eeb4>] ? btrfs_writepage_start_hook+0xbc/0xbc [btrfs] [ 7344.888431] [<ffffffff8115daad>] ? alloc_pages_current+0xa9/0xcd [ 7344.888431] [<ffffffffa037cdc9>] btrfs_readpages+0x1f/0x21 [btrfs] [ 7344.888431] [<ffffffff81128316>] __do_page_cache_readahead+0x168/0x1fc [ 7344.888431] [<ffffffff811285a0>] ondemand_readahead+0x1f6/0x207 [ 7344.888431] [<ffffffff811285a0>] ? ondemand_readahead+0x1f6/0x207 [ 7344.888431] [<ffffffff8111cf34>] ? pagecache_get_page+0x2b/0x154 [ 7344.888431] [<ffffffff8112870e>] page_cache_sync_readahead+0x3d/0x3f [ 7344.888431] [<ffffffff8111dbf7>] generic_file_read_iter+0x197/0x4e1 [ 7344.888431] [<ffffffff8117773a>] __vfs_read+0x79/0x9d [ 7344.888431] [<ffffffff81178050>] vfs_read+0x8f/0xd2 [ 7344.888431] [<ffffffff81178a38>] SyS_read+0x50/0x7e [ 7344.888431] [<ffffffff81492017>] entry_SYSCALL_64_fastpath+0x12/0x6b [ 7344.888431] Code: 8d 4d e8 45 31 c9 45 31 c0 48 8b 00 48 c1 e2 09 48 8b 80 80 fc ff ff 4c 89 65 e8 48 8b b8 f0 01 00 00 e8 1d 42 02 00 85 c0 79 02 <0f> 0b 4c 0 [ 7344.888431] RIP [<ffffffffa037c88c>] btrfs_merge_bio_hook+0x54/0x6b [btrfs] [ 7344.888431] RSP <ffff8802046878f0> [ 7344.970544] ---[ end trace eb7803b24ebab8ae ]--- Signed-off-by: Filipe Manana <fdmanana@suse.com> Reviewed-by: Josef Bacik <jbacik@fb.com> Reviewed-by: Liu Bo <bo.li.liu@oracle.com>
2016-04-26 14:39:32 +00:00
/*
* Incremented when doing extent allocations and holding a read lock
* on the space_info's groups_sem semaphore.
* Decremented when an ordered extent that represents an IO against this
* block group's range is created (after it's added to its inode's
* root's list of ordered extents) or immediately after the allocation
* if it's a metadata extent or fallocate extent (for these cases we
* don't create ordered extents).
*/
atomic_t reservations;
Btrfs: fix race between block group relocation and nocow writes Relocation of a block group waits for all existing tasks flushing dellaloc, starting direct IO writes and any ordered extents before starting the relocation process. However for direct IO writes that end up doing nocow (inode either has the flag nodatacow set or the write is against a prealloc extent) we have a short time window that allows for a race that makes relocation proceed without waiting for the direct IO write to complete first, resulting in data loss after the relocation finishes. This is illustrated by the following diagram: CPU 1 CPU 2 btrfs_relocate_block_group(bg X) direct IO write starts against an extent in block group X using nocow mode (inode has the nodatacow flag or the write is for a prealloc extent) btrfs_direct_IO() btrfs_get_blocks_direct() --> can_nocow_extent() returns 1 btrfs_inc_block_group_ro(bg X) --> turns block group into RO mode btrfs_wait_ordered_roots() --> returns and does not know about the DIO write happening at CPU 2 (the task there has not created yet an ordered extent) relocate_block_group(bg X) --> rc->stage == MOVE_DATA_EXTENTS find_next_extent() --> returns extent that the DIO write is going to write to relocate_data_extent() relocate_file_extent_cluster() --> reads the extent from disk into pages belonging to the relocation inode and dirties them --> creates DIO ordered extent btrfs_submit_direct() --> submits bio against a location on disk obtained from an extent map before the relocation started btrfs_wait_ordered_range() --> writes all the pages read before to disk (belonging to the relocation inode) relocation finishes bio completes and wrote new data to the old location of the block group So fix this by tracking the number of nocow writers for a block group and make sure relocation waits for that number to go down to 0 before starting to move the extents. The same race can also happen with buffered writes in nocow mode since the patch I recently made titled "Btrfs: don't do unnecessary delalloc flushes when relocating", because we are no longer flushing all delalloc which served as a synchonization mechanism (due to page locking) and ensured the ordered extents for nocow buffered writes were created before we called btrfs_wait_ordered_roots(). The race with direct IO writes in nocow mode existed before that patch (no pages are locked or used during direct IO) and that fixed only races with direct IO writes that do cow. Signed-off-by: Filipe Manana <fdmanana@suse.com> Reviewed-by: Josef Bacik <jbacik@fb.com>
2016-05-09 12:15:41 +00:00
/*
* Incremented while holding the spinlock *lock* by a task checking if
* it can perform a nocow write (incremented if the value for the *ro*
* field is 0). Decremented by such tasks once they create an ordered
* extent or before that if some error happens before reaching that step.
* This is to prevent races between block group relocation and nocow
* writes through direct IO.
*/
atomic_t nocow_writers;
/* Lock for free space tree operations. */
struct mutex free_space_lock;
/*
* Does the block group need to be added to the free space tree?
* Protected by free_space_lock.
*/
int needs_free_space;
/* Record locked full stripes for RAID5/6 block group */
struct btrfs_full_stripe_locks_tree full_stripe_locks_root;
};
/* delayed seq elem */
struct seq_list {
struct list_head list;
u64 seq;
};
#define SEQ_LIST_INIT(name) { .list = LIST_HEAD_INIT((name).list), .seq = 0 }
#define SEQ_LAST ((u64)-1)
enum btrfs_orphan_cleanup_state {
ORPHAN_CLEANUP_STARTED = 1,
ORPHAN_CLEANUP_DONE = 2,
};
/* used by the raid56 code to lock stripes for read/modify/write */
struct btrfs_stripe_hash {
struct list_head hash_list;
spinlock_t lock;
};
/* used by the raid56 code to lock stripes for read/modify/write */
struct btrfs_stripe_hash_table {
struct list_head stripe_cache;
spinlock_t cache_lock;
int cache_size;
struct btrfs_stripe_hash table[];
};
#define BTRFS_STRIPE_HASH_TABLE_BITS 11
Btrfs: reclaim the reserved metadata space at background Before applying this patch, the task had to reclaim the metadata space by itself if the metadata space was not enough. And When the task started the space reclamation, all the other tasks which wanted to reserve the metadata space were blocked. At some cases, they would be blocked for a long time, it made the performance fluctuate wildly. So we introduce the background metadata space reclamation, when the space is about to be exhausted, we insert a reclaim work into the workqueue, the worker of the workqueue helps us to reclaim the reserved space at the background. By this way, the tasks needn't reclaim the space by themselves at most cases, and even if the tasks have to reclaim the space or are blocked for the space reclamation, they will get enough space more quickly. Here is my test result(Tested by compilebench): Memory: 2GB CPU: 2Cores * 1CPU Partition: 40GB(SSD) Test command: # compilebench -D <mnt> -m Without this patch: intial create total runs 30 avg 54.36 MB/s (user 0.52s sys 2.44s) compile total runs 30 avg 123.72 MB/s (user 0.13s sys 1.17s) read compiled tree total runs 3 avg 81.15 MB/s (user 0.74s sys 4.89s) delete compiled tree total runs 30 avg 5.32 seconds (user 0.35s sys 4.37s) With this patch: intial create total runs 30 avg 59.80 MB/s (user 0.52s sys 2.53s) compile total runs 30 avg 151.44 MB/s (user 0.13s sys 1.11s) read compiled tree total runs 3 avg 83.25 MB/s (user 0.76s sys 4.91s) delete compiled tree total runs 30 avg 5.29 seconds (user 0.34s sys 4.34s) Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-05-14 00:29:04 +00:00
void btrfs_init_async_reclaim_work(struct work_struct *work);
/* fs_info */
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
struct reloc_control;
struct btrfs_device;
struct btrfs_fs_devices;
struct btrfs_balance_control;
btrfs: implement delayed inode items operation Changelog V5 -> V6: - Fix oom when the memory load is high, by storing the delayed nodes into the root's radix tree, and letting btrfs inodes go. Changelog V4 -> V5: - Fix the race on adding the delayed node to the inode, which is spotted by Chris Mason. - Merge Chris Mason's incremental patch into this patch. - Fix deadlock between readdir() and memory fault, which is reported by Itaru Kitayama. Changelog V3 -> V4: - Fix nested lock, which is reported by Itaru Kitayama, by updating space cache inode in time. Changelog V2 -> V3: - Fix the race between the delayed worker and the task which does delayed items balance, which is reported by Tsutomu Itoh. - Modify the patch address David Sterba's comment. - Fix the bug of the cpu recursion spinlock, reported by Chris Mason Changelog V1 -> V2: - break up the global rb-tree, use a list to manage the delayed nodes, which is created for every directory and file, and used to manage the delayed directory name index items and the delayed inode item. - introduce a worker to deal with the delayed nodes. Compare with Ext3/4, the performance of file creation and deletion on btrfs is very poor. the reason is that btrfs must do a lot of b+ tree insertions, such as inode item, directory name item, directory name index and so on. If we can do some delayed b+ tree insertion or deletion, we can improve the performance, so we made this patch which implemented delayed directory name index insertion/deletion and delayed inode update. Implementation: - introduce a delayed root object into the filesystem, that use two lists to manage the delayed nodes which are created for every file/directory. One is used to manage all the delayed nodes that have delayed items. And the other is used to manage the delayed nodes which is waiting to be dealt with by the work thread. - Every delayed node has two rb-tree, one is used to manage the directory name index which is going to be inserted into b+ tree, and the other is used to manage the directory name index which is going to be deleted from b+ tree. - introduce a worker to deal with the delayed operation. This worker is used to deal with the works of the delayed directory name index items insertion and deletion and the delayed inode update. When the delayed items is beyond the lower limit, we create works for some delayed nodes and insert them into the work queue of the worker, and then go back. When the delayed items is beyond the upper bound, we create works for all the delayed nodes that haven't been dealt with, and insert them into the work queue of the worker, and then wait for that the untreated items is below some threshold value. - When we want to insert a directory name index into b+ tree, we just add the information into the delayed inserting rb-tree. And then we check the number of the delayed items and do delayed items balance. (The balance policy is above.) - When we want to delete a directory name index from the b+ tree, we search it in the inserting rb-tree at first. If we look it up, just drop it. If not, add the key of it into the delayed deleting rb-tree. Similar to the delayed inserting rb-tree, we also check the number of the delayed items and do delayed items balance. (The same to inserting manipulation) - When we want to update the metadata of some inode, we cached the data of the inode into the delayed node. the worker will flush it into the b+ tree after dealing with the delayed insertion and deletion. - We will move the delayed node to the tail of the list after we access the delayed node, By this way, we can cache more delayed items and merge more inode updates. - If we want to commit transaction, we will deal with all the delayed node. - the delayed node will be freed when we free the btrfs inode. - Before we log the inode items, we commit all the directory name index items and the delayed inode update. I did a quick test by the benchmark tool[1] and found we can improve the performance of file creation by ~15%, and file deletion by ~20%. Before applying this patch: Create files: Total files: 50000 Total time: 1.096108 Average time: 0.000022 Delete files: Total files: 50000 Total time: 1.510403 Average time: 0.000030 After applying this patch: Create files: Total files: 50000 Total time: 0.932899 Average time: 0.000019 Delete files: Total files: 50000 Total time: 1.215732 Average time: 0.000024 [1] http://marc.info/?l=linux-btrfs&m=128212635122920&q=p3 Many thanks for Kitayama-san's help! Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Reviewed-by: David Sterba <dave@jikos.cz> Tested-by: Tsutomu Itoh <t-itoh@jp.fujitsu.com> Tested-by: Itaru Kitayama <kitayama@cl.bb4u.ne.jp> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2011-04-22 10:12:22 +00:00
struct btrfs_delayed_root;
/*
* Block group or device which contains an active swapfile. Used for preventing
* unsafe operations while a swapfile is active.
*
* These are sorted on (ptr, inode) (note that a block group or device can
* contain more than one swapfile). We compare the pointer values because we
* don't actually care what the object is, we just need a quick check whether
* the object exists in the rbtree.
*/
struct btrfs_swapfile_pin {
struct rb_node node;
void *ptr;
struct inode *inode;
/*
* If true, ptr points to a struct btrfs_block_group_cache. Otherwise,
* ptr points to a struct btrfs_device.
*/
bool is_block_group;
};
bool btrfs_pinned_by_swapfile(struct btrfs_fs_info *fs_info, void *ptr);
enum {
BTRFS_FS_BARRIER,
BTRFS_FS_CLOSING_START,
BTRFS_FS_CLOSING_DONE,
BTRFS_FS_LOG_RECOVERING,
BTRFS_FS_OPEN,
BTRFS_FS_QUOTA_ENABLED,
BTRFS_FS_UPDATE_UUID_TREE_GEN,
BTRFS_FS_CREATING_FREE_SPACE_TREE,
BTRFS_FS_BTREE_ERR,
BTRFS_FS_LOG1_ERR,
BTRFS_FS_LOG2_ERR,
BTRFS_FS_QUOTA_OVERRIDE,
/* Used to record internally whether fs has been frozen */
BTRFS_FS_FROZEN,
/*
* Indicate that a whole-filesystem exclusive operation is running
* (device replace, resize, device add/delete, balance)
*/
BTRFS_FS_EXCL_OP,
/*
* To info transaction_kthread we need an immediate commit so it
* doesn't need to wait for commit_interval
*/
BTRFS_FS_NEED_ASYNC_COMMIT,
/*
* Indicate that balance has been set up from the ioctl and is in the
* main phase. The fs_info::balance_ctl is initialized.
*/
BTRFS_FS_BALANCE_RUNNING,
/* Indicate that the cleaner thread is awake and doing something. */
BTRFS_FS_CLEANER_RUNNING,
};
struct btrfs_fs_info {
u8 chunk_tree_uuid[BTRFS_UUID_SIZE];
unsigned long flags;
struct btrfs_root *extent_root;
struct btrfs_root *tree_root;
struct btrfs_root *chunk_root;
struct btrfs_root *dev_root;
struct btrfs_root *fs_root;
Btrfs: move data checksumming into a dedicated tree Btrfs stores checksums for each data block. Until now, they have been stored in the subvolume trees, indexed by the inode that is referencing the data block. This means that when we read the inode, we've probably read in at least some checksums as well. But, this has a few problems: * The checksums are indexed by logical offset in the file. When compression is on, this means we have to do the expensive checksumming on the uncompressed data. It would be faster if we could checksum the compressed data instead. * If we implement encryption, we'll be checksumming the plain text and storing that on disk. This is significantly less secure. * For either compression or encryption, we have to get the plain text back before we can verify the checksum as correct. This makes the raid layer balancing and extent moving much more expensive. * It makes the front end caching code more complex, as we have touch the subvolume and inodes as we cache extents. * There is potentitally one copy of the checksum in each subvolume referencing an extent. The solution used here is to store the extent checksums in a dedicated tree. This allows us to index the checksums by phyiscal extent start and length. It means: * The checksum is against the data stored on disk, after any compression or encryption is done. * The checksum is stored in a central location, and can be verified without following back references, or reading inodes. This makes compression significantly faster by reducing the amount of data that needs to be checksummed. It will also allow much faster raid management code in general. The checksums are indexed by a key with a fixed objectid (a magic value in ctree.h) and offset set to the starting byte of the extent. This allows us to copy the checksum items into the fsync log tree directly (or any other tree), without having to invent a second format for them. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-12-08 21:58:54 +00:00
struct btrfs_root *csum_root;
struct btrfs_root *quota_root;
struct btrfs_root *uuid_root;
struct btrfs_root *free_space_root;
/* the log root tree is a directory of all the other log roots */
struct btrfs_root *log_root_tree;
spinlock_t fs_roots_radix_lock;
struct radix_tree_root fs_roots_radix;
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
/* block group cache stuff */
spinlock_t block_group_cache_lock;
u64 first_logical_byte;
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
struct rb_root block_group_cache_tree;
/* keep track of unallocated space */
atomic64_t free_chunk_space;
struct extent_io_tree freed_extents[2];
struct extent_io_tree *pinned_extents;
/* logical->physical extent mapping */
struct btrfs_mapping_tree mapping_tree;
btrfs: implement delayed inode items operation Changelog V5 -> V6: - Fix oom when the memory load is high, by storing the delayed nodes into the root's radix tree, and letting btrfs inodes go. Changelog V4 -> V5: - Fix the race on adding the delayed node to the inode, which is spotted by Chris Mason. - Merge Chris Mason's incremental patch into this patch. - Fix deadlock between readdir() and memory fault, which is reported by Itaru Kitayama. Changelog V3 -> V4: - Fix nested lock, which is reported by Itaru Kitayama, by updating space cache inode in time. Changelog V2 -> V3: - Fix the race between the delayed worker and the task which does delayed items balance, which is reported by Tsutomu Itoh. - Modify the patch address David Sterba's comment. - Fix the bug of the cpu recursion spinlock, reported by Chris Mason Changelog V1 -> V2: - break up the global rb-tree, use a list to manage the delayed nodes, which is created for every directory and file, and used to manage the delayed directory name index items and the delayed inode item. - introduce a worker to deal with the delayed nodes. Compare with Ext3/4, the performance of file creation and deletion on btrfs is very poor. the reason is that btrfs must do a lot of b+ tree insertions, such as inode item, directory name item, directory name index and so on. If we can do some delayed b+ tree insertion or deletion, we can improve the performance, so we made this patch which implemented delayed directory name index insertion/deletion and delayed inode update. Implementation: - introduce a delayed root object into the filesystem, that use two lists to manage the delayed nodes which are created for every file/directory. One is used to manage all the delayed nodes that have delayed items. And the other is used to manage the delayed nodes which is waiting to be dealt with by the work thread. - Every delayed node has two rb-tree, one is used to manage the directory name index which is going to be inserted into b+ tree, and the other is used to manage the directory name index which is going to be deleted from b+ tree. - introduce a worker to deal with the delayed operation. This worker is used to deal with the works of the delayed directory name index items insertion and deletion and the delayed inode update. When the delayed items is beyond the lower limit, we create works for some delayed nodes and insert them into the work queue of the worker, and then go back. When the delayed items is beyond the upper bound, we create works for all the delayed nodes that haven't been dealt with, and insert them into the work queue of the worker, and then wait for that the untreated items is below some threshold value. - When we want to insert a directory name index into b+ tree, we just add the information into the delayed inserting rb-tree. And then we check the number of the delayed items and do delayed items balance. (The balance policy is above.) - When we want to delete a directory name index from the b+ tree, we search it in the inserting rb-tree at first. If we look it up, just drop it. If not, add the key of it into the delayed deleting rb-tree. Similar to the delayed inserting rb-tree, we also check the number of the delayed items and do delayed items balance. (The same to inserting manipulation) - When we want to update the metadata of some inode, we cached the data of the inode into the delayed node. the worker will flush it into the b+ tree after dealing with the delayed insertion and deletion. - We will move the delayed node to the tail of the list after we access the delayed node, By this way, we can cache more delayed items and merge more inode updates. - If we want to commit transaction, we will deal with all the delayed node. - the delayed node will be freed when we free the btrfs inode. - Before we log the inode items, we commit all the directory name index items and the delayed inode update. I did a quick test by the benchmark tool[1] and found we can improve the performance of file creation by ~15%, and file deletion by ~20%. Before applying this patch: Create files: Total files: 50000 Total time: 1.096108 Average time: 0.000022 Delete files: Total files: 50000 Total time: 1.510403 Average time: 0.000030 After applying this patch: Create files: Total files: 50000 Total time: 0.932899 Average time: 0.000019 Delete files: Total files: 50000 Total time: 1.215732 Average time: 0.000024 [1] http://marc.info/?l=linux-btrfs&m=128212635122920&q=p3 Many thanks for Kitayama-san's help! Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Reviewed-by: David Sterba <dave@jikos.cz> Tested-by: Tsutomu Itoh <t-itoh@jp.fujitsu.com> Tested-by: Itaru Kitayama <kitayama@cl.bb4u.ne.jp> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2011-04-22 10:12:22 +00:00
/*
* block reservation for extent, checksum, root tree and
* delayed dir index item
*/
struct btrfs_block_rsv global_block_rsv;
/* block reservation for metadata operations */
struct btrfs_block_rsv trans_block_rsv;
/* block reservation for chunk tree */
struct btrfs_block_rsv chunk_block_rsv;
/* block reservation for delayed operations */
struct btrfs_block_rsv delayed_block_rsv;
btrfs: introduce delayed_refs_rsv Traditionally we've had voodoo in btrfs to account for the space that delayed refs may take up by having a global_block_rsv. This works most of the time, except when it doesn't. We've had issues reported and seen in production where sometimes the global reserve is exhausted during transaction commit before we can run all of our delayed refs, resulting in an aborted transaction. Because of this voodoo we have equally dubious flushing semantics around throttling delayed refs which we often get wrong. So instead give them their own block_rsv. This way we can always know exactly how much outstanding space we need for delayed refs. This allows us to make sure we are constantly filling that reservation up with space, and allows us to put more precise pressure on the enospc system. Instead of doing math to see if its a good time to throttle, the normal enospc code will be invoked if we have a lot of delayed refs pending, and they will be run via the normal flushing mechanism. For now the delayed_refs_rsv will hold the reservations for the delayed refs, the block group updates, and deleting csums. We could have a separate rsv for the block group updates, but the csum deletion stuff is still handled via the delayed_refs so that will stay there. Historical background: The global reserve has grown to cover everything we don't reserve space explicitly for, and we've grown a lot of weird ad-hoc heuristics to know if we're running short on space and when it's time to force a commit. A failure rate of 20-40 file systems when we run hundreds of thousands of them isn't super high, but cleaning up this code will make things less ugly and more predictible. Thus the delayed refs rsv. We always know how many delayed refs we have outstanding, and although running them generates more we can use the global reserve for that spill over, which fits better into it's desired use than a full blown reservation. This first approach is to simply take how many times we're reserving space for and multiply that by 2 in order to save enough space for the delayed refs that could be generated. This is a niave approach and will probably evolve, but for now it works. Signed-off-by: Josef Bacik <jbacik@fb.com> Reviewed-by: David Sterba <dsterba@suse.com> # high-level review [ added background notes from the cover letter ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-12-03 15:20:33 +00:00
/* block reservation for delayed refs */
struct btrfs_block_rsv delayed_refs_rsv;
struct btrfs_block_rsv empty_block_rsv;
u64 generation;
u64 last_trans_committed;
u64 avg_delayed_ref_runtime;
Btrfs: tree logging unlink/rename fixes The tree logging code allows individual files or directories to be logged without including operations on other files and directories in the FS. It tries to commit the minimal set of changes to disk in order to fsync the single file or directory that was sent to fsync or O_SYNC. The tree logging code was allowing files and directories to be unlinked if they were part of a rename operation where only one directory in the rename was in the fsync log. This patch adds a few new rules to the tree logging. 1) on rename or unlink, if the inode being unlinked isn't in the fsync log, we must force a full commit before doing an fsync of the directory where the unlink was done. The commit isn't done during the unlink, but it is forced the next time we try to log the parent directory. Solution: record transid of last unlink/rename per directory when the directory wasn't already logged. For renames this is only done when renaming to a different directory. mkdir foo/some_dir normal commit rename foo/some_dir foo2/some_dir mkdir foo/some_dir fsync foo/some_dir/some_file The fsync above will unlink the original some_dir without recording it in its new location (foo2). After a crash, some_dir will be gone unless the fsync of some_file forces a full commit 2) we must log any new names for any file or dir that is in the fsync log. This way we make sure not to lose files that are unlinked during the same transaction. 2a) we must log any new names for any file or dir during rename when the directory they are being removed from was logged. 2a is actually the more important variant. Without the extra logging a crash might unlink the old name without recreating the new one 3) after a crash, we must go through any directories with a link count of zero and redo the rm -rf mkdir f1/foo normal commit rm -rf f1/foo fsync(f1) The directory f1 was fully removed from the FS, but fsync was never called on f1, only its parent dir. After a crash the rm -rf must be replayed. This must be able to recurse down the entire directory tree. The inode link count fixup code takes care of the ugly details. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-03-24 14:24:20 +00:00
/*
* this is updated to the current trans every time a full commit
* is required instead of the faster short fsync log commits
*/
u64 last_trans_log_full_commit;
unsigned long mount_opt;
/*
* Track requests for actions that need to be done during transaction
* commit (like for some mount options).
*/
unsigned long pending_changes;
unsigned long compress_type:4;
unsigned int compress_level;
u32 commit_interval;
/*
* It is a suggestive number, the read side is safe even it gets a
* wrong number because we will write out the data into a regular
* extent. The write side(mount/remount) is under ->s_umount lock,
* so it is also safe.
*/
u64 max_inline;
struct btrfs_transaction *running_transaction;
wait_queue_head_t transaction_throttle;
wait_queue_head_t transaction_wait;
wait_queue_head_t transaction_blocked_wait;
wait_queue_head_t async_submit_wait;
/*
* Used to protect the incompat_flags, compat_flags, compat_ro_flags
* when they are updated.
*
* Because we do not clear the flags for ever, so we needn't use
* the lock on the read side.
*
* We also needn't use the lock when we mount the fs, because
* there is no other task which will update the flag.
*/
spinlock_t super_lock;
struct btrfs_super_block *super_copy;
struct btrfs_super_block *super_for_commit;
struct super_block *sb;
struct inode *btree_inode;
struct mutex tree_log_mutex;
struct mutex transaction_kthread_mutex;
struct mutex cleaner_mutex;
struct mutex chunk_mutex;
/*
* this is taken to make sure we don't set block groups ro after
* the free space cache has been allocated on them
*/
struct mutex ro_block_group_mutex;
/* this is used during read/modify/write to make sure
* no two ios are trying to mod the same stripe at the same
* time
*/
struct btrfs_stripe_hash_table *stripe_hash_table;
Btrfs: add extra flushing for renames and truncates Renames and truncates are both common ways to replace old data with new data. The filesystem can make an effort to make sure the new data is on disk before actually replacing the old data. This is especially important for rename, which many application use as though it were atomic for both the data and the metadata involved. The current btrfs code will happily replace a file that is fully on disk with one that was just created and still has pending IO. If we crash after transaction commit but before the IO is done, we'll end up replacing a good file with a zero length file. The solution used here is to create a list of inodes that need special ordering and force them to disk before the commit is done. This is similar to the ext3 style data=ordering, except it is only done on selected files. Btrfs is able to get away with this because it does not wait on commits very often, even for fsync (which use a sub-commit). For renames, we order the file when it wasn't already on disk and when it is replacing an existing file. Larger files are sent to filemap_flush right away (before the transaction handle is opened). For truncates, we order if the file goes from non-zero size down to zero size. This is a little different, because at the time of the truncate the file has no dirty bytes to order. But, we flag the inode so that it is added to the ordered list on close (via release method). We also immediately add it to the ordered list of the current transaction so that we can try to flush down any writes the application sneaks in before commit. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-03-31 17:27:11 +00:00
/*
* this protects the ordered operations list only while we are
* processing all of the entries on it. This way we make
* sure the commit code doesn't find the list temporarily empty
* because another function happens to be doing non-waiting preflush
* before jumping into the main commit.
*/
struct mutex ordered_operations_mutex;
struct rw_semaphore commit_root_sem;
Btrfs: add extra flushing for renames and truncates Renames and truncates are both common ways to replace old data with new data. The filesystem can make an effort to make sure the new data is on disk before actually replacing the old data. This is especially important for rename, which many application use as though it were atomic for both the data and the metadata involved. The current btrfs code will happily replace a file that is fully on disk with one that was just created and still has pending IO. If we crash after transaction commit but before the IO is done, we'll end up replacing a good file with a zero length file. The solution used here is to create a list of inodes that need special ordering and force them to disk before the commit is done. This is similar to the ext3 style data=ordering, except it is only done on selected files. Btrfs is able to get away with this because it does not wait on commits very often, even for fsync (which use a sub-commit). For renames, we order the file when it wasn't already on disk and when it is replacing an existing file. Larger files are sent to filemap_flush right away (before the transaction handle is opened). For truncates, we order if the file goes from non-zero size down to zero size. This is a little different, because at the time of the truncate the file has no dirty bytes to order. But, we flag the inode so that it is added to the ordered list on close (via release method). We also immediately add it to the ordered list of the current transaction so that we can try to flush down any writes the application sneaks in before commit. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-03-31 17:27:11 +00:00
struct rw_semaphore cleanup_work_sem;
struct rw_semaphore subvol_sem;
struct srcu_struct subvol_srcu;
Btrfs: kill trans_mutex We use trans_mutex for lots of things, here's a basic list 1) To serialize trans_handles joining the currently running transaction 2) To make sure that no new trans handles are started while we are committing 3) To protect the dead_roots list and the transaction lists Really the serializing trans_handles joining is not too hard, and can really get bogged down in acquiring a reference to the transaction. So replace the trans_mutex with a trans_lock spinlock and use it to do the following 1) Protect fs_info->running_transaction. All trans handles have to do is check this, and then take a reference of the transaction and keep on going. 2) Protect the fs_info->trans_list. This doesn't get used too much, basically it just holds the current transactions, which will usually just be the currently committing transaction and the currently running transaction at most. 3) Protect the dead roots list. This is only ever processed by splicing the list so this is relatively simple. 4) Protect the fs_info->reloc_ctl stuff. This is very lightweight and was using the trans_mutex before, so this is a pretty straightforward change. 5) Protect fs_info->no_trans_join. Because we don't hold the trans_lock over the entirety of the commit we need to have a way to block new people from creating a new transaction while we're doing our work. So we set no_trans_join and in join_transaction we test to see if that is set, and if it is we do a wait_on_commit. 6) Make the transaction use count atomic so we don't need to take locks to modify it when we're dropping references. 7) Add a commit_lock to the transaction to make sure multiple people trying to commit the same transaction don't race and commit at the same time. 8) Make open_ioctl_trans an atomic so we don't have to take any locks for ioctl trans. I have tested this with xfstests, but obviously it is a pretty hairy change so lots of testing is greatly appreciated. Thanks, Signed-off-by: Josef Bacik <josef@redhat.com>
2011-04-11 21:25:13 +00:00
spinlock_t trans_lock;
/*
* the reloc mutex goes with the trans lock, it is taken
* during commit to protect us from the relocation code
*/
struct mutex reloc_mutex;
struct list_head trans_list;
struct list_head dead_roots;
struct list_head caching_block_groups;
spinlock_t delayed_iput_lock;
struct list_head delayed_iputs;
atomic_t nr_delayed_iputs;
wait_queue_head_t delayed_iputs_wait;
/* this protects tree_mod_seq_list */
spinlock_t tree_mod_seq_lock;
atomic64_t tree_mod_seq;
struct list_head tree_mod_seq_list;
/* this protects tree_mod_log */
rwlock_t tree_mod_log_lock;
struct rb_root tree_mod_log;
atomic_t async_delalloc_pages;
/*
* this is used to protect the following list -- ordered_roots.
*/
spinlock_t ordered_root_lock;
Btrfs: add extra flushing for renames and truncates Renames and truncates are both common ways to replace old data with new data. The filesystem can make an effort to make sure the new data is on disk before actually replacing the old data. This is especially important for rename, which many application use as though it were atomic for both the data and the metadata involved. The current btrfs code will happily replace a file that is fully on disk with one that was just created and still has pending IO. If we crash after transaction commit but before the IO is done, we'll end up replacing a good file with a zero length file. The solution used here is to create a list of inodes that need special ordering and force them to disk before the commit is done. This is similar to the ext3 style data=ordering, except it is only done on selected files. Btrfs is able to get away with this because it does not wait on commits very often, even for fsync (which use a sub-commit). For renames, we order the file when it wasn't already on disk and when it is replacing an existing file. Larger files are sent to filemap_flush right away (before the transaction handle is opened). For truncates, we order if the file goes from non-zero size down to zero size. This is a little different, because at the time of the truncate the file has no dirty bytes to order. But, we flag the inode so that it is added to the ordered list on close (via release method). We also immediately add it to the ordered list of the current transaction so that we can try to flush down any writes the application sneaks in before commit. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-03-31 17:27:11 +00:00
/*
* all fs/file tree roots in which there are data=ordered extents
* pending writeback are added into this list.
*
Btrfs: add extra flushing for renames and truncates Renames and truncates are both common ways to replace old data with new data. The filesystem can make an effort to make sure the new data is on disk before actually replacing the old data. This is especially important for rename, which many application use as though it were atomic for both the data and the metadata involved. The current btrfs code will happily replace a file that is fully on disk with one that was just created and still has pending IO. If we crash after transaction commit but before the IO is done, we'll end up replacing a good file with a zero length file. The solution used here is to create a list of inodes that need special ordering and force them to disk before the commit is done. This is similar to the ext3 style data=ordering, except it is only done on selected files. Btrfs is able to get away with this because it does not wait on commits very often, even for fsync (which use a sub-commit). For renames, we order the file when it wasn't already on disk and when it is replacing an existing file. Larger files are sent to filemap_flush right away (before the transaction handle is opened). For truncates, we order if the file goes from non-zero size down to zero size. This is a little different, because at the time of the truncate the file has no dirty bytes to order. But, we flag the inode so that it is added to the ordered list on close (via release method). We also immediately add it to the ordered list of the current transaction so that we can try to flush down any writes the application sneaks in before commit. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-03-31 17:27:11 +00:00
* these can span multiple transactions and basically include
* every dirty data page that isn't from nodatacow
*/
struct list_head ordered_roots;
Btrfs: add extra flushing for renames and truncates Renames and truncates are both common ways to replace old data with new data. The filesystem can make an effort to make sure the new data is on disk before actually replacing the old data. This is especially important for rename, which many application use as though it were atomic for both the data and the metadata involved. The current btrfs code will happily replace a file that is fully on disk with one that was just created and still has pending IO. If we crash after transaction commit but before the IO is done, we'll end up replacing a good file with a zero length file. The solution used here is to create a list of inodes that need special ordering and force them to disk before the commit is done. This is similar to the ext3 style data=ordering, except it is only done on selected files. Btrfs is able to get away with this because it does not wait on commits very often, even for fsync (which use a sub-commit). For renames, we order the file when it wasn't already on disk and when it is replacing an existing file. Larger files are sent to filemap_flush right away (before the transaction handle is opened). For truncates, we order if the file goes from non-zero size down to zero size. This is a little different, because at the time of the truncate the file has no dirty bytes to order. But, we flag the inode so that it is added to the ordered list on close (via release method). We also immediately add it to the ordered list of the current transaction so that we can try to flush down any writes the application sneaks in before commit. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-03-31 17:27:11 +00:00
struct mutex delalloc_root_mutex;
spinlock_t delalloc_root_lock;
/* all fs/file tree roots that have delalloc inodes. */
struct list_head delalloc_roots;
/*
* there is a pool of worker threads for checksumming during writes
* and a pool for checksumming after reads. This is because readers
* can run with FS locks held, and the writers may be waiting for
* those locks. We don't want ordering in the pending list to cause
* deadlocks, and so the two are serviced separately.
*
* A third pool does submit_bio to avoid deadlocking with the other
* two
*/
struct btrfs_workqueue *workers;
struct btrfs_workqueue *delalloc_workers;
struct btrfs_workqueue *flush_workers;
struct btrfs_workqueue *endio_workers;
struct btrfs_workqueue *endio_meta_workers;
struct btrfs_workqueue *endio_raid56_workers;
struct btrfs_workqueue *endio_repair_workers;
struct btrfs_workqueue *rmw_workers;
struct btrfs_workqueue *endio_meta_write_workers;
struct btrfs_workqueue *endio_write_workers;
struct btrfs_workqueue *endio_freespace_worker;
struct btrfs_workqueue *submit_workers;
struct btrfs_workqueue *caching_workers;
struct btrfs_workqueue *readahead_workers;
/*
* fixup workers take dirty pages that didn't properly go through
* the cow mechanism and make them safe to write. It happens
* for the sys_munmap function call path
*/
struct btrfs_workqueue *fixup_workers;
struct btrfs_workqueue *delayed_workers;
/* the extent workers do delayed refs on the extent allocation tree */
struct btrfs_workqueue *extent_workers;
struct task_struct *transaction_kthread;
struct task_struct *cleaner_kthread;
u32 thread_pool_size;
struct kobject *space_info_kobj;
struct list_head pending_raid_kobjs;
spinlock_t pending_raid_kobjs_lock; /* uncontended */
u64 total_pinned;
/* used to keep from writing metadata until there is a nice batch */
struct percpu_counter dirty_metadata_bytes;
struct percpu_counter delalloc_bytes;
s32 dirty_metadata_batch;
s32 delalloc_batch;
struct list_head dirty_cowonly_roots;
struct btrfs_fs_devices *fs_devices;
/*
* The space_info list is effectively read only after initial
* setup. It is populated at mount time and cleaned up after
* all block groups are removed. RCU is used to protect it.
*/
struct list_head space_info;
struct btrfs_space_info *data_sinfo;
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
struct reloc_control *reloc_ctl;
btrfs: Do not use data_alloc_cluster in ssd mode This patch provides a band aid to improve the 'out of the box' behaviour of btrfs for disks that are detected as being an ssd. In a general purpose mixed workload scenario, the current ssd mode causes overallocation of available raw disk space for data, while leaving behind increasing amounts of unused fragmented free space. This situation leads to early ENOSPC problems which are harming user experience and adoption of btrfs as a general purpose filesystem. This patch modifies the data extent allocation behaviour of the ssd mode to make it behave identical to nossd mode. The metadata behaviour and additional ssd_spread option stay untouched so far. Recommendations for future development are to reconsider the current oversimplified nossd / ssd distinction and the broken detection mechanism based on the rotational attribute in sysfs and provide experienced users with a more flexible way to choose allocator behaviour for data and metadata, optimized for certain use cases, while keeping sane 'out of the box' default settings. The internals of the current btrfs code have more potential than what currently gets exposed to the user to choose from. The SSD story... In the first year of btrfs development, around early 2008, btrfs gained a mount option which enables specific functionality for filesystems on solid state devices. The first occurance of this functionality is in commit e18e4809, labeled "Add mount -o ssd, which includes optimizations for seek free storage". The effect on allocating free space for doing (data) writes is to 'cluster' writes together, writing them out in contiguous space, as opposed to a 'tetris' way of putting all separate writes into any free space fragment that fits (which is what the -o nossd behaviour does). A somewhat simplified explanation of what happens is that, when for example, the 'cluster' size is set to 2MiB, when we do some writes, the data allocator will search for a free space block that is 2MiB big, and put the writes in there. The ssd mode itself might allow a 2MiB cluster to be composed of multiple free space extents with some existing data in between, while the additional ssd_spread mount option kills off this option and requires fully free space. The idea behind this is (commit 536ac8ae): "The [...] clusters make it more likely a given IO will completely overwrite the ssd block, so it doesn't have to do an internal rwm cycle."; ssd block meaning nand erase block. So, effectively this means applying a "locality based algorithm" and trying to outsmart the actual ssd. Since then, various changes have been made to the involved code, but the basic idea is still present, and gets activated whenever the ssd mount option is active. This also happens by default, when the rotational flag as seen at /sys/block/<device>/queue/rotational is set to 0. However, there's a number of problems with this approach. First, what the optimization is trying to do is outsmart the ssd by assuming there is a relation between the physical address space of the block device as seen by btrfs and the actual physical storage of the ssd, and then adjusting data placement. However, since the introduction of the Flash Translation Layer (FTL) which is a part of the internal controller of an ssd, these attempts are futile. The use of good quality FTL in consumer ssd products might have been limited in 2008, but this situation has changed drastically soon after that time. Today, even the flash memory in your automatic cat feeding machine or your grandma's wheelchair has a full featured one. Second, the behaviour as described above results in the filesystem being filled up with badly fragmented free space extents because of relatively small pieces of space that are freed up by deletes, but not selected again as part of a 'cluster'. Since the algorithm prefers allocating a new chunk over going back to tetris mode, the end result is a filesystem in which all raw space is allocated, but which is composed of underutilized chunks with a 'shotgun blast' pattern of fragmented free space. Usually, the next problematic thing that happens is the filesystem wanting to allocate new space for metadata, which causes the filesystem to fail in spectacular ways. Third, the default mount options you get for an ssd ('ssd' mode enabled, 'discard' not enabled), in combination with spreading out writes over the full address space and ignoring freed up space leads to worst case behaviour in providing information to the ssd itself, since it will never learn that all the free space left behind is actually free. There are two ways to let an ssd know previously written data does not have to be preserved, which are sending explicit signals using discard or fstrim, or by simply overwriting the space with new data. The worst case behaviour is the btrfs ssd_spread mount option in combination with not having discard enabled. It has a side effect of minimizing the reuse of free space previously written in. Fourth, the rotational flag in /sys/ does not reliably indicate if the device is a locally attached ssd. For example, iSCSI or NBD displays as non-rotational, while a loop device on an ssd shows up as rotational. The combination of the second and third problem effectively means that despite all the good intentions, the btrfs ssd mode reliably causes the ssd hardware and the filesystem structures and performance to be choked to death. The clickbait version of the title of this story would have been "Btrfs ssd optimizations considered harmful for ssds". The current nossd 'tetris' mode (even still without discard) allows a pattern of overwriting much more previously used space, causing many more implicit discards to happen because of the overwrite information the ssd gets. The actual location in the physical address space, as seen from the point of view of btrfs is irrelevant, because the actual writes to the low level flash are reordered anyway thanks to the FTL. Changes made in the code 1. Make ssd mode data allocation identical to tetris mode, like nossd. 2. Adjust and clean up filesystem mount messages so that we can easily identify if a kernel has this patch applied or not, when providing support to end users. Also, make better use of the *_and_info helpers to only trigger messages on actual state changes. Backporting notes Notes for whoever wants to backport this patch to their 4.9 LTS kernel: * First apply commit 951e7966 "btrfs: drop the nossd flag when remounting with -o ssd", or fixup the differences manually. * The rest of the conflicts are because of the fs_info refactoring. So, for example, instead of using fs_info, it's root->fs_info in extent-tree.c Signed-off-by: Hans van Kranenburg <hans.van.kranenburg@mendix.com> Signed-off-by: David Sterba <dsterba@suse.com>
2017-07-28 06:31:28 +00:00
/* data_alloc_cluster is only used in ssd_spread mode */
struct btrfs_free_cluster data_alloc_cluster;
/* all metadata allocations go through this cluster */
struct btrfs_free_cluster meta_alloc_cluster;
/* auto defrag inodes go here */
spinlock_t defrag_inodes_lock;
struct rb_root defrag_inodes;
atomic_t defrag_running;
/* Used to protect avail_{data, metadata, system}_alloc_bits */
seqlock_t profiles_lock;
/*
* these three are in extended format (availability of single
* chunks is denoted by BTRFS_AVAIL_ALLOC_BIT_SINGLE bit, other
* types are denoted by corresponding BTRFS_BLOCK_GROUP_* bits)
*/
u64 avail_data_alloc_bits;
u64 avail_metadata_alloc_bits;
u64 avail_system_alloc_bits;
/* restriper state */
spinlock_t balance_lock;
struct mutex balance_mutex;
atomic_t balance_pause_req;
atomic_t balance_cancel_req;
struct btrfs_balance_control *balance_ctl;
wait_queue_head_t balance_wait_q;
u32 data_chunk_allocations;
u32 metadata_ratio;
void *bdev_holder;
/* private scrub information */
struct mutex scrub_lock;
atomic_t scrubs_running;
atomic_t scrub_pause_req;
atomic_t scrubs_paused;
atomic_t scrub_cancel_req;
wait_queue_head_t scrub_pause_wait;
/*
* The worker pointers are NULL iff the refcount is 0, ie. scrub is not
* running.
