commit 081488051d upstream.
Unmapped folios accessed through file descriptors can be underprotected.
Those folios are added to the oldest generation based on:
1. The fact that they are less costly to reclaim (no need to walk the
rmap and flush the TLB) and have less impact on performance (don't
cause major PFs and can be non-blocking if needed again).
2. The observation that they are likely to be single-use. E.g., for
client use cases like Android, its apps parse configuration files
and store the data in heap (anon); for server use cases like MySQL,
it reads from InnoDB files and holds the cached data for tables in
buffer pools (anon).
However, the oldest generation can be very short lived, and if so, it
doesn't provide the PID controller with enough time to respond to a surge
of refaults. (Note that the PID controller uses weighted refaults and
those from evicted generations only take a half of the whole weight.) In
other words, for a short lived generation, the moving average smooths out
the spike quickly.
To fix the problem:
1. For folios that are already on LRU, if they can be beyond the
tracking range of tiers, i.e., five accesses through file
descriptors, move them to the second oldest generation to give them
more time to age. (Note that tiers are used by the PID controller
to statistically determine whether folios accessed multiple times
through file descriptors are worth protecting.)
2. When adding unmapped folios to LRU, adjust the placement of them so
that they are not too close to the tail. The effect of this is
similar to the above.
On Android, launching 55 apps sequentially:
Before After Change
workingset_refault_anon 25641024 25598972 0%
workingset_refault_file 115016834 106178438 -8%
Link: https://lkml.kernel.org/r/20231208061407.2125867-1-yuzhao@google.com
Fixes: ac35a49023 ("mm: multi-gen LRU: minimal implementation")
Signed-off-by: Yu Zhao <yuzhao@google.com>
Reported-by: Charan Teja Kalla <quic_charante@quicinc.com>
Tested-by: Kalesh Singh <kaleshsingh@google.com>
Cc: T.J. Mercier <tjmercier@google.com>
Cc: Kairui Song <ryncsn@gmail.com>
Cc: Hillf Danton <hdanton@sina.com>
Cc: Jaroslav Pulchart <jaroslav.pulchart@gooddata.com>
Cc: <stable@vger.kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Greg Kroah-Hartman <gregkh@linuxfoundation.org>
To avoid confusion, the terms "promotion" and "demotion" will be applied
to the multi-gen LRU, as a new convention; the terms "activation" and
"deactivation" will be applied to the active/inactive LRU, as usual.
The aging produces young generations. Given an lruvec, it increments
max_seq when max_seq-min_seq+1 approaches MIN_NR_GENS. The aging promotes
hot pages to the youngest generation when it finds them accessed through
page tables; the demotion of cold pages happens consequently when it
increments max_seq. Promotion in the aging path does not involve any LRU
list operations, only the updates of the gen counter and
lrugen->nr_pages[]; demotion, unless as the result of the increment of
max_seq, requires LRU list operations, e.g., lru_deactivate_fn(). The
aging has the complexity O(nr_hot_pages), since it is only interested in
hot pages.
The eviction consumes old generations. Given an lruvec, it increments
min_seq when lrugen->lists[] indexed by min_seq%MAX_NR_GENS becomes empty.
A feedback loop modeled after the PID controller monitors refaults over
anon and file types and decides which type to evict when both types are
available from the same generation.
The protection of pages accessed multiple times through file descriptors
takes place in the eviction path. Each generation is divided into
multiple tiers. A page accessed N times through file descriptors is in
tier order_base_2(N). Tiers do not have dedicated lrugen->lists[], only
bits in folio->flags. The aforementioned feedback loop also monitors
refaults over all tiers and decides when to protect pages in which tiers
(N>1), using the first tier (N=0,1) as a baseline. The first tier
contains single-use unmapped clean pages, which are most likely the best
choices. In contrast to promotion in the aging path, the protection of a
page in the eviction path is achieved by moving this page to the next
generation, i.e., min_seq+1, if the feedback loop decides so. This
approach has the following advantages:
1. It removes the cost of activation in the buffered access path by
inferring whether pages accessed multiple times through file
descriptors are statistically hot and thus worth protecting in the
eviction path.
2. It takes pages accessed through page tables into account and avoids
overprotecting pages accessed multiple times through file
descriptors. (Pages accessed through page tables are in the first
tier, since N=0.)
3. More tiers provide better protection for pages accessed more than
twice through file descriptors, when under heavy buffered I/O
workloads.
Server benchmark results:
Single workload:
fio (buffered I/O): +[30, 32]%
IOPS BW
5.19-rc1: 2673k 10.2GiB/s
patch1-6: 3491k 13.3GiB/s
Single workload:
memcached (anon): -[4, 6]%
Ops/sec KB/sec
5.19-rc1: 1161501.04 45177.25
patch1-6: 1106168.46 43025.04
Configurations:
CPU: two Xeon 6154
Mem: total 256G
Node 1 was only used as a ram disk to reduce the variance in the
results.
patch drivers/block/brd.c <<EOF
99,100c99,100
< gfp_flags = GFP_NOIO | __GFP_ZERO | __GFP_HIGHMEM;
< page = alloc_page(gfp_flags);
---
> gfp_flags = GFP_NOIO | __GFP_ZERO | __GFP_HIGHMEM | __GFP_THISNODE;
> page = alloc_pages_node(1, gfp_flags, 0);
EOF
cat >>/etc/systemd/system.conf <<EOF
CPUAffinity=numa
NUMAPolicy=bind
NUMAMask=0
EOF
cat >>/etc/memcached.conf <<EOF
-m 184320
-s /var/run/memcached/memcached.sock
-a 0766
-t 36
-B binary
EOF
cat fio.sh
modprobe brd rd_nr=1 rd_size=113246208
swapoff -a
mkfs.ext4 /dev/ram0
mount -t ext4 /dev/ram0 /mnt
mkdir /sys/fs/cgroup/user.slice/test
echo 38654705664 >/sys/fs/cgroup/user.slice/test/memory.max
echo $$ >/sys/fs/cgroup/user.slice/test/cgroup.procs
fio -name=mglru --numjobs=72 --directory=/mnt --size=1408m \
--buffered=1 --ioengine=io_uring --iodepth=128 \
--iodepth_batch_submit=32 --iodepth_batch_complete=32 \
--rw=randread --random_distribution=random --norandommap \
--time_based --ramp_time=10m --runtime=5m --group_reporting
cat memcached.sh
modprobe brd rd_nr=1 rd_size=113246208
swapoff -a
mkswap /dev/ram0
swapon /dev/ram0
memtier_benchmark -S /var/run/memcached/memcached.sock \
-P memcache_binary -n allkeys --key-minimum=1 \
--key-maximum=65000000 --key-pattern=P:P -c 1 -t 36 \
--ratio 1:0 --pipeline 8 -d 2000
memtier_benchmark -S /var/run/memcached/memcached.sock \
-P memcache_binary -n allkeys --key-minimum=1 \
--key-maximum=65000000 --key-pattern=R:R -c 1 -t 36 \
--ratio 0:1 --pipeline 8 --randomize --distinct-client-seed
Client benchmark results:
kswapd profiles:
5.19-rc1
40.33% page_vma_mapped_walk (overhead)
21.80% lzo1x_1_do_compress (real work)
7.53% do_raw_spin_lock
3.95% _raw_spin_unlock_irq
2.52% vma_interval_tree_iter_next
2.37% folio_referenced_one
2.28% vma_interval_tree_subtree_search
1.97% anon_vma_interval_tree_iter_first
1.60% ptep_clear_flush
1.06% __zram_bvec_write
patch1-6
39.03% lzo1x_1_do_compress (real work)
18.47% page_vma_mapped_walk (overhead)
6.74% _raw_spin_unlock_irq
3.97% do_raw_spin_lock
2.49% ptep_clear_flush
2.48% anon_vma_interval_tree_iter_first
1.92% folio_referenced_one
1.88% __zram_bvec_write
1.48% memmove
1.31% vma_interval_tree_iter_next
Configurations:
CPU: single Snapdragon 7c
Mem: total 4G
ChromeOS MemoryPressure [1]
[1] https://chromium.googlesource.com/chromiumos/platform/tast-tests/
Link: https://lkml.kernel.org/r/20220918080010.2920238-7-yuzhao@google.com
Signed-off-by: Yu Zhao <yuzhao@google.com>
Acked-by: Brian Geffon <bgeffon@google.com>
Acked-by: Jan Alexander Steffens (heftig) <heftig@archlinux.org>
Acked-by: Oleksandr Natalenko <oleksandr@natalenko.name>
Acked-by: Steven Barrett <steven@liquorix.net>
Acked-by: Suleiman Souhlal <suleiman@google.com>
Tested-by: Daniel Byrne <djbyrne@mtu.edu>
Tested-by: Donald Carr <d@chaos-reins.com>
Tested-by: Holger Hoffstätte <holger@applied-asynchrony.com>
Tested-by: Konstantin Kharlamov <Hi-Angel@yandex.ru>
Tested-by: Shuang Zhai <szhai2@cs.rochester.edu>
Tested-by: Sofia Trinh <sofia.trinh@edi.works>
Tested-by: Vaibhav Jain <vaibhav@linux.ibm.com>
Cc: Andi Kleen <ak@linux.intel.com>
Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com>
Cc: Barry Song <baohua@kernel.org>
Cc: Catalin Marinas <catalin.marinas@arm.com>
Cc: Dave Hansen <dave.hansen@linux.intel.com>
Cc: Hillf Danton <hdanton@sina.com>
Cc: Jens Axboe <axboe@kernel.dk>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Jonathan Corbet <corbet@lwn.net>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Matthew Wilcox <willy@infradead.org>
Cc: Mel Gorman <mgorman@suse.de>
Cc: Miaohe Lin <linmiaohe@huawei.com>
Cc: Michael Larabel <Michael@MichaelLarabel.com>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: Mike Rapoport <rppt@kernel.org>
Cc: Mike Rapoport <rppt@linux.ibm.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Qi Zheng <zhengqi.arch@bytedance.com>
Cc: Tejun Heo <tj@kernel.org>
Cc: Vlastimil Babka <vbabka@suse.cz>
Cc: Will Deacon <will@kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Currently shrinkers are anonymous objects. For debugging purposes they
can be identified by count/scan function names, but it's not always
useful: e.g. for superblock's shrinkers it's nice to have at least an
idea of to which superblock the shrinker belongs.