*/
refcount_t scrub_workers_refcnt;
struct btrfs_workqueue *scrub_workers;
struct btrfs_workqueue *scrub_wr_completion_workers;
btrfs: Fix lockdep warning of wr_ctx->wr_lock in scrub_free_wr_ctx() lockdep report following warning in test: [25176.843958] ================================= [25176.844519] [ INFO: inconsistent lock state ] [25176.845047] 4.1.0-rc3 #22 Tainted: G W [25176.845591] --------------------------------- [25176.846153] inconsistent {SOFTIRQ-ON-W} -> {IN-SOFTIRQ-W} usage. [25176.846713] fsstress/26661 [HC0[0]:SC1[1]:HE1:SE0] takes: [25176.847246] (&wr_ctx->wr_lock){+.?...}, at: [<ffffffffa04cdc6d>] scrub_free_ctx+0x2d/0xf0 [btrfs] [25176.847838] {SOFTIRQ-ON-W} state was registered at: [25176.848396] [<ffffffff810bf460>] __lock_acquire+0x6a0/0xe10 [25176.848955] [<ffffffff810bfd1e>] lock_acquire+0xce/0x2c0 [25176.849491] [<ffffffff816489af>] mutex_lock_nested+0x7f/0x410 [25176.850029] [<ffffffffa04d04ff>] scrub_stripe+0x4df/0x1080 [btrfs] [25176.850575] [<ffffffffa04d11b1>] scrub_chunk.isra.19+0x111/0x130 [btrfs] [25176.851110] [<ffffffffa04d144c>] scrub_enumerate_chunks+0x27c/0x510 [btrfs] [25176.851660] [<ffffffffa04d3b87>] btrfs_scrub_dev+0x1c7/0x6c0 [btrfs] [25176.852189] [<ffffffffa04e918e>] btrfs_dev_replace_start+0x36e/0x450 [btrfs] [25176.852771] [<ffffffffa04a98e0>] btrfs_ioctl+0x1e10/0x2d20 [btrfs] [25176.853315] [<ffffffff8121c5b8>] do_vfs_ioctl+0x318/0x570 [25176.853868] [<ffffffff8121c851>] SyS_ioctl+0x41/0x80 [25176.854406] [<ffffffff8164da17>] system_call_fastpath+0x12/0x6f [25176.854935] irq event stamp: 51506 [25176.855511] hardirqs last enabled at (51506): [<ffffffff810d4ce5>] vprintk_emit+0x225/0x5e0 [25176.856059] hardirqs last disabled at (51505): [<ffffffff810d4b77>] vprintk_emit+0xb7/0x5e0 [25176.856642] softirqs last enabled at (50886): [<ffffffff81067a23>] __do_softirq+0x363/0x640 [25176.857184] softirqs last disabled at (50949): [<ffffffff8106804d>] irq_exit+0x10d/0x120 [25176.857746] other info that might help us debug this: [25176.858845] Possible unsafe locking scenario: [25176.859981] CPU0 [25176.860537] ---- [25176.861059] lock(&wr_ctx->wr_lock); [25176.861705] <Interrupt> [25176.862272] lock(&wr_ctx->wr_lock); [25176.862881] *** DEADLOCK *** Reason: Above warning is caused by: Interrupt -> bio_endio() -> ... -> scrub_put_ctx() -> scrub_free_ctx() *1 -> ... -> mutex_lock(&wr_ctx->wr_lock); scrub_put_ctx() is allowed to be called in end_bio interrupt, but in code design, it will never call scrub_free_ctx(sctx) in interrupe context(above *1), because btrfs_scrub_dev() get one additional reference of sctx->refs, which makes scrub_free_ctx() only called withine btrfs_scrub_dev(). Now the code runs out of our wish, because free sequence in scrub_pending_bio_dec() have a gap. Current code: -----------------------------------+----------------------------------- scrub_pending_bio_dec() | btrfs_scrub_dev -----------------------------------+----------------------------------- atomic_dec(&sctx->bios_in_flight); | wake_up(&sctx->list_wait); | | scrub_put_ctx() | -> atomic_dec_and_test(&sctx->refs) scrub_put_ctx(sctx); | -> atomic_dec_and_test(&sctx->refs)| -> scrub_free_ctx() | -----------------------------------+----------------------------------- We expected: -----------------------------------+----------------------------------- scrub_pending_bio_dec() | btrfs_scrub_dev -----------------------------------+----------------------------------- atomic_dec(&sctx->bios_in_flight); | wake_up(&sctx->list_wait); | scrub_put_ctx(sctx); | -> atomic_dec_and_test(&sctx->refs)| | scrub_put_ctx() | -> atomic_dec_and_test(&sctx->refs) | -> scrub_free_ctx() -----------------------------------+----------------------------------- Fix: Move scrub_pending_bio_dec() to a workqueue, to avoid this function run in interrupt context. Tested by check tracelog in debug. Changelog v1->v2: Use workqueue instead of adjust function call sequence in v1, because v1 will introduce a bug pointed out by: Filipe David Manana <fdmanana@gmail.com> Reported-by: Qu Wenruo <quwenruo@cn.fujitsu.com> Signed-off-by: Zhao Lei <zhaolei@cn.fujitsu.com> Reviewed-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2015-06-04 12:09:15 +00:00
struct btrfs_workqueue *scrub_parity_workers;
#ifdef CONFIG_BTRFS_FS_CHECK_INTEGRITY
u32 check_integrity_print_mask;
#endif
/* is qgroup tracking in a consistent state? */
u64 qgroup_flags;
/* holds configuration and tracking. Protected by qgroup_lock */
struct rb_root qgroup_tree;
spinlock_t qgroup_lock;
/*
* used to avoid frequently calling ulist_alloc()/ulist_free()
* when doing qgroup accounting, it must be protected by qgroup_lock.
*/
struct ulist *qgroup_ulist;
/* protect user change for quota operations */
struct mutex qgroup_ioctl_lock;
/* list of dirty qgroups to be written at next commit */
struct list_head dirty_qgroups;
/* used by qgroup for an efficient tree traversal */
u64 qgroup_seq;
/* qgroup rescan items */
struct mutex qgroup_rescan_lock; /* protects the progress item */
struct btrfs_key qgroup_rescan_progress;
struct btrfs_workqueue *qgroup_rescan_workers;
struct completion qgroup_rescan_completion;
Btrfs: fix qgroup rescan resume on mount When called during mount, we cannot start the rescan worker thread until open_ctree is done. This commit restuctures the qgroup rescan internals to enable a clean deferral of the rescan resume operation. First of all, the struct qgroup_rescan is removed, saving us a malloc and some initialization synchronizations problems. Its only element (the worker struct) now lives within fs_info just as the rest of the rescan code. Then setting up a rescan worker is split into several reusable stages. Currently we have three different rescan startup scenarios: (A) rescan ioctl (B) rescan resume by mount (C) rescan by quota enable Each case needs its own combination of the four following steps: (1) set the progress [A, C: zero; B: state of umount] (2) commit the transaction [A] (3) set the counters [A, C: zero; B: state of umount] (4) start worker [A, B, C] qgroup_rescan_init does step (1). There's no extra function added to commit a transaction, we've got that already. qgroup_rescan_zero_tracking does step (3). Step (4) is nothing more than a call to the generic btrfs_queue_worker. We also get rid of a double check for the rescan progress during btrfs_qgroup_account_ref, which is no longer required due to having step 2 from the list above. As a side effect, this commit prepares to move the rescan start code from btrfs_run_qgroups (which is run during commit) to a less time critical section. Signed-off-by: Jan Schmidt <list.btrfs@jan-o-sch.net> Signed-off-by: Josef Bacik <jbacik@fusionio.com>
2013-05-28 15:47:24 +00:00
struct btrfs_work qgroup_rescan_work;
bool qgroup_rescan_running; /* protected by qgroup_rescan_lock */
/* filesystem state */
unsigned long fs_state;
btrfs: implement delayed inode items operation Changelog V5 -> V6: - Fix oom when the memory load is high, by storing the delayed nodes into the root's radix tree, and letting btrfs inodes go. Changelog V4 -> V5: - Fix the race on adding the delayed node to the inode, which is spotted by Chris Mason. - Merge Chris Mason's incremental patch into this patch. - Fix deadlock between readdir() and memory fault, which is reported by Itaru Kitayama. Changelog V3 -> V4: - Fix nested lock, which is reported by Itaru Kitayama, by updating space cache inode in time. Changelog V2 -> V3: - Fix the race between the delayed worker and the task which does delayed items balance, which is reported by Tsutomu Itoh. - Modify the patch address David Sterba's comment. - Fix the bug of the cpu recursion spinlock, reported by Chris Mason Changelog V1 -> V2: - break up the global rb-tree, use a list to manage the delayed nodes, which is created for every directory and file, and used to manage the delayed directory name index items and the delayed inode item. - introduce a worker to deal with the delayed nodes. Compare with Ext3/4, the performance of file creation and deletion on btrfs is very poor. the reason is that btrfs must do a lot of b+ tree insertions, such as inode item, directory name item, directory name index and so on. If we can do some delayed b+ tree insertion or deletion, we can improve the performance, so we made this patch which implemented delayed directory name index insertion/deletion and delayed inode update. Implementation: - introduce a delayed root object into the filesystem, that use two lists to manage the delayed nodes which are created for every file/directory. One is used to manage all the delayed nodes that have delayed items. And the other is used to manage the delayed nodes which is waiting to be dealt with by the work thread. - Every delayed node has two rb-tree, one is used to manage the directory name index which is going to be inserted into b+ tree, and the other is used to manage the directory name index which is going to be deleted from b+ tree. - introduce a worker to deal with the delayed operation. This worker is used to deal with the works of the delayed directory name index items insertion and deletion and the delayed inode update. When the delayed items is beyond the lower limit, we create works for some delayed nodes and insert them into the work queue of the worker, and then go back. When the delayed items is beyond the upper bound, we create works for all the delayed nodes that haven't been dealt with, and insert them into the work queue of the worker, and then wait for that the untreated items is below some threshold value. - When we want to insert a directory name index into b+ tree, we just add the information into the delayed inserting rb-tree. And then we check the number of the delayed items and do delayed items balance. (The balance policy is above.) - When we want to delete a directory name index from the b+ tree, we search it in the inserting rb-tree at first. If we look it up, just drop it. If not, add the key of it into the delayed deleting rb-tree. Similar to the delayed inserting rb-tree, we also check the number of the delayed items and do delayed items balance. (The same to inserting manipulation) - When we want to update the metadata of some inode, we cached the data of the inode into the delayed node. the worker will flush it into the b+ tree after dealing with the delayed insertion and deletion. - We will move the delayed node to the tail of the list after we access the delayed node, By this way, we can cache more delayed items and merge more inode updates. - If we want to commit transaction, we will deal with all the delayed node. - the delayed node will be freed when we free the btrfs inode. - Before we log the inode items, we commit all the directory name index items and the delayed inode update. I did a quick test by the benchmark tool[1] and found we can improve the performance of file creation by ~15%, and file deletion by ~20%. Before applying this patch: Create files: Total files: 50000 Total time: 1.096108 Average time: 0.000022 Delete files: Total files: 50000 Total time: 1.510403 Average time: 0.000030 After applying this patch: Create files: Total files: 50000 Total time: 0.932899 Average time: 0.000019 Delete files: Total files: 50000 Total time: 1.215732 Average time: 0.000024 [1] http://marc.info/?l=linux-btrfs&m=128212635122920&q=p3 Many thanks for Kitayama-san's help! Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Reviewed-by: David Sterba <dave@jikos.cz> Tested-by: Tsutomu Itoh <t-itoh@jp.fujitsu.com> Tested-by: Itaru Kitayama <kitayama@cl.bb4u.ne.jp> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2011-04-22 10:12:22 +00:00
struct btrfs_delayed_root *delayed_root;
/* readahead tree */
spinlock_t reada_lock;
struct radix_tree_root reada_tree;
/* readahead works cnt */
atomic_t reada_works_cnt;
/* Extent buffer radix tree */
spinlock_t buffer_lock;
struct radix_tree_root buffer_radix;
/* next backup root to be overwritten */
int backup_root_index;
/* device replace state */
struct btrfs_dev_replace dev_replace;
struct semaphore uuid_tree_rescan_sem;
Btrfs: reclaim the reserved metadata space at background Before applying this patch, the task had to reclaim the metadata space by itself if the metadata space was not enough. And When the task started the space reclamation, all the other tasks which wanted to reserve the metadata space were blocked. At some cases, they would be blocked for a long time, it made the performance fluctuate wildly. So we introduce the background metadata space reclamation, when the space is about to be exhausted, we insert a reclaim work into the workqueue, the worker of the workqueue helps us to reclaim the reserved space at the background. By this way, the tasks needn't reclaim the space by themselves at most cases, and even if the tasks have to reclaim the space or are blocked for the space reclamation, they will get enough space more quickly. Here is my test result(Tested by compilebench): Memory: 2GB CPU: 2Cores * 1CPU Partition: 40GB(SSD) Test command: # compilebench -D <mnt> -m Without this patch: intial create total runs 30 avg 54.36 MB/s (user 0.52s sys 2.44s) compile total runs 30 avg 123.72 MB/s (user 0.13s sys 1.17s) read compiled tree total runs 3 avg 81.15 MB/s (user 0.74s sys 4.89s) delete compiled tree total runs 30 avg 5.32 seconds (user 0.35s sys 4.37s) With this patch: intial create total runs 30 avg 59.80 MB/s (user 0.52s sys 2.53s) compile total runs 30 avg 151.44 MB/s (user 0.13s sys 1.11s) read compiled tree total runs 3 avg 83.25 MB/s (user 0.76s sys 4.91s) delete compiled tree total runs 30 avg 5.29 seconds (user 0.34s sys 4.34s) Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-05-14 00:29:04 +00:00
/* Used to reclaim the metadata space in the background. */
struct work_struct async_reclaim_work;
spinlock_t unused_bgs_lock;
struct list_head unused_bgs;
Btrfs: fix race between transaction commit and empty block group removal Committing a transaction can race with automatic removal of empty block groups (cleaner kthread), leading to a BUG_ON() in the transaction commit code while running btrfs_finish_extent_commit(). The following sequence diagram shows how it can happen: CPU 1 CPU 2 btrfs_commit_transaction() fs_info->running_transaction = NULL btrfs_finish_extent_commit() find_first_extent_bit() -> found range for block group X in fs_info->freed_extents[] btrfs_delete_unused_bgs() -> found block group X Removed block group X's range from fs_info->freed_extents[] btrfs_remove_chunk() btrfs_remove_block_group(bg X) unpin_extent_range(bg X range) btrfs_lookup_block_group(bg X) -> returns NULL -> BUG_ON() The trace that results from the BUG_ON() is: [48665.187808] ------------[ cut here ]------------ [48665.188032] kernel BUG at fs/btrfs/extent-tree.c:5675! [48665.188032] invalid opcode: 0000 [#1] SMP DEBUG_PAGEALLOC [48665.188032] Modules linked in: dm_flakey dm_mod crc32c_generic btrfs xor raid6_pq nfsd auth_rpcgss oid_registry nfs_acl nfs lockd grace fscache sunrpc loop parport_pc evdev microcode [48665.197388] CPU: 2 PID: 31211 Comm: kworker/u32:16 Tainted: G W 3.19.0-rc5-btrfs-next-4+ #1 [48665.197388] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS rel-1.7.5-0-ge51488c-20140602_164612-nilsson.home.kraxel.org 04/01/2014 [48665.197388] Workqueue: events_unbound btrfs_async_reclaim_metadata_space [btrfs] [48665.197388] task: ffff880222011810 ti: ffff8801b56a4000 task.ti: ffff8801b56a4000 [48665.197388] RIP: 0010:[<ffffffffa0350d05>] [<ffffffffa0350d05>] unpin_extent_range+0x6a/0x1ba [btrfs] [48665.197388] RSP: 0018:ffff8801b56a7b88 EFLAGS: 00010246 [48665.197388] RAX: 0000000000000000 RBX: ffff8802143a6000 RCX: ffff8802220120c8 [48665.197388] RDX: 0000000000000001 RSI: 0000000000000001 RDI: ffff8800a3c140b0 [48665.197388] RBP: ffff8801b56a7bd8 R08: 0000000000000003 R09: 0000000000000000 [48665.197388] R10: 0000000000000000 R11: 000000000000bbac R12: 0000000012e8e000 [48665.197388] R13: ffff8800a3c14000 R14: 0000000000000000 R15: 0000000000000000 [48665.197388] FS: 0000000000000000(0000) GS:ffff88023ec40000(0000) knlGS:0000000000000000 [48665.197388] CS: 0010 DS: 0000 ES: 0000 CR0: 000000008005003b [48665.197388] CR2: 00007f065e42f270 CR3: 0000000206f70000 CR4: 00000000000006e0 [48665.197388] Stack: [48665.197388] ffff8801b56a7bd8 0000000012ea0000 01ff8800a3c14138 0000000012e9ffff [48665.197388] ffff880141df3dd8 ffff8802143a6000 ffff8800a3c14138 ffff880141df3df0 [48665.197388] ffff880141df3dd8 0000000000000000 ffff8801b56a7c08 ffffffffa0354227 [48665.197388] Call Trace: [48665.197388] [<ffffffffa0354227>] btrfs_finish_extent_commit+0xb0/0xd9 [btrfs] [48665.197388] [<ffffffffa0366b4b>] btrfs_commit_transaction+0x791/0x92c [btrfs] [48665.197388] [<ffffffffa0352432>] flush_space+0x43d/0x452 [btrfs] [48665.197388] [<ffffffff814295c3>] ? _raw_spin_unlock+0x28/0x33 [48665.197388] [<ffffffffa035255f>] btrfs_async_reclaim_metadata_space+0x118/0x164 [btrfs] [48665.197388] [<ffffffff81059917>] ? process_one_work+0x14b/0x3ab [48665.197388] [<ffffffff810599ac>] process_one_work+0x1e0/0x3ab [48665.197388] [<ffffffff81079fa9>] ? trace_hardirqs_off+0xd/0xf [48665.197388] [<ffffffff8105a55b>] worker_thread+0x210/0x2d0 [48665.197388] [<ffffffff8105a34b>] ? rescuer_thread+0x2c3/0x2c3 [48665.197388] [<ffffffff8105e5c0>] kthread+0xef/0xf7 [48665.197388] [<ffffffff81429682>] ? _raw_spin_unlock_irq+0x2d/0x39 [48665.197388] [<ffffffff8105e4d1>] ? __kthread_parkme+0xad/0xad [48665.197388] [<ffffffff81429dec>] ret_from_fork+0x7c/0xb0 [48665.197388] [<ffffffff8105e4d1>] ? __kthread_parkme+0xad/0xad [48665.197388] Code: 85 f6 74 14 49 8b 06 49 03 46 09 49 39 c4 72 1d 4c 89 f7 e8 83 ec ff ff 4c 89 e6 4c 89 ef e8 1e f1 ff ff 48 85 c0 49 89 c6 75 02 <0f> 0b 49 8b 1e 49 03 5e 09 48 8b [48665.197388] RIP [<ffffffffa0350d05>] unpin_extent_range+0x6a/0x1ba [btrfs] [48665.197388] RSP <ffff8801b56a7b88> [48665.272246] ---[ end trace b9c6ab9957521376 ]--- Fix this by ensuring that unpining the block group's range in btrfs_finish_extent_commit() is done in a synchronized fashion with removing the block group's range from freed_extents[] in btrfs_delete_unused_bgs() This race got introduced with the change: Btrfs: remove empty block groups automatically commit 47ab2a6c689913db23ccae38349714edf8365e0a Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2015-01-29 19:18:25 +00:00
struct mutex unused_bg_unpin_mutex;
Btrfs: fix race between balance and unused block group deletion We have a race between deleting an unused block group and balancing the same block group that leads to an assertion failure/BUG(), producing the following trace: [181631.208236] BTRFS: assertion failed: 0, file: fs/btrfs/volumes.c, line: 2622 [181631.220591] ------------[ cut here ]------------ [181631.222959] kernel BUG at fs/btrfs/ctree.h:4062! [181631.223932] invalid opcode: 0000 [#1] PREEMPT SMP DEBUG_PAGEALLOC [181631.224566] Modules linked in: btrfs dm_flakey dm_mod crc32c_generic xor raid6_pq nfsd auth_rpcgss oid_registry nfs_acl nfs lockd grace fscache sunrpc loop fuse acpi_cpufreq parpor$ [181631.224566] CPU: 8 PID: 17451 Comm: btrfs Tainted: G W 4.1.0-rc5-btrfs-next-10+ #1 [181631.224566] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS rel-1.8.1-0-g4adadbd-20150316_085822-nilsson.home.kraxel.org 04/01/2014 [181631.224566] task: ffff880127e09590 ti: ffff8800b5824000 task.ti: ffff8800b5824000 [181631.224566] RIP: 0010:[<ffffffffa03f19f6>] [<ffffffffa03f19f6>] assfail.constprop.50+0x1e/0x20 [btrfs] [181631.224566] RSP: 0018:ffff8800b5827ae8 EFLAGS: 00010246 [181631.224566] RAX: 0000000000000040 RBX: ffff8800109fc218 RCX: ffffffff81095dce [181631.224566] RDX: 0000000000005124 RSI: ffffffff81464819 RDI: 00000000ffffffff [181631.224566] RBP: ffff8800b5827ae8 R08: 0000000000000001 R09: 0000000000000000 [181631.224566] R10: 0000000000000000 R11: 0000000000000000 R12: ffff8800109fc200 [181631.224566] R13: ffff880020095000 R14: ffff8800b1a13f38 R15: ffff880020095000 [181631.224566] FS: 00007f70ca0b0c80(0000) GS:ffff88013ec00000(0000) knlGS:0000000000000000 [181631.224566] CS: 0010 DS: 0000 ES: 0000 CR0: 000000008005003b [181631.224566] CR2: 00007f2872ab6e68 CR3: 00000000a717c000 CR4: 00000000000006e0 [181631.224566] Stack: [181631.224566] ffff8800b5827ba8 ffffffffa03f3916 ffff8800b5827b38 ffffffffa03d080e [181631.224566] ffffffffa03d1423 ffff880020095000 ffff88001233c000 0000000000000001 [181631.224566] ffff880020095000 ffff8800b1a13f38 0000000a69c00000 0000000000000000 [181631.224566] Call Trace: [181631.224566] [<ffffffffa03f3916>] btrfs_remove_chunk+0xa4/0x6bb [btrfs] [181631.224566] [<ffffffffa03d080e>] ? join_transaction.isra.8+0xb9/0x3ba [btrfs] [181631.224566] [<ffffffffa03d1423>] ? wait_current_trans.isra.13+0x22/0xfc [btrfs] [181631.224566] [<ffffffffa03f3fbc>] btrfs_relocate_chunk.isra.29+0x8f/0xa7 [btrfs] [181631.224566] [<ffffffffa03f54df>] btrfs_balance+0xaa4/0xc52 [btrfs] [181631.224566] [<ffffffffa03fd388>] btrfs_ioctl_balance+0x23f/0x2b0 [btrfs] [181631.224566] [<ffffffff810872f9>] ? trace_hardirqs_on+0xd/0xf [181631.224566] [<ffffffffa04019a3>] btrfs_ioctl+0xfe2/0x2220 [btrfs] [181631.224566] [<ffffffff812603ed>] ? __this_cpu_preempt_check+0x13/0x15 [181631.224566] [<ffffffff81084669>] ? arch_local_irq_save+0x9/0xc [181631.224566] [<ffffffff81138def>] ? handle_mm_fault+0x834/0xcd2 [181631.224566] [<ffffffff81138def>] ? handle_mm_fault+0x834/0xcd2 [181631.224566] [<ffffffff8103e48c>] ? __do_page_fault+0x211/0x424 [181631.224566] [<ffffffff811755e6>] do_vfs_ioctl+0x3c6/0x479 (...) The sequence of steps leading to this are: CPU 0 CPU 1 btrfs_balance() btrfs_relocate_chunk() btrfs_relocate_block_group(bg X) btrfs_lookup_block_group(bg X) cleaner_kthread locks fs_info->cleaner_mutex btrfs_delete_unused_bgs() finds bg X, which became unused in the previous transaction checks bg X ->ro == 0, so it proceeds sets bg X ->ro to 1 (btrfs_set_block_group_ro(bg X)) blocks on fs_info->cleaner_mutex btrfs_remove_chunk(bg X) unlocks fs_info->cleaner_mutex acquires fs_info->cleaner_mutex relocate_block_group() --> does nothing, no extents found in the extent tree from bg X unlocks fs_info->cleaner_mutex btrfs_relocate_block_group(bg X) returns btrfs_remove_chunk(bg X) extent map not found --> ASSERT(0) Fix this by using a new mutex to make sure these 2 operations, block group relocation and removal, are serialized. This issue is reproducible by running fstests generic/038 (which stresses chunk allocation and automatic removal of unused block groups) together with the following balance loop: while true; do btrfs balance start -dusage=0 <mountpoint> ; done Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2015-06-10 23:58:53 +00:00
struct mutex delete_unused_bgs_mutex;
/* Cached block sizes */
u32 nodesize;
u32 sectorsize;
u32 stripesize;
/* Block groups and devices containing active swapfiles. */
spinlock_t swapfile_pins_lock;
struct rb_root swapfile_pins;
#ifdef CONFIG_BTRFS_FS_REF_VERIFY
spinlock_t ref_verify_lock;
struct rb_root block_tree;
#endif
};
static inline struct btrfs_fs_info *btrfs_sb(struct super_block *sb)
{
return sb->s_fs_info;
}
struct btrfs_subvolume_writers {
struct percpu_counter counter;
wait_queue_head_t wait;
};
/*
* The state of btrfs root
*/
enum {
/*
* btrfs_record_root_in_trans is a multi-step process, and it can race
* with the balancing code. But the race is very small, and only the
* first time the root is added to each transaction. So IN_TRANS_SETUP
* is used to tell us when more checks are required
*/
BTRFS_ROOT_IN_TRANS_SETUP,
BTRFS_ROOT_REF_COWS,
BTRFS_ROOT_TRACK_DIRTY,
BTRFS_ROOT_IN_RADIX,
BTRFS_ROOT_ORPHAN_ITEM_INSERTED,
BTRFS_ROOT_DEFRAG_RUNNING,
BTRFS_ROOT_FORCE_COW,
BTRFS_ROOT_MULTI_LOG_TASKS,
BTRFS_ROOT_DIRTY,
BTRFS_ROOT_DELETING,
btrfs: relocation: Delay reloc tree deletion after merge_reloc_roots Relocation code will drop btrfs_root::reloc_root as soon as merge_reloc_root() finishes. However later qgroup code will need to access btrfs_root::reloc_root after merge_reloc_root() for delayed subtree rescan. So alter the timming of resetting btrfs_root:::reloc_root, make it happens after transaction commit. With this patch, we will introduce a new btrfs_root::state, BTRFS_ROOT_DEAD_RELOC_TREE, to info part of btrfs_root::reloc_tree user that although btrfs_root::reloc_tree is still non-NULL, but still it's not used any more. The lifespan of btrfs_root::reloc tree will become: Old behavior | New ------------------------------------------------------------------------ btrfs_init_reloc_root() --- | btrfs_init_reloc_root() --- set reloc_root | | set reloc_root | | | | | | | merge_reloc_root() | | merge_reloc_root() | |- btrfs_update_reloc_root() --- | |- btrfs_update_reloc_root() -+- clear btrfs_root::reloc_root | set ROOT_DEAD_RELOC_TREE | | record root into dirty | | roots rbtree | | | | reloc_block_group() Or | | btrfs_recover_relocation() | | | After transaction commit | | |- clean_dirty_subvols() --- | clear btrfs_root::reloc_root During ROOT_DEAD_RELOC_TREE set lifespan, the only user of btrfs_root::reloc_tree should be qgroup. Since reloc root needs a longer life-span, this patch will also delay btrfs_drop_snapshot() call. Now btrfs_drop_snapshot() is called in clean_dirty_subvols(). This patch will increase the size of btrfs_root by 16 bytes. Signed-off-by: Qu Wenruo <wqu@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2019-01-23 07:15:14 +00:00
/*
* Reloc tree is orphan, only kept here for qgroup delayed subtree scan
*
* Set for the subvolume tree owning the reloc tree.
*/
BTRFS_ROOT_DEAD_RELOC_TREE,
btrfs: check for refs on snapshot delete resume There's a bug in snapshot deletion where we won't update the drop_progress key if we're in the UPDATE_BACKREF stage. This is a problem because we could drop refs for blocks we know don't belong to ours. If we crash or umount at the right time we could experience messages such as the following when snapshot deletion resumes BTRFS error (device dm-3): unable to find ref byte nr 66797568 parent 0 root 258 owner 1 offset 0 ------------[ cut here ]------------ WARNING: CPU: 3 PID: 16052 at fs/btrfs/extent-tree.c:7108 __btrfs_free_extent.isra.78+0x62c/0xb30 [btrfs] CPU: 3 PID: 16052 Comm: umount Tainted: G W OE 5.0.0-rc4+ #147 Hardware name: To Be Filled By O.E.M. To Be Filled By O.E.M./890FX Deluxe5, BIOS P1.40 05/03/2011 RIP: 0010:__btrfs_free_extent.isra.78+0x62c/0xb30 [btrfs] RSP: 0018:ffffc90005cd7b18 EFLAGS: 00010286 RAX: 0000000000000000 RBX: 0000000000000001 RCX: 0000000000000000 RDX: ffff88842fade680 RSI: ffff88842fad6b18 RDI: ffff88842fad6b18 RBP: ffffc90005cd7bc8 R08: 0000000000000000 R09: 0000000000000001 R10: 0000000000000001 R11: ffffffff822696b8 R12: 0000000003fb4000 R13: 0000000000000001 R14: 0000000000000102 R15: ffff88819c9d67e0 FS: 00007f08bb138fc0(0000) GS:ffff88842fac0000(0000) knlGS:0000000000000000 CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033 CR2: 00007f8f5d861ea0 CR3: 00000003e99fe000 CR4: 00000000000006e0 Call Trace: ? _raw_spin_unlock+0x27/0x40 ? btrfs_merge_delayed_refs+0x356/0x3e0 [btrfs] __btrfs_run_delayed_refs+0x75a/0x13c0 [btrfs] ? join_transaction+0x2b/0x460 [btrfs] btrfs_run_delayed_refs+0xf3/0x1c0 [btrfs] btrfs_commit_transaction+0x52/0xa50 [btrfs] ? start_transaction+0xa6/0x510 [btrfs] btrfs_sync_fs+0x79/0x1c0 [btrfs] sync_filesystem+0x70/0x90 generic_shutdown_super+0x27/0x120 kill_anon_super+0x12/0x30 btrfs_kill_super+0x16/0xa0 [btrfs] deactivate_locked_super+0x43/0x70 deactivate_super+0x40/0x60 cleanup_mnt+0x3f/0x80 __cleanup_mnt+0x12/0x20 task_work_run+0x8b/0xc0 exit_to_usermode_loop+0xce/0xd0 do_syscall_64+0x20b/0x210 entry_SYSCALL_64_after_hwframe+0x49/0xbe To fix this simply mark dead roots we read from disk as DEAD and then set the walk_control->restarted flag so we know we have a restarted deletion. From here whenever we try to drop refs for blocks we check to verify our ref is set on them, and if it is not we skip it. Once we find a ref that is set we unset walk_control->restarted since the tree should be in a normal state from then on, and any problems we run into from there are different issues. I tested this with an existing broken fs and my reproducer that creates a broken fs and it fixed both file systems. Reviewed-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Josef Bacik <josef@toxicpanda.com> Signed-off-by: David Sterba <dsterba@suse.com>
2019-02-06 20:46:14 +00:00
/* Mark dead root stored on device whose cleanup needs to be resumed */
BTRFS_ROOT_DEAD_TREE,
};
btrfs: qgroup: Introduce per-root swapped blocks infrastructure To allow delayed subtree swap rescan, btrfs needs to record per-root information about which tree blocks get swapped. This patch introduces the required infrastructure. The designed workflow will be: 1) Record the subtree root block that gets swapped. During subtree swap: O = Old tree blocks N = New tree blocks reloc tree subvolume tree X Root Root / \ / \ NA OB OA OB / | | \ / | | \ NC ND OE OF OC OD OE OF In this case, NA and OA are going to be swapped, record (NA, OA) into subvolume tree X. 2) After subtree swap. reloc tree subvolume tree X Root Root / \ / \ OA OB NA OB / | | \ / | | \ OC OD OE OF NC ND OE OF 3a) COW happens for OB If we are going to COW tree block OB, we check OB's bytenr against tree X's swapped_blocks structure. If it doesn't fit any, nothing will happen. 3b) COW happens for NA Check NA's bytenr against tree X's swapped_blocks, and get a hit. Then we do subtree scan on both subtrees OA and NA. Resulting 6 tree blocks to be scanned (OA, OC, OD, NA, NC, ND). Then no matter what we do to subvolume tree X, qgroup numbers will still be correct. Then NA's record gets removed from X's swapped_blocks. 4) Transaction commit Any record in X's swapped_blocks gets removed, since there is no modification to swapped subtrees, no need to trigger heavy qgroup subtree rescan for them. This will introduce 128 bytes overhead for each btrfs_root even qgroup is not enabled. This is to reduce memory allocations and potential failures. Signed-off-by: Qu Wenruo <wqu@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2019-01-23 07:15:16 +00:00
/*
* Record swapped tree blocks of a subvolume tree for delayed subtree trace
* code. For detail check comment in fs/btrfs/qgroup.c.
*/
struct btrfs_qgroup_swapped_blocks {
spinlock_t lock;
/* RM_EMPTY_ROOT() of above blocks[] */
bool swapped;
struct rb_root blocks[BTRFS_MAX_LEVEL];
};
/*
* in ram representation of the tree. extent_root is used for all allocations
* and for the extent tree extent_root root.
*/
struct btrfs_root {
struct extent_buffer *node;
struct extent_buffer *commit_root;
struct btrfs_root *log_root;
2008-09-26 14:09:34 +00:00
struct btrfs_root *reloc_root;
unsigned long state;
struct btrfs_root_item root_item;
struct btrfs_key root_key;
struct btrfs_fs_info *fs_info;
struct extent_io_tree dirty_log_pages;
struct mutex objectid_mutex;
spinlock_t accounting_lock;
struct btrfs_block_rsv *block_rsv;
/* free ino cache stuff */
struct btrfs_free_space_ctl *free_ino_ctl;
enum btrfs_caching_type ino_cache_state;
spinlock_t ino_cache_lock;
wait_queue_head_t ino_cache_wait;
struct btrfs_free_space_ctl *free_ino_pinned;
u64 ino_cache_progress;
struct inode *ino_cache_inode;
struct mutex log_mutex;
wait_queue_head_t log_writer_wait;
wait_queue_head_t log_commit_wait[2];
struct list_head log_ctxs[2];
atomic_t log_writers;
atomic_t log_commit[2];
atomic_t log_batch;
int log_transid;
/* No matter the commit succeeds or not*/
int log_transid_committed;
/* Just be updated when the commit succeeds. */
int last_log_commit;
pid_t log_start_pid;
u64 last_trans;
u32 type;
u64 highest_objectid;
u64 defrag_trans_start;
struct btrfs_key defrag_progress;
struct btrfs_key defrag_max;
/* the dirty list is only used by non-reference counted roots */
struct list_head dirty_list;
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
struct list_head root_list;
spinlock_t log_extents_lock[2];
struct list_head logged_list[2];
int orphan_cleanup_state;
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
spinlock_t inode_lock;
/* red-black tree that keeps track of in-memory inodes */
struct rb_root inode_tree;
btrfs: implement delayed inode items operation Changelog V5 -> V6: - Fix oom when the memory load is high, by storing the delayed nodes into the root's radix tree, and letting btrfs inodes go. Changelog V4 -> V5: - Fix the race on adding the delayed node to the inode, which is spotted by Chris Mason. - Merge Chris Mason's incremental patch into this patch. - Fix deadlock between readdir() and memory fault, which is reported by Itaru Kitayama. Changelog V3 -> V4: - Fix nested lock, which is reported by Itaru Kitayama, by updating space cache inode in time. Changelog V2 -> V3: - Fix the race between the delayed worker and the task which does delayed items balance, which is reported by Tsutomu Itoh. - Modify the patch address David Sterba's comment. - Fix the bug of the cpu recursion spinlock, reported by Chris Mason Changelog V1 -> V2: - break up the global rb-tree, use a list to manage the delayed nodes, which is created for every directory and file, and used to manage the delayed directory name index items and the delayed inode item. - introduce a worker to deal with the delayed nodes. Compare with Ext3/4, the performance of file creation and deletion on btrfs is very poor. the reason is that btrfs must do a lot of b+ tree insertions, such as inode item, directory name item, directory name index and so on. If we can do some delayed b+ tree insertion or deletion, we can improve the performance, so we made this patch which implemented delayed directory name index insertion/deletion and delayed inode update. Implementation: - introduce a delayed root object into the filesystem, that use two lists to manage the delayed nodes which are created for every file/directory. One is used to manage all the delayed nodes that have delayed items. And the other is used to manage the delayed nodes which is waiting to be dealt with by the work thread. - Every delayed node has two rb-tree, one is used to manage the directory name index which is going to be inserted into b+ tree, and the other is used to manage the directory name index which is going to be deleted from b+ tree. - introduce a worker to deal with the delayed operation. This worker is used to deal with the works of the delayed directory name index items insertion and deletion and the delayed inode update. When the delayed items is beyond the lower limit, we create works for some delayed nodes and insert them into the work queue of the worker, and then go back. When the delayed items is beyond the upper bound, we create works for all the delayed nodes that haven't been dealt with, and insert them into the work queue of the worker, and then wait for that the untreated items is below some threshold value. - When we want to insert a directory name index into b+ tree, we just add the information into the delayed inserting rb-tree. And then we check the number of the delayed items and do delayed items balance. (The balance policy is above.) - When we want to delete a directory name index from the b+ tree, we search it in the inserting rb-tree at first. If we look it up, just drop it. If not, add the key of it into the delayed deleting rb-tree. Similar to the delayed inserting rb-tree, we also check the number of the delayed items and do delayed items balance. (The same to inserting manipulation) - When we want to update the metadata of some inode, we cached the data of the inode into the delayed node. the worker will flush it into the b+ tree after dealing with the delayed insertion and deletion. - We will move the delayed node to the tail of the list after we access the delayed node, By this way, we can cache more delayed items and merge more inode updates. - If we want to commit transaction, we will deal with all the delayed node. - the delayed node will be freed when we free the btrfs inode. - Before we log the inode items, we commit all the directory name index items and the delayed inode update. I did a quick test by the benchmark tool[1] and found we can improve the performance of file creation by ~15%, and file deletion by ~20%. Before applying this patch: Create files: Total files: 50000 Total time: 1.096108 Average time: 0.000022 Delete files: Total files: 50000 Total time: 1.510403 Average time: 0.000030 After applying this patch: Create files: Total files: 50000 Total time: 0.932899 Average time: 0.000019 Delete files: Total files: 50000 Total time: 1.215732 Average time: 0.000024 [1] http://marc.info/?l=linux-btrfs&m=128212635122920&q=p3 Many thanks for Kitayama-san's help! Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Reviewed-by: David Sterba <dave@jikos.cz> Tested-by: Tsutomu Itoh <t-itoh@jp.fujitsu.com> Tested-by: Itaru Kitayama <kitayama@cl.bb4u.ne.jp> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2011-04-22 10:12:22 +00:00
/*
* radix tree that keeps track of delayed nodes of every inode,
* protected by inode_lock
*/
struct radix_tree_root delayed_nodes_tree;
/*
* right now this just gets used so that a root has its own devid
* for stat. It may be used for more later
*/
dev_t anon_dev;
spinlock_t root_item_lock;
refcount_t refs;
struct mutex delalloc_mutex;
spinlock_t delalloc_lock;
/*
* all of the inodes that have delalloc bytes. It is possible for
* this list to be empty even when there is still dirty data=ordered
* extents waiting to finish IO.
*/
struct list_head delalloc_inodes;
struct list_head delalloc_root;
u64 nr_delalloc_inodes;
struct mutex ordered_extent_mutex;
/*
* this is used by the balancing code to wait for all the pending
* ordered extents
*/
spinlock_t ordered_extent_lock;
/*
* all of the data=ordered extents pending writeback
* these can span multiple transactions and basically include
* every dirty data page that isn't from nodatacow
*/
struct list_head ordered_extents;
struct list_head ordered_root;
u64 nr_ordered_extents;
btrfs: relocation: Delay reloc tree deletion after merge_reloc_roots Relocation code will drop btrfs_root::reloc_root as soon as merge_reloc_root() finishes. However later qgroup code will need to access btrfs_root::reloc_root after merge_reloc_root() for delayed subtree rescan. So alter the timming of resetting btrfs_root:::reloc_root, make it happens after transaction commit. With this patch, we will introduce a new btrfs_root::state, BTRFS_ROOT_DEAD_RELOC_TREE, to info part of btrfs_root::reloc_tree user that although btrfs_root::reloc_tree is still non-NULL, but still it's not used any more. The lifespan of btrfs_root::reloc tree will become: Old behavior | New ------------------------------------------------------------------------ btrfs_init_reloc_root() --- | btrfs_init_reloc_root() --- set reloc_root | | set reloc_root | | | | | | | merge_reloc_root() | | merge_reloc_root() | |- btrfs_update_reloc_root() --- | |- btrfs_update_reloc_root() -+- clear btrfs_root::reloc_root | set ROOT_DEAD_RELOC_TREE | | record root into dirty | | roots rbtree | | | | reloc_block_group() Or | | btrfs_recover_relocation() | | | After transaction commit | | |- clean_dirty_subvols() --- | clear btrfs_root::reloc_root During ROOT_DEAD_RELOC_TREE set lifespan, the only user of btrfs_root::reloc_tree should be qgroup. Since reloc root needs a longer life-span, this patch will also delay btrfs_drop_snapshot() call. Now btrfs_drop_snapshot() is called in clean_dirty_subvols(). This patch will increase the size of btrfs_root by 16 bytes. Signed-off-by: Qu Wenruo <wqu@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2019-01-23 07:15:14 +00:00
/*
* Not empty if this subvolume root has gone through tree block swap
* (relocation)
*
* Will be used by reloc_control::dirty_subvol_roots.
*/
struct list_head reloc_dirty_list;
/*
* Number of currently running SEND ioctls to prevent
* manipulation with the read-only status via SUBVOL_SETFLAGS
*/
int send_in_progress;
struct btrfs_subvolume_writers *subv_writers;
atomic_t will_be_snapshotted;
Btrfs: fix unexpected failure of nocow buffered writes after snapshotting when low on space Commit e9894fd3e3b3 ("Btrfs: fix snapshot vs nocow writting") forced nocow writes to fallback to COW, during writeback, when a snapshot is created. This resulted in writes made before creating the snapshot to unexpectedly fail with ENOSPC during writeback when success (0) was returned to user space through the write system call. The steps leading to this problem are: 1. When it's not possible to allocate data space for a write, the buffered write path checks if a NOCOW write is possible. If it is, it will not reserve space and success (0) is returned to user space. 2. Then when a snapshot is created, the root's will_be_snapshotted atomic is incremented and writeback is triggered for all inode's that belong to the root being snapshotted. Incrementing that atomic forces all previous writes to fallback to COW during writeback (running delalloc). 3. This results in the writeback for the inodes to fail and therefore setting the ENOSPC error in their mappings, so that a subsequent fsync on them will report the error to user space. So it's not a completely silent data loss (since fsync will report ENOSPC) but it's a very unexpected and undesirable behaviour, because if a clean shutdown/unmount of the filesystem happens without previous calls to fsync, it is expected to have the data present in the files after mounting the filesystem again. So fix this by adding a new atomic named snapshot_force_cow to the root structure which prevents this behaviour and works the following way: 1. It is incremented when we start to create a snapshot after triggering writeback and before waiting for writeback to finish. 2. This new atomic is now what is used by writeback (running delalloc) to decide whether we need to fallback to COW or not. Because we incremented this new atomic after triggering writeback in the snapshot creation ioctl, we ensure that all buffered writes that happened before snapshot creation will succeed and not fallback to COW (which would make them fail with ENOSPC). 3. The existing atomic, will_be_snapshotted, is kept because it is used to force new buffered writes, that start after we started snapshotting, to reserve data space even when NOCOW is possible. This makes these writes fail early with ENOSPC when there's no available space to allocate, preventing the unexpected behaviour of writeback later failing with ENOSPC due to a fallback to COW mode. Fixes: e9894fd3e3b3 ("Btrfs: fix snapshot vs nocow writting") Signed-off-by: Robbie Ko <robbieko@synology.com> Reviewed-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2018-08-06 02:30:30 +00:00
atomic_t snapshot_force_cow;
/* For qgroup metadata reserved space */
spinlock_t qgroup_meta_rsv_lock;
u64 qgroup_meta_rsv_pertrans;
u64 qgroup_meta_rsv_prealloc;
/* Number of active swapfiles */
atomic_t nr_swapfiles;
btrfs: qgroup: Introduce per-root swapped blocks infrastructure To allow delayed subtree swap rescan, btrfs needs to record per-root information about which tree blocks get swapped. This patch introduces the required infrastructure. The designed workflow will be: 1) Record the subtree root block that gets swapped. During subtree swap: O = Old tree blocks N = New tree blocks reloc tree subvolume tree X Root Root / \ / \ NA OB OA OB / | | \ / | | \ NC ND OE OF OC OD OE OF In this case, NA and OA are going to be swapped, record (NA, OA) into subvolume tree X. 2) After subtree swap. reloc tree subvolume tree X Root Root / \ / \ OA OB NA OB / | | \ / | | \ OC OD OE OF NC ND OE OF 3a) COW happens for OB If we are going to COW tree block OB, we check OB's bytenr against tree X's swapped_blocks structure. If it doesn't fit any, nothing will happen. 3b) COW happens for NA Check NA's bytenr against tree X's swapped_blocks, and get a hit. Then we do subtree scan on both subtrees OA and NA. Resulting 6 tree blocks to be scanned (OA, OC, OD, NA, NC, ND). Then no matter what we do to subvolume tree X, qgroup numbers will still be correct. Then NA's record gets removed from X's swapped_blocks. 4) Transaction commit Any record in X's swapped_blocks gets removed, since there is no modification to swapped subtrees, no need to trigger heavy qgroup subtree rescan for them. This will introduce 128 bytes overhead for each btrfs_root even qgroup is not enabled. This is to reduce memory allocations and potential failures. Signed-off-by: Qu Wenruo <wqu@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2019-01-23 07:15:16 +00:00
/* Record pairs of swapped blocks for qgroup */
struct btrfs_qgroup_swapped_blocks swapped_blocks;
#ifdef CONFIG_BTRFS_FS_RUN_SANITY_TESTS
u64 alloc_bytenr;
#endif
};
struct btrfs_file_private {
void *filldir_buf;
};
static inline u32 btrfs_inode_sectorsize(const struct inode *inode)
{
return btrfs_sb(inode->i_sb)->sectorsize;
}
static inline u32 BTRFS_LEAF_DATA_SIZE(const struct btrfs_fs_info *info)
{
return info->nodesize - sizeof(struct btrfs_header);
}
#define BTRFS_LEAF_DATA_OFFSET offsetof(struct btrfs_leaf, items)
static inline u32 BTRFS_MAX_ITEM_SIZE(const struct btrfs_fs_info *info)
{
return BTRFS_LEAF_DATA_SIZE(info) - sizeof(struct btrfs_item);
}
static inline u32 BTRFS_NODEPTRS_PER_BLOCK(const struct btrfs_fs_info *info)
{
return BTRFS_LEAF_DATA_SIZE(info) / sizeof(struct btrfs_key_ptr);
}
#define BTRFS_FILE_EXTENT_INLINE_DATA_START \
(offsetof(struct btrfs_file_extent_item, disk_bytenr))
static inline u32 BTRFS_MAX_INLINE_DATA_SIZE(const struct btrfs_fs_info *info)
{
return BTRFS_MAX_ITEM_SIZE(info) -
BTRFS_FILE_EXTENT_INLINE_DATA_START;
}
static inline u32 BTRFS_MAX_XATTR_SIZE(const struct btrfs_fs_info *info)
{
return BTRFS_MAX_ITEM_SIZE(info) - sizeof(struct btrfs_dir_item);
}
/*
* Flags for mount options.
*
* Note: don't forget to add new options to btrfs_show_options()
*/
#define BTRFS_MOUNT_NODATASUM (1 << 0)
#define BTRFS_MOUNT_NODATACOW (1 << 1)
#define BTRFS_MOUNT_NOBARRIER (1 << 2)
#define BTRFS_MOUNT_SSD (1 << 3)
#define BTRFS_MOUNT_DEGRADED (1 << 4)
Btrfs: Add zlib compression support This is a large change for adding compression on reading and writing, both for inline and regular extents. It does some fairly large surgery to the writeback paths. Compression is off by default and enabled by mount -o compress. Even when the -o compress mount option is not used, it is possible to read compressed extents off the disk. If compression for a given set of pages fails to make them smaller, the file is flagged to avoid future compression attempts later. * While finding delalloc extents, the pages are locked before being sent down to the delalloc handler. This allows the delalloc handler to do complex things such as cleaning the pages, marking them writeback and starting IO on their behalf. * Inline extents are inserted at delalloc time now. This allows us to compress the data before inserting the inline extent, and it allows us to insert an inline extent that spans multiple pages. * All of the in-memory extent representations (extent_map.c, ordered-data.c etc) are changed to record both an in-memory size and an on disk size, as well as a flag for compression. From a disk format point of view, the extent pointers in the file are changed to record the on disk size of a given extent and some encoding flags. Space in the disk format is allocated for compression encoding, as well as encryption and a generic 'other' field. Neither the encryption or the 'other' field are currently used. In order to limit the amount of data read for a single random read in the file, the size of a compressed extent is limited to 128k. This is a software only limit, the disk format supports u64 sized compressed extents. In order to limit the ram consumed while processing extents, the uncompressed size of a compressed extent is limited to 256k. This is a software only limit and will be subject to tuning later. Checksumming is still done on compressed extents, and it is done on the uncompressed version of the data. This way additional encodings can be layered on without having to figure out which encoding to checksum. Compression happens at delalloc time, which is basically singled threaded because it is usually done by a single pdflush thread. This makes it tricky to spread the compression load across all the cpus on the box. We'll have to look at parallel pdflush walks of dirty inodes at a later time. Decompression is hooked into readpages and it does spread across CPUs nicely. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-10-29 18:49:59 +00:00
#define BTRFS_MOUNT_COMPRESS (1 << 5)
#define BTRFS_MOUNT_NOTREELOG (1 << 6)
#define BTRFS_MOUNT_FLUSHONCOMMIT (1 << 7)
#define BTRFS_MOUNT_SSD_SPREAD (1 << 8)
#define BTRFS_MOUNT_NOSSD (1 << 9)
#define BTRFS_MOUNT_DISCARD (1 << 10)
#define BTRFS_MOUNT_FORCE_COMPRESS (1 << 11)
#define BTRFS_MOUNT_SPACE_CACHE (1 << 12)
#define BTRFS_MOUNT_CLEAR_CACHE (1 << 13)
#define BTRFS_MOUNT_USER_SUBVOL_RM_ALLOWED (1 << 14)
#define BTRFS_MOUNT_ENOSPC_DEBUG (1 << 15)
#define BTRFS_MOUNT_AUTO_DEFRAG (1 << 16)
#define BTRFS_MOUNT_INODE_MAP_CACHE (1 << 17)
#define BTRFS_MOUNT_USEBACKUPROOT (1 << 18)
#define BTRFS_MOUNT_SKIP_BALANCE (1 << 19)
#define BTRFS_MOUNT_CHECK_INTEGRITY (1 << 20)
#define BTRFS_MOUNT_CHECK_INTEGRITY_INCLUDING_EXTENT_DATA (1 << 21)
#define BTRFS_MOUNT_PANIC_ON_FATAL_ERROR (1 << 22)
#define BTRFS_MOUNT_RESCAN_UUID_TREE (1 << 23)
#define BTRFS_MOUNT_FRAGMENT_DATA (1 << 24)
#define BTRFS_MOUNT_FRAGMENT_METADATA (1 << 25)
#define BTRFS_MOUNT_FREE_SPACE_TREE (1 << 26)
#define BTRFS_MOUNT_NOLOGREPLAY (1 << 27)
#define BTRFS_MOUNT_REF_VERIFY (1 << 28)
#define BTRFS_DEFAULT_COMMIT_INTERVAL (30)
#define BTRFS_DEFAULT_MAX_INLINE (2048)
#define btrfs_clear_opt(o, opt) ((o) &= ~BTRFS_MOUNT_##opt)
#define btrfs_set_opt(o, opt) ((o) |= BTRFS_MOUNT_##opt)
#define btrfs_raw_test_opt(o, opt) ((o) & BTRFS_MOUNT_##opt)
#define btrfs_test_opt(fs_info, opt) ((fs_info)->mount_opt & \
BTRFS_MOUNT_##opt)
#define btrfs_set_and_info(fs_info, opt, fmt, args...) \
{ \
if (!btrfs_test_opt(fs_info, opt)) \
btrfs_info(fs_info, fmt, ##args); \
btrfs_set_opt(fs_info->mount_opt, opt); \
}
#define btrfs_clear_and_info(fs_info, opt, fmt, args...) \
{ \
if (btrfs_test_opt(fs_info, opt)) \
btrfs_info(fs_info, fmt, ##args); \
btrfs_clear_opt(fs_info->mount_opt, opt); \
}
#ifdef CONFIG_BTRFS_DEBUG
static inline int
btrfs_should_fragment_free_space(struct btrfs_block_group_cache *block_group)
{
struct btrfs_fs_info *fs_info = block_group->fs_info;
return (btrfs_test_opt(fs_info, FRAGMENT_METADATA) &&
block_group->flags & BTRFS_BLOCK_GROUP_METADATA) ||
(btrfs_test_opt(fs_info, FRAGMENT_DATA) &&
block_group->flags & BTRFS_BLOCK_GROUP_DATA);
}
#endif
/*
* Requests for changes that need to be done during transaction commit.
*
* Internal mount options that are used for special handling of the real
* mount options (eg. cannot be set during remount and have to be set during
* transaction commit)
*/
#define BTRFS_PENDING_SET_INODE_MAP_CACHE (0)
#define BTRFS_PENDING_CLEAR_INODE_MAP_CACHE (1)
#define BTRFS_PENDING_COMMIT (2)
#define btrfs_test_pending(info, opt) \
test_bit(BTRFS_PENDING_##opt, &(info)->pending_changes)
#define btrfs_set_pending(info, opt) \
set_bit(BTRFS_PENDING_##opt, &(info)->pending_changes)
#define btrfs_clear_pending(info, opt) \
clear_bit(BTRFS_PENDING_##opt, &(info)->pending_changes)
/*
* Helpers for setting pending mount option changes.
*
* Expects corresponding macros
* BTRFS_PENDING_SET_ and CLEAR_ + short mount option name
*/
#define btrfs_set_pending_and_info(info, opt, fmt, args...) \
do { \
if (!btrfs_raw_test_opt((info)->mount_opt, opt)) { \
btrfs_info((info), fmt, ##args); \
btrfs_set_pending((info), SET_##opt); \
btrfs_clear_pending((info), CLEAR_##opt); \
} \
} while(0)
#define btrfs_clear_pending_and_info(info, opt, fmt, args...) \
do { \
if (btrfs_raw_test_opt((info)->mount_opt, opt)) { \
btrfs_info((info), fmt, ##args); \
btrfs_set_pending((info), CLEAR_##opt); \
btrfs_clear_pending((info), SET_##opt); \
} \
} while(0)
/*
* Inode flags
*/
#define BTRFS_INODE_NODATASUM (1 << 0)
#define BTRFS_INODE_NODATACOW (1 << 1)
#define BTRFS_INODE_READONLY (1 << 2)
Btrfs: Add zlib compression support This is a large change for adding compression on reading and writing, both for inline and regular extents. It does some fairly large surgery to the writeback paths. Compression is off by default and enabled by mount -o compress. Even when the -o compress mount option is not used, it is possible to read compressed extents off the disk. If compression for a given set of pages fails to make them smaller, the file is flagged to avoid future compression attempts later. * While finding delalloc extents, the pages are locked before being sent down to the delalloc handler. This allows the delalloc handler to do complex things such as cleaning the pages, marking them writeback and starting IO on their behalf. * Inline extents are inserted at delalloc time now. This allows us to compress the data before inserting the inline extent, and it allows us to insert an inline extent that spans multiple pages. * All of the in-memory extent representations (extent_map.c, ordered-data.c etc) are changed to record both an in-memory size and an on disk size, as well as a flag for compression. From a disk format point of view, the extent pointers in the file are changed to record the on disk size of a given extent and some encoding flags. Space in the disk format is allocated for compression encoding, as well as encryption and a generic 'other' field. Neither the encryption or the 'other' field are currently used. In order to limit the amount of data read for a single random read in the file, the size of a compressed extent is limited to 128k. This is a software only limit, the disk format supports u64 sized compressed extents. In order to limit the ram consumed while processing extents, the uncompressed size of a compressed extent is limited to 256k. This is a software only limit and will be subject to tuning later. Checksumming is still done on compressed extents, and it is done on the uncompressed version of the data. This way additional encodings can be layered on without having to figure out which encoding to checksum. Compression happens at delalloc time, which is basically singled threaded because it is usually done by a single pdflush thread. This makes it tricky to spread the compression load across all the cpus on the box. We'll have to look at parallel pdflush walks of dirty inodes at a later time. Decompression is hooked into readpages and it does spread across CPUs nicely. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-10-29 18:49:59 +00:00
#define BTRFS_INODE_NOCOMPRESS (1 << 3)
#define BTRFS_INODE_PREALLOC (1 << 4)
#define BTRFS_INODE_SYNC (1 << 5)
#define BTRFS_INODE_IMMUTABLE (1 << 6)
#define BTRFS_INODE_APPEND (1 << 7)
#define BTRFS_INODE_NODUMP (1 << 8)
#define BTRFS_INODE_NOATIME (1 << 9)
#define BTRFS_INODE_DIRSYNC (1 << 10)
#define BTRFS_INODE_COMPRESS (1 << 11)
#define BTRFS_INODE_ROOT_ITEM_INIT (1 << 31)
#define BTRFS_INODE_FLAG_MASK \
(BTRFS_INODE_NODATASUM | \
BTRFS_INODE_NODATACOW | \
BTRFS_INODE_READONLY | \
BTRFS_INODE_NOCOMPRESS | \
BTRFS_INODE_PREALLOC | \
BTRFS_INODE_SYNC | \
BTRFS_INODE_IMMUTABLE | \
BTRFS_INODE_APPEND | \
BTRFS_INODE_NODUMP | \
BTRFS_INODE_NOATIME | \
BTRFS_INODE_DIRSYNC | \
BTRFS_INODE_COMPRESS | \
BTRFS_INODE_ROOT_ITEM_INIT)
struct btrfs_map_token {
const struct extent_buffer *eb;
char *kaddr;
unsigned long offset;
};
#define BTRFS_BYTES_TO_BLKS(fs_info, bytes) \
((bytes) >> (fs_info)->sb->s_blocksize_bits)
static inline void btrfs_init_map_token (struct btrfs_map_token *token)
{
token->kaddr = NULL;
}
/* some macros to generate set/get functions for the struct fields. This
* assumes there is a lefoo_to_cpu for every type, so lets make a simple
* one for u8:
*/
#define le8_to_cpu(v) (v)
#define cpu_to_le8(v) (v)
#define __le8 u8
#define read_eb_member(eb, ptr, type, member, result) (\
read_extent_buffer(eb, (char *)(result), \
((unsigned long)(ptr)) + \
offsetof(type, member), \
sizeof(((type *)0)->member)))
#define write_eb_member(eb, ptr, type, member, result) (\
write_extent_buffer(eb, (char *)(result), \
((unsigned long)(ptr)) + \
offsetof(type, member), \
sizeof(((type *)0)->member)))
#define DECLARE_BTRFS_SETGET_BITS(bits) \
u##bits btrfs_get_token_##bits(const struct extent_buffer *eb, \
const void *ptr, unsigned long off, \
struct btrfs_map_token *token); \
void btrfs_set_token_##bits(struct extent_buffer *eb, const void *ptr, \
unsigned long off, u##bits val, \
struct btrfs_map_token *token); \
static inline u##bits btrfs_get_##bits(const struct extent_buffer *eb, \
const void *ptr, \
unsigned long off) \
{ \
return btrfs_get_token_##bits(eb, ptr, off, NULL); \
} \
static inline void btrfs_set_##bits(struct extent_buffer *eb, void *ptr,\
unsigned long off, u##bits val) \
{ \
btrfs_set_token_##bits(eb, ptr, off, val, NULL); \
}
DECLARE_BTRFS_SETGET_BITS(8)
DECLARE_BTRFS_SETGET_BITS(16)
DECLARE_BTRFS_SETGET_BITS(32)
DECLARE_BTRFS_SETGET_BITS(64)
#define BTRFS_SETGET_FUNCS(name, type, member, bits) \
static inline u##bits btrfs_##name(const struct extent_buffer *eb, \
const type *s) \
{ \
BUILD_BUG_ON(sizeof(u##bits) != sizeof(((type *)0))->member); \
return btrfs_get_##bits(eb, s, offsetof(type, member)); \
} \
static inline void btrfs_set_##name(struct extent_buffer *eb, type *s, \
u##bits val) \
{ \
BUILD_BUG_ON(sizeof(u##bits) != sizeof(((type *)0))->member); \
btrfs_set_##bits(eb, s, offsetof(type, member), val); \
} \
static inline u##bits btrfs_token_##name(const struct extent_buffer *eb,\
const type *s, \
struct btrfs_map_token *token) \
{ \
BUILD_BUG_ON(sizeof(u##bits) != sizeof(((type *)0))->member); \
return btrfs_get_token_##bits(eb, s, offsetof(type, member), token); \
} \
static inline void btrfs_set_token_##name(struct extent_buffer *eb, \
type *s, u##bits val, \
struct btrfs_map_token *token) \
{ \
BUILD_BUG_ON(sizeof(u##bits) != sizeof(((type *)0))->member); \
btrfs_set_token_##bits(eb, s, offsetof(type, member), val, token); \
}
#define BTRFS_SETGET_HEADER_FUNCS(name, type, member, bits) \
static inline u##bits btrfs_##name(const struct extent_buffer *eb) \
{ \
const type *p = page_address(eb->pages[0]); \
u##bits res = le##bits##_to_cpu(p->member); \
return res; \
} \
static inline void btrfs_set_##name(struct extent_buffer *eb, \
u##bits val) \
{ \
type *p = page_address(eb->pages[0]); \
p->member = cpu_to_le##bits(val); \
}
#define BTRFS_SETGET_STACK_FUNCS(name, type, member, bits) \
static inline u##bits btrfs_##name(const type *s) \
{ \
return le##bits##_to_cpu(s->member); \
} \
static inline void btrfs_set_##name(type *s, u##bits val) \
{ \
s->member = cpu_to_le##bits(val); \
}
static inline u64 btrfs_device_total_bytes(struct extent_buffer *eb,
struct btrfs_dev_item *s)
{
BUILD_BUG_ON(sizeof(u64) !=
sizeof(((struct btrfs_dev_item *)0))->total_bytes);
return btrfs_get_64(eb, s, offsetof(struct btrfs_dev_item,
total_bytes));
}
static inline void btrfs_set_device_total_bytes(struct extent_buffer *eb,
struct btrfs_dev_item *s,
u64 val)
{
BUILD_BUG_ON(sizeof(u64) !=
sizeof(((struct btrfs_dev_item *)0))->total_bytes);
WARN_ON(!IS_ALIGNED(val, eb->fs_info->sectorsize));
btrfs_set_64(eb, s, offsetof(struct btrfs_dev_item, total_bytes), val);
}
BTRFS_SETGET_FUNCS(device_type, struct btrfs_dev_item, type, 64);
BTRFS_SETGET_FUNCS(device_bytes_used, struct btrfs_dev_item, bytes_used, 64);
BTRFS_SETGET_FUNCS(device_io_align, struct btrfs_dev_item, io_align, 32);
BTRFS_SETGET_FUNCS(device_io_width, struct btrfs_dev_item, io_width, 32);
BTRFS_SETGET_FUNCS(device_start_offset, struct btrfs_dev_item,
start_offset, 64);
BTRFS_SETGET_FUNCS(device_sector_size, struct btrfs_dev_item, sector_size, 32);
BTRFS_SETGET_FUNCS(device_id, struct btrfs_dev_item, devid, 64);
BTRFS_SETGET_FUNCS(device_group, struct btrfs_dev_item, dev_group, 32);
BTRFS_SETGET_FUNCS(device_seek_speed, struct btrfs_dev_item, seek_speed, 8);
BTRFS_SETGET_FUNCS(device_bandwidth, struct btrfs_dev_item, bandwidth, 8);
BTRFS_SETGET_FUNCS(device_generation, struct btrfs_dev_item, generation, 64);
BTRFS_SETGET_STACK_FUNCS(stack_device_type, struct btrfs_dev_item, type, 64);
BTRFS_SETGET_STACK_FUNCS(stack_device_total_bytes, struct btrfs_dev_item,
total_bytes, 64);
BTRFS_SETGET_STACK_FUNCS(stack_device_bytes_used, struct btrfs_dev_item,
bytes_used, 64);
BTRFS_SETGET_STACK_FUNCS(stack_device_io_align, struct btrfs_dev_item,
io_align, 32);
BTRFS_SETGET_STACK_FUNCS(stack_device_io_width, struct btrfs_dev_item,
io_width, 32);
BTRFS_SETGET_STACK_FUNCS(stack_device_sector_size, struct btrfs_dev_item,
sector_size, 32);
BTRFS_SETGET_STACK_FUNCS(stack_device_id, struct btrfs_dev_item, devid, 64);
BTRFS_SETGET_STACK_FUNCS(stack_device_group, struct btrfs_dev_item,
dev_group, 32);
BTRFS_SETGET_STACK_FUNCS(stack_device_seek_speed, struct btrfs_dev_item,
seek_speed, 8);
BTRFS_SETGET_STACK_FUNCS(stack_device_bandwidth, struct btrfs_dev_item,
bandwidth, 8);
BTRFS_SETGET_STACK_FUNCS(stack_device_generation, struct btrfs_dev_item,
generation, 64);
static inline unsigned long btrfs_device_uuid(struct btrfs_dev_item *d)
{
return (unsigned long)d + offsetof(struct btrfs_dev_item, uuid);
}
static inline unsigned long btrfs_device_fsid(struct btrfs_dev_item *d)
{
return (unsigned long)d + offsetof(struct btrfs_dev_item, fsid);
}
BTRFS_SETGET_FUNCS(chunk_length, struct btrfs_chunk, length, 64);
BTRFS_SETGET_FUNCS(chunk_owner, struct btrfs_chunk, owner, 64);
BTRFS_SETGET_FUNCS(chunk_stripe_len, struct btrfs_chunk, stripe_len, 64);
BTRFS_SETGET_FUNCS(chunk_io_align, struct btrfs_chunk, io_align, 32);
BTRFS_SETGET_FUNCS(chunk_io_width, struct btrfs_chunk, io_width, 32);
BTRFS_SETGET_FUNCS(chunk_sector_size, struct btrfs_chunk, sector_size, 32);
BTRFS_SETGET_FUNCS(chunk_type, struct btrfs_chunk, type, 64);
BTRFS_SETGET_FUNCS(chunk_num_stripes, struct btrfs_chunk, num_stripes, 16);
BTRFS_SETGET_FUNCS(chunk_sub_stripes, struct btrfs_chunk, sub_stripes, 16);
BTRFS_SETGET_FUNCS(stripe_devid, struct btrfs_stripe, devid, 64);
BTRFS_SETGET_FUNCS(stripe_offset, struct btrfs_stripe, offset, 64);
static inline char *btrfs_stripe_dev_uuid(struct btrfs_stripe *s)
{
return (char *)s + offsetof(struct btrfs_stripe, dev_uuid);
}
BTRFS_SETGET_STACK_FUNCS(stack_chunk_length, struct btrfs_chunk, length, 64);
BTRFS_SETGET_STACK_FUNCS(stack_chunk_owner, struct btrfs_chunk, owner, 64);
BTRFS_SETGET_STACK_FUNCS(stack_chunk_stripe_len, struct btrfs_chunk,
stripe_len, 64);
BTRFS_SETGET_STACK_FUNCS(stack_chunk_io_align, struct btrfs_chunk,
io_align, 32);
BTRFS_SETGET_STACK_FUNCS(stack_chunk_io_width, struct btrfs_chunk,
io_width, 32);
BTRFS_SETGET_STACK_FUNCS(stack_chunk_sector_size, struct btrfs_chunk,
sector_size, 32);
BTRFS_SETGET_STACK_FUNCS(stack_chunk_type, struct btrfs_chunk, type, 64);
BTRFS_SETGET_STACK_FUNCS(stack_chunk_num_stripes, struct btrfs_chunk,
num_stripes, 16);
BTRFS_SETGET_STACK_FUNCS(stack_chunk_sub_stripes, struct btrfs_chunk,
sub_stripes, 16);
BTRFS_SETGET_STACK_FUNCS(stack_stripe_devid, struct btrfs_stripe, devid, 64);
BTRFS_SETGET_STACK_FUNCS(stack_stripe_offset, struct btrfs_stripe, offset, 64);
static inline struct btrfs_stripe *btrfs_stripe_nr(struct btrfs_chunk *c,
int nr)
{
unsigned long offset = (unsigned long)c;
offset += offsetof(struct btrfs_chunk, stripe);
offset += nr * sizeof(struct btrfs_stripe);
return (struct btrfs_stripe *)offset;
}
static inline char *btrfs_stripe_dev_uuid_nr(struct btrfs_chunk *c, int nr)
{
return btrfs_stripe_dev_uuid(btrfs_stripe_nr(c, nr));
}
static inline u64 btrfs_stripe_offset_nr(struct extent_buffer *eb,
struct btrfs_chunk *c, int nr)
{
return btrfs_stripe_offset(eb, btrfs_stripe_nr(c, nr));
}
static inline u64 btrfs_stripe_devid_nr(struct extent_buffer *eb,
struct btrfs_chunk *c, int nr)
{
return btrfs_stripe_devid(eb, btrfs_stripe_nr(c, nr));
}
/* struct btrfs_block_group_item */
BTRFS_SETGET_STACK_FUNCS(block_group_used, struct btrfs_block_group_item,
used, 64);
BTRFS_SETGET_FUNCS(disk_block_group_used, struct btrfs_block_group_item,
used, 64);
BTRFS_SETGET_STACK_FUNCS(block_group_chunk_objectid,
struct btrfs_block_group_item, chunk_objectid, 64);
BTRFS_SETGET_FUNCS(disk_block_group_chunk_objectid,
struct btrfs_block_group_item, chunk_objectid, 64);
BTRFS_SETGET_FUNCS(disk_block_group_flags,
struct btrfs_block_group_item, flags, 64);
BTRFS_SETGET_STACK_FUNCS(block_group_flags,
struct btrfs_block_group_item, flags, 64);
/* struct btrfs_free_space_info */
BTRFS_SETGET_FUNCS(free_space_extent_count, struct btrfs_free_space_info,
extent_count, 32);
BTRFS_SETGET_FUNCS(free_space_flags, struct btrfs_free_space_info, flags, 32);
/* struct btrfs_inode_ref */
BTRFS_SETGET_FUNCS(inode_ref_name_len, struct btrfs_inode_ref, name_len, 16);
BTRFS_SETGET_FUNCS(inode_ref_index, struct btrfs_inode_ref, index, 64);
/* struct btrfs_inode_extref */
BTRFS_SETGET_FUNCS(inode_extref_parent, struct btrfs_inode_extref,
parent_objectid, 64);
BTRFS_SETGET_FUNCS(inode_extref_name_len, struct btrfs_inode_extref,
name_len, 16);
BTRFS_SETGET_FUNCS(inode_extref_index, struct btrfs_inode_extref, index, 64);
/* struct btrfs_inode_item */
BTRFS_SETGET_FUNCS(inode_generation, struct btrfs_inode_item, generation, 64);
BTRFS_SETGET_FUNCS(inode_sequence, struct btrfs_inode_item, sequence, 64);
BTRFS_SETGET_FUNCS(inode_transid, struct btrfs_inode_item, transid, 64);
BTRFS_SETGET_FUNCS(inode_size, struct btrfs_inode_item, size, 64);
BTRFS_SETGET_FUNCS(inode_nbytes, struct btrfs_inode_item, nbytes, 64);
BTRFS_SETGET_FUNCS(inode_block_group, struct btrfs_inode_item, block_group, 64);
BTRFS_SETGET_FUNCS(inode_nlink, struct btrfs_inode_item, nlink, 32);
BTRFS_SETGET_FUNCS(inode_uid, struct btrfs_inode_item, uid, 32);
BTRFS_SETGET_FUNCS(inode_gid, struct btrfs_inode_item, gid, 32);
BTRFS_SETGET_FUNCS(inode_mode, struct btrfs_inode_item, mode, 32);
BTRFS_SETGET_FUNCS(inode_rdev, struct btrfs_inode_item, rdev, 64);
BTRFS_SETGET_FUNCS(inode_flags, struct btrfs_inode_item, flags, 64);
BTRFS_SETGET_STACK_FUNCS(stack_inode_generation, struct btrfs_inode_item,
generation, 64);
BTRFS_SETGET_STACK_FUNCS(stack_inode_sequence, struct btrfs_inode_item,
sequence, 64);
BTRFS_SETGET_STACK_FUNCS(stack_inode_transid, struct btrfs_inode_item,
transid, 64);
BTRFS_SETGET_STACK_FUNCS(stack_inode_size, struct btrfs_inode_item, size, 64);
BTRFS_SETGET_STACK_FUNCS(stack_inode_nbytes, struct btrfs_inode_item,
nbytes, 64);
BTRFS_SETGET_STACK_FUNCS(stack_inode_block_group, struct btrfs_inode_item,
block_group, 64);
BTRFS_SETGET_STACK_FUNCS(stack_inode_nlink, struct btrfs_inode_item, nlink, 32);
BTRFS_SETGET_STACK_FUNCS(stack_inode_uid, struct btrfs_inode_item, uid, 32);
BTRFS_SETGET_STACK_FUNCS(stack_inode_gid, struct btrfs_inode_item, gid, 32);
BTRFS_SETGET_STACK_FUNCS(stack_inode_mode, struct btrfs_inode_item, mode, 32);
BTRFS_SETGET_STACK_FUNCS(stack_inode_rdev, struct btrfs_inode_item, rdev, 64);
BTRFS_SETGET_STACK_FUNCS(stack_inode_flags, struct btrfs_inode_item, flags, 64);
BTRFS_SETGET_FUNCS(timespec_sec, struct btrfs_timespec, sec, 64);
BTRFS_SETGET_FUNCS(timespec_nsec, struct btrfs_timespec, nsec, 32);
BTRFS_SETGET_STACK_FUNCS(stack_timespec_sec, struct btrfs_timespec, sec, 64);
BTRFS_SETGET_STACK_FUNCS(stack_timespec_nsec, struct btrfs_timespec, nsec, 32);
/* struct btrfs_dev_extent */
BTRFS_SETGET_FUNCS(dev_extent_chunk_tree, struct btrfs_dev_extent,
chunk_tree, 64);
BTRFS_SETGET_FUNCS(dev_extent_chunk_objectid, struct btrfs_dev_extent,
chunk_objectid, 64);
BTRFS_SETGET_FUNCS(dev_extent_chunk_offset, struct btrfs_dev_extent,
chunk_offset, 64);
BTRFS_SETGET_FUNCS(dev_extent_length, struct btrfs_dev_extent, length, 64);
static inline unsigned long btrfs_dev_extent_chunk_tree_uuid(struct btrfs_dev_extent *dev)
{
unsigned long ptr = offsetof(struct btrfs_dev_extent, chunk_tree_uuid);
return (unsigned long)dev + ptr;
}
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
BTRFS_SETGET_FUNCS(extent_refs, struct btrfs_extent_item, refs, 64);
BTRFS_SETGET_FUNCS(extent_generation, struct btrfs_extent_item,
generation, 64);
BTRFS_SETGET_FUNCS(extent_flags, struct btrfs_extent_item, flags, 64);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
BTRFS_SETGET_FUNCS(extent_refs_v0, struct btrfs_extent_item_v0, refs, 32);
BTRFS_SETGET_FUNCS(tree_block_level, struct btrfs_tree_block_info, level, 8);
static inline void btrfs_tree_block_key(struct extent_buffer *eb,
struct btrfs_tree_block_info *item,
struct btrfs_disk_key *key)
{
read_eb_member(eb, item, struct btrfs_tree_block_info, key, key);
}
static inline void btrfs_set_tree_block_key(struct extent_buffer *eb,
struct btrfs_tree_block_info *item,
struct btrfs_disk_key *key)
{
write_eb_member(eb, item, struct btrfs_tree_block_info, key, key);
}
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
BTRFS_SETGET_FUNCS(extent_data_ref_root, struct btrfs_extent_data_ref,
root, 64);
BTRFS_SETGET_FUNCS(extent_data_ref_objectid, struct btrfs_extent_data_ref,
objectid, 64);
BTRFS_SETGET_FUNCS(extent_data_ref_offset, struct btrfs_extent_data_ref,
offset, 64);
BTRFS_SETGET_FUNCS(extent_data_ref_count, struct btrfs_extent_data_ref,
count, 32);
BTRFS_SETGET_FUNCS(shared_data_ref_count, struct btrfs_shared_data_ref,
count, 32);
BTRFS_SETGET_FUNCS(extent_inline_ref_type, struct btrfs_extent_inline_ref,
type, 8);
BTRFS_SETGET_FUNCS(extent_inline_ref_offset, struct btrfs_extent_inline_ref,
offset, 64);
static inline u32 btrfs_extent_inline_ref_size(int type)
{
if (type == BTRFS_TREE_BLOCK_REF_KEY ||
type == BTRFS_SHARED_BLOCK_REF_KEY)
return sizeof(struct btrfs_extent_inline_ref);
if (type == BTRFS_SHARED_DATA_REF_KEY)
return sizeof(struct btrfs_shared_data_ref) +
sizeof(struct btrfs_extent_inline_ref);
if (type == BTRFS_EXTENT_DATA_REF_KEY)
return sizeof(struct btrfs_extent_data_ref) +
offsetof(struct btrfs_extent_inline_ref, offset);
return 0;
}
BTRFS_SETGET_FUNCS(ref_root_v0, struct btrfs_extent_ref_v0, root, 64);
BTRFS_SETGET_FUNCS(ref_generation_v0, struct btrfs_extent_ref_v0,
generation, 64);
BTRFS_SETGET_FUNCS(ref_objectid_v0, struct btrfs_extent_ref_v0, objectid, 64);
BTRFS_SETGET_FUNCS(ref_count_v0, struct btrfs_extent_ref_v0, count, 32);
/* struct btrfs_node */
BTRFS_SETGET_FUNCS(key_blockptr, struct btrfs_key_ptr, blockptr, 64);
BTRFS_SETGET_FUNCS(key_generation, struct btrfs_key_ptr, generation, 64);
BTRFS_SETGET_STACK_FUNCS(stack_key_blockptr, struct btrfs_key_ptr,
blockptr, 64);
BTRFS_SETGET_STACK_FUNCS(stack_key_generation, struct btrfs_key_ptr,
generation, 64);
static inline u64 btrfs_node_blockptr(struct extent_buffer *eb, int nr)
{
unsigned long ptr;
ptr = offsetof(struct btrfs_node, ptrs) +
sizeof(struct btrfs_key_ptr) * nr;
return btrfs_key_blockptr(eb, (struct btrfs_key_ptr *)ptr);
}
static inline void btrfs_set_node_blockptr(struct extent_buffer *eb,
int nr, u64 val)
{
unsigned long ptr;
ptr = offsetof(struct btrfs_node, ptrs) +
sizeof(struct btrfs_key_ptr) * nr;
btrfs_set_key_blockptr(eb, (struct btrfs_key_ptr *)ptr, val);
}
static inline u64 btrfs_node_ptr_generation(struct extent_buffer *eb, int nr)
{
unsigned long ptr;
ptr = offsetof(struct btrfs_node, ptrs) +
sizeof(struct btrfs_key_ptr) * nr;
return btrfs_key_generation(eb, (struct btrfs_key_ptr *)ptr);
}
static inline void btrfs_set_node_ptr_generation(struct extent_buffer *eb,
int nr, u64 val)
{
unsigned long ptr;
ptr = offsetof(struct btrfs_node, ptrs) +
sizeof(struct btrfs_key_ptr) * nr;
btrfs_set_key_generation(eb, (struct btrfs_key_ptr *)ptr, val);
}
static inline unsigned long btrfs_node_key_ptr_offset(int nr)
{
return offsetof(struct btrfs_node, ptrs) +
sizeof(struct btrfs_key_ptr) * nr;
}
void btrfs_node_key(const struct extent_buffer *eb,
struct btrfs_disk_key *disk_key, int nr);
static inline void btrfs_set_node_key(struct extent_buffer *eb,
struct btrfs_disk_key *disk_key, int nr)
{
unsigned long ptr;
ptr = btrfs_node_key_ptr_offset(nr);
write_eb_member(eb, (struct btrfs_key_ptr *)ptr,
struct btrfs_key_ptr, key, disk_key);
}
/* struct btrfs_item */
BTRFS_SETGET_FUNCS(item_offset, struct btrfs_item, offset, 32);
BTRFS_SETGET_FUNCS(item_size, struct btrfs_item, size, 32);
BTRFS_SETGET_STACK_FUNCS(stack_item_offset, struct btrfs_item, offset, 32);
BTRFS_SETGET_STACK_FUNCS(stack_item_size, struct btrfs_item, size, 32);
static inline unsigned long btrfs_item_nr_offset(int nr)
{
return offsetof(struct btrfs_leaf, items) +
sizeof(struct btrfs_item) * nr;
}
static inline struct btrfs_item *btrfs_item_nr(int nr)
{
return (struct btrfs_item *)btrfs_item_nr_offset(nr);
}
static inline u32 btrfs_item_end(const struct extent_buffer *eb,
struct btrfs_item *item)
{
return btrfs_item_offset(eb, item) + btrfs_item_size(eb, item);
}
static inline u32 btrfs_item_end_nr(const struct extent_buffer *eb, int nr)
{
return btrfs_item_end(eb, btrfs_item_nr(nr));
}
static inline u32 btrfs_item_offset_nr(const struct extent_buffer *eb, int nr)
{
return btrfs_item_offset(eb, btrfs_item_nr(nr));
}
static inline u32 btrfs_item_size_nr(const struct extent_buffer *eb, int nr)
{
return btrfs_item_size(eb, btrfs_item_nr(nr));
}
static inline void btrfs_item_key(const struct extent_buffer *eb,
struct btrfs_disk_key *disk_key, int nr)
{
struct btrfs_item *item = btrfs_item_nr(nr);
read_eb_member(eb, item, struct btrfs_item, key, disk_key);
}
static inline void btrfs_set_item_key(struct extent_buffer *eb,
struct btrfs_disk_key *disk_key, int nr)
{
struct btrfs_item *item = btrfs_item_nr(nr);
write_eb_member(eb, item, struct btrfs_item, key, disk_key);
}
BTRFS_SETGET_FUNCS(dir_log_end, struct btrfs_dir_log_item, end, 64);
/*
* struct btrfs_root_ref
*/
BTRFS_SETGET_FUNCS(root_ref_dirid, struct btrfs_root_ref, dirid, 64);
BTRFS_SETGET_FUNCS(root_ref_sequence, struct btrfs_root_ref, sequence, 64);
BTRFS_SETGET_FUNCS(root_ref_name_len, struct btrfs_root_ref, name_len, 16);
/* struct btrfs_dir_item */
BTRFS_SETGET_FUNCS(dir_data_len, struct btrfs_dir_item, data_len, 16);
BTRFS_SETGET_FUNCS(dir_type, struct btrfs_dir_item, type, 8);
BTRFS_SETGET_FUNCS(dir_name_len, struct btrfs_dir_item, name_len, 16);
BTRFS_SETGET_FUNCS(dir_transid, struct btrfs_dir_item, transid, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dir_type, struct btrfs_dir_item, type, 8);
BTRFS_SETGET_STACK_FUNCS(stack_dir_data_len, struct btrfs_dir_item,
data_len, 16);
BTRFS_SETGET_STACK_FUNCS(stack_dir_name_len, struct btrfs_dir_item,
name_len, 16);
BTRFS_SETGET_STACK_FUNCS(stack_dir_transid, struct btrfs_dir_item,
transid, 64);
static inline void btrfs_dir_item_key(const struct extent_buffer *eb,
const struct btrfs_dir_item *item,
struct btrfs_disk_key *key)
{
read_eb_member(eb, item, struct btrfs_dir_item, location, key);
}
static inline void btrfs_set_dir_item_key(struct extent_buffer *eb,
struct btrfs_dir_item *item,
const struct btrfs_disk_key *key)
{