This commit adds names to shrinkers. register_shrinker() and
prealloc_shrinker() functions are extended to take a format and arguments
to master a name.
In some cases it's not possible to determine a good name at the time when
a shrinker is allocated. For such cases shrinker_debugfs_rename() is
provided.
The expected format is:
<subsystem>-<shrinker_type>[:<instance>]-<id>
For some shrinkers an instance can be encoded as (MAJOR:MINOR) pair.
After this change the shrinker debugfs directory looks like:
$ cd /sys/kernel/debug/shrinker/
$ ls
dquota-cache-16 sb-devpts-28 sb-proc-47 sb-tmpfs-42
mm-shadow-18 sb-devtmpfs-5 sb-proc-48 sb-tmpfs-43
mm-zspool:zram0-34 sb-hugetlbfs-17 sb-pstore-31 sb-tmpfs-44
rcu-kfree-0 sb-hugetlbfs-33 sb-rootfs-2 sb-tmpfs-49
sb-aio-20 sb-iomem-12 sb-securityfs-6 sb-tracefs-13
sb-anon_inodefs-15 sb-mqueue-21 sb-selinuxfs-22 sb-xfs:vda1-36
sb-bdev-3 sb-nsfs-4 sb-sockfs-8 sb-zsmalloc-19
sb-bpf-32 sb-pipefs-14 sb-sysfs-26 thp-deferred_split-10
sb-btrfs:vda2-24 sb-proc-25 sb-tmpfs-1 thp-zero-9
sb-cgroup2-30 sb-proc-39 sb-tmpfs-27 xfs-buf:vda1-37
sb-configfs-23 sb-proc-41 sb-tmpfs-29 xfs-inodegc:vda1-38
sb-dax-11 sb-proc-45 sb-tmpfs-35
sb-debugfs-7 sb-proc-46 sb-tmpfs-40
[roman.gushchin@linux.dev: fix build warnings]
Link: https://lkml.kernel.org/r/Yr+ZTnLb9lJk6fJO@castle
Reported-by: kernel test robot <lkp@intel.com>
Link: https://lkml.kernel.org/r/20220601032227.4076670-4-roman.gushchin@linux.dev
Signed-off-by: Roman Gushchin <roman.gushchin@linux.dev>
Cc: Christophe JAILLET <christophe.jaillet@wanadoo.fr>
Cc: Dave Chinner <dchinner@redhat.com>
Cc: Hillf Danton <hdanton@sina.com>
Cc: Kent Overstreet <kent.overstreet@gmail.com>
Cc: Muchun Song <songmuchun@bytedance.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Daniel Dao has reported [1] a regression on workloads that may trigger a
lot of refaults (anon and file). The underlying issue is that flushing
rstat is expensive. Although rstat flush are batched with (nr_cpus *
MEMCG_BATCH) stat updates, it seems like there are workloads which
genuinely do stat updates larger than batch value within short amount of
time. Since the rstat flush can happen in the performance critical
codepaths like page faults, such workload can suffer greatly.
This patch fixes this regression by making the rstat flushing
conditional in the performance critical codepaths. More specifically,
the kernel relies on the async periodic rstat flusher to flush the stats
and only if the periodic flusher is delayed by more than twice the
amount of its normal time window then the kernel allows rstat flushing
from the performance critical codepaths.
Now the question: what are the side-effects of this change? The worst
that can happen is the refault codepath will see 4sec old lruvec stats
and may cause false (or missed) activations of the refaulted page which
may under-or-overestimate the workingset size. Though that is not very
concerning as the kernel can already miss or do false activations.
There are two more codepaths whose flushing behavior is not changed by
this patch and we may need to come to them in future. One is the
writeback stats used by dirty throttling and second is the deactivation
heuristic in the reclaim. For now keeping an eye on them and if there
is report of regression due to these codepaths, we will reevaluate then.
Link: https://lore.kernel.org/all/CA+wXwBSyO87ZX5PVwdHm-=dBjZYECGmfnydUicUyrQqndgX2MQ@mail.gmail.com [1]
Link: https://lkml.kernel.org/r/20220304184040.1304781-1-shakeelb@google.com
Fixes: 1f828223b7 ("memcg: flush lruvec stats in the refault")
Signed-off-by: Shakeel Butt <shakeelb@google.com>
Reported-by: Daniel Dao <dqminh@cloudflare.com>
Tested-by: Ivan Babrou <ivan@cloudflare.com>
Cc: Michal Hocko <mhocko@suse.com>
Cc: Roman Gushchin <roman.gushchin@linux.dev>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Michal Koutný <mkoutny@suse.com>
Cc: Frank Hofmann <fhofmann@cloudflare.com>
Cc: <stable@vger.kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
- Rewrite how munlock works to massively reduce the contention
on i_mmap_rwsem (Hugh Dickins):
https://lore.kernel.org/linux-mm/8e4356d-9622-a7f0-b2c-f116b5f2efea@google.com/
- Sort out the page refcount mess for ZONE_DEVICE pages (Christoph Hellwig):
https://lore.kernel.org/linux-mm/20220210072828.2930359-1-hch@lst.de/
- Convert GUP to use folios and make pincount available for order-1
pages. (Matthew Wilcox)
- Convert a few more truncation functions to use folios (Matthew Wilcox)
- Convert page_vma_mapped_walk to use PFNs instead of pages (Matthew Wilcox)
- Convert rmap_walk to use folios (Matthew Wilcox)
- Convert most of shrink_page_list() to use a folio (Matthew Wilcox)
- Add support for creating large folios in readahead (Matthew Wilcox)
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Merge tag 'folio-5.18c' of git://git.infradead.org/users/willy/pagecache
Pull folio updates from Matthew Wilcox:
- Rewrite how munlock works to massively reduce the contention on
i_mmap_rwsem (Hugh Dickins):
https://lore.kernel.org/linux-mm/8e4356d-9622-a7f0-b2c-f116b5f2efea@google.com/
- Sort out the page refcount mess for ZONE_DEVICE pages (Christoph
Hellwig):
https://lore.kernel.org/linux-mm/20220210072828.2930359-1-hch@lst.de/
- Convert GUP to use folios and make pincount available for order-1
pages. (Matthew Wilcox)
- Convert a few more truncation functions to use folios (Matthew
Wilcox)
- Convert page_vma_mapped_walk to use PFNs instead of pages (Matthew
Wilcox)
- Convert rmap_walk to use folios (Matthew Wilcox)
- Convert most of shrink_page_list() to use a folio (Matthew Wilcox)
- Add support for creating large folios in readahead (Matthew Wilcox)
* tag 'folio-5.18c' of git://git.infradead.org/users/willy/pagecache: (114 commits)
mm/damon: minor cleanup for damon_pa_young
selftests/vm/transhuge-stress: Support file-backed PMD folios
mm/filemap: Support VM_HUGEPAGE for file mappings
mm/readahead: Switch to page_cache_ra_order
mm/readahead: Align file mappings for non-DAX
mm/readahead: Add large folio readahead
mm: Support arbitrary THP sizes
mm: Make large folios depend on THP
mm: Fix READ_ONLY_THP warning
mm/filemap: Allow large folios to be added to the page cache
mm: Turn can_split_huge_page() into can_split_folio()
mm/vmscan: Convert pageout() to take a folio
mm/vmscan: Turn page_check_references() into folio_check_references()
mm/vmscan: Account large folios correctly
mm/vmscan: Optimise shrink_page_list for non-PMD-sized folios
mm/vmscan: Free non-shmem folios without splitting them
mm/rmap: Constify the rmap_walk_control argument
mm/rmap: Convert rmap_walk() to take a folio
mm: Turn page_anon_vma() into folio_anon_vma()
mm/rmap: Turn page_lock_anon_vma_read() into folio_lock_anon_vma_read()
...
Commit 68d48e6a2d ("mm: workingset: add vmstat counter for shadow
nodes") introduced an IRQ-off check to ensure that a lock is held which
also disabled interrupts. This does not work the same way on PREEMPT_RT
because none of the locks, that are held, disable interrupts.
Replace this check with a lockdep assert which ensures that the lock is
held.
Link: https://lkml.kernel.org/r/20220301122143.1521823-3-bigeasy@linutronix.de
Signed-off-by: Sebastian Andrzej Siewior <bigeasy@linutronix.de>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Tejun Heo <tj@kernel.org>
Cc: Zefan Li <lizefan.x@bytedance.com>
Cc: Thomas Gleixner <tglx@linutronix.de>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The workingset will add the xa_node to the shadow_nodes list. So the
allocation of xa_node should be done by kmem_cache_alloc_lru(). Using
xas_set_lru() to pass the list_lru which we want to insert xa_node into to
set up the xa_node reclaim context correctly.
Link: https://lkml.kernel.org/r/20220228122126.37293-9-songmuchun@bytedance.com
Signed-off-by: Muchun Song <songmuchun@bytedance.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Roman Gushchin <roman.gushchin@linux.dev>
Cc: Alex Shi <alexs@kernel.org>
Cc: Anna Schumaker <Anna.Schumaker@Netapp.com>
Cc: Chao Yu <chao@kernel.org>
Cc: Dave Chinner <david@fromorbit.com>
Cc: Fam Zheng <fam.zheng@bytedance.com>
Cc: Jaegeuk Kim <jaegeuk@kernel.org>
Cc: Kari Argillander <kari.argillander@gmail.com>
Cc: Matthew Wilcox (Oracle) <willy@infradead.org>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: Qi Zheng <zhengqi.arch@bytedance.com>
Cc: Shakeel Butt <shakeelb@google.com>
Cc: Theodore Ts'o <tytso@mit.edu>
Cc: Trond Myklebust <trond.myklebust@hammerspace.com>
Cc: Vladimir Davydov <vdavydov.dev@gmail.com>
Cc: Vlastimil Babka <vbabka@suse.cz>
Cc: Wei Yang <richard.weiyang@gmail.com>
Cc: Xiongchun Duan <duanxiongchun@bytedance.com>
Cc: Yang Shi <shy828301@gmail.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This removes an assumption that THPs are the only kind of compound
pages and removes a few hidden calls to compound_head().
Signed-off-by: Matthew Wilcox (Oracle) <willy@infradead.org>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Merge more updates from Andrew Morton:
"87 patches.
Subsystems affected by this patch series: mm (pagecache and hugetlb),
procfs, misc, MAINTAINERS, lib, checkpatch, binfmt, kallsyms, ramfs,
init, codafs, nilfs2, hfs, crash_dump, signals, seq_file, fork,
sysvfs, kcov, gdb, resource, selftests, and ipc"
* emailed patches from Andrew Morton <akpm@linux-foundation.org>: (87 commits)
ipc/ipc_sysctl.c: remove fallback for !CONFIG_PROC_SYSCTL
ipc: check checkpoint_restore_ns_capable() to modify C/R proc files
selftests/kselftest/runner/run_one(): allow running non-executable files
virtio-mem: disallow mapping virtio-mem memory via /dev/mem
kernel/resource: disallow access to exclusive system RAM regions
kernel/resource: clean up and optimize iomem_is_exclusive()
scripts/gdb: handle split debug for vmlinux
kcov: replace local_irq_save() with a local_lock_t
kcov: avoid enable+disable interrupts if !in_task()
kcov: allocate per-CPU memory on the relevant node
Documentation/kcov: define `ip' in the example
Documentation/kcov: include types.h in the example
sysv: use BUILD_BUG_ON instead of runtime check
kernel/fork.c: unshare(): use swap() to make code cleaner
seq_file: fix passing wrong private data
seq_file: move seq_escape() to a header
signal: remove duplicate include in signal.h
crash_dump: remove duplicate include in crash_dump.h
crash_dump: fix boolreturn.cocci warning
hfs/hfsplus: use WARN_ON for sanity check
...
Historically (pre-2.5), the inode shrinker used to reclaim only empty
inodes and skip over those that still contained page cache. This caused
problems on highmem hosts: struct inode could put fill lowmem zones
before the cache was getting reclaimed in the highmem zones.
To address this, the inode shrinker started to strip page cache to
facilitate reclaiming lowmem. However, this comes with its own set of
problems: the shrinkers may drop actively used page cache just because
the inodes are not currently open or dirty - think working with a large
git tree. It further doesn't respect cgroup memory protection settings
and can cause priority inversions between containers.
Nowadays, the page cache also holds non-resident info for evicted cache
pages in order to detect refaults. We've come to rely heavily on this
data inside reclaim for protecting the cache workingset and driving swap
behavior. We also use it to quantify and report workload health through
psi. The latter in turn is used for fleet health monitoring, as well as
driving automated memory sizing of workloads and containers, proactive
reclaim and memory offloading schemes.
The consequences of dropping page cache prematurely is that we're seeing
subtle and not-so-subtle failures in all of the above-mentioned
scenarios, with the workload generally entering unexpected thrashing
states while losing the ability to reliably detect it.
To fix this on non-highmem systems at least, going back to rotating
inodes on the LRU isn't feasible. We've tried (commit a76cf1a474
("mm: don't reclaim inodes with many attached pages")) and failed
(commit 69056ee6a8 ("Revert "mm: don't reclaim inodes with many
attached pages"")).
The issue is mostly that shrinker pools attract pressure based on their
size, and when objects get skipped the shrinkers remember this as
deferred reclaim work. This accumulates excessive pressure on the
remaining inodes, and we can quickly eat into heavily used ones, or
dirty ones that require IO to reclaim, when there potentially is plenty
of cold, clean cache around still.
Instead, this patch keeps populated inodes off the inode LRU in the
first place - just like an open file or dirty state would. An otherwise
clean and unused inode then gets queued when the last cache entry
disappears. This solves the problem without reintroducing the reclaim
issues, and generally is a bit more scalable than having to wade through
potentially hundreds of thousands of busy inodes.
Locking is a bit tricky because the locks protecting the inode state
(i_lock) and the inode LRU (lru_list.lock) don't nest inside the
irq-safe page cache lock (i_pages.xa_lock). Page cache deletions are
serialized through i_lock, taken before the i_pages lock, to make sure
depopulated inodes are queued reliably. Additions may race with
deletions, but we'll check again in the shrinker. If additions race
with the shrinker itself, we're protected by the i_lock: if find_inode()
or iput() win, the shrinker will bail on the elevated i_count or
I_REFERENCED; if the shrinker wins and goes ahead with the inode, it
will set I_FREEING and inhibit further igets(), which will cause the
other side to create a new instance of the inode instead.
Link: https://lkml.kernel.org/r/20210614211904.14420-4-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Cc: Roman Gushchin <guro@fb.com>
Cc: Tejun Heo <tj@kernel.org>
Cc: Dave Chinner <david@fromorbit.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This nets us 178 bytes of savings from removing calls to compound_head.
The three callers all grow a little, but each of them will be converted
to use folios soon, so that's fine.
Signed-off-by: Matthew Wilcox (Oracle) <willy@infradead.org>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: David Howells <dhowells@redhat.com>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
This function already assumed it was being passed a head page. No real
change here, except that thp_nr_pages() compiles away on kernels with
THP compiled out while folio_nr_pages() is always present. Also convert
page_memcg_rcu() to folio_memcg_rcu().