write_eb_member(eb, item, struct btrfs_dir_item, location, key);
}
BTRFS_SETGET_FUNCS(free_space_entries, struct btrfs_free_space_header,
num_entries, 64);
BTRFS_SETGET_FUNCS(free_space_bitmaps, struct btrfs_free_space_header,
num_bitmaps, 64);
BTRFS_SETGET_FUNCS(free_space_generation, struct btrfs_free_space_header,
generation, 64);
static inline void btrfs_free_space_key(const struct extent_buffer *eb,
const struct btrfs_free_space_header *h,
struct btrfs_disk_key *key)
{
read_eb_member(eb, h, struct btrfs_free_space_header, location, key);
}
static inline void btrfs_set_free_space_key(struct extent_buffer *eb,
struct btrfs_free_space_header *h,
const struct btrfs_disk_key *key)
{
write_eb_member(eb, h, struct btrfs_free_space_header, location, key);
}
/* struct btrfs_disk_key */
BTRFS_SETGET_STACK_FUNCS(disk_key_objectid, struct btrfs_disk_key,
objectid, 64);
BTRFS_SETGET_STACK_FUNCS(disk_key_offset, struct btrfs_disk_key, offset, 64);
BTRFS_SETGET_STACK_FUNCS(disk_key_type, struct btrfs_disk_key, type, 8);
static inline void btrfs_disk_key_to_cpu(struct btrfs_key *cpu,
const struct btrfs_disk_key *disk)
{
cpu->offset = le64_to_cpu(disk->offset);
cpu->type = disk->type;
cpu->objectid = le64_to_cpu(disk->objectid);
}
static inline void btrfs_cpu_key_to_disk(struct btrfs_disk_key *disk,
const struct btrfs_key *cpu)
{
disk->offset = cpu_to_le64(cpu->offset);
disk->type = cpu->type;
disk->objectid = cpu_to_le64(cpu->objectid);
}
static inline void btrfs_node_key_to_cpu(const struct extent_buffer *eb,
struct btrfs_key *key, int nr)
{
struct btrfs_disk_key disk_key;
btrfs_node_key(eb, &disk_key, nr);
btrfs_disk_key_to_cpu(key, &disk_key);
}
static inline void btrfs_item_key_to_cpu(const struct extent_buffer *eb,
struct btrfs_key *key, int nr)
{
struct btrfs_disk_key disk_key;
btrfs_item_key(eb, &disk_key, nr);
btrfs_disk_key_to_cpu(key, &disk_key);
}
static inline void btrfs_dir_item_key_to_cpu(const struct extent_buffer *eb,
const struct btrfs_dir_item *item,
struct btrfs_key *key)
{
struct btrfs_disk_key disk_key;
btrfs_dir_item_key(eb, item, &disk_key);
btrfs_disk_key_to_cpu(key, &disk_key);
}
static inline u8 btrfs_key_type(const struct btrfs_key *key)
{
return key->type;
}
static inline void btrfs_set_key_type(struct btrfs_key *key, u8 val)
{
key->type = val;
}
/* struct btrfs_header */
BTRFS_SETGET_HEADER_FUNCS(header_bytenr, struct btrfs_header, bytenr, 64);
BTRFS_SETGET_HEADER_FUNCS(header_generation, struct btrfs_header,
generation, 64);
BTRFS_SETGET_HEADER_FUNCS(header_owner, struct btrfs_header, owner, 64);
BTRFS_SETGET_HEADER_FUNCS(header_nritems, struct btrfs_header, nritems, 32);
BTRFS_SETGET_HEADER_FUNCS(header_flags, struct btrfs_header, flags, 64);
BTRFS_SETGET_HEADER_FUNCS(header_level, struct btrfs_header, level, 8);
BTRFS_SETGET_STACK_FUNCS(stack_header_generation, struct btrfs_header,
generation, 64);
BTRFS_SETGET_STACK_FUNCS(stack_header_owner, struct btrfs_header, owner, 64);
BTRFS_SETGET_STACK_FUNCS(stack_header_nritems, struct btrfs_header,
nritems, 32);
BTRFS_SETGET_STACK_FUNCS(stack_header_bytenr, struct btrfs_header, bytenr, 64);
static inline int btrfs_header_flag(const struct extent_buffer *eb, u64 flag)
{
return (btrfs_header_flags(eb) & flag) == flag;
}
static inline void btrfs_set_header_flag(struct extent_buffer *eb, u64 flag)
{
u64 flags = btrfs_header_flags(eb);
btrfs_set_header_flags(eb, flags | flag);
}
static inline void btrfs_clear_header_flag(struct extent_buffer *eb, u64 flag)
{
u64 flags = btrfs_header_flags(eb);
btrfs_set_header_flags(eb, flags & ~flag);
}
static inline int btrfs_header_backref_rev(const struct extent_buffer *eb)
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
{
u64 flags = btrfs_header_flags(eb);
return flags >> BTRFS_BACKREF_REV_SHIFT;
}
static inline void btrfs_set_header_backref_rev(struct extent_buffer *eb,
int rev)
{
u64 flags = btrfs_header_flags(eb);
flags &= ~BTRFS_BACKREF_REV_MASK;
flags |= (u64)rev << BTRFS_BACKREF_REV_SHIFT;
btrfs_set_header_flags(eb, flags);
}
static inline unsigned long btrfs_header_fsid(void)
{
return offsetof(struct btrfs_header, fsid);
}
static inline unsigned long btrfs_header_chunk_tree_uuid(const struct extent_buffer *eb)
{
return offsetof(struct btrfs_header, chunk_tree_uuid);
}
static inline int btrfs_is_leaf(const struct extent_buffer *eb)
{
return btrfs_header_level(eb) == 0;
}
/* struct btrfs_root_item */
BTRFS_SETGET_FUNCS(disk_root_generation, struct btrfs_root_item,
generation, 64);
BTRFS_SETGET_FUNCS(disk_root_refs, struct btrfs_root_item, refs, 32);
BTRFS_SETGET_FUNCS(disk_root_bytenr, struct btrfs_root_item, bytenr, 64);
BTRFS_SETGET_FUNCS(disk_root_level, struct btrfs_root_item, level, 8);
BTRFS_SETGET_STACK_FUNCS(root_generation, struct btrfs_root_item,
generation, 64);
BTRFS_SETGET_STACK_FUNCS(root_bytenr, struct btrfs_root_item, bytenr, 64);
BTRFS_SETGET_STACK_FUNCS(root_level, struct btrfs_root_item, level, 8);
BTRFS_SETGET_STACK_FUNCS(root_dirid, struct btrfs_root_item, root_dirid, 64);
BTRFS_SETGET_STACK_FUNCS(root_refs, struct btrfs_root_item, refs, 32);
BTRFS_SETGET_STACK_FUNCS(root_flags, struct btrfs_root_item, flags, 64);
BTRFS_SETGET_STACK_FUNCS(root_used, struct btrfs_root_item, bytes_used, 64);
BTRFS_SETGET_STACK_FUNCS(root_limit, struct btrfs_root_item, byte_limit, 64);
BTRFS_SETGET_STACK_FUNCS(root_last_snapshot, struct btrfs_root_item,
last_snapshot, 64);
BTRFS_SETGET_STACK_FUNCS(root_generation_v2, struct btrfs_root_item,
generation_v2, 64);
BTRFS_SETGET_STACK_FUNCS(root_ctransid, struct btrfs_root_item,
ctransid, 64);
BTRFS_SETGET_STACK_FUNCS(root_otransid, struct btrfs_root_item,
otransid, 64);
BTRFS_SETGET_STACK_FUNCS(root_stransid, struct btrfs_root_item,
stransid, 64);
BTRFS_SETGET_STACK_FUNCS(root_rtransid, struct btrfs_root_item,
rtransid, 64);
static inline bool btrfs_root_readonly(const struct btrfs_root *root)
{
return (root->root_item.flags & cpu_to_le64(BTRFS_ROOT_SUBVOL_RDONLY)) != 0;
}
static inline bool btrfs_root_dead(const struct btrfs_root *root)
{
return (root->root_item.flags & cpu_to_le64(BTRFS_ROOT_SUBVOL_DEAD)) != 0;
}
/* struct btrfs_root_backup */
BTRFS_SETGET_STACK_FUNCS(backup_tree_root, struct btrfs_root_backup,
tree_root, 64);
BTRFS_SETGET_STACK_FUNCS(backup_tree_root_gen, struct btrfs_root_backup,
tree_root_gen, 64);
BTRFS_SETGET_STACK_FUNCS(backup_tree_root_level, struct btrfs_root_backup,
tree_root_level, 8);
BTRFS_SETGET_STACK_FUNCS(backup_chunk_root, struct btrfs_root_backup,
chunk_root, 64);
BTRFS_SETGET_STACK_FUNCS(backup_chunk_root_gen, struct btrfs_root_backup,
chunk_root_gen, 64);
BTRFS_SETGET_STACK_FUNCS(backup_chunk_root_level, struct btrfs_root_backup,
chunk_root_level, 8);
BTRFS_SETGET_STACK_FUNCS(backup_extent_root, struct btrfs_root_backup,
extent_root, 64);
BTRFS_SETGET_STACK_FUNCS(backup_extent_root_gen, struct btrfs_root_backup,
extent_root_gen, 64);
BTRFS_SETGET_STACK_FUNCS(backup_extent_root_level, struct btrfs_root_backup,
extent_root_level, 8);
BTRFS_SETGET_STACK_FUNCS(backup_fs_root, struct btrfs_root_backup,
fs_root, 64);
BTRFS_SETGET_STACK_FUNCS(backup_fs_root_gen, struct btrfs_root_backup,
fs_root_gen, 64);
BTRFS_SETGET_STACK_FUNCS(backup_fs_root_level, struct btrfs_root_backup,
fs_root_level, 8);
BTRFS_SETGET_STACK_FUNCS(backup_dev_root, struct btrfs_root_backup,
dev_root, 64);
BTRFS_SETGET_STACK_FUNCS(backup_dev_root_gen, struct btrfs_root_backup,
dev_root_gen, 64);
BTRFS_SETGET_STACK_FUNCS(backup_dev_root_level, struct btrfs_root_backup,
dev_root_level, 8);
BTRFS_SETGET_STACK_FUNCS(backup_csum_root, struct btrfs_root_backup,
csum_root, 64);
BTRFS_SETGET_STACK_FUNCS(backup_csum_root_gen, struct btrfs_root_backup,
csum_root_gen, 64);
BTRFS_SETGET_STACK_FUNCS(backup_csum_root_level, struct btrfs_root_backup,
csum_root_level, 8);
BTRFS_SETGET_STACK_FUNCS(backup_total_bytes, struct btrfs_root_backup,
total_bytes, 64);
BTRFS_SETGET_STACK_FUNCS(backup_bytes_used, struct btrfs_root_backup,
bytes_used, 64);
BTRFS_SETGET_STACK_FUNCS(backup_num_devices, struct btrfs_root_backup,
num_devices, 64);
/* struct btrfs_balance_item */
BTRFS_SETGET_FUNCS(balance_flags, struct btrfs_balance_item, flags, 64);
static inline void btrfs_balance_data(const struct extent_buffer *eb,
const struct btrfs_balance_item *bi,
struct btrfs_disk_balance_args *ba)
{
read_eb_member(eb, bi, struct btrfs_balance_item, data, ba);
}
static inline void btrfs_set_balance_data(struct extent_buffer *eb,
struct btrfs_balance_item *bi,
const struct btrfs_disk_balance_args *ba)
{
write_eb_member(eb, bi, struct btrfs_balance_item, data, ba);
}
static inline void btrfs_balance_meta(const struct extent_buffer *eb,
const struct btrfs_balance_item *bi,
struct btrfs_disk_balance_args *ba)
{
read_eb_member(eb, bi, struct btrfs_balance_item, meta, ba);
}
static inline void btrfs_set_balance_meta(struct extent_buffer *eb,
struct btrfs_balance_item *bi,
const struct btrfs_disk_balance_args *ba)
{
write_eb_member(eb, bi, struct btrfs_balance_item, meta, ba);
}
static inline void btrfs_balance_sys(const struct extent_buffer *eb,
const struct btrfs_balance_item *bi,
struct btrfs_disk_balance_args *ba)
{
read_eb_member(eb, bi, struct btrfs_balance_item, sys, ba);
}
static inline void btrfs_set_balance_sys(struct extent_buffer *eb,
struct btrfs_balance_item *bi,
const struct btrfs_disk_balance_args *ba)
{
write_eb_member(eb, bi, struct btrfs_balance_item, sys, ba);
}
static inline void
btrfs_disk_balance_args_to_cpu(struct btrfs_balance_args *cpu,
const struct btrfs_disk_balance_args *disk)
{
memset(cpu, 0, sizeof(*cpu));
cpu->profiles = le64_to_cpu(disk->profiles);
cpu->usage = le64_to_cpu(disk->usage);
cpu->devid = le64_to_cpu(disk->devid);
cpu->pstart = le64_to_cpu(disk->pstart);
cpu->pend = le64_to_cpu(disk->pend);
cpu->vstart = le64_to_cpu(disk->vstart);
cpu->vend = le64_to_cpu(disk->vend);
cpu->target = le64_to_cpu(disk->target);
cpu->flags = le64_to_cpu(disk->flags);
cpu->limit = le64_to_cpu(disk->limit);
cpu->stripes_min = le32_to_cpu(disk->stripes_min);
cpu->stripes_max = le32_to_cpu(disk->stripes_max);
}
static inline void
btrfs_cpu_balance_args_to_disk(struct btrfs_disk_balance_args *disk,
const struct btrfs_balance_args *cpu)
{
memset(disk, 0, sizeof(*disk));
disk->profiles = cpu_to_le64(cpu->profiles);
disk->usage = cpu_to_le64(cpu->usage);
disk->devid = cpu_to_le64(cpu->devid);
disk->pstart = cpu_to_le64(cpu->pstart);
disk->pend = cpu_to_le64(cpu->pend);
disk->vstart = cpu_to_le64(cpu->vstart);
disk->vend = cpu_to_le64(cpu->vend);
disk->target = cpu_to_le64(cpu->target);
disk->flags = cpu_to_le64(cpu->flags);
disk->limit = cpu_to_le64(cpu->limit);
disk->stripes_min = cpu_to_le32(cpu->stripes_min);
disk->stripes_max = cpu_to_le32(cpu->stripes_max);
}
/* struct btrfs_super_block */
BTRFS_SETGET_STACK_FUNCS(super_bytenr, struct btrfs_super_block, bytenr, 64);
BTRFS_SETGET_STACK_FUNCS(super_flags, struct btrfs_super_block, flags, 64);
BTRFS_SETGET_STACK_FUNCS(super_generation, struct btrfs_super_block,
generation, 64);
BTRFS_SETGET_STACK_FUNCS(super_root, struct btrfs_super_block, root, 64);
BTRFS_SETGET_STACK_FUNCS(super_sys_array_size,
struct btrfs_super_block, sys_chunk_array_size, 32);
BTRFS_SETGET_STACK_FUNCS(super_chunk_root_generation,
struct btrfs_super_block, chunk_root_generation, 64);
BTRFS_SETGET_STACK_FUNCS(super_root_level, struct btrfs_super_block,
root_level, 8);
BTRFS_SETGET_STACK_FUNCS(super_chunk_root, struct btrfs_super_block,
chunk_root, 64);
BTRFS_SETGET_STACK_FUNCS(super_chunk_root_level, struct btrfs_super_block,
chunk_root_level, 8);
BTRFS_SETGET_STACK_FUNCS(super_log_root, struct btrfs_super_block,
log_root, 64);
BTRFS_SETGET_STACK_FUNCS(super_log_root_transid, struct btrfs_super_block,
log_root_transid, 64);
BTRFS_SETGET_STACK_FUNCS(super_log_root_level, struct btrfs_super_block,
log_root_level, 8);
BTRFS_SETGET_STACK_FUNCS(super_total_bytes, struct btrfs_super_block,
total_bytes, 64);
BTRFS_SETGET_STACK_FUNCS(super_bytes_used, struct btrfs_super_block,
bytes_used, 64);
BTRFS_SETGET_STACK_FUNCS(super_sectorsize, struct btrfs_super_block,
sectorsize, 32);
BTRFS_SETGET_STACK_FUNCS(super_nodesize, struct btrfs_super_block,
nodesize, 32);
BTRFS_SETGET_STACK_FUNCS(super_stripesize, struct btrfs_super_block,
stripesize, 32);
BTRFS_SETGET_STACK_FUNCS(super_root_dir, struct btrfs_super_block,
root_dir_objectid, 64);
BTRFS_SETGET_STACK_FUNCS(super_num_devices, struct btrfs_super_block,
num_devices, 64);
BTRFS_SETGET_STACK_FUNCS(super_compat_flags, struct btrfs_super_block,
compat_flags, 64);
BTRFS_SETGET_STACK_FUNCS(super_compat_ro_flags, struct btrfs_super_block,
compat_ro_flags, 64);
BTRFS_SETGET_STACK_FUNCS(super_incompat_flags, struct btrfs_super_block,
incompat_flags, 64);
BTRFS_SETGET_STACK_FUNCS(super_csum_type, struct btrfs_super_block,
csum_type, 16);
BTRFS_SETGET_STACK_FUNCS(super_cache_generation, struct btrfs_super_block,
cache_generation, 64);
BTRFS_SETGET_STACK_FUNCS(super_magic, struct btrfs_super_block, magic, 64);
BTRFS_SETGET_STACK_FUNCS(super_uuid_tree_generation, struct btrfs_super_block,
uuid_tree_generation, 64);
static inline int btrfs_super_csum_size(const struct btrfs_super_block *s)
{
u16 t = btrfs_super_csum_type(s);
/*
* csum type is validated at mount time
*/
return btrfs_csum_sizes[t];
}
/*
* The leaf data grows from end-to-front in the node.
* this returns the address of the start of the last item,
* which is the stop of the leaf data stack
*/
static inline unsigned int leaf_data_end(const struct extent_buffer *leaf)
{
u32 nr = btrfs_header_nritems(leaf);
if (nr == 0)
return BTRFS_LEAF_DATA_SIZE(leaf->fs_info);
return btrfs_item_offset_nr(leaf, nr - 1);
}
/* struct btrfs_file_extent_item */
BTRFS_SETGET_FUNCS(file_extent_type, struct btrfs_file_extent_item, type, 8);
BTRFS_SETGET_STACK_FUNCS(stack_file_extent_disk_bytenr,
struct btrfs_file_extent_item, disk_bytenr, 64);
BTRFS_SETGET_STACK_FUNCS(stack_file_extent_offset,
struct btrfs_file_extent_item, offset, 64);
BTRFS_SETGET_STACK_FUNCS(stack_file_extent_generation,
struct btrfs_file_extent_item, generation, 64);
BTRFS_SETGET_STACK_FUNCS(stack_file_extent_num_bytes,
struct btrfs_file_extent_item, num_bytes, 64);
BTRFS_SETGET_STACK_FUNCS(stack_file_extent_disk_num_bytes,
struct btrfs_file_extent_item, disk_num_bytes, 64);
BTRFS_SETGET_STACK_FUNCS(stack_file_extent_compression,
struct btrfs_file_extent_item, compression, 8);
static inline unsigned long
btrfs_file_extent_inline_start(const struct btrfs_file_extent_item *e)
{
return (unsigned long)e + BTRFS_FILE_EXTENT_INLINE_DATA_START;
}
static inline u32 btrfs_file_extent_calc_inline_size(u32 datasize)
{
return BTRFS_FILE_EXTENT_INLINE_DATA_START + datasize;
}
BTRFS_SETGET_FUNCS(file_extent_disk_bytenr, struct btrfs_file_extent_item,
disk_bytenr, 64);
BTRFS_SETGET_FUNCS(file_extent_generation, struct btrfs_file_extent_item,
generation, 64);
BTRFS_SETGET_FUNCS(file_extent_disk_num_bytes, struct btrfs_file_extent_item,
disk_num_bytes, 64);
BTRFS_SETGET_FUNCS(file_extent_offset, struct btrfs_file_extent_item,
offset, 64);
BTRFS_SETGET_FUNCS(file_extent_num_bytes, struct btrfs_file_extent_item,
num_bytes, 64);
Btrfs: Add zlib compression support This is a large change for adding compression on reading and writing, both for inline and regular extents. It does some fairly large surgery to the writeback paths. Compression is off by default and enabled by mount -o compress. Even when the -o compress mount option is not used, it is possible to read compressed extents off the disk. If compression for a given set of pages fails to make them smaller, the file is flagged to avoid future compression attempts later. * While finding delalloc extents, the pages are locked before being sent down to the delalloc handler. This allows the delalloc handler to do complex things such as cleaning the pages, marking them writeback and starting IO on their behalf. * Inline extents are inserted at delalloc time now. This allows us to compress the data before inserting the inline extent, and it allows us to insert an inline extent that spans multiple pages. * All of the in-memory extent representations (extent_map.c, ordered-data.c etc) are changed to record both an in-memory size and an on disk size, as well as a flag for compression. From a disk format point of view, the extent pointers in the file are changed to record the on disk size of a given extent and some encoding flags. Space in the disk format is allocated for compression encoding, as well as encryption and a generic 'other' field. Neither the encryption or the 'other' field are currently used. In order to limit the amount of data read for a single random read in the file, the size of a compressed extent is limited to 128k. This is a software only limit, the disk format supports u64 sized compressed extents. In order to limit the ram consumed while processing extents, the uncompressed size of a compressed extent is limited to 256k. This is a software only limit and will be subject to tuning later. Checksumming is still done on compressed extents, and it is done on the uncompressed version of the data. This way additional encodings can be layered on without having to figure out which encoding to checksum. Compression happens at delalloc time, which is basically singled threaded because it is usually done by a single pdflush thread. This makes it tricky to spread the compression load across all the cpus on the box. We'll have to look at parallel pdflush walks of dirty inodes at a later time. Decompression is hooked into readpages and it does spread across CPUs nicely. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-10-29 18:49:59 +00:00
BTRFS_SETGET_FUNCS(file_extent_ram_bytes, struct btrfs_file_extent_item,
ram_bytes, 64);
BTRFS_SETGET_FUNCS(file_extent_compression, struct btrfs_file_extent_item,
compression, 8);
BTRFS_SETGET_FUNCS(file_extent_encryption, struct btrfs_file_extent_item,
encryption, 8);
BTRFS_SETGET_FUNCS(file_extent_other_encoding, struct btrfs_file_extent_item,
other_encoding, 16);
/*
* this returns the number of bytes used by the item on disk, minus the
* size of any extent headers. If a file is compressed on disk, this is
* the compressed size
*/
static inline u32 btrfs_file_extent_inline_item_len(
const struct extent_buffer *eb,
struct btrfs_item *e)
Btrfs: Add zlib compression support This is a large change for adding compression on reading and writing, both for inline and regular extents. It does some fairly large surgery to the writeback paths. Compression is off by default and enabled by mount -o compress. Even when the -o compress mount option is not used, it is possible to read compressed extents off the disk. If compression for a given set of pages fails to make them smaller, the file is flagged to avoid future compression attempts later. * While finding delalloc extents, the pages are locked before being sent down to the delalloc handler. This allows the delalloc handler to do complex things such as cleaning the pages, marking them writeback and starting IO on their behalf. * Inline extents are inserted at delalloc time now. This allows us to compress the data before inserting the inline extent, and it allows us to insert an inline extent that spans multiple pages. * All of the in-memory extent representations (extent_map.c, ordered-data.c etc) are changed to record both an in-memory size and an on disk size, as well as a flag for compression. From a disk format point of view, the extent pointers in the file are changed to record the on disk size of a given extent and some encoding flags. Space in the disk format is allocated for compression encoding, as well as encryption and a generic 'other' field. Neither the encryption or the 'other' field are currently used. In order to limit the amount of data read for a single random read in the file, the size of a compressed extent is limited to 128k. This is a software only limit, the disk format supports u64 sized compressed extents. In order to limit the ram consumed while processing extents, the uncompressed size of a compressed extent is limited to 256k. This is a software only limit and will be subject to tuning later. Checksumming is still done on compressed extents, and it is done on the uncompressed version of the data. This way additional encodings can be layered on without having to figure out which encoding to checksum. Compression happens at delalloc time, which is basically singled threaded because it is usually done by a single pdflush thread. This makes it tricky to spread the compression load across all the cpus on the box. We'll have to look at parallel pdflush walks of dirty inodes at a later time. Decompression is hooked into readpages and it does spread across CPUs nicely. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-10-29 18:49:59 +00:00
{
return btrfs_item_size(eb, e) - BTRFS_FILE_EXTENT_INLINE_DATA_START;
Btrfs: Add zlib compression support This is a large change for adding compression on reading and writing, both for inline and regular extents. It does some fairly large surgery to the writeback paths. Compression is off by default and enabled by mount -o compress. Even when the -o compress mount option is not used, it is possible to read compressed extents off the disk. If compression for a given set of pages fails to make them smaller, the file is flagged to avoid future compression attempts later. * While finding delalloc extents, the pages are locked before being sent down to the delalloc handler. This allows the delalloc handler to do complex things such as cleaning the pages, marking them writeback and starting IO on their behalf. * Inline extents are inserted at delalloc time now. This allows us to compress the data before inserting the inline extent, and it allows us to insert an inline extent that spans multiple pages. * All of the in-memory extent representations (extent_map.c, ordered-data.c etc) are changed to record both an in-memory size and an on disk size, as well as a flag for compression. From a disk format point of view, the extent pointers in the file are changed to record the on disk size of a given extent and some encoding flags. Space in the disk format is allocated for compression encoding, as well as encryption and a generic 'other' field. Neither the encryption or the 'other' field are currently used. In order to limit the amount of data read for a single random read in the file, the size of a compressed extent is limited to 128k. This is a software only limit, the disk format supports u64 sized compressed extents. In order to limit the ram consumed while processing extents, the uncompressed size of a compressed extent is limited to 256k. This is a software only limit and will be subject to tuning later. Checksumming is still done on compressed extents, and it is done on the uncompressed version of the data. This way additional encodings can be layered on without having to figure out which encoding to checksum. Compression happens at delalloc time, which is basically singled threaded because it is usually done by a single pdflush thread. This makes it tricky to spread the compression load across all the cpus on the box. We'll have to look at parallel pdflush walks of dirty inodes at a later time. Decompression is hooked into readpages and it does spread across CPUs nicely. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-10-29 18:49:59 +00:00
}
/* btrfs_dev_stats_item */
static inline u64 btrfs_dev_stats_value(const struct extent_buffer *eb,
const struct btrfs_dev_stats_item *ptr,
int index)
{
u64 val;
read_extent_buffer(eb, &val,
offsetof(struct btrfs_dev_stats_item, values) +
((unsigned long)ptr) + (index * sizeof(u64)),
sizeof(val));
return val;
}
static inline void btrfs_set_dev_stats_value(struct extent_buffer *eb,
struct btrfs_dev_stats_item *ptr,
int index, u64 val)
{
write_extent_buffer(eb, &val,
offsetof(struct btrfs_dev_stats_item, values) +
((unsigned long)ptr) + (index * sizeof(u64)),
sizeof(val));
}
/* btrfs_qgroup_status_item */
BTRFS_SETGET_FUNCS(qgroup_status_generation, struct btrfs_qgroup_status_item,
generation, 64);
BTRFS_SETGET_FUNCS(qgroup_status_version, struct btrfs_qgroup_status_item,
version, 64);
BTRFS_SETGET_FUNCS(qgroup_status_flags, struct btrfs_qgroup_status_item,
flags, 64);
BTRFS_SETGET_FUNCS(qgroup_status_rescan, struct btrfs_qgroup_status_item,
rescan, 64);
/* btrfs_qgroup_info_item */
BTRFS_SETGET_FUNCS(qgroup_info_generation, struct btrfs_qgroup_info_item,
generation, 64);
BTRFS_SETGET_FUNCS(qgroup_info_rfer, struct btrfs_qgroup_info_item, rfer, 64);
BTRFS_SETGET_FUNCS(qgroup_info_rfer_cmpr, struct btrfs_qgroup_info_item,
rfer_cmpr, 64);
BTRFS_SETGET_FUNCS(qgroup_info_excl, struct btrfs_qgroup_info_item, excl, 64);
BTRFS_SETGET_FUNCS(qgroup_info_excl_cmpr, struct btrfs_qgroup_info_item,
excl_cmpr, 64);
BTRFS_SETGET_STACK_FUNCS(stack_qgroup_info_generation,
struct btrfs_qgroup_info_item, generation, 64);
BTRFS_SETGET_STACK_FUNCS(stack_qgroup_info_rfer, struct btrfs_qgroup_info_item,
rfer, 64);
BTRFS_SETGET_STACK_FUNCS(stack_qgroup_info_rfer_cmpr,
struct btrfs_qgroup_info_item, rfer_cmpr, 64);
BTRFS_SETGET_STACK_FUNCS(stack_qgroup_info_excl, struct btrfs_qgroup_info_item,
excl, 64);
BTRFS_SETGET_STACK_FUNCS(stack_qgroup_info_excl_cmpr,
struct btrfs_qgroup_info_item, excl_cmpr, 64);
/* btrfs_qgroup_limit_item */
BTRFS_SETGET_FUNCS(qgroup_limit_flags, struct btrfs_qgroup_limit_item,
flags, 64);
BTRFS_SETGET_FUNCS(qgroup_limit_max_rfer, struct btrfs_qgroup_limit_item,
max_rfer, 64);
BTRFS_SETGET_FUNCS(qgroup_limit_max_excl, struct btrfs_qgroup_limit_item,
max_excl, 64);
BTRFS_SETGET_FUNCS(qgroup_limit_rsv_rfer, struct btrfs_qgroup_limit_item,
rsv_rfer, 64);
BTRFS_SETGET_FUNCS(qgroup_limit_rsv_excl, struct btrfs_qgroup_limit_item,
rsv_excl, 64);
/* btrfs_dev_replace_item */
BTRFS_SETGET_FUNCS(dev_replace_src_devid,
struct btrfs_dev_replace_item, src_devid, 64);
BTRFS_SETGET_FUNCS(dev_replace_cont_reading_from_srcdev_mode,
struct btrfs_dev_replace_item, cont_reading_from_srcdev_mode,
64);
BTRFS_SETGET_FUNCS(dev_replace_replace_state, struct btrfs_dev_replace_item,
replace_state, 64);
BTRFS_SETGET_FUNCS(dev_replace_time_started, struct btrfs_dev_replace_item,
time_started, 64);
BTRFS_SETGET_FUNCS(dev_replace_time_stopped, struct btrfs_dev_replace_item,
time_stopped, 64);
BTRFS_SETGET_FUNCS(dev_replace_num_write_errors, struct btrfs_dev_replace_item,
num_write_errors, 64);
BTRFS_SETGET_FUNCS(dev_replace_num_uncorrectable_read_errors,
struct btrfs_dev_replace_item, num_uncorrectable_read_errors,
64);
BTRFS_SETGET_FUNCS(dev_replace_cursor_left, struct btrfs_dev_replace_item,
cursor_left, 64);
BTRFS_SETGET_FUNCS(dev_replace_cursor_right, struct btrfs_dev_replace_item,
cursor_right, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_src_devid,
struct btrfs_dev_replace_item, src_devid, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_cont_reading_from_srcdev_mode,
struct btrfs_dev_replace_item,
cont_reading_from_srcdev_mode, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_replace_state,
struct btrfs_dev_replace_item, replace_state, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_time_started,
struct btrfs_dev_replace_item, time_started, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_time_stopped,
struct btrfs_dev_replace_item, time_stopped, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_num_write_errors,
struct btrfs_dev_replace_item, num_write_errors, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_num_uncorrectable_read_errors,
struct btrfs_dev_replace_item,
num_uncorrectable_read_errors, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_cursor_left,
struct btrfs_dev_replace_item, cursor_left, 64);
BTRFS_SETGET_STACK_FUNCS(stack_dev_replace_cursor_right,
struct btrfs_dev_replace_item, cursor_right, 64);
/* helper function to cast into the data area of the leaf. */
#define btrfs_item_ptr(leaf, slot, type) \
((type *)(BTRFS_LEAF_DATA_OFFSET + \
btrfs_item_offset_nr(leaf, slot)))
#define btrfs_item_ptr_offset(leaf, slot) \
((unsigned long)(BTRFS_LEAF_DATA_OFFSET + \
btrfs_item_offset_nr(leaf, slot)))
btrfs: Remove custom crc32c init code The custom crc32 init code was introduced in 14a958e678cd ("Btrfs: fix btrfs boot when compiled as built-in") to enable using btrfs as a built-in. However, later as pointed out by 60efa5eb2e88 ("Btrfs: use late_initcall instead of module_init") this wasn't enough and finally btrfs was switched to late_initcall which comes after the generic crc32c implementation is initiliased. The latter commit superseeded the former. Now that we don't have to maintain our own code let's just remove it and switch to using the generic implementation. Despite touching a lot of files the patch is really simple. Here is the gist of the changes: 1. Select LIBCRC32C rather than the low-level modules. 2. s/btrfs_crc32c/crc32c/g 3. replace hash.h with linux/crc32c.h 4. Move the btrfs namehash funcs to ctree.h and change the tree accordingly. I've tested this with btrfs being both a module and a built-in and xfstest doesn't complain. Does seem to fix the longstanding problem of not automatically selectiong the crc32c module when btrfs is used. Possibly there is a workaround in dracut. The modinfo confirms that now all the module dependencies are there: before: depends: zstd_compress,zstd_decompress,raid6_pq,xor,zlib_deflate after: depends: libcrc32c,zstd_compress,zstd_decompress,raid6_pq,xor,zlib_deflate Signed-off-by: Nikolay Borisov <nborisov@suse.com> Reviewed-by: David Sterba <dsterba@suse.com> [ add more info to changelog from mails ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-01-08 09:45:05 +00:00
static inline u64 btrfs_name_hash(const char *name, int len)
{
return crc32c((u32)~1, name, len);
}
/*
* Figure the key offset of an extended inode ref
*/
static inline u64 btrfs_extref_hash(u64 parent_objectid, const char *name,
int len)
{
return (u64) crc32c(parent_objectid, name, len);
}
static inline bool btrfs_mixed_space_info(struct btrfs_space_info *space_info)
{
return ((space_info->flags & BTRFS_BLOCK_GROUP_METADATA) &&
(space_info->flags & BTRFS_BLOCK_GROUP_DATA));
}
static inline gfp_t btrfs_alloc_write_mask(struct address_space *mapping)
{
return mapping_gfp_constraint(mapping, ~__GFP_FS);
}
/* extent-tree.c */
enum btrfs_inline_ref_type {
BTRFS_REF_TYPE_INVALID,
BTRFS_REF_TYPE_BLOCK,
BTRFS_REF_TYPE_DATA,
BTRFS_REF_TYPE_ANY,
};
int btrfs_get_extent_inline_ref_type(const struct extent_buffer *eb,
struct btrfs_extent_inline_ref *iref,
enum btrfs_inline_ref_type is_data);
u64 btrfs_csum_bytes_to_leaves(struct btrfs_fs_info *fs_info, u64 csum_bytes);
static inline u64 btrfs_calc_trans_metadata_size(struct btrfs_fs_info *fs_info,
unsigned num_items)
btrfs: implement delayed inode items operation Changelog V5 -> V6: - Fix oom when the memory load is high, by storing the delayed nodes into the root's radix tree, and letting btrfs inodes go. Changelog V4 -> V5: - Fix the race on adding the delayed node to the inode, which is spotted by Chris Mason. - Merge Chris Mason's incremental patch into this patch. - Fix deadlock between readdir() and memory fault, which is reported by Itaru Kitayama. Changelog V3 -> V4: - Fix nested lock, which is reported by Itaru Kitayama, by updating space cache inode in time. Changelog V2 -> V3: - Fix the race between the delayed worker and the task which does delayed items balance, which is reported by Tsutomu Itoh. - Modify the patch address David Sterba's comment. - Fix the bug of the cpu recursion spinlock, reported by Chris Mason Changelog V1 -> V2: - break up the global rb-tree, use a list to manage the delayed nodes, which is created for every directory and file, and used to manage the delayed directory name index items and the delayed inode item. - introduce a worker to deal with the delayed nodes. Compare with Ext3/4, the performance of file creation and deletion on btrfs is very poor. the reason is that btrfs must do a lot of b+ tree insertions, such as inode item, directory name item, directory name index and so on. If we can do some delayed b+ tree insertion or deletion, we can improve the performance, so we made this patch which implemented delayed directory name index insertion/deletion and delayed inode update. Implementation: - introduce a delayed root object into the filesystem, that use two lists to manage the delayed nodes which are created for every file/directory. One is used to manage all the delayed nodes that have delayed items. And the other is used to manage the delayed nodes which is waiting to be dealt with by the work thread. - Every delayed node has two rb-tree, one is used to manage the directory name index which is going to be inserted into b+ tree, and the other is used to manage the directory name index which is going to be deleted from b+ tree. - introduce a worker to deal with the delayed operation. This worker is used to deal with the works of the delayed directory name index items insertion and deletion and the delayed inode update. When the delayed items is beyond the lower limit, we create works for some delayed nodes and insert them into the work queue of the worker, and then go back. When the delayed items is beyond the upper bound, we create works for all the delayed nodes that haven't been dealt with, and insert them into the work queue of the worker, and then wait for that the untreated items is below some threshold value. - When we want to insert a directory name index into b+ tree, we just add the information into the delayed inserting rb-tree. And then we check the number of the delayed items and do delayed items balance. (The balance policy is above.) - When we want to delete a directory name index from the b+ tree, we search it in the inserting rb-tree at first. If we look it up, just drop it. If not, add the key of it into the delayed deleting rb-tree. Similar to the delayed inserting rb-tree, we also check the number of the delayed items and do delayed items balance. (The same to inserting manipulation) - When we want to update the metadata of some inode, we cached the data of the inode into the delayed node. the worker will flush it into the b+ tree after dealing with the delayed insertion and deletion. - We will move the delayed node to the tail of the list after we access the delayed node, By this way, we can cache more delayed items and merge more inode updates. - If we want to commit transaction, we will deal with all the delayed node. - the delayed node will be freed when we free the btrfs inode. - Before we log the inode items, we commit all the directory name index items and the delayed inode update. I did a quick test by the benchmark tool[1] and found we can improve the performance of file creation by ~15%, and file deletion by ~20%. Before applying this patch: Create files: Total files: 50000 Total time: 1.096108 Average time: 0.000022 Delete files: Total files: 50000 Total time: 1.510403 Average time: 0.000030 After applying this patch: Create files: Total files: 50000 Total time: 0.932899 Average time: 0.000019 Delete files: Total files: 50000 Total time: 1.215732 Average time: 0.000024 [1] http://marc.info/?l=linux-btrfs&m=128212635122920&q=p3 Many thanks for Kitayama-san's help! Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Reviewed-by: David Sterba <dave@jikos.cz> Tested-by: Tsutomu Itoh <t-itoh@jp.fujitsu.com> Tested-by: Itaru Kitayama <kitayama@cl.bb4u.ne.jp> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2011-04-22 10:12:22 +00:00
{
Btrfs: fix delalloc accounting leak caused by u32 overflow btrfs_calc_trans_metadata_size() does an unsigned 32-bit multiplication, which can overflow if num_items >= 4 GB / (nodesize * BTRFS_MAX_LEVEL * 2). For a nodesize of 16kB, this overflow happens at 16k items. Usually, num_items is a small constant passed to btrfs_start_transaction(), but we also use btrfs_calc_trans_metadata_size() for metadata reservations for extent items in btrfs_delalloc_{reserve,release}_metadata(). In drop_outstanding_extents(), num_items is calculated as inode->reserved_extents - inode->outstanding_extents. The difference between these two counters is usually small, but if many delalloc extents are reserved and then the outstanding extents are merged in btrfs_merge_extent_hook(), the difference can become large enough to overflow in btrfs_calc_trans_metadata_size(). The overflow manifests itself as a leak of a multiple of 4 GB in delalloc_block_rsv and the metadata bytes_may_use counter. This in turn can cause early ENOSPC errors. Additionally, these WARN_ONs in extent-tree.c will be hit when unmounting: WARN_ON(fs_info->delalloc_block_rsv.size > 0); WARN_ON(fs_info->delalloc_block_rsv.reserved > 0); WARN_ON(space_info->bytes_pinned > 0 || space_info->bytes_reserved > 0 || space_info->bytes_may_use > 0); Fix it by casting nodesize to a u64 so that btrfs_calc_trans_metadata_size() does a full 64-bit multiplication. While we're here, do the same in btrfs_calc_trunc_metadata_size(); this can't overflow with any existing uses, but it's better to be safe here than have another hard-to-debug problem later on. Cc: stable@vger.kernel.org Signed-off-by: Omar Sandoval <osandov@fb.com> Reviewed-by: David Sterba <dsterba@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2017-06-02 08:20:01 +00:00
return (u64)fs_info->nodesize * BTRFS_MAX_LEVEL * 2 * num_items;
}
/*
* Doing a truncate won't result in new nodes or leaves, just what we need for
* COW.
*/
static inline u64 btrfs_calc_trunc_metadata_size(struct btrfs_fs_info *fs_info,
unsigned num_items)
{
Btrfs: fix delalloc accounting leak caused by u32 overflow btrfs_calc_trans_metadata_size() does an unsigned 32-bit multiplication, which can overflow if num_items >= 4 GB / (nodesize * BTRFS_MAX_LEVEL * 2). For a nodesize of 16kB, this overflow happens at 16k items. Usually, num_items is a small constant passed to btrfs_start_transaction(), but we also use btrfs_calc_trans_metadata_size() for metadata reservations for extent items in btrfs_delalloc_{reserve,release}_metadata(). In drop_outstanding_extents(), num_items is calculated as inode->reserved_extents - inode->outstanding_extents. The difference between these two counters is usually small, but if many delalloc extents are reserved and then the outstanding extents are merged in btrfs_merge_extent_hook(), the difference can become large enough to overflow in btrfs_calc_trans_metadata_size(). The overflow manifests itself as a leak of a multiple of 4 GB in delalloc_block_rsv and the metadata bytes_may_use counter. This in turn can cause early ENOSPC errors. Additionally, these WARN_ONs in extent-tree.c will be hit when unmounting: WARN_ON(fs_info->delalloc_block_rsv.size > 0); WARN_ON(fs_info->delalloc_block_rsv.reserved > 0); WARN_ON(space_info->bytes_pinned > 0 || space_info->bytes_reserved > 0 || space_info->bytes_may_use > 0); Fix it by casting nodesize to a u64 so that btrfs_calc_trans_metadata_size() does a full 64-bit multiplication. While we're here, do the same in btrfs_calc_trunc_metadata_size(); this can't overflow with any existing uses, but it's better to be safe here than have another hard-to-debug problem later on. Cc: stable@vger.kernel.org Signed-off-by: Omar Sandoval <osandov@fb.com> Reviewed-by: David Sterba <dsterba@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2017-06-02 08:20:01 +00:00
return (u64)fs_info->nodesize * BTRFS_MAX_LEVEL * num_items;
btrfs: implement delayed inode items operation Changelog V5 -> V6: - Fix oom when the memory load is high, by storing the delayed nodes into the root's radix tree, and letting btrfs inodes go. Changelog V4 -> V5: - Fix the race on adding the delayed node to the inode, which is spotted by Chris Mason. - Merge Chris Mason's incremental patch into this patch. - Fix deadlock between readdir() and memory fault, which is reported by Itaru Kitayama. Changelog V3 -> V4: - Fix nested lock, which is reported by Itaru Kitayama, by updating space cache inode in time. Changelog V2 -> V3: - Fix the race between the delayed worker and the task which does delayed items balance, which is reported by Tsutomu Itoh. - Modify the patch address David Sterba's comment. - Fix the bug of the cpu recursion spinlock, reported by Chris Mason Changelog V1 -> V2: - break up the global rb-tree, use a list to manage the delayed nodes, which is created for every directory and file, and used to manage the delayed directory name index items and the delayed inode item. - introduce a worker to deal with the delayed nodes. Compare with Ext3/4, the performance of file creation and deletion on btrfs is very poor. the reason is that btrfs must do a lot of b+ tree insertions, such as inode item, directory name item, directory name index and so on. If we can do some delayed b+ tree insertion or deletion, we can improve the performance, so we made this patch which implemented delayed directory name index insertion/deletion and delayed inode update. Implementation: - introduce a delayed root object into the filesystem, that use two lists to manage the delayed nodes which are created for every file/directory. One is used to manage all the delayed nodes that have delayed items. And the other is used to manage the delayed nodes which is waiting to be dealt with by the work thread. - Every delayed node has two rb-tree, one is used to manage the directory name index which is going to be inserted into b+ tree, and the other is used to manage the directory name index which is going to be deleted from b+ tree. - introduce a worker to deal with the delayed operation. This worker is used to deal with the works of the delayed directory name index items insertion and deletion and the delayed inode update. When the delayed items is beyond the lower limit, we create works for some delayed nodes and insert them into the work queue of the worker, and then go back. When the delayed items is beyond the upper bound, we create works for all the delayed nodes that haven't been dealt with, and insert them into the work queue of the worker, and then wait for that the untreated items is below some threshold value. - When we want to insert a directory name index into b+ tree, we just add the information into the delayed inserting rb-tree. And then we check the number of the delayed items and do delayed items balance. (The balance policy is above.) - When we want to delete a directory name index from the b+ tree, we search it in the inserting rb-tree at first. If we look it up, just drop it. If not, add the key of it into the delayed deleting rb-tree. Similar to the delayed inserting rb-tree, we also check the number of the delayed items and do delayed items balance. (The same to inserting manipulation) - When we want to update the metadata of some inode, we cached the data of the inode into the delayed node. the worker will flush it into the b+ tree after dealing with the delayed insertion and deletion. - We will move the delayed node to the tail of the list after we access the delayed node, By this way, we can cache more delayed items and merge more inode updates. - If we want to commit transaction, we will deal with all the delayed node. - the delayed node will be freed when we free the btrfs inode. - Before we log the inode items, we commit all the directory name index items and the delayed inode update. I did a quick test by the benchmark tool[1] and found we can improve the performance of file creation by ~15%, and file deletion by ~20%. Before applying this patch: Create files: Total files: 50000 Total time: 1.096108 Average time: 0.000022 Delete files: Total files: 50000 Total time: 1.510403 Average time: 0.000030 After applying this patch: Create files: Total files: 50000 Total time: 0.932899 Average time: 0.000019 Delete files: Total files: 50000 Total time: 1.215732 Average time: 0.000024 [1] http://marc.info/?l=linux-btrfs&m=128212635122920&q=p3 Many thanks for Kitayama-san's help! Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Reviewed-by: David Sterba <dave@jikos.cz> Tested-by: Tsutomu Itoh <t-itoh@jp.fujitsu.com> Tested-by: Itaru Kitayama <kitayama@cl.bb4u.ne.jp> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2011-04-22 10:12:22 +00:00
}
int btrfs_should_throttle_delayed_refs(struct btrfs_trans_handle *trans);
bool btrfs_check_space_for_delayed_refs(struct btrfs_fs_info *fs_info);
Btrfs: don't do unnecessary delalloc flushes when relocating Before we start the actual relocation process of a block group, we do calls to flush delalloc of all inodes and then wait for ordered extents to complete. However we do these flush calls just to make sure we don't race with concurrent tasks that have actually already started to run delalloc and have allocated an extent from the block group we want to relocate, right before we set it to readonly mode, but have not yet created the respective ordered extents. The flush calls make us wait for such concurrent tasks because they end up calling filemap_fdatawrite_range() (through btrfs_start_delalloc_roots() -> __start_delalloc_inodes() -> btrfs_alloc_delalloc_work() -> btrfs_run_delalloc_work()) which ends up serializing us with those tasks due to attempts to lock the same pages (and the delalloc flush procedure calls the allocator and creates the ordered extents before unlocking the pages). These flushing calls not only make us waste time (cpu, IO) but also reduce the chances of writing larger extents (applications might be writing to contiguous ranges and we flush before they finish dirtying the whole ranges). So make sure we don't flush delalloc and just wait for concurrent tasks that have already started flushing delalloc and have allocated an extent from the block group we are about to relocate. This change also ends up fixing a race with direct IO writes that makes relocation not wait for direct IO ordered extents. This race is illustrated by the following diagram: CPU 1 CPU 2 btrfs_relocate_block_group(bg X) starts direct IO write, target inode currently has no ordered extents ongoing nor dirty pages (delalloc regions), therefore the root for our inode is not in the list fs_info->ordered_roots btrfs_direct_IO() __blockdev_direct_IO() btrfs_get_blocks_direct() btrfs_lock_extent_direct() locks range in the io tree btrfs_new_extent_direct() btrfs_reserve_extent() --> extent allocated from bg X btrfs_inc_block_group_ro(bg X) btrfs_start_delalloc_roots() __start_delalloc_inodes() --> does nothing, no dealloc ranges in the inode's io tree so the inode's root is not in the list fs_info->delalloc_roots btrfs_wait_ordered_roots() --> does not find the inode's root in the list fs_info->ordered_roots --> ends up not waiting for the direct IO write started by the task at CPU 2 relocate_block_group(rc->stage == MOVE_DATA_EXTENTS) prepare_to_relocate() btrfs_commit_transaction() iterates the extent tree, using its commit root and moves extents into new locations btrfs_add_ordered_extent_dio() --> now a ordered extent is created and added to the list root->ordered_extents and the root added to the list fs_info->ordered_roots --> this is too late and the task at CPU 1 already started the relocation btrfs_commit_transaction() btrfs_finish_ordered_io() btrfs_alloc_reserved_file_extent() --> adds delayed data reference for the extent allocated from bg X relocate_block_group(rc->stage == UPDATE_DATA_PTRS) prepare_to_relocate() btrfs_commit_transaction() --> delayed refs are run, so an extent item for the allocated extent from bg X is added to extent tree --> commit roots are switched, so the next scan in the extent tree will see the extent item sees the extent in the extent tree When this happens the relocation produces the following warning when it finishes: [ 7260.832836] ------------[ cut here ]------------ [ 7260.834653] WARNING: CPU: 5 PID: 6765 at fs/btrfs/relocation.c:4318 btrfs_relocate_block_group+0x245/0x2a1 [btrfs]() [ 7260.838268] Modules linked in: btrfs crc32c_generic xor ppdev raid6_pq psmouse sg acpi_cpufreq evdev i2c_piix4 tpm_tis serio_raw tpm i2c_core pcspkr parport_pc [ 7260.850935] CPU: 5 PID: 6765 Comm: btrfs Not tainted 4.5.0-rc6-btrfs-next-28+ #1 [ 7260.852998] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS by qemu-project.org 04/01/2014 [ 7260.852998] 0000000000000000 ffff88020bf57bc0 ffffffff812648b3 0000000000000000 [ 7260.852998] 0000000000000009 ffff88020bf57bf8 ffffffff81051608 ffffffffa03c1b2d [ 7260.852998] ffff8800b2bbb800 0000000000000000 ffff8800b17bcc58 ffff8800399dd000 [ 7260.852998] Call Trace: [ 7260.852998] [<ffffffff812648b3>] dump_stack+0x67/0x90 [ 7260.852998] [<ffffffff81051608>] warn_slowpath_common+0x99/0xb2 [ 7260.852998] [<ffffffffa03c1b2d>] ? btrfs_relocate_block_group+0x245/0x2a1 [btrfs] [ 7260.852998] [<ffffffff810516d4>] warn_slowpath_null+0x1a/0x1c [ 7260.852998] [<ffffffffa03c1b2d>] btrfs_relocate_block_group+0x245/0x2a1 [btrfs] [ 7260.852998] [<ffffffffa039d9de>] btrfs_relocate_chunk.isra.29+0x66/0xdb [btrfs] [ 7260.852998] [<ffffffffa039f314>] btrfs_balance+0xde1/0xe4e [btrfs] [ 7260.852998] [<ffffffff8127d671>] ? debug_smp_processor_id+0x17/0x19 [ 7260.852998] [<ffffffffa03a9583>] btrfs_ioctl_balance+0x255/0x2d3 [btrfs] [ 7260.852998] [<ffffffffa03ac96a>] btrfs_ioctl+0x11e0/0x1dff [btrfs] [ 7260.852998] [<ffffffff811451df>] ? handle_mm_fault+0x443/0xd63 [ 7260.852998] [<ffffffff81491817>] ? _raw_spin_unlock+0x31/0x44 [ 7260.852998] [<ffffffff8108b36a>] ? arch_local_irq_save+0x9/0xc [ 7260.852998] [<ffffffff811876ab>] vfs_ioctl+0x18/0x34 [ 7260.852998] [<ffffffff81187cb2>] do_vfs_ioctl+0x550/0x5be [ 7260.852998] [<ffffffff81190c30>] ? __fget_light+0x4d/0x71 [ 7260.852998] [<ffffffff81187d77>] SyS_ioctl+0x57/0x79 [ 7260.852998] [<ffffffff81492017>] entry_SYSCALL_64_fastpath+0x12/0x6b [ 7260.893268] ---[ end trace eb7803b24ebab8ad ]--- This is because at the end of the first stage, in relocate_block_group(), we commit the current transaction, which makes delayed refs run, the commit roots are switched and so the second stage will find the extent item that the ordered extent added to the delayed refs. But this extent was not moved (ordered extent completed after first stage finished), so at the end of the relocation our block group item still has a positive used bytes counter, triggering a warning at the end of btrfs_relocate_block_group(). Later on when trying to read the extent contents from disk we hit a BUG_ON() due to the inability to map a block with a logical address that belongs to the block group we relocated and is no longer valid, resulting in the following trace: [ 7344.885290] BTRFS critical (device sdi): unable to find logical 12845056 len 4096 [ 7344.887518] ------------[ cut here ]------------ [ 7344.888431] kernel BUG at fs/btrfs/inode.c:1833! [ 7344.888431] invalid opcode: 0000 [#1] PREEMPT SMP DEBUG_PAGEALLOC [ 7344.888431] Modules linked in: btrfs crc32c_generic xor ppdev raid6_pq psmouse sg acpi_cpufreq evdev i2c_piix4 tpm_tis serio_raw tpm i2c_core pcspkr parport_pc [ 7344.888431] CPU: 0 PID: 6831 Comm: od Tainted: G W 4.5.0-rc6-btrfs-next-28+ #1 [ 7344.888431] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS by qemu-project.org 04/01/2014 [ 7344.888431] task: ffff880215818600 ti: ffff880204684000 task.ti: ffff880204684000 [ 7344.888431] RIP: 0010:[<ffffffffa037c88c>] [<ffffffffa037c88c>] btrfs_merge_bio_hook+0x54/0x6b [btrfs] [ 7344.888431] RSP: 0018:ffff8802046878f0 EFLAGS: 00010282 [ 7344.888431] RAX: 00000000ffffffea RBX: 0000000000001000 RCX: 0000000000000001 [ 7344.888431] RDX: ffff88023ec0f950 RSI: ffffffff8183b638 RDI: 00000000ffffffff [ 7344.888431] RBP: ffff880204687908 R08: 0000000000000001 R09: 0000000000000000 [ 7344.888431] R10: ffff880204687770 R11: ffffffff82f2d52d R12: 0000000000001000 [ 7344.888431] R13: ffff88021afbfee8 R14: 0000000000006208 R15: ffff88006cd199b0 [ 7344.888431] FS: 00007f1f9e1d6700(0000) GS:ffff88023ec00000(0000) knlGS:0000000000000000 [ 7344.888431] CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033 [ 7344.888431] CR2: 00007f1f9dc8cb60 CR3: 000000023e3b6000 CR4: 00000000000006f0 [ 7344.888431] Stack: [ 7344.888431] 0000000000001000 0000000000001000 ffff880204687b98 ffff880204687950 [ 7344.888431] ffffffffa0395c8f ffffea0004d64d48 0000000000000000 0000000000001000 [ 7344.888431] ffffea0004d64d48 0000000000001000 0000000000000000 0000000000000000 [ 7344.888431] Call Trace: [ 7344.888431] [<ffffffffa0395c8f>] submit_extent_page+0xf5/0x16f [btrfs] [ 7344.888431] [<ffffffffa03970ac>] __do_readpage+0x4a0/0x4f1 [btrfs] [ 7344.888431] [<ffffffffa039680d>] ? btrfs_create_repair_bio+0xcb/0xcb [btrfs] [ 7344.888431] [<ffffffffa037eeb4>] ? btrfs_writepage_start_hook+0xbc/0xbc [btrfs] [ 7344.888431] [<ffffffff8108df55>] ? trace_hardirqs_on+0xd/0xf [ 7344.888431] [<ffffffffa039728c>] __do_contiguous_readpages.constprop.26+0xc2/0xe4 [btrfs] [ 7344.888431] [<ffffffffa037eeb4>] ? btrfs_writepage_start_hook+0xbc/0xbc [btrfs] [ 7344.888431] [<ffffffffa039739b>] __extent_readpages.constprop.25+0xed/0x100 [btrfs] [ 7344.888431] [<ffffffff81129d24>] ? lru_cache_add+0xe/0x10 [ 7344.888431] [<ffffffffa0397ea8>] extent_readpages+0x160/0x1aa [btrfs] [ 7344.888431] [<ffffffffa037eeb4>] ? btrfs_writepage_start_hook+0xbc/0xbc [btrfs] [ 7344.888431] [<ffffffff8115daad>] ? alloc_pages_current+0xa9/0xcd [ 7344.888431] [<ffffffffa037cdc9>] btrfs_readpages+0x1f/0x21 [btrfs] [ 7344.888431] [<ffffffff81128316>] __do_page_cache_readahead+0x168/0x1fc [ 7344.888431] [<ffffffff811285a0>] ondemand_readahead+0x1f6/0x207 [ 7344.888431] [<ffffffff811285a0>] ? ondemand_readahead+0x1f6/0x207 [ 7344.888431] [<ffffffff8111cf34>] ? pagecache_get_page+0x2b/0x154 [ 7344.888431] [<ffffffff8112870e>] page_cache_sync_readahead+0x3d/0x3f [ 7344.888431] [<ffffffff8111dbf7>] generic_file_read_iter+0x197/0x4e1 [ 7344.888431] [<ffffffff8117773a>] __vfs_read+0x79/0x9d [ 7344.888431] [<ffffffff81178050>] vfs_read+0x8f/0xd2 [ 7344.888431] [<ffffffff81178a38>] SyS_read+0x50/0x7e [ 7344.888431] [<ffffffff81492017>] entry_SYSCALL_64_fastpath+0x12/0x6b [ 7344.888431] Code: 8d 4d e8 45 31 c9 45 31 c0 48 8b 00 48 c1 e2 09 48 8b 80 80 fc ff ff 4c 89 65 e8 48 8b b8 f0 01 00 00 e8 1d 42 02 00 85 c0 79 02 <0f> 0b 4c 0 [ 7344.888431] RIP [<ffffffffa037c88c>] btrfs_merge_bio_hook+0x54/0x6b [btrfs] [ 7344.888431] RSP <ffff8802046878f0> [ 7344.970544] ---[ end trace eb7803b24ebab8ae ]--- Signed-off-by: Filipe Manana <fdmanana@suse.com> Reviewed-by: Josef Bacik <jbacik@fb.com> Reviewed-by: Liu Bo <bo.li.liu@oracle.com>
2016-04-26 14:39:32 +00:00
void btrfs_dec_block_group_reservations(struct btrfs_fs_info *fs_info,
const u64 start);
void btrfs_wait_block_group_reservations(struct btrfs_block_group_cache *bg);
Btrfs: fix race between block group relocation and nocow writes Relocation of a block group waits for all existing tasks flushing dellaloc, starting direct IO writes and any ordered extents before starting the relocation process. However for direct IO writes that end up doing nocow (inode either has the flag nodatacow set or the write is against a prealloc extent) we have a short time window that allows for a race that makes relocation proceed without waiting for the direct IO write to complete first, resulting in data loss after the relocation finishes. This is illustrated by the following diagram: CPU 1 CPU 2 btrfs_relocate_block_group(bg X) direct IO write starts against an extent in block group X using nocow mode (inode has the nodatacow flag or the write is for a prealloc extent) btrfs_direct_IO() btrfs_get_blocks_direct() --> can_nocow_extent() returns 1 btrfs_inc_block_group_ro(bg X) --> turns block group into RO mode btrfs_wait_ordered_roots() --> returns and does not know about the DIO write happening at CPU 2 (the task there has not created yet an ordered extent) relocate_block_group(bg X) --> rc->stage == MOVE_DATA_EXTENTS find_next_extent() --> returns extent that the DIO write is going to write to relocate_data_extent() relocate_file_extent_cluster() --> reads the extent from disk into pages belonging to the relocation inode and dirties them --> creates DIO ordered extent btrfs_submit_direct() --> submits bio against a location on disk obtained from an extent map before the relocation started btrfs_wait_ordered_range() --> writes all the pages read before to disk (belonging to the relocation inode) relocation finishes bio completes and wrote new data to the old location of the block group So fix this by tracking the number of nocow writers for a block group and make sure relocation waits for that number to go down to 0 before starting to move the extents. The same race can also happen with buffered writes in nocow mode since the patch I recently made titled "Btrfs: don't do unnecessary delalloc flushes when relocating", because we are no longer flushing all delalloc which served as a synchonization mechanism (due to page locking) and ensured the ordered extents for nocow buffered writes were created before we called btrfs_wait_ordered_roots(). The race with direct IO writes in nocow mode existed before that patch (no pages are locked or used during direct IO) and that fixed only races with direct IO writes that do cow. Signed-off-by: Filipe Manana <fdmanana@suse.com> Reviewed-by: Josef Bacik <jbacik@fb.com>
2016-05-09 12:15:41 +00:00
bool btrfs_inc_nocow_writers(struct btrfs_fs_info *fs_info, u64 bytenr);
void btrfs_dec_nocow_writers(struct btrfs_fs_info *fs_info, u64 bytenr);
void btrfs_wait_nocow_writers(struct btrfs_block_group_cache *bg);
void btrfs_put_block_group(struct btrfs_block_group_cache *cache);
Btrfs: do extent allocation and reference count updates in the background The extent allocation tree maintains a reference count and full back reference information for every extent allocated in the filesystem. For subvolume and snapshot trees, every time a block goes through COW, the new copy of the block adds a reference on every block it points to. If a btree node points to 150 leaves, then the COW code needs to go and add backrefs on 150 different extents, which might be spread all over the extent allocation tree. These updates currently happen during btrfs_cow_block, and most COWs happen during btrfs_search_slot. btrfs_search_slot has locks held on both the parent and the node we are COWing, and so we really want to avoid IO during the COW if we can. This commit adds an rbtree of pending reference count updates and extent allocations. The tree is ordered by byte number of the extent and byte number of the parent for the back reference. The tree allows us to: 1) Modify back references in something close to disk order, reducing seeks 2) Significantly reduce the number of modifications made as block pointers are balanced around 3) Do all of the extent insertion and back reference modifications outside of the performance critical btrfs_search_slot code. #3 has the added benefit of greatly reducing the btrfs stack footprint. The extent allocation tree modifications are done without the deep (and somewhat recursive) call chains used in the past. These delayed back reference updates must be done before the transaction commits, and so the rbtree is tied to the transaction. Throttling is implemented to help keep the queue of backrefs at a reasonable size. Since there was a similar mechanism in place for the extent tree extents, that is removed and replaced by the delayed reference tree. Yan Zheng <yan.zheng@oracle.com> helped review and fixup this code. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-03-13 14:10:06 +00:00
int btrfs_run_delayed_refs(struct btrfs_trans_handle *trans,
unsigned long count);
int btrfs_async_run_delayed_refs(struct btrfs_fs_info *fs_info,
unsigned long count, u64 transid, int wait);
btrfs: handle delayed ref head accounting cleanup in abort We weren't doing any of the accounting cleanup when we aborted transactions. Fix this by making cleanup_ref_head_accounting global and calling it from the abort code, this fixes the issue where our accounting was all wrong after the fs aborts. The test generic/475 on a 2G VM can trigger the problems eg.: [ 8502.136957] WARNING: CPU: 0 PID: 11064 at fs/btrfs/extent-tree.c:5986 btrfs_free_block_grou +ps+0x3dc/0x410 [btrfs] [ 8502.148372] CPU: 0 PID: 11064 Comm: umount Not tainted 5.0.0-rc1-default+ #394 [ 8502.150807] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS rel-1.11.2-0-gf9626 +cc-prebuilt.qemu-project.org 04/01/2014 [ 8502.154317] RIP: 0010:btrfs_free_block_groups+0x3dc/0x410 [btrfs] [ 8502.160623] RSP: 0018:ffffb1ab84b93de8 EFLAGS: 00010206 [ 8502.161906] RAX: 0000000001000000 RBX: ffff9f34b1756400 RCX: 0000000000000000 [ 8502.163448] RDX: 0000000000000002 RSI: 0000000000000001 RDI: ffff9f34b1755400 [ 8502.164906] RBP: ffff9f34b7e8c000 R08: 0000000000000001 R09: 0000000000000000 [ 8502.166716] R10: 0000000000000000 R11: 0000000000000001 R12: ffff9f34b7e8c108 [ 8502.168498] R13: ffff9f34b7e8c158 R14: 0000000000000000 R15: dead000000000100 [ 8502.170296] FS: 00007fb1cf15ffc0(0000) GS:ffff9f34bd400000(0000) knlGS:0000000000000000 [ 8502.172439] CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033 [ 8502.173669] CR2: 00007fb1ced507b0 CR3: 000000002f7a6000 CR4: 00000000000006f0 [ 8502.175094] Call Trace: [ 8502.175759] close_ctree+0x17f/0x350 [btrfs] [ 8502.176721] generic_shutdown_super+0x64/0x100 [ 8502.177702] kill_anon_super+0x14/0x30 [ 8502.178607] btrfs_kill_super+0x12/0xa0 [btrfs] [ 8502.179602] deactivate_locked_super+0x29/0x60 [ 8502.180595] cleanup_mnt+0x3b/0x70 [ 8502.181406] task_work_run+0x98/0xc0 [ 8502.182255] exit_to_usermode_loop+0x83/0x90 [ 8502.183113] do_syscall_64+0x15b/0x180 [ 8502.183919] entry_SYSCALL_64_after_hwframe+0x49/0xbe Corresponding to release_global_block_rsv() { ... WARN_ON(fs_info->delayed_refs_rsv.reserved > 0); CC: stable@vger.kernel.org Signed-off-by: Josef Bacik <josef@toxicpanda.com> [ add log dump ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-11-21 19:05:41 +00:00
void btrfs_cleanup_ref_head_accounting(struct btrfs_fs_info *fs_info,
struct btrfs_delayed_ref_root *delayed_refs,
struct btrfs_delayed_ref_head *head);
int btrfs_lookup_data_extent(struct btrfs_fs_info *fs_info, u64 start, u64 len);
int btrfs_lookup_extent_info(struct btrfs_trans_handle *trans,
struct btrfs_fs_info *fs_info, u64 bytenr,
u64 offset, int metadata, u64 *refs, u64 *flags);
int btrfs_pin_extent(struct btrfs_fs_info *fs_info,
u64 bytenr, u64 num, int reserved);
int btrfs_pin_extent_for_log_replay(struct btrfs_fs_info *fs_info,
u64 bytenr, u64 num_bytes);
int btrfs_exclude_logged_extents(struct extent_buffer *eb);
int btrfs_cross_ref_exist(struct btrfs_root *root,
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
u64 objectid, u64 offset, u64 bytenr);
struct btrfs_block_group_cache *btrfs_lookup_block_group(
struct btrfs_fs_info *info,
u64 bytenr);
void btrfs_get_block_group(struct btrfs_block_group_cache *cache);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
void btrfs_put_block_group(struct btrfs_block_group_cache *cache);
struct extent_buffer *btrfs_alloc_tree_block(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
u64 parent, u64 root_objectid,
const struct btrfs_disk_key *key,
int level, u64 hint,
u64 empty_size);
void btrfs_free_tree_block(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct extent_buffer *buf,
u64 parent, int last_ref);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int btrfs_alloc_reserved_file_extent(struct btrfs_trans_handle *trans,
struct btrfs_root *root, u64 owner,
u64 offset, u64 ram_bytes,
struct btrfs_key *ins);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int btrfs_alloc_logged_file_extent(struct btrfs_trans_handle *trans,
u64 root_objectid, u64 owner, u64 offset,
struct btrfs_key *ins);
btrfs: update btrfs_space_info's bytes_may_use timely This patch can fix some false ENOSPC errors, below test script can reproduce one false ENOSPC error: #!/bin/bash dd if=/dev/zero of=fs.img bs=$((1024*1024)) count=128 dev=$(losetup --show -f fs.img) mkfs.btrfs -f -M $dev mkdir /tmp/mntpoint mount $dev /tmp/mntpoint cd /tmp/mntpoint xfs_io -f -c "falloc 0 $((64*1024*1024))" testfile Above script will fail for ENOSPC reason, but indeed fs still has free space to satisfy this request. Please see call graph: btrfs_fallocate() |-> btrfs_alloc_data_chunk_ondemand() | bytes_may_use += 64M |-> btrfs_prealloc_file_range() |-> btrfs_reserve_extent() |-> btrfs_add_reserved_bytes() | alloc_type is RESERVE_ALLOC_NO_ACCOUNT, so it does not | change bytes_may_use, and bytes_reserved += 64M. Now | bytes_may_use + bytes_reserved == 128M, which is greater | than btrfs_space_info's total_bytes, false enospc occurs. | Note, the bytes_may_use decrease operation will be done in | end of btrfs_fallocate(), which is too late. Here is another simple case for buffered write: CPU 1 | CPU 2 | |-> cow_file_range() |-> __btrfs_buffered_write() |-> btrfs_reserve_extent() | | | | | | | | | ..... | |-> btrfs_check_data_free_space() | | | | |-> extent_clear_unlock_delalloc() | In CPU 1, btrfs_reserve_extent()->find_free_extent()-> btrfs_add_reserved_bytes() do not decrease bytes_may_use, the decrease operation will be delayed to be done in extent_clear_unlock_delalloc(). Assume in this case, btrfs_reserve_extent() reserved 128MB data, CPU2's btrfs_check_data_free_space() tries to reserve 100MB data space. If 100MB > data_sinfo->total_bytes - data_sinfo->bytes_used - data_sinfo->bytes_reserved - data_sinfo->bytes_pinned - data_sinfo->bytes_readonly - data_sinfo->bytes_may_use btrfs_check_data_free_space() will try to allcate new data chunk or call btrfs_start_delalloc_roots(), or commit current transaction in order to reserve some free space, obviously a lot of work. But indeed it's not necessary as long as decreasing bytes_may_use timely, we still have free space, decreasing 128M from bytes_may_use. To fix this issue, this patch chooses to update bytes_may_use for both data and metadata in btrfs_add_reserved_bytes(). For compress path, real extent length may not be equal to file content length, so introduce a ram_bytes argument for btrfs_reserve_extent(), find_free_extent() and btrfs_add_reserved_bytes(), it's becasue bytes_may_use is increased by file content length. Then compress path can update bytes_may_use correctly. Also now we can discard RESERVE_ALLOC_NO_ACCOUNT, RESERVE_ALLOC and RESERVE_FREE. As we know, usually EXTENT_DO_ACCOUNTING is used for error path. In run_delalloc_nocow(), for inode marked as NODATACOW or extent marked as PREALLOC, we also need to update bytes_may_use, but can not pass EXTENT_DO_ACCOUNTING, because it also clears metadata reservation, so here we introduce EXTENT_CLEAR_DATA_RESV flag to indicate btrfs_clear_bit_hook() to update btrfs_space_info's bytes_may_use. Meanwhile __btrfs_prealloc_file_range() will call btrfs_free_reserved_data_space() internally for both sucessful and failed path, btrfs_prealloc_file_range()'s callers does not need to call btrfs_free_reserved_data_space() any more. Signed-off-by: Wang Xiaoguang <wangxg.fnst@cn.fujitsu.com> Reviewed-by: Josef Bacik <jbacik@fb.com> Signed-off-by: David Sterba <dsterba@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2016-07-25 07:51:40 +00:00
int btrfs_reserve_extent(struct btrfs_root *root, u64 ram_bytes, u64 num_bytes,
u64 min_alloc_size, u64 empty_size, u64 hint_byte,
Btrfs: fix broken free space cache after the system crashed When we mounted the filesystem after the crash, we got the following message: BTRFS error (device xxx): block group xxxx has wrong amount of free space BTRFS error (device xxx): failed to load free space cache for block group xxx It is because we didn't update the metadata of the allocated space (in extent tree) until the file data was written into the disk. During this time, there was no information about the allocated spaces in either the extent tree nor the free space cache. when we wrote out the free space cache at this time (commit transaction), those spaces were lost. In fact, only the free space that is used to store the file data had this problem, the others didn't because the metadata of them is updated in the same transaction context. There are many methods which can fix the above problem - track the allocated space, and write it out when we write out the free space cache - account the size of the allocated space that is used to store the file data, if the size is not zero, don't write out the free space cache. The first one is complex and may make the performance drop down. This patch chose the second method, we use a per-block-group variant to account the size of that allocated space. Besides that, we also introduce a per-block-group read-write semaphore to avoid the race between the allocation and the free space cache write out. Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-06-19 02:42:50 +00:00
struct btrfs_key *ins, int is_data, int delalloc);
int btrfs_inc_ref(struct btrfs_trans_handle *trans, struct btrfs_root *root,
struct extent_buffer *buf, int full_backref);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int btrfs_dec_ref(struct btrfs_trans_handle *trans, struct btrfs_root *root,
struct extent_buffer *buf, int full_backref);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int btrfs_set_disk_extent_flags(struct btrfs_trans_handle *trans,
struct btrfs_fs_info *fs_info,
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
u64 bytenr, u64 num_bytes, u64 flags,
int level, int is_data);
int btrfs_free_extent(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
u64 bytenr, u64 num_bytes, u64 parent, u64 root_objectid,
Btrfs: fix regression running delayed references when using qgroups In the kernel 4.2 merge window we had a big changes to the implementation of delayed references and qgroups which made the no_quota field of delayed references not used anymore. More specifically the no_quota field is not used anymore as of: commit 0ed4792af0e8 ("btrfs: qgroup: Switch to new extent-oriented qgroup mechanism.") Leaving the no_quota field actually prevents delayed references from getting merged, which in turn cause the following BUG_ON(), at fs/btrfs/extent-tree.c, to be hit when qgroups are enabled: static int run_delayed_tree_ref(...) { (...) BUG_ON(node->ref_mod != 1); (...) } This happens on a scenario like the following: 1) Ref1 bytenr X, action = BTRFS_ADD_DELAYED_REF, no_quota = 1, added. 2) Ref2 bytenr X, action = BTRFS_DROP_DELAYED_REF, no_quota = 0, added. It's not merged with Ref1 because Ref1->no_quota != Ref2->no_quota. 3) Ref3 bytenr X, action = BTRFS_ADD_DELAYED_REF, no_quota = 1, added. It's not merged with the reference at the tail of the list of refs for bytenr X because the reference at the tail, Ref2 is incompatible due to Ref2->no_quota != Ref3->no_quota. 4) Ref4 bytenr X, action = BTRFS_DROP_DELAYED_REF, no_quota = 0, added. It's not merged with the reference at the tail of the list of refs for bytenr X because the reference at the tail, Ref3 is incompatible due to Ref3->no_quota != Ref4->no_quota. 5) We run delayed references, trigger merging of delayed references, through __btrfs_run_delayed_refs() -> btrfs_merge_delayed_refs(). 6) Ref1 and Ref3 are merged as Ref1->no_quota = Ref3->no_quota and all other conditions are satisfied too. So Ref1 gets a ref_mod value of 2. 7) Ref2 and Ref4 are merged as Ref2->no_quota = Ref4->no_quota and all other conditions are satisfied too. So Ref2 gets a ref_mod value of 2. 8) Ref1 and Ref2 aren't merged, because they have different values for their no_quota field. 9) Delayed reference Ref1 is picked for running (select_delayed_ref() always prefers references with an action == BTRFS_ADD_DELAYED_REF). So run_delayed_tree_ref() is called for Ref1 which triggers the BUG_ON because Ref1->red_mod != 1 (equals 2). So fix this by removing the no_quota field, as it's not used anymore as of commit 0ed4792af0e8 ("btrfs: qgroup: Switch to new extent-oriented qgroup mechanism."). The use of no_quota was also buggy in at least two places: 1) At delayed-refs.c:btrfs_add_delayed_tree_ref() - we were setting no_quota to 0 instead of 1 when the following condition was true: is_fstree(ref_root) || !fs_info->quota_enabled 2) At extent-tree.c:__btrfs_inc_extent_ref() - we were attempting to reset a node's no_quota when the condition "!is_fstree(root_objectid) || !root->fs_info->quota_enabled" was true but we did it only in an unused local stack variable, that is, we never reset the no_quota value in the node itself. This fixes the remainder of problems several people have been having when running delayed references, mostly while a balance is running in parallel, on a 4.2+ kernel. Very special thanks to Stéphane Lesimple for helping debugging this issue and testing this fix on his multi terabyte filesystem (which took more than one day to balance alone, plus fsck, etc). Also, this fixes deadlock issue when using the clone ioctl with qgroups enabled, as reported by Elias Probst in the mailing list. The deadlock happens because after calling btrfs_insert_empty_item we have our path holding a write lock on a leaf of the fs/subvol tree and then before releasing the path we called check_ref() which did backref walking, when qgroups are enabled, and tried to read lock the same leaf. The trace for this case is the following: INFO: task systemd-nspawn:6095 blocked for more than 120 seconds. (...) Call Trace: [<ffffffff86999201>] schedule+0x74/0x83 [<ffffffff863ef64c>] btrfs_tree_read_lock+0xc0/0xea [<ffffffff86137ed7>] ? wait_woken+0x74/0x74 [<ffffffff8639f0a7>] btrfs_search_old_slot+0x51a/0x810 [<ffffffff863a129b>] btrfs_next_old_leaf+0xdf/0x3ce [<ffffffff86413a00>] ? ulist_add_merge+0x1b/0x127 [<ffffffff86411688>] __resolve_indirect_refs+0x62a/0x667 [<ffffffff863ef546>] ? btrfs_clear_lock_blocking_rw+0x78/0xbe [<ffffffff864122d3>] find_parent_nodes+0xaf3/0xfc6 [<ffffffff86412838>] __btrfs_find_all_roots+0x92/0xf0 [<ffffffff864128f2>] btrfs_find_all_roots+0x45/0x65 [<ffffffff8639a75b>] ? btrfs_get_tree_mod_seq+0x2b/0x88 [<ffffffff863e852e>] check_ref+0x64/0xc4 [<ffffffff863e9e01>] btrfs_clone+0x66e/0xb5d [<ffffffff863ea77f>] btrfs_ioctl_clone+0x48f/0x5bb [<ffffffff86048a68>] ? native_sched_clock+0x28/0x77 [<ffffffff863ed9b0>] btrfs_ioctl+0xabc/0x25cb (...) The problem goes away by eleminating check_ref(), which no longer is needed as its purpose was to get a value for the no_quota field of a delayed reference (this patch removes the no_quota field as mentioned earlier). Reported-by: Stéphane Lesimple <stephane_btrfs@lesimple.fr> Tested-by: Stéphane Lesimple <stephane_btrfs@lesimple.fr> Reported-by: Elias Probst <mail@eliasprobst.eu> Reported-by: Peter Becker <floyd.net@gmail.com> Reported-by: Malte Schröder <malte@tnxip.de> Reported-by: Derek Dongray <derek@valedon.co.uk> Reported-by: Erkki Seppala <flux-btrfs@inside.org> Cc: stable@vger.kernel.org # 4.2+ Signed-off-by: Filipe Manana <fdmanana@suse.com> Reviewed-by: Qu Wenruo <quwenruo@cn.fujitsu.com>
2015-10-23 06:52:54 +00:00
u64 owner, u64 offset);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int btrfs_free_reserved_extent(struct btrfs_fs_info *fs_info,
u64 start, u64 len, int delalloc);
int btrfs_free_and_pin_reserved_extent(struct btrfs_fs_info *fs_info,
u64 start, u64 len);
void btrfs_prepare_extent_commit(struct btrfs_fs_info *fs_info);
int btrfs_finish_extent_commit(struct btrfs_trans_handle *trans);
int btrfs_inc_extent_ref(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
u64 bytenr, u64 num_bytes, u64 parent,