Signed-off-by: Matthew Wilcox (Oracle) <willy@infradead.org>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: David Howells <dhowells@redhat.com>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
This replaces mem_cgroup_page_lruvec(). All callers converted.
Signed-off-by: Matthew Wilcox (Oracle) <willy@infradead.org>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Acked-by: Mike Rapoport <rppt@linux.ibm.com>
Reviewed-by: David Howells <dhowells@redhat.com>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Prior to the commit 7e1c0d6f58 ("memcg: switch lruvec stats to rstat")
and the commit aa48e47e39 ("memcg: infrastructure to flush memcg
stats"), each lruvec memcg stats can be off by (nr_cgroups * nr_cpus *
32) at worst and for unbounded amount of time. The commit aa48e47e39
moved the lruvec stats to rstat infrastructure and the commit
7e1c0d6f58 bounded the error for all the lruvec stats to (nr_cpus *
32) at worst for at most 2 seconds. More specifically it decoupled the
number of stats and the number of cgroups from the error rate.
However this reduction in error comes with the cost of triggering the
slowpath of stats update more frequently. Previously in the slowpath
the kernel adds the stats up the memcg tree. After aa48e47e39, the
kernel triggers the asyn lruvec stats flush through queue_work(). This
causes regression reports from 0day kernel bot [1] as well as from
phoronix test suite [2].
We tried two options to fix the regression:
1) Increase the threshold to trigger the slowpath in lruvec stats
update codepath from 32 to 512.
2) Remove the slowpath from lruvec stats update codepath and instead
flush the stats in the page refault codepath. The assumption is that
the kernel timely flush the stats, so, the update tree would be
small in the refault codepath to not cause the preformance impact.
Following are the results of will-it-scale/page_fault[1|2|3] benchmark
on four settings i.e. (1) 5.15-rc1 as baseline (2) 5.15-rc1 with
aa48e47e39 and 7e1c0d6f58 reverted (3) 5.15-rc1 with option-1
(4) 5.15-rc1 with option-2.
test (1) (2) (3) (4)
pg_f1 368563 406277 (10.23%) 399693 (8.44%) 416398 (12.97%)
pg_f2 338399 372133 (9.96%) 369180 (9.09%) 381024 (12.59%)
pg_f3 500853 575399 (14.88%) 570388 (13.88%) 576083 (15.02%)
From the above result, it seems like the option-2 not only solves the
regression but also improves the performance for at least these
benchmarks.
Feng Tang (intel) ran the aim7 benchmark with these two options and
confirms that option-1 reduces the regression but option-2 removes the
regression.
Michael Larabel (phoronix) ran multiple benchmarks with these options
and reported the results at [3] and it shows for most benchmarks
option-2 removes the regression introduced by the commit aa48e47e39
("memcg: infrastructure to flush memcg stats").
Based on the experiment results, this patch proposed the option-2 as the
solution to resolve the regression.
Link: https://lore.kernel.org/all/20210726022421.GB21872@xsang-OptiPlex-9020 [1]
Link: https://www.phoronix.com/scan.php?page=article&item=linux515-compile-regress [2]
Link: https://openbenchmarking.org/result/2109226-DEBU-LINUX5104 [3]
Fixes: aa48e47e39 ("memcg: infrastructure to flush memcg stats")
Signed-off-by: Shakeel Butt <shakeelb@google.com>
Tested-by: Michael Larabel <Michael@phoronix.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Roman Gushchin <guro@fb.com>
Cc: Feng Tang <feng.tang@intel.com>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: Hillf Danton <hdanton@sina.com>,
Cc: Michal Koutný <mkoutny@suse.com>
Cc: Andrew Morton <akpm@linux-foundation.org>,
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Use the documented kernel-doc format to prevent kernel-doc warnings.
mm/workingset.c:256: warning: No description found for return value of 'workingset_eviction'
mm/workingset.c:285: warning: Function parameter or member 'folio' not described in 'workingset_refault'
mm/workingset.c:285: warning: Excess function parameter 'page' description in 'workingset_refault'
Link: https://lkml.kernel.org/r/20210808203153.10678-1-rdunlap@infradead.org
Signed-off-by: Randy Dunlap <rdunlap@infradead.org>
Cc: Matthew Wilcox (Oracle) <willy@infradead.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The magic number 1 is used in several places in workingset.c. Define a
macro WORKINGSET_SHIFT for it to improve code readability.
Link: https://lkml.kernel.org/r/20210624122307.1759342-1-linmiaohe@huawei.com
Signed-off-by: Miaohe Lin <linmiaohe@huawei.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
All the callers of mem_cgroup_page_lruvec() just pass page_pgdat(page) as
the 2nd parameter to it (except isolate_migratepages_block()). But for
isolate_migratepages_block(), the page_pgdat(page) is also equal to the
local variable of @pgdat. So mem_cgroup_page_lruvec() do not need the
pgdat parameter. Just remove it to simplify the code.
Link: https://lkml.kernel.org/r/20210417043538.9793-4-songmuchun@bytedance.com
Signed-off-by: Muchun Song <songmuchun@bytedance.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Reviewed-by: Shakeel Butt <shakeelb@google.com>
Acked-by: Roman Gushchin <guro@fb.com>
Acked-by: Michal Hocko <mhocko@suse.com>
Cc: Vladimir Davydov <vdavydov.dev@gmail.com>
Cc: Xiongchun Duan <duanxiongchun@bytedance.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
We no longer need to keep track of how many shadow entries are present in
a mapping. This saves a few writes to the inode and memory barriers.
Link: https://lkml.kernel.org/r/20201026151849.24232-3-willy@infradead.org
Signed-off-by: Matthew Wilcox (Oracle) <willy@infradead.org>
Tested-by: Vishal Verma <vishal.l.verma@intel.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The premise of the refault distance is that it can be seen as a deficit of
the inactive list space, so that if the inactive list would have had (R -
E) more slots, the page would not have been evicted but promoted to the
active list instead.
However, the way the code is ordered right now set us to be off by one, so
the real number of slots would be (R - E) + 1. I stumbled upon this when
trying to understand the code and it puzzled me that the comments did not
match what the code did.
This it not an issue at all since evictions and refaults tend to happen in
a number large enough that being off-by-one does not have any impact - and
since the compiler and CPUs are free to rearrange the execution sequence
anyway.
But as Johannes says, it is better to re-arrange the code in the proper
order since otherwise would be misleading to somebody who is actively
reading and trying to understand the logic of the code - like it happened
to me.
Link: https://lkml.kernel.org/r/20210201060651.3781-1-osalvador@suse.de
Signed-off-by: Oscar Salvador <osalvador@suse.de>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
If list_lru_shrink_count is 0, we always return SHRINK_EMPTY regardless of
the value of max_nodes. So we can return early if nodes == 0 to save some
cpu cycles of approximating a reasonable limit for the nodes.
Link: https://lkml.kernel.org/r/20210123073825.46709-1-linmiaohe@huawei.com
Signed-off-by: Miaohe Lin <linmiaohe@huawei.com>
Reviewed-by: Andrew Morton <akpm@linux-foundation.org>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Merge more updates from Andrew Morton:
"More MM work: a memcg scalability improvememt"
* emailed patches from Andrew Morton <akpm@linux-foundation.org>:
mm/lru: revise the comments of lru_lock
mm/lru: introduce relock_page_lruvec()
mm/lru: replace pgdat lru_lock with lruvec lock
mm/swap.c: serialize memcg changes in pagevec_lru_move_fn
mm/compaction: do page isolation first in compaction
mm/lru: introduce TestClearPageLRU()
mm/mlock: remove __munlock_isolate_lru_page()
mm/mlock: remove lru_lock on TestClearPageMlocked
mm/vmscan: remove lruvec reget in move_pages_to_lru
mm/lru: move lock into lru_note_cost
mm/swap.c: fold vm event PGROTATED into pagevec_move_tail_fn
mm/memcg: add debug checking in lock_page_memcg
mm: page_idle_get_page() does not need lru_lock
mm/rmap: stop store reordering issue on page->mapping
mm/vmscan: remove unnecessary lruvec adding
mm/thp: narrow lru locking
mm/thp: simplify lru_add_page_tail()
mm/thp: use head for head page in lru_add_page_tail()
mm/thp: move lru_add_page_tail() to huge_memory.c
We have to move lru_lock into lru_note_cost, since it cycle up on memcg
tree, for future per lruvec lru_lock replace. It's a bit ugly and may
cost a bit more locking, but benefit from multiple memcg locking could
cover the lost.