Btrfs: fix regression running delayed references when using qgroups In the kernel 4.2 merge window we had a big changes to the implementation of delayed references and qgroups which made the no_quota field of delayed references not used anymore. More specifically the no_quota field is not used anymore as of: commit 0ed4792af0e8 ("btrfs: qgroup: Switch to new extent-oriented qgroup mechanism.") Leaving the no_quota field actually prevents delayed references from getting merged, which in turn cause the following BUG_ON(), at fs/btrfs/extent-tree.c, to be hit when qgroups are enabled: static int run_delayed_tree_ref(...) { (...) BUG_ON(node->ref_mod != 1); (...) } This happens on a scenario like the following: 1) Ref1 bytenr X, action = BTRFS_ADD_DELAYED_REF, no_quota = 1, added. 2) Ref2 bytenr X, action = BTRFS_DROP_DELAYED_REF, no_quota = 0, added. It's not merged with Ref1 because Ref1->no_quota != Ref2->no_quota. 3) Ref3 bytenr X, action = BTRFS_ADD_DELAYED_REF, no_quota = 1, added. It's not merged with the reference at the tail of the list of refs for bytenr X because the reference at the tail, Ref2 is incompatible due to Ref2->no_quota != Ref3->no_quota. 4) Ref4 bytenr X, action = BTRFS_DROP_DELAYED_REF, no_quota = 0, added. It's not merged with the reference at the tail of the list of refs for bytenr X because the reference at the tail, Ref3 is incompatible due to Ref3->no_quota != Ref4->no_quota. 5) We run delayed references, trigger merging of delayed references, through __btrfs_run_delayed_refs() -> btrfs_merge_delayed_refs(). 6) Ref1 and Ref3 are merged as Ref1->no_quota = Ref3->no_quota and all other conditions are satisfied too. So Ref1 gets a ref_mod value of 2. 7) Ref2 and Ref4 are merged as Ref2->no_quota = Ref4->no_quota and all other conditions are satisfied too. So Ref2 gets a ref_mod value of 2. 8) Ref1 and Ref2 aren't merged, because they have different values for their no_quota field. 9) Delayed reference Ref1 is picked for running (select_delayed_ref() always prefers references with an action == BTRFS_ADD_DELAYED_REF). So run_delayed_tree_ref() is called for Ref1 which triggers the BUG_ON because Ref1->red_mod != 1 (equals 2). So fix this by removing the no_quota field, as it's not used anymore as of commit 0ed4792af0e8 ("btrfs: qgroup: Switch to new extent-oriented qgroup mechanism."). The use of no_quota was also buggy in at least two places: 1) At delayed-refs.c:btrfs_add_delayed_tree_ref() - we were setting no_quota to 0 instead of 1 when the following condition was true: is_fstree(ref_root) || !fs_info->quota_enabled 2) At extent-tree.c:__btrfs_inc_extent_ref() - we were attempting to reset a node's no_quota when the condition "!is_fstree(root_objectid) || !root->fs_info->quota_enabled" was true but we did it only in an unused local stack variable, that is, we never reset the no_quota value in the node itself. This fixes the remainder of problems several people have been having when running delayed references, mostly while a balance is running in parallel, on a 4.2+ kernel. Very special thanks to Stéphane Lesimple for helping debugging this issue and testing this fix on his multi terabyte filesystem (which took more than one day to balance alone, plus fsck, etc). Also, this fixes deadlock issue when using the clone ioctl with qgroups enabled, as reported by Elias Probst in the mailing list. The deadlock happens because after calling btrfs_insert_empty_item we have our path holding a write lock on a leaf of the fs/subvol tree and then before releasing the path we called check_ref() which did backref walking, when qgroups are enabled, and tried to read lock the same leaf. The trace for this case is the following: INFO: task systemd-nspawn:6095 blocked for more than 120 seconds. (...) Call Trace: [<ffffffff86999201>] schedule+0x74/0x83 [<ffffffff863ef64c>] btrfs_tree_read_lock+0xc0/0xea [<ffffffff86137ed7>] ? wait_woken+0x74/0x74 [<ffffffff8639f0a7>] btrfs_search_old_slot+0x51a/0x810 [<ffffffff863a129b>] btrfs_next_old_leaf+0xdf/0x3ce [<ffffffff86413a00>] ? ulist_add_merge+0x1b/0x127 [<ffffffff86411688>] __resolve_indirect_refs+0x62a/0x667 [<ffffffff863ef546>] ? btrfs_clear_lock_blocking_rw+0x78/0xbe [<ffffffff864122d3>] find_parent_nodes+0xaf3/0xfc6 [<ffffffff86412838>] __btrfs_find_all_roots+0x92/0xf0 [<ffffffff864128f2>] btrfs_find_all_roots+0x45/0x65 [<ffffffff8639a75b>] ? btrfs_get_tree_mod_seq+0x2b/0x88 [<ffffffff863e852e>] check_ref+0x64/0xc4 [<ffffffff863e9e01>] btrfs_clone+0x66e/0xb5d [<ffffffff863ea77f>] btrfs_ioctl_clone+0x48f/0x5bb [<ffffffff86048a68>] ? native_sched_clock+0x28/0x77 [<ffffffff863ed9b0>] btrfs_ioctl+0xabc/0x25cb (...) The problem goes away by eleminating check_ref(), which no longer is needed as its purpose was to get a value for the no_quota field of a delayed reference (this patch removes the no_quota field as mentioned earlier). Reported-by: Stéphane Lesimple <stephane_btrfs@lesimple.fr> Tested-by: Stéphane Lesimple <stephane_btrfs@lesimple.fr> Reported-by: Elias Probst <mail@eliasprobst.eu> Reported-by: Peter Becker <floyd.net@gmail.com> Reported-by: Malte Schröder <malte@tnxip.de> Reported-by: Derek Dongray <derek@valedon.co.uk> Reported-by: Erkki Seppala <flux-btrfs@inside.org> Cc: stable@vger.kernel.org # 4.2+ Signed-off-by: Filipe Manana <fdmanana@suse.com> Reviewed-by: Qu Wenruo <quwenruo@cn.fujitsu.com>
2015-10-23 06:52:54 +00:00
u64 root_objectid, u64 owner, u64 offset);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int btrfs_start_dirty_block_groups(struct btrfs_trans_handle *trans);
int btrfs_write_dirty_block_groups(struct btrfs_trans_handle *trans,
struct btrfs_fs_info *fs_info);
int btrfs_setup_space_cache(struct btrfs_trans_handle *trans,
struct btrfs_fs_info *fs_info);
int btrfs_extent_readonly(struct btrfs_fs_info *fs_info, u64 bytenr);
int btrfs_free_block_groups(struct btrfs_fs_info *info);
int btrfs_read_block_groups(struct btrfs_fs_info *info);
int btrfs_can_relocate(struct btrfs_fs_info *fs_info, u64 bytenr);
int btrfs_make_block_group(struct btrfs_trans_handle *trans,
u64 bytes_used, u64 type, u64 chunk_offset,
u64 size);
void btrfs_add_raid_kobjects(struct btrfs_fs_info *fs_info);
struct btrfs_trans_handle *btrfs_start_trans_remove_block_group(
Btrfs: fix the number of transaction units needed to remove a block group We were using only 1 transaction unit when attempting to delete an unused block group but in reality we need 3 + N units, where N corresponds to the number of stripes. We were accounting only for the addition of the orphan item (for the block group's free space cache inode) but we were not accounting that we need to delete one block group item from the extent tree, one free space item from the tree of tree roots and N device extent items from the device tree. While one unit is not enough, it worked most of the time because for each single unit we are too pessimistic and assume an entire tree path, with the highest possible heigth (8), needs to be COWed with eventual node splits at every possible level in the tree, so there was usually enough reserved space for removing all the items and adding the orphan item. However after adding the orphan item, writepages() can by called by the VM subsystem against the btree inode when we are under memory pressure, which causes writeback to start for the nodes we COWed before, this forces the operation to remove the free space item to COW again some (or all of) the same nodes (in the tree of tree roots). Even without writepages() being called, we could fail with ENOSPC because these items are located in multiple trees and one of them might have a higher heigth and require node/leaf splits at many levels, exhausting all the reserved space before removing all the items and adding the orphan. In the kernel 4.0 release, commit 3d84be799194 ("Btrfs: fix BUG_ON in btrfs_orphan_add() when delete unused block group"), we attempted to fix a BUG_ON due to ENOSPC when trying to add the orphan item by making the cleaner kthread reserve one transaction unit before attempting to remove the block group, but this was not enough. We had a couple user reports still hitting the same BUG_ON after 4.0, like Stefan Priebe's report on a 4.2-rc6 kernel for example: http://www.spinics.net/lists/linux-btrfs/msg46070.html So fix this by reserving all the necessary units of metadata. Reported-by: Stefan Priebe <s.priebe@profihost.ag> Fixes: 3d84be799194 ("Btrfs: fix BUG_ON in btrfs_orphan_add() when delete unused block group") Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2015-11-13 23:57:17 +00:00
struct btrfs_fs_info *fs_info,
const u64 chunk_offset);
2008-09-26 14:09:34 +00:00
int btrfs_remove_block_group(struct btrfs_trans_handle *trans,
u64 group_start, struct extent_map *em);
void btrfs_delete_unused_bgs(struct btrfs_fs_info *fs_info);
void btrfs_get_block_group_trimming(struct btrfs_block_group_cache *cache);
void btrfs_put_block_group_trimming(struct btrfs_block_group_cache *cache);
void btrfs_create_pending_block_groups(struct btrfs_trans_handle *trans);
u64 btrfs_data_alloc_profile(struct btrfs_fs_info *fs_info);
u64 btrfs_metadata_alloc_profile(struct btrfs_fs_info *fs_info);
u64 btrfs_system_alloc_profile(struct btrfs_fs_info *fs_info);
void btrfs_clear_space_info_full(struct btrfs_fs_info *info);
enum btrfs_reserve_flush_enum {
/* If we are in the transaction, we can't flush anything.*/
BTRFS_RESERVE_NO_FLUSH,
/*
* Flushing delalloc may cause deadlock somewhere, in this
* case, use FLUSH LIMIT
*/
BTRFS_RESERVE_FLUSH_LIMIT,
BTRFS_RESERVE_FLUSH_ALL,
};
enum btrfs_flush_state {
FLUSH_DELAYED_ITEMS_NR = 1,
FLUSH_DELAYED_ITEMS = 2,
FLUSH_DELAYED_REFS_NR = 3,
FLUSH_DELAYED_REFS = 4,
FLUSH_DELALLOC = 5,
FLUSH_DELALLOC_WAIT = 6,
ALLOC_CHUNK = 7,
ALLOC_CHUNK_FORCE = 8,
COMMIT_TRANS = 9,
};
int btrfs_alloc_data_chunk_ondemand(struct btrfs_inode *inode, u64 bytes);
int btrfs_check_data_free_space(struct inode *inode,
struct extent_changeset **reserved, u64 start, u64 len);
btrfs: qgroup: Fix qgroup reserved space underflow by only freeing reserved ranges [BUG] For the following case, btrfs can underflow qgroup reserved space at an error path: (Page size 4K, function name without "btrfs_" prefix) Task A | Task B ---------------------------------------------------------------------- Buffered_write [0, 2K) | |- check_data_free_space() | | |- qgroup_reserve_data() | | Range aligned to page | | range [0, 4K) <<< | | 4K bytes reserved <<< | |- copy pages to page cache | | Buffered_write [2K, 4K) | |- check_data_free_space() | | |- qgroup_reserved_data() | | Range alinged to page | | range [0, 4K) | | Already reserved by A <<< | | 0 bytes reserved <<< | |- delalloc_reserve_metadata() | | And it *FAILED* (Maybe EQUOTA) | |- free_reserved_data_space() |- qgroup_free_data() Range aligned to page range [0, 4K) Freeing 4K (Special thanks to Chandan for the detailed report and analyse) [CAUSE] Above Task B is freeing reserved data range [0, 4K) which is actually reserved by Task A. And at writeback time, page dirty by Task A will go through writeback routine, which will free 4K reserved data space at file extent insert time, causing the qgroup underflow. [FIX] For btrfs_qgroup_free_data(), add @reserved parameter to only free data ranges reserved by previous btrfs_qgroup_reserve_data(). So in above case, Task B will try to free 0 byte, so no underflow. Reported-by: Chandan Rajendra <chandan@linux.vnet.ibm.com> Signed-off-by: Qu Wenruo <quwenruo@cn.fujitsu.com> Reviewed-by: Chandan Rajendra <chandan@linux.vnet.ibm.com> Tested-by: Chandan Rajendra <chandan@linux.vnet.ibm.com> Signed-off-by: David Sterba <dsterba@suse.com>
2017-02-27 07:10:39 +00:00
void btrfs_free_reserved_data_space(struct inode *inode,
struct extent_changeset *reserved, u64 start, u64 len);
void btrfs_delalloc_release_space(struct inode *inode,
btrfs: qgroup: Use separate meta reservation type for delalloc Before this patch, btrfs qgroup is mixing per-transcation meta rsv with preallocated meta rsv, making it quite easy to underflow qgroup meta reservation. Since we have the new qgroup meta rsv types, apply it to delalloc reservation. Now for delalloc, most of its reserved space will use META_PREALLOC qgroup rsv type. And for callers reducing outstanding extent like btrfs_finish_ordered_io(), they will convert corresponding META_PREALLOC reservation to META_PERTRANS. This is mainly due to the fact that current qgroup numbers will only be updated in btrfs_commit_transaction(), that's to say if we don't keep such placeholder reservation, we can exceed qgroup limitation. And for callers freeing outstanding extent in error handler, we will just free META_PREALLOC bytes. This behavior makes callers of btrfs_qgroup_release_meta() or btrfs_qgroup_convert_meta() to be aware of which type they are. So in this patch, btrfs_delalloc_release_metadata() and its callers get an extra parameter to info qgroup to do correct meta convert/release. The good news is, even we use the wrong type (convert or free), it won't cause obvious bug, as prealloc type is always in good shape, and the type only affects how per-trans meta is increased or not. So the worst case will be at most metadata limitation can be sometimes exceeded (no convert at all) or metadata limitation is reached too soon (no free at all). Signed-off-by: Qu Wenruo <wqu@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2017-12-12 07:34:32 +00:00
struct extent_changeset *reserved,
u64 start, u64 len, bool qgroup_free);
void btrfs_free_reserved_data_space_noquota(struct inode *inode, u64 start,
u64 len);
Btrfs: fix -ENOSPC when finishing block group creation While creating a block group, we often end up getting ENOSPC while updating the chunk tree, which leads to a transaction abortion that produces a trace like the following: [30670.116368] WARNING: CPU: 4 PID: 20735 at fs/btrfs/super.c:260 __btrfs_abort_transaction+0x52/0x106 [btrfs]() [30670.117777] BTRFS: Transaction aborted (error -28) (...) [30670.163567] Call Trace: [30670.163906] [<ffffffff8142fa46>] dump_stack+0x4f/0x7b [30670.164522] [<ffffffff8108b6a2>] ? console_unlock+0x361/0x3ad [30670.165171] [<ffffffff81045ea5>] warn_slowpath_common+0xa1/0xbb [30670.166323] [<ffffffffa035daa7>] ? __btrfs_abort_transaction+0x52/0x106 [btrfs] [30670.167213] [<ffffffff81045f05>] warn_slowpath_fmt+0x46/0x48 [30670.167862] [<ffffffffa035daa7>] __btrfs_abort_transaction+0x52/0x106 [btrfs] [30670.169116] [<ffffffffa03743d7>] btrfs_create_pending_block_groups+0x101/0x130 [btrfs] [30670.170593] [<ffffffffa038426a>] __btrfs_end_transaction+0x84/0x366 [btrfs] [30670.171960] [<ffffffffa038455c>] btrfs_end_transaction+0x10/0x12 [btrfs] [30670.174649] [<ffffffffa036eb6b>] btrfs_check_data_free_space+0x11f/0x27c [btrfs] [30670.176092] [<ffffffffa039450d>] btrfs_fallocate+0x7c8/0xb96 [btrfs] [30670.177218] [<ffffffff812459f2>] ? __this_cpu_preempt_check+0x13/0x15 [30670.178622] [<ffffffff81152447>] vfs_fallocate+0x14c/0x1de [30670.179642] [<ffffffff8116b915>] ? __fget_light+0x2d/0x4f [30670.180692] [<ffffffff81152863>] SyS_fallocate+0x47/0x62 [30670.186737] [<ffffffff81435b32>] system_call_fastpath+0x12/0x17 [30670.187792] ---[ end trace 0373e6b491c4a8cc ]--- This is because we don't do proper space reservation for the chunk block reserve when we have multiple tasks allocating chunks in parallel. So block group creation has 2 phases, and the first phase essentially checks if there is enough space in the system space_info, allocating a new system chunk if there isn't, while the second phase updates the device, extent and chunk trees. However, because the updates to the chunk tree happen in the second phase, if we have N tasks, each with its own transaction handle, allocating new chunks in parallel and if there is only enough space in the system space_info to allocate M chunks, where M < N, none of the tasks ends up allocating a new system chunk in the first phase and N - M tasks will get -ENOSPC when attempting to update the chunk tree in phase 2 if they need to COW any nodes/leafs from the chunk tree. Fix this by doing proper reservation in the chunk block reserve. The issue could be reproduced by running fstests generic/038 in a loop, which eventually triggered the problem. Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: Chris Mason <clm@fb.com>
2015-05-20 13:01:54 +00:00
void btrfs_trans_release_chunk_metadata(struct btrfs_trans_handle *trans);
int btrfs_subvolume_reserve_metadata(struct btrfs_root *root,
struct btrfs_block_rsv *rsv,
int nitems, bool use_global_rsv);
void btrfs_subvolume_release_metadata(struct btrfs_fs_info *fs_info,
struct btrfs_block_rsv *rsv);
btrfs: qgroup: Use separate meta reservation type for delalloc Before this patch, btrfs qgroup is mixing per-transcation meta rsv with preallocated meta rsv, making it quite easy to underflow qgroup meta reservation. Since we have the new qgroup meta rsv types, apply it to delalloc reservation. Now for delalloc, most of its reserved space will use META_PREALLOC qgroup rsv type. And for callers reducing outstanding extent like btrfs_finish_ordered_io(), they will convert corresponding META_PREALLOC reservation to META_PERTRANS. This is mainly due to the fact that current qgroup numbers will only be updated in btrfs_commit_transaction(), that's to say if we don't keep such placeholder reservation, we can exceed qgroup limitation. And for callers freeing outstanding extent in error handler, we will just free META_PREALLOC bytes. This behavior makes callers of btrfs_qgroup_release_meta() or btrfs_qgroup_convert_meta() to be aware of which type they are. So in this patch, btrfs_delalloc_release_metadata() and its callers get an extra parameter to info qgroup to do correct meta convert/release. The good news is, even we use the wrong type (convert or free), it won't cause obvious bug, as prealloc type is always in good shape, and the type only affects how per-trans meta is increased or not. So the worst case will be at most metadata limitation can be sometimes exceeded (no convert at all) or metadata limitation is reached too soon (no free at all). Signed-off-by: Qu Wenruo <wqu@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2017-12-12 07:34:32 +00:00
void btrfs_delalloc_release_extents(struct btrfs_inode *inode, u64 num_bytes,
bool qgroup_free);
Btrfs: rework outstanding_extents Right now we do a lot of weird hoops around outstanding_extents in order to keep the extent count consistent. This is because we logically transfer the outstanding_extent count from the initial reservation through the set_delalloc_bits. This makes it pretty difficult to get a handle on how and when we need to mess with outstanding_extents. Fix this by revamping the rules of how we deal with outstanding_extents. Now instead everybody that is holding on to a delalloc extent is required to increase the outstanding extents count for itself. This means we'll have something like this btrfs_delalloc_reserve_metadata - outstanding_extents = 1 btrfs_set_extent_delalloc - outstanding_extents = 2 btrfs_release_delalloc_extents - outstanding_extents = 1 for an initial file write. Now take the append write where we extend an existing delalloc range but still under the maximum extent size btrfs_delalloc_reserve_metadata - outstanding_extents = 2 btrfs_set_extent_delalloc btrfs_set_bit_hook - outstanding_extents = 3 btrfs_merge_extent_hook - outstanding_extents = 2 btrfs_delalloc_release_extents - outstanding_extnets = 1 In order to make the ordered extent transition we of course must now make ordered extents carry their own outstanding_extent reservation, so for cow_file_range we end up with btrfs_add_ordered_extent - outstanding_extents = 2 clear_extent_bit - outstanding_extents = 1 btrfs_remove_ordered_extent - outstanding_extents = 0 This makes all manipulations of outstanding_extents much more explicit. Every successful call to btrfs_delalloc_reserve_metadata _must_ now be combined with btrfs_release_delalloc_extents, even in the error case, as that is the only function that actually modifies the outstanding_extents counter. The drawback to this is now we are much more likely to have transient cases where outstanding_extents is much larger than it actually should be. This could happen before as we manipulated the delalloc bits, but now it happens basically at every write. This may put more pressure on the ENOSPC flushing code, but I think making this code simpler is worth the cost. I have another change coming to mitigate this side-effect somewhat. I also added trace points for the counter manipulation. These were used by a bpf script I wrote to help track down leak issues. Signed-off-by: Josef Bacik <jbacik@fb.com> Signed-off-by: David Sterba <dsterba@suse.com>
2017-10-19 18:15:55 +00:00
int btrfs_delalloc_reserve_metadata(struct btrfs_inode *inode, u64 num_bytes);
btrfs: qgroup: Use separate meta reservation type for delalloc Before this patch, btrfs qgroup is mixing per-transcation meta rsv with preallocated meta rsv, making it quite easy to underflow qgroup meta reservation. Since we have the new qgroup meta rsv types, apply it to delalloc reservation. Now for delalloc, most of its reserved space will use META_PREALLOC qgroup rsv type. And for callers reducing outstanding extent like btrfs_finish_ordered_io(), they will convert corresponding META_PREALLOC reservation to META_PERTRANS. This is mainly due to the fact that current qgroup numbers will only be updated in btrfs_commit_transaction(), that's to say if we don't keep such placeholder reservation, we can exceed qgroup limitation. And for callers freeing outstanding extent in error handler, we will just free META_PREALLOC bytes. This behavior makes callers of btrfs_qgroup_release_meta() or btrfs_qgroup_convert_meta() to be aware of which type they are. So in this patch, btrfs_delalloc_release_metadata() and its callers get an extra parameter to info qgroup to do correct meta convert/release. The good news is, even we use the wrong type (convert or free), it won't cause obvious bug, as prealloc type is always in good shape, and the type only affects how per-trans meta is increased or not. So the worst case will be at most metadata limitation can be sometimes exceeded (no convert at all) or metadata limitation is reached too soon (no free at all). Signed-off-by: Qu Wenruo <wqu@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2017-12-12 07:34:32 +00:00
void btrfs_delalloc_release_metadata(struct btrfs_inode *inode, u64 num_bytes,
bool qgroup_free);
int btrfs_delalloc_reserve_space(struct inode *inode,
struct extent_changeset **reserved, u64 start, u64 len);
void btrfs_init_block_rsv(struct btrfs_block_rsv *rsv, unsigned short type);
struct btrfs_block_rsv *btrfs_alloc_block_rsv(struct btrfs_fs_info *fs_info,
unsigned short type);
void btrfs_init_metadata_block_rsv(struct btrfs_fs_info *fs_info,
struct btrfs_block_rsv *rsv,
unsigned short type);
void btrfs_free_block_rsv(struct btrfs_fs_info *fs_info,
struct btrfs_block_rsv *rsv);
int btrfs_block_rsv_add(struct btrfs_root *root,
struct btrfs_block_rsv *block_rsv, u64 num_bytes,
enum btrfs_reserve_flush_enum flush);
int btrfs_block_rsv_check(struct btrfs_block_rsv *block_rsv, int min_factor);
int btrfs_block_rsv_refill(struct btrfs_root *root,
struct btrfs_block_rsv *block_rsv, u64 min_reserved,
enum btrfs_reserve_flush_enum flush);
int btrfs_block_rsv_migrate(struct btrfs_block_rsv *src_rsv,
struct btrfs_block_rsv *dst_rsv, u64 num_bytes,
bool update_size);
int btrfs_cond_migrate_bytes(struct btrfs_fs_info *fs_info,
struct btrfs_block_rsv *dest, u64 num_bytes,
int min_factor);
void btrfs_block_rsv_release(struct btrfs_fs_info *fs_info,
struct btrfs_block_rsv *block_rsv,
u64 num_bytes);
btrfs: introduce delayed_refs_rsv Traditionally we've had voodoo in btrfs to account for the space that delayed refs may take up by having a global_block_rsv. This works most of the time, except when it doesn't. We've had issues reported and seen in production where sometimes the global reserve is exhausted during transaction commit before we can run all of our delayed refs, resulting in an aborted transaction. Because of this voodoo we have equally dubious flushing semantics around throttling delayed refs which we often get wrong. So instead give them their own block_rsv. This way we can always know exactly how much outstanding space we need for delayed refs. This allows us to make sure we are constantly filling that reservation up with space, and allows us to put more precise pressure on the enospc system. Instead of doing math to see if its a good time to throttle, the normal enospc code will be invoked if we have a lot of delayed refs pending, and they will be run via the normal flushing mechanism. For now the delayed_refs_rsv will hold the reservations for the delayed refs, the block group updates, and deleting csums. We could have a separate rsv for the block group updates, but the csum deletion stuff is still handled via the delayed_refs so that will stay there. Historical background: The global reserve has grown to cover everything we don't reserve space explicitly for, and we've grown a lot of weird ad-hoc heuristics to know if we're running short on space and when it's time to force a commit. A failure rate of 20-40 file systems when we run hundreds of thousands of them isn't super high, but cleaning up this code will make things less ugly and more predictible. Thus the delayed refs rsv. We always know how many delayed refs we have outstanding, and although running them generates more we can use the global reserve for that spill over, which fits better into it's desired use than a full blown reservation. This first approach is to simply take how many times we're reserving space for and multiply that by 2 in order to save enough space for the delayed refs that could be generated. This is a niave approach and will probably evolve, but for now it works. Signed-off-by: Josef Bacik <jbacik@fb.com> Reviewed-by: David Sterba <dsterba@suse.com> # high-level review [ added background notes from the cover letter ] Signed-off-by: David Sterba <dsterba@suse.com>
2018-12-03 15:20:33 +00:00
void btrfs_delayed_refs_rsv_release(struct btrfs_fs_info *fs_info, int nr);
void btrfs_update_delayed_refs_rsv(struct btrfs_trans_handle *trans);
int btrfs_delayed_refs_rsv_refill(struct btrfs_fs_info *fs_info,
enum btrfs_reserve_flush_enum flush);
void btrfs_migrate_to_delayed_refs_rsv(struct btrfs_fs_info *fs_info,
struct btrfs_block_rsv *src,
u64 num_bytes);
int btrfs_inc_block_group_ro(struct btrfs_block_group_cache *cache);
void btrfs_dec_block_group_ro(struct btrfs_block_group_cache *cache);
void btrfs_put_block_group_cache(struct btrfs_fs_info *info);
btrfs: fix wrong free space information of btrfs When we store data by raid profile in btrfs with two or more different size disks, df command shows there is some free space in the filesystem, but the user can not write any data in fact, df command shows the wrong free space information of btrfs. # mkfs.btrfs -d raid1 /dev/sda9 /dev/sda10 # btrfs-show Label: none uuid: a95cd49e-6e33-45b8-8741-a36153ce4b64 Total devices 2 FS bytes used 28.00KB devid 1 size 5.01GB used 2.03GB path /dev/sda9 devid 2 size 10.00GB used 2.01GB path /dev/sda10 # btrfs device scan /dev/sda9 /dev/sda10 # mount /dev/sda9 /mnt # dd if=/dev/zero of=tmpfile0 bs=4K count=9999999999 (fill the filesystem) # sync # df -TH Filesystem Type Size Used Avail Use% Mounted on /dev/sda9 btrfs 17G 8.6G 5.4G 62% /mnt # btrfs-show Label: none uuid: a95cd49e-6e33-45b8-8741-a36153ce4b64 Total devices 2 FS bytes used 3.99GB devid 1 size 5.01GB used 5.01GB path /dev/sda9 devid 2 size 10.00GB used 4.99GB path /dev/sda10 It is because btrfs cannot allocate chunks when one of the pairing disks has no space, the free space on the other disks can not be used for ever, and should be subtracted from the total space, but btrfs doesn't subtract this space from the total. It is strange to the user. This patch fixes it by calcing the free space that can be used to allocate chunks. Implementation: 1. get all the devices free space, and align them by stripe length. 2. sort the devices by the free space. 3. check the free space of the devices, 3.1. if it is not zero, and then check the number of the devices that has more free space than this device, if the number of the devices is beyond the min stripe number, the free space can be used, and add into total free space. if the number of the devices is below the min stripe number, we can not use the free space, the check ends. 3.2. if the free space is zero, check the next devices, goto 3.1 This implementation is just likely fake chunk allocation. After appling this patch, df can show correct space information: # df -TH Filesystem Type Size Used Avail Use% Mounted on /dev/sda9 btrfs 17G 8.6G 0 100% /mnt Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2011-01-05 10:07:31 +00:00
u64 btrfs_account_ro_block_groups_free_space(struct btrfs_space_info *sinfo);
int btrfs_error_unpin_extent_range(struct btrfs_fs_info *fs_info,
u64 start, u64 end);
int btrfs_discard_extent(struct btrfs_fs_info *fs_info, u64 bytenr,
u64 num_bytes, u64 *actual_bytes);
int btrfs_force_chunk_alloc(struct btrfs_trans_handle *trans, u64 type);
int btrfs_trim_fs(struct btrfs_fs_info *fs_info, struct fstrim_range *range);
int btrfs_init_space_info(struct btrfs_fs_info *fs_info);
int btrfs_delayed_refs_qgroup_accounting(struct btrfs_trans_handle *trans,
struct btrfs_fs_info *fs_info);
int btrfs_start_write_no_snapshotting(struct btrfs_root *root);
void btrfs_end_write_no_snapshotting(struct btrfs_root *root);
void btrfs_wait_for_snapshot_creation(struct btrfs_root *root);
void check_system_chunk(struct btrfs_trans_handle *trans, const u64 type);
u64 add_new_free_space(struct btrfs_block_group_cache *block_group,
u64 start, u64 end);
void btrfs_mark_bg_unused(struct btrfs_block_group_cache *bg);
/* ctree.c */
int btrfs_bin_search(struct extent_buffer *eb, const struct btrfs_key *key,
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int level, int *slot);
int btrfs_comp_cpu_keys(const struct btrfs_key *k1, const struct btrfs_key *k2);
int btrfs_previous_item(struct btrfs_root *root,
struct btrfs_path *path, u64 min_objectid,
int type);
int btrfs_previous_extent_item(struct btrfs_root *root,
struct btrfs_path *path, u64 min_objectid);
void btrfs_set_item_key_safe(struct btrfs_fs_info *fs_info,
struct btrfs_path *path,
const struct btrfs_key *new_key);
struct extent_buffer *btrfs_root_node(struct btrfs_root *root);
struct extent_buffer *btrfs_lock_root_node(struct btrfs_root *root);
struct extent_buffer *btrfs_read_lock_root_node(struct btrfs_root *root);
int btrfs_find_next_key(struct btrfs_root *root, struct btrfs_path *path,
struct btrfs_key *key, int lowest_level,
u64 min_trans);
int btrfs_search_forward(struct btrfs_root *root, struct btrfs_key *min_key,
struct btrfs_path *path,
u64 min_trans);
enum btrfs_compare_tree_result {
BTRFS_COMPARE_TREE_NEW,
BTRFS_COMPARE_TREE_DELETED,
BTRFS_COMPARE_TREE_CHANGED,
Btrfs: fix send issues related to inode number reuse If you are sending a snapshot and specifying a parent snapshot we will walk the trees and figure out where they differ and send the differences only. The way we check for differences are if the leaves aren't the same and if the keys are not the same within the leaves. So if neither leaf is the same (ie the leaf has been cow'ed from the parent snapshot) we walk each item in the send root and check it against the parent root. If the items match exactly then we don't do anything. This doesn't quite work for inode refs, since they will just have the name and the parent objectid. If you move the file from a directory and then remove that directory and re-create a directory with the same inode number as the old directory and then move that file back into that directory we will assume that nothing changed and you will get errors when you try to receive. In order to fix this we need to do extra checking to see if the inode ref really is the same or not. So do this by passing down BTRFS_COMPARE_TREE_SAME if the items match. Then if the key type is an inode ref we can do some extra checking, otherwise we just keep processing. The extra checking is to look up the generation of the directory in the parent volume and compare it to the generation of the send volume. If they match then they are the same directory and we are good to go. If they don't we have to add them to the changed refs list. This means we have to track the generation of the ref we're trying to lookup when we iterate all the refs for a particular inode. So in the case of looking for new refs we have to get the generation from the parent volume, and in the case of looking for deleted refs we have to get the generation from the send volume to compare with. There was also the issue of using a ulist to keep track of the directories we needed to check. Because we can get a deleted ref and a new ref for the same inode number the ulist won't work since it indexes based on the value. So instead just dup any directory ref we find and add it to a local list, and then process that list as normal and do away with using a ulist for this altogether. Before we would fail all of the tests in the far-progs that related to moving directories (test group 32). With this patch we now pass these tests, and all of the tests in the far-progs send testing suite. Thanks, Signed-off-by: Josef Bacik <jbacik@fusionio.com> Signed-off-by: Chris Mason <chris.mason@fusionio.com>
2013-08-16 20:52:55 +00:00
BTRFS_COMPARE_TREE_SAME,
};
typedef int (*btrfs_changed_cb_t)(struct btrfs_path *left_path,
struct btrfs_path *right_path,
struct btrfs_key *key,
enum btrfs_compare_tree_result result,
void *ctx);
int btrfs_compare_trees(struct btrfs_root *left_root,
struct btrfs_root *right_root,
btrfs_changed_cb_t cb, void *ctx);
int btrfs_cow_block(struct btrfs_trans_handle *trans,
struct btrfs_root *root, struct extent_buffer *buf,
struct extent_buffer *parent, int parent_slot,
struct extent_buffer **cow_ret);
int btrfs_copy_root(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct extent_buffer *buf,
struct extent_buffer **cow_ret, u64 new_root_objectid);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int btrfs_block_can_be_shared(struct btrfs_root *root,
struct extent_buffer *buf);
void btrfs_extend_item(struct btrfs_fs_info *fs_info, struct btrfs_path *path,
u32 data_size);
void btrfs_truncate_item(struct btrfs_fs_info *fs_info,
struct btrfs_path *path, u32 new_size, int from_end);
int btrfs_split_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path,
const struct btrfs_key *new_key,
unsigned long split_offset);
int btrfs_duplicate_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path,
const struct btrfs_key *new_key);
int btrfs_find_item(struct btrfs_root *fs_root, struct btrfs_path *path,
u64 inum, u64 ioff, u8 key_type, struct btrfs_key *found_key);
int btrfs_search_slot(struct btrfs_trans_handle *trans, struct btrfs_root *root,
const struct btrfs_key *key, struct btrfs_path *p,
int ins_len, int cow);
int btrfs_search_old_slot(struct btrfs_root *root, const struct btrfs_key *key,
struct btrfs_path *p, u64 time_seq);
int btrfs_search_slot_for_read(struct btrfs_root *root,
const struct btrfs_key *key,
struct btrfs_path *p, int find_higher,
int return_any);
int btrfs_realloc_node(struct btrfs_trans_handle *trans,
struct btrfs_root *root, struct extent_buffer *parent,
int start_slot, u64 *last_ret,
struct btrfs_key *progress);
void btrfs_release_path(struct btrfs_path *p);
struct btrfs_path *btrfs_alloc_path(void);
void btrfs_free_path(struct btrfs_path *p);
Btrfs: Change btree locking to use explicit blocking points Most of the btrfs metadata operations can be protected by a spinlock, but some operations still need to schedule. So far, btrfs has been using a mutex along with a trylock loop, most of the time it is able to avoid going for the full mutex, so the trylock loop is a big performance gain. This commit is step one for getting rid of the blocking locks entirely. btrfs_tree_lock takes a spinlock, and the code explicitly switches to a blocking lock when it starts an operation that can schedule. We'll be able get rid of the blocking locks in smaller pieces over time. Tracing allows us to find the most common cause of blocking, so we can start with the hot spots first. The basic idea is: btrfs_tree_lock() returns with the spin lock held btrfs_set_lock_blocking() sets the EXTENT_BUFFER_BLOCKING bit in the extent buffer flags, and then drops the spin lock. The buffer is still considered locked by all of the btrfs code. If btrfs_tree_lock gets the spinlock but finds the blocking bit set, it drops the spin lock and waits on a wait queue for the blocking bit to go away. Much of the code that needs to set the blocking bit finishes without actually blocking a good percentage of the time. So, an adaptive spin is still used against the blocking bit to avoid very high context switch rates. btrfs_clear_lock_blocking() clears the blocking bit and returns with the spinlock held again. btrfs_tree_unlock() can be called on either blocking or spinning locks, it does the right thing based on the blocking bit. ctree.c has a helper function to set/clear all the locked buffers in a path as blocking. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-02-04 14:25:08 +00:00
void btrfs_set_path_blocking(struct btrfs_path *p);
void btrfs_unlock_up_safe(struct btrfs_path *p, int level);
int btrfs_del_items(struct btrfs_trans_handle *trans, struct btrfs_root *root,
struct btrfs_path *path, int slot, int nr);
static inline int btrfs_del_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path)
{
return btrfs_del_items(trans, root, path, path->slots[0], 1);
}
void setup_items_for_insert(struct btrfs_root *root, struct btrfs_path *path,
const struct btrfs_key *cpu_key, u32 *data_size,
u32 total_data, u32 total_size, int nr);
int btrfs_insert_item(struct btrfs_trans_handle *trans, struct btrfs_root *root,
const struct btrfs_key *key, void *data, u32 data_size);
int btrfs_insert_empty_items(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path,
const struct btrfs_key *cpu_key, u32 *data_size,
int nr);
static inline int btrfs_insert_empty_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path,
const struct btrfs_key *key,
u32 data_size)
{
return btrfs_insert_empty_items(trans, root, path, key, &data_size, 1);
}
int btrfs_next_leaf(struct btrfs_root *root, struct btrfs_path *path);
int btrfs_prev_leaf(struct btrfs_root *root, struct btrfs_path *path);
int btrfs_next_old_leaf(struct btrfs_root *root, struct btrfs_path *path,
u64 time_seq);
static inline int btrfs_next_old_item(struct btrfs_root *root,
struct btrfs_path *p, u64 time_seq)
{
++p->slots[0];
if (p->slots[0] >= btrfs_header_nritems(p->nodes[0]))
return btrfs_next_old_leaf(root, p, time_seq);
return 0;
}
static inline int btrfs_next_item(struct btrfs_root *root, struct btrfs_path *p)
{
return btrfs_next_old_item(root, p, 0);
}
int btrfs_leaf_free_space(struct extent_buffer *leaf);
int __must_check btrfs_drop_snapshot(struct btrfs_root *root,
struct btrfs_block_rsv *block_rsv,
int update_ref, int for_reloc);
int btrfs_drop_subtree(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct extent_buffer *node,
struct extent_buffer *parent);
static inline int btrfs_fs_closing(struct btrfs_fs_info *fs_info)
{
/*
* Do it this way so we only ever do one test_bit in the normal case.