Link: https://lkml.kernel.org/r/1604566549-62481-11-git-send-email-alex.shi@linux.alibaba.com
Signed-off-by: Alex Shi <alex.shi@linux.alibaba.com>
Acked-by: Hugh Dickins <hughd@google.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Alexander Duyck <alexander.duyck@gmail.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Andrey Ryabinin <aryabinin@virtuozzo.com>
Cc: "Chen, Rong A" <rong.a.chen@intel.com>
Cc: Daniel Jordan <daniel.m.jordan@oracle.com>
Cc: "Huang, Ying" <ying.huang@intel.com>
Cc: Jann Horn <jannh@google.com>
Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com>
Cc: Kirill A. Shutemov <kirill@shutemov.name>
Cc: Konstantin Khlebnikov <khlebnikov@yandex-team.ru>
Cc: Matthew Wilcox (Oracle) <willy@infradead.org>
Cc: Mel Gorman <mgorman@techsingularity.net>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: Michal Hocko <mhocko@suse.com>
Cc: Mika Penttilä <mika.penttila@nextfour.com>
Cc: Minchan Kim <minchan@kernel.org>
Cc: Shakeel Butt <shakeelb@google.com>
Cc: Tejun Heo <tj@kernel.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: Vladimir Davydov <vdavydov.dev@gmail.com>
Cc: Vlastimil Babka <vbabka@suse.cz>
Cc: Wei Yang <richard.weiyang@gmail.com>
Cc: Yang Shi <yang.shi@linux.alibaba.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Core:
- support "prefer busy polling" NAPI operation mode, where we defer softirq
for some time expecting applications to periodically busy poll
- AF_XDP: improve efficiency by more batching and hindering
the adjacency cache prefetcher
- af_packet: make packet_fanout.arr size configurable up to 64K
- tcp: optimize TCP zero copy receive in presence of partial or unaligned
reads making zero copy a performance win for much smaller messages
- XDP: add bulk APIs for returning / freeing frames
- sched: support fragmenting IP packets as they come out of conntrack
- net: allow virtual netdevs to forward UDP L4 and fraglist GSO skbs
BPF:
- BPF switch from crude rlimit-based to memcg-based memory accounting
- BPF type format information for kernel modules and related tracing
enhancements
- BPF implement task local storage for BPF LSM
- allow the FENTRY/FEXIT/RAW_TP tracing programs to use bpf_sk_storage
Protocols:
- mptcp: improve multiple xmit streams support, memory accounting and
many smaller improvements
- TLS: support CHACHA20-POLY1305 cipher
- seg6: add support for SRv6 End.DT4/DT6 behavior
- sctp: Implement RFC 6951: UDP Encapsulation of SCTP
- ppp_generic: add ability to bridge channels directly
- bridge: Connectivity Fault Management (CFM) support as is defined in
IEEE 802.1Q section 12.14.
Drivers:
- mlx5: make use of the new auxiliary bus to organize the driver internals
- mlx5: more accurate port TX timestamping support
- mlxsw:
- improve the efficiency of offloaded next hop updates by using
the new nexthop object API
- support blackhole nexthops
- support IEEE 802.1ad (Q-in-Q) bridging
- rtw88: major bluetooth co-existance improvements
- iwlwifi: support new 6 GHz frequency band
- ath11k: Fast Initial Link Setup (FILS)
- mt7915: dual band concurrent (DBDC) support
- net: ipa: add basic support for IPA v4.5
Refactor:
- a few pieces of in_interrupt() cleanup work from Sebastian Andrzej Siewior
- phy: add support for shared interrupts; get rid of multiple driver
APIs and have the drivers write a full IRQ handler, slight growth
of driver code should be compensated by the simpler API which
also allows shared IRQs
- add common code for handling netdev per-cpu counters
- move TX packet re-allocation from Ethernet switch tag drivers to
a central place
- improve efficiency and rename nla_strlcpy
- number of W=1 warning cleanups as we now catch those in a patchwork
build bot
Old code removal:
- wan: delete the DLCI / SDLA drivers
- wimax: move to staging
- wifi: remove old WDS wifi bridging support
Signed-off-by: Jakub Kicinski <kuba@kernel.org>
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Merge tag 'net-next-5.11' of git://git.kernel.org/pub/scm/linux/kernel/git/netdev/net-next
Pull networking updates from Jakub Kicinski:
"Core:
- support "prefer busy polling" NAPI operation mode, where we defer
softirq for some time expecting applications to periodically busy
poll
- AF_XDP: improve efficiency by more batching and hindering the
adjacency cache prefetcher
- af_packet: make packet_fanout.arr size configurable up to 64K
- tcp: optimize TCP zero copy receive in presence of partial or
unaligned reads making zero copy a performance win for much smaller
messages
- XDP: add bulk APIs for returning / freeing frames
- sched: support fragmenting IP packets as they come out of conntrack
- net: allow virtual netdevs to forward UDP L4 and fraglist GSO skbs
BPF:
- BPF switch from crude rlimit-based to memcg-based memory accounting
- BPF type format information for kernel modules and related tracing
enhancements
- BPF implement task local storage for BPF LSM
- allow the FENTRY/FEXIT/RAW_TP tracing programs to use
bpf_sk_storage
Protocols:
- mptcp: improve multiple xmit streams support, memory accounting and
many smaller improvements
- TLS: support CHACHA20-POLY1305 cipher
- seg6: add support for SRv6 End.DT4/DT6 behavior
- sctp: Implement RFC 6951: UDP Encapsulation of SCTP
- ppp_generic: add ability to bridge channels directly
- bridge: Connectivity Fault Management (CFM) support as is defined
in IEEE 802.1Q section 12.14.
Drivers:
- mlx5: make use of the new auxiliary bus to organize the driver
internals
- mlx5: more accurate port TX timestamping support
- mlxsw:
- improve the efficiency of offloaded next hop updates by using
the new nexthop object API
- support blackhole nexthops
- support IEEE 802.1ad (Q-in-Q) bridging
- rtw88: major bluetooth co-existance improvements
- iwlwifi: support new 6 GHz frequency band
- ath11k: Fast Initial Link Setup (FILS)
- mt7915: dual band concurrent (DBDC) support
- net: ipa: add basic support for IPA v4.5
Refactor:
- a few pieces of in_interrupt() cleanup work from Sebastian Andrzej
Siewior
- phy: add support for shared interrupts; get rid of multiple driver
APIs and have the drivers write a full IRQ handler, slight growth
of driver code should be compensated by the simpler API which also
allows shared IRQs
- add common code for handling netdev per-cpu counters
- move TX packet re-allocation from Ethernet switch tag drivers to a
central place
- improve efficiency and rename nla_strlcpy
- number of W=1 warning cleanups as we now catch those in a patchwork
build bot
Old code removal:
- wan: delete the DLCI / SDLA drivers
- wimax: move to staging
- wifi: remove old WDS wifi bridging support"
* tag 'net-next-5.11' of git://git.kernel.org/pub/scm/linux/kernel/git/netdev/net-next: (1922 commits)
net: hns3: fix expression that is currently always true
net: fix proc_fs init handling in af_packet and tls
nfc: pn533: convert comma to semicolon
af_vsock: Assign the vsock transport considering the vsock address flags
af_vsock: Set VMADDR_FLAG_TO_HOST flag on the receive path
vsock_addr: Check for supported flag values
vm_sockets: Add VMADDR_FLAG_TO_HOST vsock flag
vm_sockets: Add flags field in the vsock address data structure
net: Disable NETIF_F_HW_TLS_TX when HW_CSUM is disabled
tcp: Add logic to check for SYN w/ data in tcp_simple_retransmit
net: mscc: ocelot: install MAC addresses in .ndo_set_rx_mode from process context
nfc: s3fwrn5: Release the nfc firmware
net: vxget: clean up sparse warnings
mlxsw: spectrum_router: Use eXtended mezzanine to offload IPv4 router
mlxsw: spectrum: Set KVH XLT cache mode for Spectrum2/3
mlxsw: spectrum_router_xm: Introduce basic XM cache flushing
mlxsw: reg: Add Router LPM Cache Enable Register
mlxsw: reg: Add Router LPM Cache ML Delete Register
mlxsw: spectrum_router_xm: Implement L-value tracking for M-index
mlxsw: reg: Add XM Router M Table Register
...
The *_lruvec_slab_state is also suitable for pages allocated from buddy,
not just for the slab objects. But the function name seems to tell us
that only slab object is applicable. So we can rename the keyword of slab
to kmem.