*/
if (test_bit(BTRFS_FS_CLOSING_START, &fs_info->flags)) {
if (test_bit(BTRFS_FS_CLOSING_DONE, &fs_info->flags))
return 2;
return 1;
}
return 0;
}
/*
* If we remount the fs to be R/O or umount the fs, the cleaner needn't do
* anything except sleeping. This function is used to check the status of
* the fs.
*/
static inline int btrfs_need_cleaner_sleep(struct btrfs_fs_info *fs_info)
{
Rename superblock flags (MS_xyz -> SB_xyz) This is a pure automated search-and-replace of the internal kernel superblock flags. The s_flags are now called SB_*, with the names and the values for the moment mirroring the MS_* flags that they're equivalent to. Note how the MS_xyz flags are the ones passed to the mount system call, while the SB_xyz flags are what we then use in sb->s_flags. The script to do this was: # places to look in; re security/*: it generally should *not* be # touched (that stuff parses mount(2) arguments directly), but # there are two places where we really deal with superblock flags. FILES="drivers/mtd drivers/staging/lustre fs ipc mm \ include/linux/fs.h include/uapi/linux/bfs_fs.h \ security/apparmor/apparmorfs.c security/apparmor/include/lib.h" # the list of MS_... constants SYMS="RDONLY NOSUID NODEV NOEXEC SYNCHRONOUS REMOUNT MANDLOCK \ DIRSYNC NOATIME NODIRATIME BIND MOVE REC VERBOSE SILENT \ POSIXACL UNBINDABLE PRIVATE SLAVE SHARED RELATIME KERNMOUNT \ I_VERSION STRICTATIME LAZYTIME SUBMOUNT NOREMOTELOCK NOSEC BORN \ ACTIVE NOUSER" SED_PROG= for i in $SYMS; do SED_PROG="$SED_PROG -e s/MS_$i/SB_$i/g"; done # we want files that contain at least one of MS_..., # with fs/namespace.c and fs/pnode.c excluded. L=$(for i in $SYMS; do git grep -w -l MS_$i $FILES; done| sort|uniq|grep -v '^fs/namespace.c'|grep -v '^fs/pnode.c') for f in $L; do sed -i $f $SED_PROG; done Requested-by: Al Viro <viro@zeniv.linux.org.uk> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-11-27 21:05:09 +00:00
return fs_info->sb->s_flags & SB_RDONLY || btrfs_fs_closing(fs_info);
}
static inline void free_fs_info(struct btrfs_fs_info *fs_info)
{
kfree(fs_info->balance_ctl);
kfree(fs_info->delayed_root);
kfree(fs_info->extent_root);
kfree(fs_info->tree_root);
kfree(fs_info->chunk_root);
kfree(fs_info->dev_root);
kfree(fs_info->csum_root);
kfree(fs_info->quota_root);
kfree(fs_info->uuid_root);
kfree(fs_info->free_space_root);
kfree(fs_info->super_copy);
kfree(fs_info->super_for_commit);
kvfree(fs_info);
}
/* tree mod log functions from ctree.c */
u64 btrfs_get_tree_mod_seq(struct btrfs_fs_info *fs_info,
struct seq_list *elem);
void btrfs_put_tree_mod_seq(struct btrfs_fs_info *fs_info,
struct seq_list *elem);
int btrfs_old_root_level(struct btrfs_root *root, u64 time_seq);
/* root-item.c */
int btrfs_add_root_ref(struct btrfs_trans_handle *trans, u64 root_id,
u64 ref_id, u64 dirid, u64 sequence, const char *name,
int name_len);
int btrfs_del_root_ref(struct btrfs_trans_handle *trans, u64 root_id,
u64 ref_id, u64 dirid, u64 *sequence, const char *name,
int name_len);
int btrfs_del_root(struct btrfs_trans_handle *trans,
const struct btrfs_key *key);
int btrfs_insert_root(struct btrfs_trans_handle *trans, struct btrfs_root *root,
const struct btrfs_key *key,
struct btrfs_root_item *item);
int __must_check btrfs_update_root(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_key *key,
struct btrfs_root_item *item);
int btrfs_find_root(struct btrfs_root *root, const struct btrfs_key *search_key,
struct btrfs_path *path, struct btrfs_root_item *root_item,
struct btrfs_key *root_key);
int btrfs_find_orphan_roots(struct btrfs_fs_info *fs_info);
void btrfs_set_root_node(struct btrfs_root_item *item,
struct extent_buffer *node);
void btrfs_check_and_init_root_item(struct btrfs_root_item *item);
void btrfs_update_root_times(struct btrfs_trans_handle *trans,
struct btrfs_root *root);
Btrfs: introduce a tree for items that map UUIDs to something Mapping UUIDs to subvolume IDs is an operation with a high effort today. Today, the algorithm even has quadratic effort (based on the number of existing subvolumes), which means, that it takes minutes to send/receive a single subvolume if 10,000 subvolumes exist. But even linear effort would be too much since it is a waste. And these data structures to allow mapping UUIDs to subvolume IDs are created every time a btrfs send/receive instance is started. It is much more efficient to maintain a searchable persistent data structure in the filesystem, one that is updated whenever a subvolume/snapshot is created and deleted, and when the received subvolume UUID is set by the btrfs-receive tool. Therefore kernel code is added with this commit that is able to maintain data structures in the filesystem that allow to quickly search for a given UUID and to retrieve data that is assigned to this UUID, like which subvolume ID is related to this UUID. This commit adds a new tree to hold UUID-to-data mapping items. The key of the items is the full UUID plus the key type BTRFS_UUID_KEY. Multiple data blocks can be stored for a given UUID, a type/length/ value scheme is used. Now follows the lengthy justification, why a new tree was added instead of using the existing root tree: The first approach was to not create another tree that holds UUID items. Instead, the items should just go into the top root tree. Unfortunately this confused the algorithm to assign the objectid of subvolumes and snapshots. The reason is that btrfs_find_free_objectid() calls btrfs_find_highest_objectid() for the first created subvol or snapshot after mounting a filesystem, and this function simply searches for the largest used objectid in the root tree keys to pick the next objectid to assign. Of course, the UUID keys have always been the ones with the highest offset value, and the next assigned subvol ID was wastefully huge. To use any other existing tree did not look proper. To apply a workaround such as setting the objectid to zero in the UUID item key and to implement collision handling would either add limitations (in case of a btrfs_extend_item() approach to handle the collisions) or a lot of complexity and source code (in case a key would be looked up that is free of collisions). Adding new code that introduces limitations is not good, and adding code that is complex and lengthy for no good reason is also not good. That's the justification why a completely new tree was introduced. Signed-off-by: Stefan Behrens <sbehrens@giantdisaster.de> Signed-off-by: Josef Bacik <jbacik@fusionio.com> Signed-off-by: Chris Mason <chris.mason@fusionio.com>
2013-08-15 15:11:17 +00:00
/* uuid-tree.c */
int btrfs_uuid_tree_add(struct btrfs_trans_handle *trans, u8 *uuid, u8 type,
Btrfs: introduce a tree for items that map UUIDs to something Mapping UUIDs to subvolume IDs is an operation with a high effort today. Today, the algorithm even has quadratic effort (based on the number of existing subvolumes), which means, that it takes minutes to send/receive a single subvolume if 10,000 subvolumes exist. But even linear effort would be too much since it is a waste. And these data structures to allow mapping UUIDs to subvolume IDs are created every time a btrfs send/receive instance is started. It is much more efficient to maintain a searchable persistent data structure in the filesystem, one that is updated whenever a subvolume/snapshot is created and deleted, and when the received subvolume UUID is set by the btrfs-receive tool. Therefore kernel code is added with this commit that is able to maintain data structures in the filesystem that allow to quickly search for a given UUID and to retrieve data that is assigned to this UUID, like which subvolume ID is related to this UUID. This commit adds a new tree to hold UUID-to-data mapping items. The key of the items is the full UUID plus the key type BTRFS_UUID_KEY. Multiple data blocks can be stored for a given UUID, a type/length/ value scheme is used. Now follows the lengthy justification, why a new tree was added instead of using the existing root tree: The first approach was to not create another tree that holds UUID items. Instead, the items should just go into the top root tree. Unfortunately this confused the algorithm to assign the objectid of subvolumes and snapshots. The reason is that btrfs_find_free_objectid() calls btrfs_find_highest_objectid() for the first created subvol or snapshot after mounting a filesystem, and this function simply searches for the largest used objectid in the root tree keys to pick the next objectid to assign. Of course, the UUID keys have always been the ones with the highest offset value, and the next assigned subvol ID was wastefully huge. To use any other existing tree did not look proper. To apply a workaround such as setting the objectid to zero in the UUID item key and to implement collision handling would either add limitations (in case of a btrfs_extend_item() approach to handle the collisions) or a lot of complexity and source code (in case a key would be looked up that is free of collisions). Adding new code that introduces limitations is not good, and adding code that is complex and lengthy for no good reason is also not good. That's the justification why a completely new tree was introduced. Signed-off-by: Stefan Behrens <sbehrens@giantdisaster.de> Signed-off-by: Josef Bacik <jbacik@fusionio.com> Signed-off-by: Chris Mason <chris.mason@fusionio.com>
2013-08-15 15:11:17 +00:00
u64 subid);
int btrfs_uuid_tree_remove(struct btrfs_trans_handle *trans, u8 *uuid, u8 type,
Btrfs: introduce a tree for items that map UUIDs to something Mapping UUIDs to subvolume IDs is an operation with a high effort today. Today, the algorithm even has quadratic effort (based on the number of existing subvolumes), which means, that it takes minutes to send/receive a single subvolume if 10,000 subvolumes exist. But even linear effort would be too much since it is a waste. And these data structures to allow mapping UUIDs to subvolume IDs are created every time a btrfs send/receive instance is started. It is much more efficient to maintain a searchable persistent data structure in the filesystem, one that is updated whenever a subvolume/snapshot is created and deleted, and when the received subvolume UUID is set by the btrfs-receive tool. Therefore kernel code is added with this commit that is able to maintain data structures in the filesystem that allow to quickly search for a given UUID and to retrieve data that is assigned to this UUID, like which subvolume ID is related to this UUID. This commit adds a new tree to hold UUID-to-data mapping items. The key of the items is the full UUID plus the key type BTRFS_UUID_KEY. Multiple data blocks can be stored for a given UUID, a type/length/ value scheme is used. Now follows the lengthy justification, why a new tree was added instead of using the existing root tree: The first approach was to not create another tree that holds UUID items. Instead, the items should just go into the top root tree. Unfortunately this confused the algorithm to assign the objectid of subvolumes and snapshots. The reason is that btrfs_find_free_objectid() calls btrfs_find_highest_objectid() for the first created subvol or snapshot after mounting a filesystem, and this function simply searches for the largest used objectid in the root tree keys to pick the next objectid to assign. Of course, the UUID keys have always been the ones with the highest offset value, and the next assigned subvol ID was wastefully huge. To use any other existing tree did not look proper. To apply a workaround such as setting the objectid to zero in the UUID item key and to implement collision handling would either add limitations (in case of a btrfs_extend_item() approach to handle the collisions) or a lot of complexity and source code (in case a key would be looked up that is free of collisions). Adding new code that introduces limitations is not good, and adding code that is complex and lengthy for no good reason is also not good. That's the justification why a completely new tree was introduced. Signed-off-by: Stefan Behrens <sbehrens@giantdisaster.de> Signed-off-by: Josef Bacik <jbacik@fusionio.com> Signed-off-by: Chris Mason <chris.mason@fusionio.com>
2013-08-15 15:11:17 +00:00
u64 subid);
int btrfs_uuid_tree_iterate(struct btrfs_fs_info *fs_info,
int (*check_func)(struct btrfs_fs_info *, u8 *, u8,
u64));
Btrfs: introduce a tree for items that map UUIDs to something Mapping UUIDs to subvolume IDs is an operation with a high effort today. Today, the algorithm even has quadratic effort (based on the number of existing subvolumes), which means, that it takes minutes to send/receive a single subvolume if 10,000 subvolumes exist. But even linear effort would be too much since it is a waste. And these data structures to allow mapping UUIDs to subvolume IDs are created every time a btrfs send/receive instance is started. It is much more efficient to maintain a searchable persistent data structure in the filesystem, one that is updated whenever a subvolume/snapshot is created and deleted, and when the received subvolume UUID is set by the btrfs-receive tool. Therefore kernel code is added with this commit that is able to maintain data structures in the filesystem that allow to quickly search for a given UUID and to retrieve data that is assigned to this UUID, like which subvolume ID is related to this UUID. This commit adds a new tree to hold UUID-to-data mapping items. The key of the items is the full UUID plus the key type BTRFS_UUID_KEY. Multiple data blocks can be stored for a given UUID, a type/length/ value scheme is used. Now follows the lengthy justification, why a new tree was added instead of using the existing root tree: The first approach was to not create another tree that holds UUID items. Instead, the items should just go into the top root tree. Unfortunately this confused the algorithm to assign the objectid of subvolumes and snapshots. The reason is that btrfs_find_free_objectid() calls btrfs_find_highest_objectid() for the first created subvol or snapshot after mounting a filesystem, and this function simply searches for the largest used objectid in the root tree keys to pick the next objectid to assign. Of course, the UUID keys have always been the ones with the highest offset value, and the next assigned subvol ID was wastefully huge. To use any other existing tree did not look proper. To apply a workaround such as setting the objectid to zero in the UUID item key and to implement collision handling would either add limitations (in case of a btrfs_extend_item() approach to handle the collisions) or a lot of complexity and source code (in case a key would be looked up that is free of collisions). Adding new code that introduces limitations is not good, and adding code that is complex and lengthy for no good reason is also not good. That's the justification why a completely new tree was introduced. Signed-off-by: Stefan Behrens <sbehrens@giantdisaster.de> Signed-off-by: Josef Bacik <jbacik@fusionio.com> Signed-off-by: Chris Mason <chris.mason@fusionio.com>
2013-08-15 15:11:17 +00:00
/* dir-item.c */
int btrfs_check_dir_item_collision(struct btrfs_root *root, u64 dir,
const char *name, int name_len);
int btrfs_insert_dir_item(struct btrfs_trans_handle *trans, const char *name,
int name_len, struct btrfs_inode *dir,
struct btrfs_key *location, u8 type, u64 index);
struct btrfs_dir_item *btrfs_lookup_dir_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path, u64 dir,
const char *name, int name_len,
int mod);
struct btrfs_dir_item *
btrfs_lookup_dir_index_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path, u64 dir,
u64 objectid, const char *name, int name_len,
int mod);
struct btrfs_dir_item *
btrfs_search_dir_index_item(struct btrfs_root *root,
struct btrfs_path *path, u64 dirid,
const char *name, int name_len);
int btrfs_delete_one_dir_name(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path,
struct btrfs_dir_item *di);
int btrfs_insert_xattr_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path, u64 objectid,
const char *name, u16 name_len,
const void *data, u16 data_len);
struct btrfs_dir_item *btrfs_lookup_xattr(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path, u64 dir,
const char *name, u16 name_len,
int mod);
struct btrfs_dir_item *btrfs_match_dir_item_name(struct btrfs_fs_info *fs_info,
struct btrfs_path *path,
const char *name,
int name_len);
/* orphan.c */
int btrfs_insert_orphan_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root, u64 offset);
int btrfs_del_orphan_item(struct btrfs_trans_handle *trans,
struct btrfs_root *root, u64 offset);
int btrfs_find_orphan_item(struct btrfs_root *root, u64 offset);
/* inode-item.c */
int btrfs_insert_inode_ref(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
const char *name, int name_len,
u64 inode_objectid, u64 ref_objectid, u64 index);
int btrfs_del_inode_ref(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
const char *name, int name_len,
u64 inode_objectid, u64 ref_objectid, u64 *index);
int btrfs_insert_empty_inode(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path, u64 objectid);
int btrfs_lookup_inode(struct btrfs_trans_handle *trans, struct btrfs_root
*root, struct btrfs_path *path,
struct btrfs_key *location, int mod);
struct btrfs_inode_extref *
btrfs_lookup_inode_extref(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path,
const char *name, int name_len,
u64 inode_objectid, u64 ref_objectid, int ins_len,
int cow);
Btrfs: fix log replay failure after unlink and link combination If we have a file with 2 (or more) hard links in the same directory, remove one of the hard links, create a new file (or link an existing file) in the same directory with the name of the removed hard link, and then finally fsync the new file, we end up with a log that fails to replay, causing a mount failure. Example: $ mkfs.btrfs -f /dev/sdb $ mount /dev/sdb /mnt $ mkdir /mnt/testdir $ touch /mnt/testdir/foo $ ln /mnt/testdir/foo /mnt/testdir/bar $ sync $ unlink /mnt/testdir/bar $ touch /mnt/testdir/bar $ xfs_io -c "fsync" /mnt/testdir/bar <power failure> $ mount /dev/sdb /mnt mount: mount(2) failed: /mnt: No such file or directory When replaying the log, for that example, we also see the following in dmesg/syslog: [71813.671307] BTRFS info (device dm-0): failed to delete reference to bar, inode 258 parent 257 [71813.674204] ------------[ cut here ]------------ [71813.675694] BTRFS: Transaction aborted (error -2) [71813.677236] WARNING: CPU: 1 PID: 13231 at fs/btrfs/inode.c:4128 __btrfs_unlink_inode+0x17b/0x355 [btrfs] [71813.679669] Modules linked in: btrfs xfs f2fs dm_flakey dm_mod dax ghash_clmulni_intel ppdev pcbc aesni_intel aes_x86_64 crypto_simd cryptd glue_helper evdev psmouse i2c_piix4 parport_pc i2c_core pcspkr sg serio_raw parport button sunrpc loop autofs4 ext4 crc16 mbcache jbd2 zstd_decompress zstd_compress xxhash raid10 raid456 async_raid6_recov async_memcpy async_pq async_xor async_tx xor raid6_pq libcrc32c crc32c_generic raid1 raid0 multipath linear md_mod ata_generic sd_mod virtio_scsi ata_piix libata virtio_pci virtio_ring crc32c_intel floppy virtio e1000 scsi_mod [last unloaded: btrfs] [71813.679669] CPU: 1 PID: 13231 Comm: mount Tainted: G W 4.15.0-rc9-btrfs-next-56+ #1 [71813.679669] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS rel-1.10.2-0-g5f4c7b1-prebuilt.qemu-project.org 04/01/2014 [71813.679669] RIP: 0010:__btrfs_unlink_inode+0x17b/0x355 [btrfs] [71813.679669] RSP: 0018:ffffc90001cef738 EFLAGS: 00010286 [71813.679669] RAX: 0000000000000025 RBX: ffff880217ce4708 RCX: 0000000000000001 [71813.679669] RDX: 0000000000000000 RSI: ffffffff81c14bae RDI: 00000000ffffffff [71813.679669] RBP: ffffc90001cef7c0 R08: 0000000000000001 R09: 0000000000000001 [71813.679669] R10: ffffc90001cef5e0 R11: ffffffff8343f007 R12: ffff880217d474c8 [71813.679669] R13: 00000000fffffffe R14: ffff88021ccf1548 R15: 0000000000000101 [71813.679669] FS: 00007f7cee84c480(0000) GS:ffff88023fc80000(0000) knlGS:0000000000000000 [71813.679669] CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033 [71813.679669] CR2: 00007f7cedc1abf9 CR3: 00000002354b4003 CR4: 00000000001606e0 [71813.679669] Call Trace: [71813.679669] btrfs_unlink_inode+0x17/0x41 [btrfs] [71813.679669] drop_one_dir_item+0xfa/0x131 [btrfs] [71813.679669] add_inode_ref+0x71e/0x851 [btrfs] [71813.679669] ? __lock_is_held+0x39/0x71 [71813.679669] ? replay_one_buffer+0x53/0x53a [btrfs] [71813.679669] replay_one_buffer+0x4a4/0x53a [btrfs] [71813.679669] ? rcu_read_unlock+0x3a/0x57 [71813.679669] ? __lock_is_held+0x39/0x71 [71813.679669] walk_up_log_tree+0x101/0x1d2 [btrfs] [71813.679669] walk_log_tree+0xad/0x188 [btrfs] [71813.679669] btrfs_recover_log_trees+0x1fa/0x31e [btrfs] [71813.679669] ? replay_one_extent+0x544/0x544 [btrfs] [71813.679669] open_ctree+0x1cf6/0x2209 [btrfs] [71813.679669] btrfs_mount_root+0x368/0x482 [btrfs] [71813.679669] ? trace_hardirqs_on_caller+0x14c/0x1a6 [71813.679669] ? __lockdep_init_map+0x176/0x1c2 [71813.679669] ? mount_fs+0x64/0x10b [71813.679669] mount_fs+0x64/0x10b [71813.679669] vfs_kern_mount+0x68/0xce [71813.679669] btrfs_mount+0x13e/0x772 [btrfs] [71813.679669] ? trace_hardirqs_on_caller+0x14c/0x1a6 [71813.679669] ? __lockdep_init_map+0x176/0x1c2 [71813.679669] ? mount_fs+0x64/0x10b [71813.679669] mount_fs+0x64/0x10b [71813.679669] vfs_kern_mount+0x68/0xce [71813.679669] do_mount+0x6e5/0x973 [71813.679669] ? memdup_user+0x3e/0x5c [71813.679669] SyS_mount+0x72/0x98 [71813.679669] entry_SYSCALL_64_fastpath+0x1e/0x8b [71813.679669] RIP: 0033:0x7f7cedf150ba [71813.679669] RSP: 002b:00007ffca71da688 EFLAGS: 00000206 [71813.679669] Code: 7f a0 e8 51 0c fd ff 48 8b 43 50 f0 0f ba a8 30 2c 00 00 02 72 17 41 83 fd fb 74 11 44 89 ee 48 c7 c7 7d 11 7f a0 e8 38 f5 8d e0 <0f> ff 44 89 e9 ba 20 10 00 00 eb 4d 48 8b 4d b0 48 8b 75 88 4c [71813.679669] ---[ end trace 83bd473fc5b4663b ]--- [71813.854764] BTRFS: error (device dm-0) in __btrfs_unlink_inode:4128: errno=-2 No such entry [71813.886994] BTRFS: error (device dm-0) in btrfs_replay_log:2307: errno=-2 No such entry (Failed to recover log tree) [71813.903357] BTRFS error (device dm-0): cleaner transaction attach returned -30 [71814.128078] BTRFS error (device dm-0): open_ctree failed This happens because the log has inode reference items for both inode 258 (the first file we created) and inode 259 (the second file created), and when processing the reference item for inode 258, we replace the corresponding item in the subvolume tree (which has two names, "foo" and "bar") witht he one in the log (which only has one name, "foo") without removing the corresponding dir index keys from the parent directory. Later, when processing the inode reference item for inode 259, which has a name of "bar" associated to it, we notice that dir index entries exist for that name and for a different inode, so we attempt to unlink that name, which fails because the inode reference item for inode 258 no longer has the name "bar" associated to it, making a call to btrfs_unlink_inode() fail with a -ENOENT error. Fix this by unlinking all the names in an inode reference item from a subvolume tree that are not present in the inode reference item found in the log tree, before overwriting it with the item from the log tree. Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2018-02-28 15:56:10 +00:00
int btrfs_find_name_in_backref(struct extent_buffer *leaf, int slot,
const char *name,
int name_len, struct btrfs_inode_ref **ref_ret);
int btrfs_find_name_in_ext_backref(struct extent_buffer *leaf, int slot,
u64 ref_objectid, const char *name,
int name_len,
struct btrfs_inode_extref **extref_ret);
/* file-item.c */
struct btrfs_dio_private;
int btrfs_del_csums(struct btrfs_trans_handle *trans,
struct btrfs_fs_info *fs_info, u64 bytenr, u64 len);
blk_status_t btrfs_lookup_bio_sums(struct inode *inode, struct bio *bio, u32 *dst);
blk_status_t btrfs_lookup_bio_sums_dio(struct inode *inode, struct bio *bio,
u64 logical_offset);
int btrfs_insert_file_extent(struct btrfs_trans_handle *trans,
Btrfs: Add zlib compression support This is a large change for adding compression on reading and writing, both for inline and regular extents. It does some fairly large surgery to the writeback paths. Compression is off by default and enabled by mount -o compress. Even when the -o compress mount option is not used, it is possible to read compressed extents off the disk. If compression for a given set of pages fails to make them smaller, the file is flagged to avoid future compression attempts later. * While finding delalloc extents, the pages are locked before being sent down to the delalloc handler. This allows the delalloc handler to do complex things such as cleaning the pages, marking them writeback and starting IO on their behalf. * Inline extents are inserted at delalloc time now. This allows us to compress the data before inserting the inline extent, and it allows us to insert an inline extent that spans multiple pages. * All of the in-memory extent representations (extent_map.c, ordered-data.c etc) are changed to record both an in-memory size and an on disk size, as well as a flag for compression. From a disk format point of view, the extent pointers in the file are changed to record the on disk size of a given extent and some encoding flags. Space in the disk format is allocated for compression encoding, as well as encryption and a generic 'other' field. Neither the encryption or the 'other' field are currently used. In order to limit the amount of data read for a single random read in the file, the size of a compressed extent is limited to 128k. This is a software only limit, the disk format supports u64 sized compressed extents. In order to limit the ram consumed while processing extents, the uncompressed size of a compressed extent is limited to 256k. This is a software only limit and will be subject to tuning later. Checksumming is still done on compressed extents, and it is done on the uncompressed version of the data. This way additional encodings can be layered on without having to figure out which encoding to checksum. Compression happens at delalloc time, which is basically singled threaded because it is usually done by a single pdflush thread. This makes it tricky to spread the compression load across all the cpus on the box. We'll have to look at parallel pdflush walks of dirty inodes at a later time. Decompression is hooked into readpages and it does spread across CPUs nicely. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-10-29 18:49:59 +00:00
struct btrfs_root *root,
u64 objectid, u64 pos,
u64 disk_offset, u64 disk_num_bytes,
u64 num_bytes, u64 offset, u64 ram_bytes,
u8 compression, u8 encryption, u16 other_encoding);
int btrfs_lookup_file_extent(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_path *path, u64 objectid,
u64 bytenr, int mod);
int btrfs_csum_file_blocks(struct btrfs_trans_handle *trans,
Btrfs: move data checksumming into a dedicated tree Btrfs stores checksums for each data block. Until now, they have been stored in the subvolume trees, indexed by the inode that is referencing the data block. This means that when we read the inode, we've probably read in at least some checksums as well. But, this has a few problems: * The checksums are indexed by logical offset in the file. When compression is on, this means we have to do the expensive checksumming on the uncompressed data. It would be faster if we could checksum the compressed data instead. * If we implement encryption, we'll be checksumming the plain text and storing that on disk. This is significantly less secure. * For either compression or encryption, we have to get the plain text back before we can verify the checksum as correct. This makes the raid layer balancing and extent moving much more expensive. * It makes the front end caching code more complex, as we have touch the subvolume and inodes as we cache extents. * There is potentitally one copy of the checksum in each subvolume referencing an extent. The solution used here is to store the extent checksums in a dedicated tree. This allows us to index the checksums by phyiscal extent start and length. It means: * The checksum is against the data stored on disk, after any compression or encryption is done. * The checksum is stored in a central location, and can be verified without following back references, or reading inodes. This makes compression significantly faster by reducing the amount of data that needs to be checksummed. It will also allow much faster raid management code in general. The checksums are indexed by a key with a fixed objectid (a magic value in ctree.h) and offset set to the starting byte of the extent. This allows us to copy the checksum items into the fsync log tree directly (or any other tree), without having to invent a second format for them. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-12-08 21:58:54 +00:00
struct btrfs_root *root,
struct btrfs_ordered_sum *sums);
blk_status_t btrfs_csum_one_bio(struct inode *inode, struct bio *bio,
u64 file_start, int contig);
int btrfs_lookup_csums_range(struct btrfs_root *root, u64 start, u64 end,
struct list_head *list, int search_commit);
void btrfs_extent_item_to_extent_map(struct btrfs_inode *inode,
const struct btrfs_path *path,
struct btrfs_file_extent_item *fi,
const bool new_inline,
struct extent_map *em);
/* inode.c */
struct extent_map *btrfs_get_extent_fiemap(struct btrfs_inode *inode,
u64 start, u64 len);
noinline int can_nocow_extent(struct inode *inode, u64 offset, u64 *len,
u64 *orig_start, u64 *orig_block_len,
u64 *ram_bytes);
void __btrfs_del_delalloc_inode(struct btrfs_root *root,
struct btrfs_inode *inode);
struct inode *btrfs_lookup_dentry(struct inode *dir, struct dentry *dentry);
int btrfs_set_inode_index(struct btrfs_inode *dir, u64 *index);
int btrfs_unlink_inode(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct btrfs_inode *dir, struct btrfs_inode *inode,
const char *name, int name_len);
int btrfs_add_link(struct btrfs_trans_handle *trans,
struct btrfs_inode *parent_inode, struct btrfs_inode *inode,
const char *name, int name_len, int add_backref, u64 index);
int btrfs_delete_subvolume(struct inode *dir, struct dentry *dentry);
int btrfs_truncate_block(struct inode *inode, loff_t from, loff_t len,
int front);
int btrfs_truncate_inode_items(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct inode *inode, u64 new_size,
u32 min_type);
btrfs: use tagged writepage to mitigate livelock of snapshot Snapshot is expected to be fast. But if there are writers steadily creating dirty pages in our subvolume, the snapshot may take a very long time to complete. To fix the problem, we use tagged writepage for snapshot flusher as we do in the generic write_cache_pages(), so we can omit pages dirtied after the snapshot command. This does not change the semantics regarding which data get to the snapshot, if there are pages being dirtied during the snapshotting operation. There's a sync called before snapshot is taken in old/new case, any IO in flight just after that may be in the snapshot but this depends on other system effects that might still sync the IO. We do a simple snapshot speed test on a Intel D-1531 box: fio --ioengine=libaio --iodepth=32 --bs=4k --rw=write --size=64G --direct=0 --thread=1 --numjobs=1 --time_based --runtime=120 --filename=/mnt/sub/testfile --name=job1 --group_reporting & sleep 5; time btrfs sub snap -r /mnt/sub /mnt/snap; killall fio original: 1m58sec patched: 6.54sec This is the best case for this patch since for a sequential write case, we omit nearly all pages dirtied after the snapshot command. For a multi writers, random write test: fio --ioengine=libaio --iodepth=32 --bs=4k --rw=randwrite --size=64G --direct=0 --thread=1 --numjobs=4 --time_based --runtime=120 --filename=/mnt/sub/testfile --name=job1 --group_reporting & sleep 5; time btrfs sub snap -r /mnt/sub /mnt/snap; killall fio original: 15.83sec patched: 10.35sec The improvement is smaller compared to the sequential write case, since we omit only half of the pages dirtied after snapshot command. Reviewed-by: Nikolay Borisov <nborisov@suse.com> Signed-off-by: Ethan Lien <ethanlien@synology.com> Reviewed-by: David Sterba <dsterba@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2018-11-01 06:49:03 +00:00
int btrfs_start_delalloc_snapshot(struct btrfs_root *root);
int btrfs_start_delalloc_roots(struct btrfs_fs_info *fs_info, int nr);
int btrfs_set_extent_delalloc(struct inode *inode, u64 start, u64 end,
Btrfs: fix reported number of inode blocks after buffered append writes The patch from commit a7e3b975a0f9 ("Btrfs: fix reported number of inode blocks") introduced a regression where if we do a buffered write starting at position equal to or greater than the file's size and then stat(2) the file before writeback is triggered, the number of used blocks does not change (unless there's a prealloc/unwritten extent). Example: $ xfs_io -f -c "pwrite -S 0xab 0 64K" foobar $ du -h foobar 0 foobar $ sync $ du -h foobar 64K foobar The first version of that patch didn't had this regression and the second version, which was the one committed, was made only to address some performance regression detected by the intel test robots using fs_mark. This fixes the regression by setting the new delaloc bit in the range, and doing it at btrfs_dirty_pages() while setting the regular dealloc bit as well, so that this way we set both bits at once avoiding navigation of the inode's io tree twice. Doing it at btrfs_dirty_pages() is also the most meaninful place, as we should set the new dellaloc bit when if we set the delalloc bit, which happens only if we copied bytes into the pages at __btrfs_buffered_write(). This was making some of LTP's du tests fail, which can be quickly run using a command line like the following: $ ./runltp -q -p -l /ltp.log -f commands -s du -d /mnt Fixes: a7e3b975a0f9 ("Btrfs: fix reported number of inode blocks") Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2017-11-04 00:16:59 +00:00
unsigned int extra_bits,
struct extent_state **cached_state, int dedupe);
int btrfs_create_subvol_root(struct btrfs_trans_handle *trans,
Btrfs: add support for inode properties This change adds infrastructure to allow for generic properties for inodes. Properties are name/value pairs that can be associated with inodes for different purposes. They are stored as xattrs with the prefix "btrfs." Properties can be inherited - this means when a directory inode has inheritable properties set, these are added to new inodes created under that directory. Further, subvolumes can also have properties associated with them, and they can be inherited from their parent subvolume. Naturally, directory properties have priority over subvolume properties (in practice a subvolume property is just a regular property associated with the root inode, objectid 256, of the subvolume's fs tree). This change also adds one specific property implementation, named "compression", whose values can be "lzo" or "zlib" and it's an inheritable property. The corresponding changes to btrfs-progs were also implemented. A patch with xfstests for this feature will follow once there's agreement on this change/feature. Further, the script at the bottom of this commit message was used to do some benchmarks to measure any performance penalties of this feature. Basically the tests correspond to: Test 1 - create a filesystem and mount it with compress-force=lzo, then sequentially create N files of 64Kb each, measure how long it took to create the files, unmount the filesystem, mount the filesystem and perform an 'ls -lha' against the test directory holding the N files, and report the time the command took. Test 2 - create a filesystem and don't use any compression option when mounting it - instead set the compression property of the subvolume's root to 'lzo'. Then create N files of 64Kb, and report the time it took. The unmount the filesystem, mount it again and perform an 'ls -lha' like in the former test. This means every single file ends up with a property (xattr) associated to it. Test 3 - same as test 2, but uses 4 properties - 3 are duplicates of the compression property, have no real effect other than adding more work when inheriting properties and taking more btree leaf space. Test 4 - same as test 3 but with 10 properties per file. Results (in seconds, and averages of 5 runs each), for different N numbers of files follow. * Without properties (test 1) file creation time ls -lha time 10 000 files 3.49 0.76 100 000 files 47.19 8.37 1 000 000 files 518.51 107.06 * With 1 property (compression property set to lzo - test 2) file creation time ls -lha time 10 000 files 3.63 0.93 100 000 files 48.56 9.74 1 000 000 files 537.72 125.11 * With 4 properties (test 3) file creation time ls -lha time 10 000 files 3.94 1.20 100 000 files 52.14 11.48 1 000 000 files 572.70 142.13 * With 10 properties (test 4) file creation time ls -lha time 10 000 files 4.61 1.35 100 000 files 58.86 13.83 1 000 000 files 656.01 177.61 The increased latencies with properties are essencialy because of: *) When creating an inode, we now synchronously write 1 more item (an xattr item) for each property inherited from the parent dir (or subvolume). This could be done in an asynchronous way such as we do for dir intex items (delayed-inode.c), which could help reduce the file creation latency; *) With properties, we now have larger fs trees. For this particular test each xattr item uses 75 bytes of leaf space in the fs tree. This could be less by using a new item for xattr items, instead of the current btrfs_dir_item, since we could cut the 'location' and 'type' fields (saving 18 bytes) and maybe 'transid' too (saving a total of 26 bytes per xattr item) from the btrfs_dir_item type. Also tried batching the xattr insertions (ignoring proper hash collision handling, since it didn't exist) when creating files that inherit properties from their parent inode/subvolume, but the end results were (surprisingly) essentially the same. Test script: $ cat test.pl #!/usr/bin/perl -w use strict; use Time::HiRes qw(time); use constant NUM_FILES => 10_000; use constant FILE_SIZES => (64 * 1024); use constant DEV => '/dev/sdb4'; use constant MNT_POINT => '/home/fdmanana/btrfs-tests/dev'; use constant TEST_DIR => (MNT_POINT . '/testdir'); system("mkfs.btrfs", "-l", "16384", "-f", DEV) == 0 or die "mkfs.btrfs failed!"; # following line for testing without properties #system("mount", "-o", "compress-force=lzo", DEV, MNT_POINT) == 0 or die "mount failed!"; # following 2 lines for testing with properties system("mount", DEV, MNT_POINT) == 0 or die "mount failed!"; system("btrfs", "prop", "set", MNT_POINT, "compression", "lzo") == 0 or die "set prop failed!"; system("mkdir", TEST_DIR) == 0 or die "mkdir failed!"; my ($t1, $t2); $t1 = time(); for (my $i = 1; $i <= NUM_FILES; $i++) { my $p = TEST_DIR . '/file_' . $i; open(my $f, '>', $p) or die "Error opening file!"; $f->autoflush(1); for (my $j = 0; $j < FILE_SIZES; $j += 4096) { print $f ('A' x 4096) or die "Error writing to file!"; } close($f); } $t2 = time(); print "Time to create " . NUM_FILES . ": " . ($t2 - $t1) . " seconds.\n"; system("umount", DEV) == 0 or die "umount failed!"; system("mount", DEV, MNT_POINT) == 0 or die "mount failed!"; $t1 = time(); system("bash -c 'ls -lha " . TEST_DIR . " > /dev/null'") == 0 or die "ls failed!"; $t2 = time(); print "Time to ls -lha all files: " . ($t2 - $t1) . " seconds.\n"; system("umount", DEV) == 0 or die "umount failed!"; Signed-off-by: Filipe David Borba Manana <fdmanana@gmail.com> Signed-off-by: Josef Bacik <jbacik@fb.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-01-07 11:47:46 +00:00
struct btrfs_root *new_root,
struct btrfs_root *parent_root,
u64 new_dirid);
void btrfs_set_delalloc_extent(struct inode *inode, struct extent_state *state,
unsigned *bits);
void btrfs_clear_delalloc_extent(struct inode *inode,
struct extent_state *state, unsigned *bits);
void btrfs_merge_delalloc_extent(struct inode *inode, struct extent_state *new,
struct extent_state *other);
void btrfs_split_delalloc_extent(struct inode *inode,
struct extent_state *orig, u64 split);
int btrfs_bio_fits_in_stripe(struct page *page, size_t size, struct bio *bio,
unsigned long bio_flags);
void btrfs_set_range_writeback(struct extent_io_tree *tree, u64 start, u64 end);
vm_fault_t btrfs_page_mkwrite(struct vm_fault *vmf);
int btrfs_readpage(struct file *file, struct page *page);
void btrfs_evict_inode(struct inode *inode);
int btrfs_write_inode(struct inode *inode, struct writeback_control *wbc);
struct inode *btrfs_alloc_inode(struct super_block *sb);
void btrfs_destroy_inode(struct inode *inode);
int btrfs_drop_inode(struct inode *inode);
int __init btrfs_init_cachep(void);
void __cold btrfs_destroy_cachep(void);