Link: https://lkml.kernel.org/r/20201117085249.24319-1-songmuchun@bytedance.com
Signed-off-by: Muchun Song <songmuchun@bytedance.com>
Acked-by: Roman Gushchin <guro@fb.com>
Reviewed-by: Shakeel Butt <shakeelb@google.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Patch series "mm: allow mapping accounted kernel pages to userspace", v6.
Currently a non-slab kernel page which has been charged to a memory cgroup
can't be mapped to userspace. The underlying reason is simple: PageKmemcg
flag is defined as a page type (like buddy, offline, etc), so it takes a
bit from a page->mapped counter. Pages with a type set can't be mapped to
userspace.
But in general the kmemcg flag has nothing to do with mapping to
userspace. It only means that the page has been accounted by the page
allocator, so it has to be properly uncharged on release.
Some bpf maps are mapping the vmalloc-based memory to userspace, and their
memory can't be accounted because of this implementation detail.
This patchset removes this limitation by moving the PageKmemcg flag into
one of the free bits of the page->mem_cgroup pointer. Also it formalizes
accesses to the page->mem_cgroup and page->obj_cgroups using new helpers,
adds several checks and removes a couple of obsolete functions. As the
result the code became more robust with fewer open-coded bit tricks.
This patch (of 4):
Currently there are many open-coded reads of the page->mem_cgroup pointer,
as well as a couple of read helpers, which are barely used.
It creates an obstacle on a way to reuse some bits of the pointer for
storing additional bits of information. In fact, we already do this for
slab pages, where the last bit indicates that a pointer has an attached
vector of objcg pointers instead of a regular memcg pointer.
This commits uses 2 existing helpers and introduces a new helper to
converts all read sides to calls of these helpers:
struct mem_cgroup *page_memcg(struct page *page);
struct mem_cgroup *page_memcg_rcu(struct page *page);
struct mem_cgroup *page_memcg_check(struct page *page);
page_memcg_check() is intended to be used in cases when the page can be a
slab page and have a memcg pointer pointing at objcg vector. It does
check the lowest bit, and if set, returns NULL. page_memcg() contains a
VM_BUG_ON_PAGE() check for the page not being a slab page.
To make sure nobody uses a direct access, struct page's
mem_cgroup/obj_cgroups is converted to unsigned long memcg_data.
Signed-off-by: Roman Gushchin <guro@fb.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Alexei Starovoitov <ast@kernel.org>
Reviewed-by: Shakeel Butt <shakeelb@google.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Michal Hocko <mhocko@suse.com>
Link: https://lkml.kernel.org/r/20201027001657.3398190-1-guro@fb.com
Link: https://lkml.kernel.org/r/20201027001657.3398190-2-guro@fb.com
Link: https://lore.kernel.org/bpf/20201201215900.3569844-2-guro@fb.com
- Fix the test suite after introduction of the local_lock
- Fix a bug in the IDA spotted by Coverity
- Change the API that allows the workingset code to delete a node
- Fix xas_reload() when dealing with entries that occupy multiple indices
- Add a few more tests to the test suite
- Fix an unsigned int being shifted into an unsigned long
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Merge tag 'xarray-5.9' of git://git.infradead.org/users/willy/xarray
Pull XArray updates from Matthew Wilcox:
- Fix the test suite after introduction of the local_lock
- Fix a bug in the IDA spotted by Coverity
- Change the API that allows the workingset code to delete a node
- Fix xas_reload() when dealing with entries that occupy multiple
indices
- Add a few more tests to the test suite
- Fix an unsigned int being shifted into an unsigned long
* tag 'xarray-5.9' of git://git.infradead.org/users/willy/xarray:
XArray: Fix xas_create_range for ranges above 4 billion
radix-tree: fix the comment of radix_tree_next_slot()
XArray: Fix xas_reload for multi-index entries
XArray: Add private interface for workingset node deletion
XArray: Fix xas_for_each_conflict documentation
XArray: Test marked multiorder iterations
XArray: Test two more things about xa_cmpxchg
ida: Free allocated bitmap in error path
radix tree test suite: Fix compilation
Fix following warnings caused by mismatch bewteen function parameters and
comments.
mm/workingset.c:228: warning: Function parameter or member 'lruvec' not described in 'workingset_age_nonresident'
mm/workingset.c:228: warning: Excess function parameter 'memcg' description in 'workingset_age_nonresident'
Signed-off-by: Xiaofei Tan <tanxiaofei@huawei.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Link: https://lkml.kernel.org/r/1600485913-11192-1-git-send-email-tanxiaofei@huawei.com
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Move the tricky bits of dealing with the XArray from the workingset
code to the XArray. Make it clear in the documentation that this is a
private interface, and only export it for the benefit of the test suite.
Signed-off-by: Matthew Wilcox (Oracle) <willy@infradead.org>
The thp prefix is more frequently used than hpage and we should be
consistent between the various functions.
[akpm@linux-foundation.org: fix mm/migrate.c]
Signed-off-by: Matthew Wilcox (Oracle) <willy@infradead.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Reviewed-by: William Kucharski <william.kucharski@oracle.com>
Reviewed-by: Zi Yan <ziy@nvidia.com>
Cc: Mike Kravetz <mike.kravetz@oracle.com>
Cc: David Hildenbrand <david@redhat.com>
Cc: "Kirill A. Shutemov" <kirill.shutemov@linux.intel.com>
Link: http://lkml.kernel.org/r/20200629151959.15779-6-willy@infradead.org
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This patch implements workingset detection for anonymous LRU. All the
infrastructure is implemented by the previous patches so this patch just
activates the workingset detection by installing/retrieving the shadow
entry and adding refault calculation.
Signed-off-by: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Cc: Hugh Dickins <hughd@google.com>
Cc: Matthew Wilcox <willy@infradead.org>
Cc: Mel Gorman <mgorman@techsingularity.net>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: Minchan Kim <minchan@kernel.org>
Link: http://lkml.kernel.org/r/1595490560-15117-6-git-send-email-iamjoonsoo.kim@lge.com
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
To prepare the workingset detection for anon LRU, this patch splits
workingset event counters for refault, activate and restore into anon and
file variants, as well as the refaults counter in struct lruvec.
Signed-off-by: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Cc: Hugh Dickins <hughd@google.com>
Cc: Matthew Wilcox <willy@infradead.org>
Cc: Mel Gorman <mgorman@techsingularity.net>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: Minchan Kim <minchan@kernel.org>
Link: http://lkml.kernel.org/r/1595490560-15117-4-git-send-email-iamjoonsoo.kim@lge.com
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
In order to prepare for per-object slab memory accounting, convert
NR_SLAB_RECLAIMABLE and NR_SLAB_UNRECLAIMABLE vmstat items to bytes.
To make it obvious, rename them to NR_SLAB_RECLAIMABLE_B and
NR_SLAB_UNRECLAIMABLE_B (similar to NR_KERNEL_STACK_KB).
Internally global and per-node counters are stored in pages, however memcg
and lruvec counters are stored in bytes. This scheme may look weird, but
only for now. As soon as slab pages will be shared between multiple
cgroups, global and node counters will reflect the total number of slab
pages. However memcg and lruvec counters will be used for per-memcg slab
memory tracking, which will take separate kernel objects in the account.
Keeping global and node counters in pages helps to avoid additional
overhead.
The size of slab memory shouldn't exceed 4Gb on 32-bit machines, so it
will fit into atomic_long_t we use for vmstats.
Signed-off-by: Roman Gushchin <guro@fb.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Reviewed-by: Shakeel Butt <shakeelb@google.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Cc: Christoph Lameter <cl@linux.com>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: Tejun Heo <tj@kernel.org>
Link: http://lkml.kernel.org/r/20200623174037.3951353-4-guro@fb.com
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Patch series "fix for "mm: balance LRU lists based on relative
thrashing" patchset"
This patchset fixes some problems of the patchset, "mm: balance LRU
lists based on relative thrashing", which is now merged on the mainline.
Patch "mm: workingset: let cache workingset challenge anon fix" is the
result of discussion with Johannes. See following link.
http://lkml.kernel.org/r/20200520232525.798933-6-hannes@cmpxchg.org
And, the other two are minor things which are found when I try to rebase
my patchset.
This patch (of 3):
After ("mm: workingset: let cache workingset challenge anon fix"), we
compare refault distances to active_file + anon. But age of the
non-resident information is only driven by the file LRU. As a result,
we may overestimate the recency of any incoming refaults and activate
them too eagerly, causing unnecessary LRU churn in certain situations.
Make anon aging drive nonresident age as well to address that.
Link: http://lkml.kernel.org/r/1592288204-27734-1-git-send-email-iamjoonsoo.kim@lge.com
Link: http://lkml.kernel.org/r/1592288204-27734-2-git-send-email-iamjoonsoo.kim@lge.com
Fixes: 34e58cac6d ("mm: workingset: let cache workingset challenge anon")
Reported-by: Joonsoo Kim <js1304@gmail.com>
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Signed-off-by: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Cc: Rik van Riel <riel@surriel.com>
Cc: Minchan Kim <minchan.kim@gmail.com>
Cc: Michal Hocko <mhocko@suse.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The VM tries to balance reclaim pressure between anon and file so as to
reduce the amount of IO incurred due to the memory shortage. It already
counts refaults and swapins, but in addition it should also count
writepage calls during reclaim.
For swap, this is obvious: it's IO that wouldn't have occurred if the
anonymous memory hadn't been under memory pressure. From a relative
balancing point of view this makes sense as well: even if anon is cold and
reclaimable, a cache that isn't thrashing may have equally cold pages that
don't require IO to reclaim.
For file writeback, it's trickier: some of the reclaim writepage IO would
have likely occurred anyway due to dirty expiration. But not all of it -
premature writeback reduces batching and generates additional writes.
Since the flushers are already woken up by the time the VM starts writing
cache pages one by one, let's assume that we'e likely causing writes that
wouldn't have happened without memory pressure. In addition, the per-page
cost of IO would have probably been much cheaper if written in larger
batches from the flusher thread rather than the single-page-writes from
kswapd.
For our purposes - getting the trend right to accelerate convergence on a
stable state that doesn't require paging at all - this is sufficiently
accurate. If we later wanted to optimize for sustained thrashing, we can
still refine the measurements.
Count all writepage calls from kswapd as IO cost toward the LRU that the
page belongs to.
Why do this dynamically? Don't we know in advance that anon pages require
IO to reclaim, and so could build in a static bias?
First, scanning is not the same as reclaiming. If all the anon pages are
referenced, we may not swap for a while just because we're scanning the
anon list. During this time, however, it's important that we age
anonymous memory and the page cache at the same rate so that their
hot-cold gradients are comparable. Everything else being equal, we still
want to reclaim the coldest memory overall.
Second, we keep copies in swap unless the page changes. If there is
swap-backed data that's mostly read (tmpfs file) and has been swapped out
before, we can reclaim it without incurring additional IO.
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Cc: Michal Hocko <mhocko@suse.com>
Cc: Minchan Kim <minchan@kernel.org>
Cc: Rik van Riel <riel@surriel.com>
Link: http://lkml.kernel.org/r/20200520232525.798933-14-hannes@cmpxchg.org
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Since the LRUs were split into anon and file lists, the VM has been
balancing between page cache and anonymous pages based on per-list ratios
of scanned vs. rotated pages. In most cases that tips page reclaim
towards the list that is easier to reclaim and has the fewest actively
used pages, but there are a few problems with it:
1. Refaults and LRU rotations are weighted the same way, even though
one costs IO and the other costs a bit of CPU.
2. The less we scan an LRU list based on already observed rotations,
the more we increase the sampling interval for new references, and
rotations become even more likely on that list. This can enter a
death spiral in which we stop looking at one list completely until
the other one is all but annihilated by page reclaim.
Since commit a528910e12 ("mm: thrash detection-based file cache sizing")
we have refault detection for the page cache. Along with swapin events,
they are good indicators of when the file or anon list, respectively, is
too small for its workingset and needs to grow.
For example, if the page cache is thrashing, the cache pages need more
time in memory, while there may be colder pages on the anonymous list.
Likewise, if swapped pages are faulting back in, it indicates that we
reclaim anonymous pages too aggressively and should back off.
Replace LRU rotations with refaults and swapins as the basis for relative
reclaim cost of the two LRUs. This will have the VM target list balances
that incur the least amount of IO on aggregate.
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Cc: Michal Hocko <mhocko@suse.com>
Cc: Minchan Kim <minchan@kernel.org>
Cc: Rik van Riel <riel@surriel.com>
Link: http://lkml.kernel.org/r/20200520232525.798933-12-hannes@cmpxchg.org
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
We activate cache refaults with reuse distances in pages smaller than the
size of the total cache. This allows new pages with competitive access
frequencies to establish themselves, as well as challenge and potentially
displace pages on the active list that have gone cold.
However, that assumes that active cache can only replace other active
cache in a competition for the hottest memory. This is not a great
default assumption. The page cache might be thrashing while there are
enough completely cold and unused anonymous pages sitting around that we'd
only have to write to swap once to stop all IO from the cache.
Activate cache refaults when their reuse distance in pages is smaller than
the total userspace workingset, including anonymous pages.
Reclaim can still decide how to balance pressure among the two LRUs
depending on the IO situation. Rotational drives will prefer avoiding
random IO from swap and go harder after cache. But fundamentally, hot
cache should be able to compete with anon pages for a place in RAM.
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Cc: Michal Hocko <mhocko@suse.com>
Cc: Minchan Kim <minchan@kernel.org>
Cc: Rik van Riel <riel@surriel.com>
Link: http://lkml.kernel.org/r/20200520232525.798933-6-hannes@cmpxchg.org
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
We use refault information to determine whether the cache workingset is
stable or transitioning, and dynamically adjust the inactive:active file
LRU ratio so as to maximize protection from one-off cache during stable
periods, and minimize IO during transitions.
With cgroups and their nested LRU lists, we currently don't do this
correctly. While recursive cgroup reclaim establishes a relative LRU
order among the pages of all involved cgroups, refaults only affect the
local LRU order in the cgroup in which they are occuring. As a result,
cache transitions can take longer in a cgrouped system as the active pages
of sibling cgroups aren't challenged when they should be.
[ Right now, this is somewhat theoretical, because the siblings, under
continued regular reclaim pressure, should eventually run out of
inactive pages - and since inactive:active *size* balancing is also
done on a cgroup-local level, we will challenge the active pages
eventually in most cases. But the next patch will move that relative
size enforcement to the reclaim root as well, and then this patch
here will be necessary to propagate refault pressure to siblings. ]
This patch moves refault detection to the root of reclaim. Instead of
remembering the cgroup owner of an evicted page, remember the cgroup that
caused the reclaim to happen. When refaults later occur, they'll
correctly influence the cross-cgroup LRU order that reclaim follows.
I.e. if global reclaim kicked out pages in some subgroup A/B/C, the
refault of those pages will challenge the global LRU order, and not just
the local order down inside C.
[hannes@cmpxchg.org: use page_memcg() instead of another lookup]
Link: http://lkml.kernel.org/r/20191115160722.GA309754@cmpxchg.org
Link: http://lkml.kernel.org/r/20191107205334.158354-3-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Reviewed-by: Suren Baghdasaryan <surenb@google.com>
Cc: Andrey Ryabinin <aryabinin@virtuozzo.com>
Cc: Michal Hocko <mhocko@suse.com>
Cc: Rik van Riel <riel@surriel.com>
Cc: Shakeel Butt <shakeelb@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
There is a per-memcg lruvec and a NUMA node lruvec. Which one is being
used is somewhat confusing right now, and it's easy to make mistakes -
especially when it comes to global reclaim.
How it works: when memory cgroups are enabled, we always use the
root_mem_cgroup's per-node lruvecs. When memory cgroups are not compiled
in or disabled at runtime, we use pgdat->lruvec.
Document that in a comment.
Due to the way the reclaim code is generalized, all lookups use the
mem_cgroup_lruvec() helper function, and nobody should have to find the
right lruvec manually right now. But to avoid future mistakes, rename the
pgdat->lruvec member to pgdat->__lruvec and delete the convenience wrapper
that suggests it's a commonly accessed member.
While in this area, swap the mem_cgroup_lruvec() argument order. The name
suggests a memcg operation, yet it takes a pgdat first and a memcg second.
I have to double take every time I call this. Fix that.
Link: http://lkml.kernel.org/r/20191022144803.302233-3-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Michal Hocko <mhocko@suse.com>
Reviewed-by: Shakeel Butt <shakeelb@google.com>
Cc: Roman Gushchin <guro@fb.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Memcg counters for shadow nodes are broken because the memcg pointer is
obtained in a wrong way. The following approach is used:
virt_to_page(xa_node)->mem_cgroup
Since commit 4d96ba3530 ("mm: memcg/slab: stop setting
page->mem_cgroup pointer for slab pages") page->mem_cgroup pointer isn't
set for slab pages, so memcg_from_slab_page() should be used instead.