Btrfs: fix deadlock on tree root leaf when finding free extent When we are writing out a free space cache, during the transaction commit phase, we can end up in a deadlock which results in a stack trace like the following: schedule+0x28/0x80 btrfs_tree_read_lock+0x8e/0x120 [btrfs] ? finish_wait+0x80/0x80 btrfs_read_lock_root_node+0x2f/0x40 [btrfs] btrfs_search_slot+0xf6/0x9f0 [btrfs] ? evict_refill_and_join+0xd0/0xd0 [btrfs] ? inode_insert5+0x119/0x190 btrfs_lookup_inode+0x3a/0xc0 [btrfs] ? kmem_cache_alloc+0x166/0x1d0 btrfs_iget+0x113/0x690 [btrfs] __lookup_free_space_inode+0xd8/0x150 [btrfs] lookup_free_space_inode+0x5b/0xb0 [btrfs] load_free_space_cache+0x7c/0x170 [btrfs] ? cache_block_group+0x72/0x3b0 [btrfs] cache_block_group+0x1b3/0x3b0 [btrfs] ? finish_wait+0x80/0x80 find_free_extent+0x799/0x1010 [btrfs] btrfs_reserve_extent+0x9b/0x180 [btrfs] btrfs_alloc_tree_block+0x1b3/0x4f0 [btrfs] __btrfs_cow_block+0x11d/0x500 [btrfs] btrfs_cow_block+0xdc/0x180 [btrfs] btrfs_search_slot+0x3bd/0x9f0 [btrfs] btrfs_lookup_inode+0x3a/0xc0 [btrfs] ? kmem_cache_alloc+0x166/0x1d0 btrfs_update_inode_item+0x46/0x100 [btrfs] cache_save_setup+0xe4/0x3a0 [btrfs] btrfs_start_dirty_block_groups+0x1be/0x480 [btrfs] btrfs_commit_transaction+0xcb/0x8b0 [btrfs] At cache_save_setup() we need to update the inode item of a block group's cache which is located in the tree root (fs_info->tree_root), which means that it may result in COWing a leaf from that tree. If that happens we need to find a free metadata extent and while looking for one, if we find a block group which was not cached yet we attempt to load its cache by calling cache_block_group(). However this function will try to load the inode of the free space cache, which requires finding the matching inode item in the tree root - if that inode item is located in the same leaf as the inode item of the space cache we are updating at cache_save_setup(), we end up in a deadlock, since we try to obtain a read lock on the same extent buffer that we previously write locked. So fix this by using the tree root's commit root when searching for a block group's free space cache inode item when we are attempting to load a free space cache. This is safe since block groups once loaded stay in memory forever, as well as their caches, so after they are first loaded we will never need to read their inode items again. For new block groups, once they are created they get their ->cached field set to BTRFS_CACHE_FINISHED meaning we will not need to read their inode item. Reported-by: Andrew Nelson <andrew.s.nelson@gmail.com> Link: https://lore.kernel.org/linux-btrfs/CAPTELenq9x5KOWuQ+fa7h1r3nsJG8vyiTH8+ifjURc_duHh2Wg@mail.gmail.com/ Fixes: 9d66e233c704 ("Btrfs: load free space cache if it exists") Tested-by: Andrew Nelson <andrew.s.nelson@gmail.com> Reviewed-by: Josef Bacik <josef@toxicpanda.com> Signed-off-by: Filipe Manana <fdmanana@suse.com> Signed-off-by: David Sterba <dsterba@suse.com>
2018-10-24 09:13:03 +00:00
struct inode *btrfs_iget_path(struct super_block *s, struct btrfs_key *location,
struct btrfs_root *root, int *new,
struct btrfs_path *path);
struct inode *btrfs_iget(struct super_block *s, struct btrfs_key *location,
Btrfs: change how we mount subvolumes This work is in preperation for being able to set a different root as the default mounting root. There is currently a problem with how we mount subvolumes. We cannot currently mount a subvolume of a subvolume, you can only mount subvolumes/snapshots of the default subvolume. So say you take a snapshot of the default subvolume and call it snap1, and then take a snapshot of snap1 and call it snap2, so now you have / /snap1 /snap1/snap2 as your available volumes. Currently you can only mount / and /snap1, you cannot mount /snap1/snap2. To fix this problem instead of passing subvolid=<name> you must pass in subvolid=<treeid>, where <treeid> is the tree id that gets spit out via the subvolume listing you get from the subvolume listing patches (btrfs filesystem list). This allows us to mount /, /snap1 and /snap1/snap2 as the root volume. In addition to the above, we also now read the default dir item in the tree root to get the root key that it points to. For now this just points at what has always been the default subvolme, but later on I plan to change it to point at whatever root you want to be the new default root, so you can just set the default mount and not have to mount with -o subvolid=<treeid>. I tested this out with the above scenario and it worked perfectly. Thanks, mount -o subvol operates inside the selected subvolid. For example: mount -o subvol=snap1,subvolid=256 /dev/xxx /mnt /mnt will have the snap1 directory for the subvolume with id 256. mount -o subvol=snap /dev/xxx /mnt /mnt will be the snap directory of whatever the default subvolume is. Signed-off-by: Josef Bacik <josef@redhat.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-12-04 17:38:27 +00:00
struct btrfs_root *root, int *was_new);
struct extent_map *btrfs_get_extent(struct btrfs_inode *inode,
struct page *page, size_t pg_offset,
u64 start, u64 end, int create);
int btrfs_update_inode(struct btrfs_trans_handle *trans,
struct btrfs_root *root,
struct inode *inode);
int btrfs_update_inode_fallback(struct btrfs_trans_handle *trans,
struct btrfs_root *root, struct inode *inode);
int btrfs_orphan_add(struct btrfs_trans_handle *trans,
struct btrfs_inode *inode);
int btrfs_orphan_cleanup(struct btrfs_root *root);
int btrfs_cont_expand(struct inode *inode, loff_t oldsize, loff_t size);
void btrfs_add_delayed_iput(struct inode *inode);
void btrfs_run_delayed_iputs(struct btrfs_fs_info *fs_info);
int btrfs_wait_on_delayed_iputs(struct btrfs_fs_info *fs_info);
int btrfs_prealloc_file_range(struct inode *inode, int mode,
u64 start, u64 num_bytes, u64 min_size,
loff_t actual_len, u64 *alloc_hint);
int btrfs_prealloc_file_range_trans(struct inode *inode,
struct btrfs_trans_handle *trans, int mode,
u64 start, u64 num_bytes, u64 min_size,
loff_t actual_len, u64 *alloc_hint);
int btrfs_run_delalloc_range(struct inode *inode, struct page *locked_page,
u64 start, u64 end, int *page_started, unsigned long *nr_written,
struct writeback_control *wbc);
int btrfs_writepage_cow_fixup(struct page *page, u64 start, u64 end);
void btrfs_writepage_endio_finish_ordered(struct page *page, u64 start,
u64 end, int uptodate);
extern const struct dentry_operations btrfs_dentry_operations;
/* ioctl.c */
long btrfs_ioctl(struct file *file, unsigned int cmd, unsigned long arg);
long btrfs_compat_ioctl(struct file *file, unsigned int cmd, unsigned long arg);
int btrfs_ioctl_get_supported_features(void __user *arg);
void btrfs_sync_inode_flags_to_i_flags(struct inode *inode);
int btrfs_is_empty_uuid(u8 *uuid);
int btrfs_defrag_file(struct inode *inode, struct file *file,
struct btrfs_ioctl_defrag_range_args *range,
u64 newer_than, unsigned long max_pages);
void btrfs_get_block_group_info(struct list_head *groups_list,
struct btrfs_ioctl_space_info *space);
void btrfs_update_ioctl_balance_args(struct btrfs_fs_info *fs_info,
struct btrfs_ioctl_balance_args *bargs);
/* file.c */
int __init btrfs_auto_defrag_init(void);
void __cold btrfs_auto_defrag_exit(void);
int btrfs_add_inode_defrag(struct btrfs_trans_handle *trans,
struct btrfs_inode *inode);
int btrfs_run_defrag_inodes(struct btrfs_fs_info *fs_info);
void btrfs_cleanup_defrag_inodes(struct btrfs_fs_info *fs_info);
int btrfs_sync_file(struct file *file, loff_t start, loff_t end, int datasync);
void btrfs_drop_extent_cache(struct btrfs_inode *inode, u64 start, u64 end,
int skip_pinned);
extern const struct file_operations btrfs_file_operations;
Btrfs: turbo charge fsync At least for the vm workload. Currently on fsync we will 1) Truncate all items in the log tree for the given inode if they exist and 2) Copy all items for a given inode into the log The problem with this is that for things like VMs you can have lots of extents from the fragmented writing behavior, and worst yet you may have only modified a few extents, not the entire thing. This patch fixes this problem by tracking which transid modified our extent, and then when we do the tree logging we find all of the extents we've modified in our current transaction, sort them and commit them. We also only truncate up to the xattrs of the inode and copy that stuff in normally, and then just drop any extents in the range we have that exist in the log already. Here are some numbers of a 50 meg fio job that does random writes and fsync()s after every write Original Patched SATA drive 82KB/s 140KB/s Fusion drive 431KB/s 2532KB/s So around 2-6 times faster depending on your hardware. There are a few corner cases, for example if you truncate at all we have to do it the old way since there is no way to be sure what is in the log is ok. This probably could be done smarter, but if you write-fsync-truncate-write-fsync you deserve what you get. All this work is in RAM of course so if your inode gets evicted from cache and you read it in and fsync it we'll do it the slow way if we are still in the same transaction that we last modified the inode in. The biggest cool part of this is that it requires no changes to the recovery code, so if you fsync with this patch and crash and load an old kernel, it will run the recovery and be a-ok. I have tested this pretty thoroughly with an fsync tester and everything comes back fine, as well as xfstests. Thanks, Signed-off-by: Josef Bacik <jbacik@fusionio.com>
2012-08-17 17:14:17 +00:00
int __btrfs_drop_extents(struct btrfs_trans_handle *trans,
struct btrfs_root *root, struct inode *inode,
struct btrfs_path *path, u64 start, u64 end,
Btrfs: faster file extent item replace operations When writing to a file we drop existing file extent items that cover the write range and then add a new file extent item that represents that write range. Before this change we were doing a tree lookup to remove the file extent items, and then after we did another tree lookup to insert the new file extent item. Most of the time all the file extent items we need to drop are located within a single leaf - this is the leaf where our new file extent item ends up at. Therefore, in this common case just combine these 2 operations into a single one. By avoiding the second btree navigation for insertion of the new file extent item, we reduce btree node/leaf lock acquisitions/releases, btree block/leaf COW operations, CPU time on btree node/leaf key binary searches, etc. Besides for file writes, this is an operation that happens for file fsync's as well. However log btrees are much less likely to big as big as regular fs btrees, therefore the impact of this change is smaller. The following benchmark was performed against an SSD drive and a HDD drive, both for random and sequential writes: sysbench --test=fileio --file-num=4096 --file-total-size=8G \ --file-test-mode=[rndwr|seqwr] --num-threads=512 \ --file-block-size=8192 \ --max-requests=1000000 \ --file-fsync-freq=0 --file-io-mode=sync [prepare|run] All results below are averages of 10 runs of the respective test. ** SSD sequential writes Before this change: 225.88 Mb/sec After this change: 277.26 Mb/sec ** SSD random writes Before this change: 49.91 Mb/sec After this change: 56.39 Mb/sec ** HDD sequential writes Before this change: 68.53 Mb/sec After this change: 69.87 Mb/sec ** HDD random writes Before this change: 13.04 Mb/sec After this change: 14.39 Mb/sec Signed-off-by: Filipe David Borba Manana <fdmanana@gmail.com> Signed-off-by: Josef Bacik <jbacik@fb.com> Signed-off-by: Chris Mason <clm@fb.com>
2014-01-07 11:42:27 +00:00
u64 *drop_end, int drop_cache,
int replace_extent,
u32 extent_item_size,
int *key_inserted);
Btrfs: turbo charge fsync At least for the vm workload. Currently on fsync we will 1) Truncate all items in the log tree for the given inode if they exist and 2) Copy all items for a given inode into the log The problem with this is that for things like VMs you can have lots of extents from the fragmented writing behavior, and worst yet you may have only modified a few extents, not the entire thing. This patch fixes this problem by tracking which transid modified our extent, and then when we do the tree logging we find all of the extents we've modified in our current transaction, sort them and commit them. We also only truncate up to the xattrs of the inode and copy that stuff in normally, and then just drop any extents in the range we have that exist in the log already. Here are some numbers of a 50 meg fio job that does random writes and fsync()s after every write Original Patched SATA drive 82KB/s 140KB/s Fusion drive 431KB/s 2532KB/s So around 2-6 times faster depending on your hardware. There are a few corner cases, for example if you truncate at all we have to do it the old way since there is no way to be sure what is in the log is ok. This probably could be done smarter, but if you write-fsync-truncate-write-fsync you deserve what you get. All this work is in RAM of course so if your inode gets evicted from cache and you read it in and fsync it we'll do it the slow way if we are still in the same transaction that we last modified the inode in. The biggest cool part of this is that it requires no changes to the recovery code, so if you fsync with this patch and crash and load an old kernel, it will run the recovery and be a-ok. I have tested this pretty thoroughly with an fsync tester and everything comes back fine, as well as xfstests. Thanks, Signed-off-by: Josef Bacik <jbacik@fusionio.com>
2012-08-17 17:14:17 +00:00
int btrfs_drop_extents(struct btrfs_trans_handle *trans,
struct btrfs_root *root, struct inode *inode, u64 start,
u64 end, int drop_cache);
int btrfs_mark_extent_written(struct btrfs_trans_handle *trans,
struct btrfs_inode *inode, u64 start, u64 end);
int btrfs_release_file(struct inode *inode, struct file *file);
int btrfs_dirty_pages(struct inode *inode, struct page **pages,
size_t num_pages, loff_t pos, size_t write_bytes,
struct extent_state **cached);
int btrfs_fdatawrite_range(struct inode *inode, loff_t start, loff_t end);
loff_t btrfs_remap_file_range(struct file *file_in, loff_t pos_in,
struct file *file_out, loff_t pos_out,
loff_t len, unsigned int remap_flags);
/* tree-defrag.c */
int btrfs_defrag_leaves(struct btrfs_trans_handle *trans,
struct btrfs_root *root);
/* sysfs.c */
int __init btrfs_init_sysfs(void);
void __cold btrfs_exit_sysfs(void);
int btrfs_sysfs_add_mounted(struct btrfs_fs_info *fs_info);
void btrfs_sysfs_remove_mounted(struct btrfs_fs_info *fs_info);
/* super.c */
int btrfs_parse_options(struct btrfs_fs_info *info, char *options,
unsigned long new_flags);
int btrfs_sync_fs(struct super_block *sb, int wait);
static inline __printf(2, 3) __cold
void btrfs_no_printk(const struct btrfs_fs_info *fs_info, const char *fmt, ...)
{
}
#ifdef CONFIG_PRINTK
__printf(2, 3)
__cold
void btrfs_printk(const struct btrfs_fs_info *fs_info, const char *fmt, ...);
#else
#define btrfs_printk(fs_info, fmt, args...) \
btrfs_no_printk(fs_info, fmt, ##args)
#endif
#define btrfs_emerg(fs_info, fmt, args...) \
btrfs_printk(fs_info, KERN_EMERG fmt, ##args)
#define btrfs_alert(fs_info, fmt, args...) \
btrfs_printk(fs_info, KERN_ALERT fmt, ##args)
#define btrfs_crit(fs_info, fmt, args...) \
btrfs_printk(fs_info, KERN_CRIT fmt, ##args)
#define btrfs_err(fs_info, fmt, args...) \
btrfs_printk(fs_info, KERN_ERR fmt, ##args)
#define btrfs_warn(fs_info, fmt, args...) \
btrfs_printk(fs_info, KERN_WARNING fmt, ##args)
#define btrfs_notice(fs_info, fmt, args...) \
btrfs_printk(fs_info, KERN_NOTICE fmt, ##args)
#define btrfs_info(fs_info, fmt, args...) \
btrfs_printk(fs_info, KERN_INFO fmt, ##args)
/*
* Wrappers that use printk_in_rcu
*/
#define btrfs_emerg_in_rcu(fs_info, fmt, args...) \
btrfs_printk_in_rcu(fs_info, KERN_EMERG fmt, ##args)
#define btrfs_alert_in_rcu(fs_info, fmt, args...) \
btrfs_printk_in_rcu(fs_info, KERN_ALERT fmt, ##args)
#define btrfs_crit_in_rcu(fs_info, fmt, args...) \
btrfs_printk_in_rcu(fs_info, KERN_CRIT fmt, ##args)
#define btrfs_err_in_rcu(fs_info, fmt, args...) \
btrfs_printk_in_rcu(fs_info, KERN_ERR fmt, ##args)
#define btrfs_warn_in_rcu(fs_info, fmt, args...) \
btrfs_printk_in_rcu(fs_info, KERN_WARNING fmt, ##args)
#define btrfs_notice_in_rcu(fs_info, fmt, args...) \
btrfs_printk_in_rcu(fs_info, KERN_NOTICE fmt, ##args)
#define btrfs_info_in_rcu(fs_info, fmt, args...) \
btrfs_printk_in_rcu(fs_info, KERN_INFO fmt, ##args)
/*
* Wrappers that use a ratelimited printk_in_rcu
*/
#define btrfs_emerg_rl_in_rcu(fs_info, fmt, args...) \
btrfs_printk_rl_in_rcu(fs_info, KERN_EMERG fmt, ##args)
#define btrfs_alert_rl_in_rcu(fs_info, fmt, args...) \
btrfs_printk_rl_in_rcu(fs_info, KERN_ALERT fmt, ##args)
#define btrfs_crit_rl_in_rcu(fs_info, fmt, args...) \
btrfs_printk_rl_in_rcu(fs_info, KERN_CRIT fmt, ##args)
#define btrfs_err_rl_in_rcu(fs_info, fmt, args...) \
btrfs_printk_rl_in_rcu(fs_info, KERN_ERR fmt, ##args)
#define btrfs_warn_rl_in_rcu(fs_info, fmt, args...) \
btrfs_printk_rl_in_rcu(fs_info, KERN_WARNING fmt, ##args)
#define btrfs_notice_rl_in_rcu(fs_info, fmt, args...) \
btrfs_printk_rl_in_rcu(fs_info, KERN_NOTICE fmt, ##args)
#define btrfs_info_rl_in_rcu(fs_info, fmt, args...) \
btrfs_printk_rl_in_rcu(fs_info, KERN_INFO fmt, ##args)
/*
* Wrappers that use a ratelimited printk
*/
#define btrfs_emerg_rl(fs_info, fmt, args...) \
btrfs_printk_ratelimited(fs_info, KERN_EMERG fmt, ##args)
#define btrfs_alert_rl(fs_info, fmt, args...) \
btrfs_printk_ratelimited(fs_info, KERN_ALERT fmt, ##args)
#define btrfs_crit_rl(fs_info, fmt, args...) \
btrfs_printk_ratelimited(fs_info, KERN_CRIT fmt, ##args)
#define btrfs_err_rl(fs_info, fmt, args...) \
btrfs_printk_ratelimited(fs_info, KERN_ERR fmt, ##args)
#define btrfs_warn_rl(fs_info, fmt, args...) \
btrfs_printk_ratelimited(fs_info, KERN_WARNING fmt, ##args)
#define btrfs_notice_rl(fs_info, fmt, args...) \
btrfs_printk_ratelimited(fs_info, KERN_NOTICE fmt, ##args)
#define btrfs_info_rl(fs_info, fmt, args...) \
btrfs_printk_ratelimited(fs_info, KERN_INFO fmt, ##args)
#if defined(CONFIG_DYNAMIC_DEBUG)
#define btrfs_debug(fs_info, fmt, args...) \
_dynamic_func_call_no_desc(fmt, btrfs_printk, \
fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_in_rcu(fs_info, fmt, args...) \
_dynamic_func_call_no_desc(fmt, btrfs_printk_in_rcu, \
fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_rl_in_rcu(fs_info, fmt, args...) \
_dynamic_func_call_no_desc(fmt, btrfs_printk_rl_in_rcu, \
fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_rl(fs_info, fmt, args...) \
_dynamic_func_call_no_desc(fmt, btrfs_printk_ratelimited, \
fs_info, KERN_DEBUG fmt, ##args)
#elif defined(DEBUG)
#define btrfs_debug(fs_info, fmt, args...) \
btrfs_printk(fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_in_rcu(fs_info, fmt, args...) \
btrfs_printk_in_rcu(fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_rl_in_rcu(fs_info, fmt, args...) \
btrfs_printk_rl_in_rcu(fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_rl(fs_info, fmt, args...) \
btrfs_printk_ratelimited(fs_info, KERN_DEBUG fmt, ##args)
#else
#define btrfs_debug(fs_info, fmt, args...) \
btrfs_no_printk(fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_in_rcu(fs_info, fmt, args...) \
btrfs_no_printk_in_rcu(fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_rl_in_rcu(fs_info, fmt, args...) \
btrfs_no_printk_in_rcu(fs_info, KERN_DEBUG fmt, ##args)
#define btrfs_debug_rl(fs_info, fmt, args...) \
btrfs_no_printk(fs_info, KERN_DEBUG fmt, ##args)
#endif
#define btrfs_printk_in_rcu(fs_info, fmt, args...) \
do { \
rcu_read_lock(); \
btrfs_printk(fs_info, fmt, ##args); \
rcu_read_unlock(); \
} while (0)
#define btrfs_no_printk_in_rcu(fs_info, fmt, args...) \
do { \
rcu_read_lock(); \
btrfs_no_printk(fs_info, fmt, ##args); \
rcu_read_unlock(); \
} while (0)
#define btrfs_printk_ratelimited(fs_info, fmt, args...) \
do { \
static DEFINE_RATELIMIT_STATE(_rs, \
DEFAULT_RATELIMIT_INTERVAL, \
DEFAULT_RATELIMIT_BURST); \
if (__ratelimit(&_rs)) \
btrfs_printk(fs_info, fmt, ##args); \
} while (0)
#define btrfs_printk_rl_in_rcu(fs_info, fmt, args...) \
do { \
rcu_read_lock(); \
btrfs_printk_ratelimited(fs_info, fmt, ##args); \
rcu_read_unlock(); \
} while (0)
__cold
static inline void assfail(const char *expr, const char *file, int line)
{
if (IS_ENABLED(CONFIG_BTRFS_ASSERT)) {
pr_err("assertion failed: %s, file: %s, line: %d\n",
expr, file, line);
BUG();
}
}
#define ASSERT(expr) \
(likely(expr) ? (void)0 : assfail(#expr, __FILE__, __LINE__))
/*
* Use that for functions that are conditionally exported for sanity tests but
* otherwise static
*/
#ifndef CONFIG_BTRFS_FS_RUN_SANITY_TESTS
#define EXPORT_FOR_TESTS static
#else
#define EXPORT_FOR_TESTS
#endif
__cold
static inline void btrfs_print_v0_err(struct btrfs_fs_info *fs_info)
{
btrfs_err(fs_info,
"Unsupported V0 extent filesystem detected. Aborting. Please re-create your filesystem with a newer kernel");
}
__printf(5, 6)
__cold
void __btrfs_handle_fs_error(struct btrfs_fs_info *fs_info, const char *function,
unsigned int line, int errno, const char *fmt, ...);
const char *btrfs_decode_error(int errno);
__cold
void __btrfs_abort_transaction(struct btrfs_trans_handle *trans,
const char *function,
unsigned int line, int errno);
/*
* Call btrfs_abort_transaction as early as possible when an error condition is
* detected, that way the exact line number is reported.
*/
#define btrfs_abort_transaction(trans, errno) \
do { \
/* Report first abort since mount */ \
if (!test_and_set_bit(BTRFS_FS_STATE_TRANS_ABORTED, \
&((trans)->fs_info->fs_state))) { \
if ((errno) != -EIO) { \
WARN(1, KERN_DEBUG \
"BTRFS: Transaction aborted (error %d)\n", \
(errno)); \
} else { \
btrfs_debug((trans)->fs_info, \
"Transaction aborted (error %d)", \
(errno)); \
} \
} \
__btrfs_abort_transaction((trans), __func__, \
__LINE__, (errno)); \
} while (0)
#define btrfs_handle_fs_error(fs_info, errno, fmt, args...) \
do { \
__btrfs_handle_fs_error((fs_info), __func__, __LINE__, \
(errno), fmt, ##args); \
} while (0)
__printf(5, 6)
__cold
void __btrfs_panic(struct btrfs_fs_info *fs_info, const char *function,
unsigned int line, int errno, const char *fmt, ...);
/*
* If BTRFS_MOUNT_PANIC_ON_FATAL_ERROR is in mount_opt, __btrfs_panic
* will panic(). Otherwise we BUG() here.
*/
#define btrfs_panic(fs_info, errno, fmt, args...) \
do { \
__btrfs_panic(fs_info, __func__, __LINE__, errno, fmt, ##args); \
BUG(); \
} while (0)
/* compatibility and incompatibility defines */
#define btrfs_set_fs_incompat(__fs_info, opt) \
__btrfs_set_fs_incompat((__fs_info), BTRFS_FEATURE_INCOMPAT_##opt)
static inline void __btrfs_set_fs_incompat(struct btrfs_fs_info *fs_info,
u64 flag)
{
struct btrfs_super_block *disk_super;
u64 features;
disk_super = fs_info->super_copy;
features = btrfs_super_incompat_flags(disk_super);
if (!(features & flag)) {
spin_lock(&fs_info->super_lock);
features = btrfs_super_incompat_flags(disk_super);
if (!(features & flag)) {
features |= flag;
btrfs_set_super_incompat_flags(disk_super, features);
btrfs_info(fs_info, "setting %llu feature flag",
flag);
}
spin_unlock(&fs_info->super_lock);
}
}
#define btrfs_clear_fs_incompat(__fs_info, opt) \
__btrfs_clear_fs_incompat((__fs_info), BTRFS_FEATURE_INCOMPAT_##opt)
static inline void __btrfs_clear_fs_incompat(struct btrfs_fs_info *fs_info,
u64 flag)
{
struct btrfs_super_block *disk_super;
u64 features;
disk_super = fs_info->super_copy;
features = btrfs_super_incompat_flags(disk_super);
if (features & flag) {
spin_lock(&fs_info->super_lock);
features = btrfs_super_incompat_flags(disk_super);
if (features & flag) {
features &= ~flag;
btrfs_set_super_incompat_flags(disk_super, features);
btrfs_info(fs_info, "clearing %llu feature flag",
flag);
}
spin_unlock(&fs_info->super_lock);
}
}
#define btrfs_fs_incompat(fs_info, opt) \
__btrfs_fs_incompat((fs_info), BTRFS_FEATURE_INCOMPAT_##opt)
static inline bool __btrfs_fs_incompat(struct btrfs_fs_info *fs_info, u64 flag)
{
struct btrfs_super_block *disk_super;
disk_super = fs_info->super_copy;
return !!(btrfs_super_incompat_flags(disk_super) & flag);
}
#define btrfs_set_fs_compat_ro(__fs_info, opt) \
__btrfs_set_fs_compat_ro((__fs_info), BTRFS_FEATURE_COMPAT_RO_##opt)
static inline void __btrfs_set_fs_compat_ro(struct btrfs_fs_info *fs_info,
u64 flag)
{
struct btrfs_super_block *disk_super;
u64 features;
disk_super = fs_info->super_copy;
features = btrfs_super_compat_ro_flags(disk_super);
if (!(features & flag)) {
spin_lock(&fs_info->super_lock);
features = btrfs_super_compat_ro_flags(disk_super);
if (!(features & flag)) {
features |= flag;
btrfs_set_super_compat_ro_flags(disk_super, features);
btrfs_info(fs_info, "setting %llu ro feature flag",
flag);
}
spin_unlock(&fs_info->super_lock);
}
}
#define btrfs_clear_fs_compat_ro(__fs_info, opt) \
__btrfs_clear_fs_compat_ro((__fs_info), BTRFS_FEATURE_COMPAT_RO_##opt)
static inline void __btrfs_clear_fs_compat_ro(struct btrfs_fs_info *fs_info,
u64 flag)
{
struct btrfs_super_block *disk_super;
u64 features;
disk_super = fs_info->super_copy;
features = btrfs_super_compat_ro_flags(disk_super);
if (features & flag) {
spin_lock(&fs_info->super_lock);
features = btrfs_super_compat_ro_flags(disk_super);
if (features & flag) {
features &= ~flag;
btrfs_set_super_compat_ro_flags(disk_super, features);
btrfs_info(fs_info, "clearing %llu ro feature flag",
flag);
}
spin_unlock(&fs_info->super_lock);
}
}
#define btrfs_fs_compat_ro(fs_info, opt) \
__btrfs_fs_compat_ro((fs_info), BTRFS_FEATURE_COMPAT_RO_##opt)
static inline int __btrfs_fs_compat_ro(struct btrfs_fs_info *fs_info, u64 flag)
{
struct btrfs_super_block *disk_super;
disk_super = fs_info->super_copy;
return !!(btrfs_super_compat_ro_flags(disk_super) & flag);
}
/* acl.c */
#ifdef CONFIG_BTRFS_FS_POSIX_ACL
struct posix_acl *btrfs_get_acl(struct inode *inode, int type);
int btrfs_set_acl(struct inode *inode, struct posix_acl *acl, int type);
int btrfs_init_acl(struct btrfs_trans_handle *trans,
struct inode *inode, struct inode *dir);
#else
#define btrfs_get_acl NULL
#define btrfs_set_acl NULL
static inline int btrfs_init_acl(struct btrfs_trans_handle *trans,
struct inode *inode, struct inode *dir)
{
return 0;
}
#endif
Btrfs: free space accounting redo 1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: Chris Mason <chris.mason@oracle.com>
2008-09-23 17:14:11 +00:00
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
/* relocation.c */
int btrfs_relocate_block_group(struct btrfs_fs_info *fs_info, u64 group_start);
Btrfs: Mixed back reference (FORWARD ROLLING FORMAT CHANGE) This commit introduces a new kind of back reference for btrfs metadata. Once a filesystem has been mounted with this commit, IT WILL NO LONGER BE MOUNTABLE BY OLDER KERNELS. When a tree block in subvolume tree is cow'd, the reference counts of all extents it points to are increased by one. At transaction commit time, the old root of the subvolume is recorded in a "dead root" data structure, and the btree it points to is later walked, dropping reference counts and freeing any blocks where the reference count goes to 0. The increments done during cow and decrements done after commit cancel out, and the walk is a very expensive way to go about freeing the blocks that are no longer referenced by the new btree root. This commit reduces the transaction overhead by avoiding the need for dead root records. When a non-shared tree block is cow'd, we free the old block at once, and the new block inherits old block's references. When a tree block with reference count > 1 is cow'd, we increase the reference counts of all extents the new block points to by one, and decrease the old block's reference count by one. This dead tree avoidance code removes the need to modify the reference counts of lower level extents when a non-shared tree block is cow'd. But we still need to update back ref for all pointers in the block. This is because the location of the block is recorded in the back ref item. We can solve this by introducing a new type of back ref. The new back ref provides information about pointer's key, level and in which tree the pointer lives. This information allow us to find the pointer by searching the tree. The shortcoming of the new back ref is that it only works for pointers in tree blocks referenced by their owner trees. This is mostly a problem for snapshots, where resolving one of these fuzzy back references would be O(number_of_snapshots) and quite slow. The solution used here is to use the fuzzy back references in the common case where a given tree block is only referenced by one root, and use the full back references when multiple roots have a reference on a given block. This commit adds per subvolume red-black tree to keep trace of cached inodes. The red-black tree helps the balancing code to find cached inodes whose inode numbers within a given range. This commit improves the balancing code by introducing several data structures to keep the state of balancing. The most important one is the back ref cache. It caches how the upper level tree blocks are referenced. This greatly reduce the overhead of checking back ref. The improved balancing code scales significantly better with a large number of snapshots. This is a very large commit and was written in a number of pieces. But, they depend heavily on the disk format change and were squashed together to make sure git bisect didn't end up in a bad state wrt space balancing or the format change. Signed-off-by: Yan Zheng <zheng.yan@oracle.com> Signed-off-by: Chris Mason <chris.mason@oracle.com>
2009-06-10 14:45:14 +00:00
int btrfs_init_reloc_root(struct btrfs_trans_handle *trans,
struct btrfs_root *root);
int btrfs_update_reloc_root(struct btrfs_trans_handle *trans,
struct btrfs_root *root);
int btrfs_recover_relocation(struct btrfs_root *root);
int btrfs_reloc_clone_csums(struct inode *inode, u64 file_pos, u64 len);
int btrfs_reloc_cow_block(struct btrfs_trans_handle *trans,
struct btrfs_root *root, struct extent_buffer *buf,
struct extent_buffer *cow);
void btrfs_reloc_pre_snapshot(struct btrfs_pending_snapshot *pending,
u64 *bytes_to_reserve);
int btrfs_reloc_post_snapshot(struct btrfs_trans_handle *trans,
struct btrfs_pending_snapshot *pending);
/* scrub.c */
int btrfs_scrub_dev(struct btrfs_fs_info *fs_info, u64 devid, u64 start,
u64 end, struct btrfs_scrub_progress *progress,
int readonly, int is_dev_replace);
void btrfs_scrub_pause(struct btrfs_fs_info *fs_info);
void btrfs_scrub_continue(struct btrfs_fs_info *fs_info);
int btrfs_scrub_cancel(struct btrfs_fs_info *info);
int btrfs_scrub_cancel_dev(struct btrfs_fs_info *info,
struct btrfs_device *dev);
int btrfs_scrub_progress(struct btrfs_fs_info *fs_info, u64 devid,
struct btrfs_scrub_progress *progress);
static inline void btrfs_init_full_stripe_locks_tree(
struct btrfs_full_stripe_locks_tree *locks_root)
{
locks_root->root = RB_ROOT;
mutex_init(&locks_root->lock);
}
Btrfs: fix use-after-free in the finishing procedure of the device replace During device replace test, we hit a null pointer deference (It was very easy to reproduce it by running xfstests' btrfs/011 on the devices with the virtio scsi driver). There were two bugs that caused this problem: - We might allocate new chunks on the replaced device after we updated the mapping tree. And we forgot to replace the source device in those mapping of the new chunks. - We might get the mapping information which including the source device before the mapping information update. And then submit the bio which was based on that mapping information after we freed the source device. For the first bug, we can fix it by doing mapping tree update and source device remove in the same context of the chunk mutex. The chunk mutex is used to protect the allocable device list, the above method can avoid the new chunk allocation, and after we remove the source device, all the new chunks will be allocated on the new device. So it can fix the first bug. For the second bug, we need make sure all flighting bios are finished and no new bios are produced during we are removing the source device. To fix this problem, we introduced a global @bio_counter, we not only inc/dec @bio_counter outsize of map_blocks, but also inc it before submitting bio and dec @bio_counter when ending bios. Since Raid56 is a little different and device replace dosen't support raid56 yet, it is not addressed in the patch and I add comments to make sure we will fix it in the future. Reported-by: Qu Wenruo <quwenruo@cn.fujitsu.com> Signed-off-by: Wang Shilong <wangsl.fnst@cn.fujitsu.com> Signed-off-by: Miao Xie <miaox@cn.fujitsu.com> Signed-off-by: Josef Bacik <jbacik@fb.com>
2014-01-30 08:46:55 +00:00
/* dev-replace.c */
void btrfs_bio_counter_inc_blocked(struct btrfs_fs_info *fs_info);
void btrfs_bio_counter_inc_noblocked(struct btrfs_fs_info *fs_info);
void btrfs_bio_counter_sub(struct btrfs_fs_info *fs_info, s64 amount);
static inline void btrfs_bio_counter_dec(struct btrfs_fs_info *fs_info)
{
btrfs_bio_counter_sub(fs_info, 1);
}
btrfs: initial readahead code and prototypes This is the implementation for the generic read ahead framework. To trigger a readahead, btrfs_reada_add must be called. It will start a read ahead for the given range [start, end) on tree root. The returned handle can either be used to wait on the readahead to finish (btrfs_reada_wait), or to send it to the background (btrfs_reada_detach). The read ahead works as follows: On btrfs_reada_add, the root of the tree is inserted into a radix_tree. reada_start_machine will then search for extents to prefetch and trigger some reads. When a read finishes for a node, all contained node/leaf pointers that lie in the given range will also be enqueued. The reads will be triggered in sequential order, thus giving a big win over a naive enumeration. It will also make use of multi-device layouts. Each disk will have its on read pointer and all disks will by utilized in parallel. Also will no two disks read both sides of a mirror simultaneously, as this would waste seeking capacity. Instead both disks will read different parts of the filesystem. Any number of readaheads can be started in parallel. The read order will be determined globally, i.e. 2 parallel readaheads will normally finish faster than the 2 started one after another. Changes v2: - protect root->node by transaction instead of node_lock - fix missed branches: The readahead had a too simple check to determine if a branch from a node should be checked or not. It now also records the upper bound of each node to see if the requested RA range lies within. - use KERN_CONT to debug output, to avoid line breaks - defer reada_start_machine to worker to avoid deadlock Changes v3: - protect root->node by rcu Changes v5: - changed EIO-semantics of reada_tree_block_flagged - remove spin_lock from reada_control and make elems an atomic_t - remove unused read_total from reada_control - kill reada_key_cmp, use btrfs_comp_cpu_keys instead - use kref-style release functions where possible - return struct reada_control * instead of void * from btrfs_reada_add Signed-off-by: Arne Jansen <sensille@gmx.net>
2011-05-23 12:33:49 +00:00
/* reada.c */
struct reada_control {
struct btrfs_fs_info *fs_info; /* tree to prefetch */
btrfs: initial readahead code and prototypes This is the implementation for the generic read ahead framework. To trigger a readahead, btrfs_reada_add must be called. It will start a read ahead for the given range [start, end) on tree root. The returned handle can either be used to wait on the readahead to finish (btrfs_reada_wait), or to send it to the background (btrfs_reada_detach). The read ahead works as follows: On btrfs_reada_add, the root of the tree is inserted into a radix_tree. reada_start_machine will then search for extents to prefetch and trigger some reads. When a read finishes for a node, all contained node/leaf pointers that lie in the given range will also be enqueued. The reads will be triggered in sequential order, thus giving a big win over a naive enumeration. It will also make use of multi-device layouts. Each disk will have its on read pointer and all disks will by utilized in parallel. Also will no two disks read both sides of a mirror simultaneously, as this would waste seeking capacity. Instead both disks will read different parts of the filesystem. Any number of readaheads can be started in parallel. The read order will be determined globally, i.e. 2 parallel readaheads will normally finish faster than the 2 started one after another. Changes v2: - protect root->node by transaction instead of node_lock - fix missed branches: The readahead had a too simple check to determine if a branch from a node should be checked or not. It now also records the upper bound of each node to see if the requested RA range lies within. - use KERN_CONT to debug output, to avoid line breaks - defer reada_start_machine to worker to avoid deadlock Changes v3: - protect root->node by rcu Changes v5: - changed EIO-semantics of reada_tree_block_flagged - remove spin_lock from reada_control and make elems an atomic_t - remove unused read_total from reada_control - kill reada_key_cmp, use btrfs_comp_cpu_keys instead - use kref-style release functions where possible - return struct reada_control * instead of void * from btrfs_reada_add Signed-off-by: Arne Jansen <sensille@gmx.net>
2011-05-23 12:33:49 +00:00
struct btrfs_key key_start;
struct btrfs_key key_end; /* exclusive */
atomic_t elems;
struct kref refcnt;
wait_queue_head_t wait;
};
struct reada_control *btrfs_reada_add(struct btrfs_root *root,
struct btrfs_key *start, struct btrfs_key *end);
int btrfs_reada_wait(void *handle);
void btrfs_reada_detach(void *handle);
int btree_readahead_hook(struct extent_buffer *eb, int err);
btrfs: initial readahead code and prototypes This is the implementation for the generic read ahead framework. To trigger a readahead, btrfs_reada_add must be called. It will start a read ahead for the given range [start, end) on tree root. The returned handle can either be used to wait on the readahead to finish (btrfs_reada_wait), or to send it to the background (btrfs_reada_detach). The read ahead works as follows: On btrfs_reada_add, the root of the tree is inserted into a radix_tree. reada_start_machine will then search for extents to prefetch and trigger some reads. When a read finishes for a node, all contained node/leaf pointers that lie in the given range will also be enqueued. The reads will be triggered in sequential order, thus giving a big win over a naive enumeration. It will also make use of multi-device layouts. Each disk will have its on read pointer and all disks will by utilized in parallel. Also will no two disks read both sides of a mirror simultaneously, as this would waste seeking capacity. Instead both disks will read different parts of the filesystem. Any number of readaheads can be started in parallel. The read order will be determined globally, i.e. 2 parallel readaheads will normally finish faster than the 2 started one after another. Changes v2: - protect root->node by transaction instead of node_lock - fix missed branches: The readahead had a too simple check to determine if a branch from a node should be checked or not. It now also records the upper bound of each node to see if the requested RA range lies within. - use KERN_CONT to debug output, to avoid line breaks - defer reada_start_machine to worker to avoid deadlock Changes v3: - protect root->node by rcu Changes v5: - changed EIO-semantics of reada_tree_block_flagged - remove spin_lock from reada_control and make elems an atomic_t - remove unused read_total from reada_control - kill reada_key_cmp, use btrfs_comp_cpu_keys instead - use kref-style release functions where possible - return struct reada_control * instead of void * from btrfs_reada_add Signed-off-by: Arne Jansen <sensille@gmx.net>
2011-05-23 12:33:49 +00:00
static inline int is_fstree(u64 rootid)
{
if (rootid == BTRFS_FS_TREE_OBJECTID ||
((s64)rootid >= (s64)BTRFS_FIRST_FREE_OBJECTID &&
!btrfs_qgroup_level(rootid)))
return 1;
return 0;
}
static inline int btrfs_defrag_cancelled(struct btrfs_fs_info *fs_info)
{
return signal_pending(current);
}
/* Sanity test specific functions */
#ifdef CONFIG_BTRFS_FS_RUN_SANITY_TESTS
void btrfs_test_inode_set_ops(struct inode *inode);
void btrfs_test_destroy_inode(struct inode *inode);
static inline int btrfs_is_testing(struct btrfs_fs_info *fs_info)
{
return test_bit(BTRFS_FS_STATE_DUMMY_FS_INFO, &fs_info->fs_state);
}
#else
static inline int btrfs_is_testing(struct btrfs_fs_info *fs_info)
{
return 0;
}
#endif
static inline void cond_wake_up(struct wait_queue_head *wq)
{
/*
* This implies a full smp_mb barrier, see comments for
* waitqueue_active why.
*/
if (wq_has_sleeper(wq))
wake_up(wq);
}
static inline void cond_wake_up_nomb(struct wait_queue_head *wq)
{
/*
* Special case for conditional wakeup where the barrier required for
* waitqueue_active is implied by some of the preceding code. Eg. one
* of such atomic operations (atomic_dec_and_return, ...), or a
* unlock/lock sequence, etc.
*/
if (waitqueue_active(wq))
wake_up(wq);
}
#endif