Also I doubt that it ever worked correctly: virt_to_head_page() should
be used instead of virt_to_page(). Otherwise objects residing on tail
pages are not accounted, because only the head page contains a valid
mem_cgroup pointer. That was a case since the introduction of these
counters by the commit 68d48e6a2d ("mm: workingset: add vmstat counter
for shadow nodes").
Link: http://lkml.kernel.org/r/20190801233532.138743-1-guro@fb.com
Fixes: 4d96ba3530 ("mm: memcg/slab: stop setting page->mem_cgroup pointer for slab pages")
Signed-off-by: Roman Gushchin <guro@fb.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Cc: Vladimir Davydov <vdavydov.dev@gmail.com>
Cc: Shakeel Butt <shakeelb@google.com>
Cc: Michal Hocko <mhocko@suse.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Patch series "mm: memcontrol: memory.stat cost & correctness".
The cgroup memory.stat file holds recursive statistics for the entire
subtree. The current implementation does this tree walk on-demand
whenever the file is read. This is giving us problems in production.
1. The cost of aggregating the statistics on-demand is high. A lot of
system service cgroups are mostly idle and their stats don't change
between reads, yet we always have to check them. There are also always
some lazily-dying cgroups sitting around that are pinned by a handful
of remaining page cache; the same applies to them.
In an application that periodically monitors memory.stat in our
fleet, we have seen the aggregation consume up to 5% CPU time.
2. When cgroups die and disappear from the cgroup tree, so do their
accumulated vm events. The result is that the event counters at
higher-level cgroups can go backwards and confuse some of our
automation, let alone people looking at the graphs over time.
To address both issues, this patch series changes the stat
implementation to spill counts upwards when the counters change.
The upward spilling is batched using the existing per-cpu cache. In a
sparse file stress test with 5 level cgroup nesting, the additional cost
of the flushing was negligible (a little under 1% of CPU at 100% CPU
utilization, compared to the 5% of reading memory.stat during regular
operation).
This patch (of 4):
memcg_page_state(), lruvec_page_state(), memcg_sum_events() are
currently returning the state of the local memcg or lruvec, not the
recursive state.
In practice there is a demand for both versions, although the callers
that want the recursive counts currently sum them up by hand.
Per default, cgroups are considered recursive entities and generally we
expect more users of the recursive counters, with the local counts being
special cases. To reflect that in the name, add a _local suffix to the
current implementations.
The following patch will re-incarnate these functions with recursive
semantics, but with an O(1) implementation.
[hannes@cmpxchg.org: fix bisection hole]
Link: http://lkml.kernel.org/r/20190417160347.GC23013@cmpxchg.org
Link: http://lkml.kernel.org/r/20190412151507.2769-2-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Reviewed-by: Shakeel Butt <shakeelb@google.com>
Reviewed-by: Roman Gushchin <guro@fb.com>
Cc: Michal Hocko <mhocko@kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
mem_cgroup_node_nr_lru_pages() is just a convenience wrapper around
lruvec_page_state() that takes bitmasks of lru indexes and aggregates the
counts for those.
Replace callsites where the bitmask is simple enough with direct
lruvec_page_state() calls.
This removes the last extern user of mem_cgroup_node_nr_lru_pages(), so
make that function private again, too.
Link: http://lkml.kernel.org/r/20190228163020.24100-5-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Reviewed-by: Roman Gushchin <guro@fb.com>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: Tejun Heo <tj@kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
workingset_eviction() doesn't use and never did use the @mapping
argument. Remove it.
Link: http://lkml.kernel.org/r/20190228083329.31892-1-aryabinin@virtuozzo.com
Signed-off-by: Andrey Ryabinin <aryabinin@virtuozzo.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Rik van Riel <riel@surriel.com>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Acked-by: Mel Gorman <mgorman@techsingularity.net>
Cc: Michal Hocko <mhocko@kernel.org>
Cc: William Kucharski <william.kucharski@oracle.com>
Cc: John Hubbard <jhubbard@nvidia.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
totalram_pages and totalhigh_pages are made static inline function.
Main motivation was that managed_page_count_lock handling was complicating
things. It was discussed in length here,
https://lore.kernel.org/patchwork/patch/995739/#1181785 So it seemes
better to remove the lock and convert variables to atomic, with preventing
poteintial store-to-read tearing as a bonus.
[akpm@linux-foundation.org: coding style fixes]
Link: http://lkml.kernel.org/r/1542090790-21750-4-git-send-email-arunks@codeaurora.org
Signed-off-by: Arun KS <arunks@codeaurora.org>
Suggested-by: Michal Hocko <mhocko@suse.com>
Suggested-by: Vlastimil Babka <vbabka@suse.cz>
Reviewed-by: Konstantin Khlebnikov <khlebnikov@yandex-team.ru>
Reviewed-by: Pavel Tatashin <pasha.tatashin@soleen.com>
Acked-by: Michal Hocko <mhocko@suse.com>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Cc: David Hildenbrand <david@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Pull XArray conversion from Matthew Wilcox:
"The XArray provides an improved interface to the radix tree data
structure, providing locking as part of the API, specifying GFP flags
at allocation time, eliminating preloading, less re-walking the tree,
more efficient iterations and not exposing RCU-protected pointers to
its users.
This patch set
1. Introduces the XArray implementation
2. Converts the pagecache to use it
3. Converts memremap to use it
The page cache is the most complex and important user of the radix
tree, so converting it was most important. Converting the memremap
code removes the only other user of the multiorder code, which allows
us to remove the radix tree code that supported it.
I have 40+ followup patches to convert many other users of the radix
tree over to the XArray, but I'd like to get this part in first. The
other conversions haven't been in linux-next and aren't suitable for
applying yet, but you can see them in the xarray-conv branch if you're
interested"
* 'xarray' of git://git.infradead.org/users/willy/linux-dax: (90 commits)
radix tree: Remove multiorder support
radix tree test: Convert multiorder tests to XArray
radix tree tests: Convert item_delete_rcu to XArray
radix tree tests: Convert item_kill_tree to XArray
radix tree tests: Move item_insert_order
radix tree test suite: Remove multiorder benchmarking
radix tree test suite: Remove __item_insert
memremap: Convert to XArray
xarray: Add range store functionality
xarray: Move multiorder_check to in-kernel tests
xarray: Move multiorder_shrink to kernel tests
xarray: Move multiorder account test in-kernel
radix tree test suite: Convert iteration test to XArray
radix tree test suite: Convert tag_tagged_items to XArray
radix tree: Remove radix_tree_clear_tags
radix tree: Remove radix_tree_maybe_preload_order
radix tree: Remove split/join code
radix tree: Remove radix_tree_update_node_t
page cache: Finish XArray conversion
dax: Convert page fault handlers to XArray
...
The page cache and most shrinkable slab caches hold data that has been
read from disk, but there are some caches that only cache CPU work, such
as the dentry and inode caches of procfs and sysfs, as well as the subset
of radix tree nodes that track non-resident page cache.
Currently, all these are shrunk at the same rate: using DEFAULT_SEEKS for
the shrinker's seeks setting tells the reclaim algorithm that for every
two page cache pages scanned it should scan one slab object.
This is a bogus setting. A virtual inode that required no IO to create is
not twice as valuable as a page cache page; shadow cache entries with
eviction distances beyond the size of memory aren't either.
In most cases, the behavior in practice is still fine. Such virtual
caches don't tend to grow and assert themselves aggressively, and usually
get picked up before they cause problems. But there are scenarios where
that's not true.
Our database workloads suffer from two of those. For one, their file
workingset is several times bigger than available memory, which has the
kernel aggressively create shadow page cache entries for the non-resident
parts of it. The workingset code does tell the VM that most of these are
expendable, but the VM ends up balancing them 2:1 to cache pages as per
the seeks setting. This is a huge waste of memory.
These workloads also deal with tens of thousands of open files and use
/proc for introspection, which ends up growing the proc_inode_cache to
absurdly large sizes - again at the cost of valuable cache space, which
isn't a reasonable trade-off, given that proc inodes can be re-created
without involving the disk.
This patch implements a "zero-seek" setting for shrinkers that results in
a target ratio of 0:1 between their objects and IO-backed caches. This
allows such virtual caches to grow when memory is available (they do
cache/avoid CPU work after all), but effectively disables them as soon as
IO-backed objects are under pressure.
It then switches the shrinkers for procfs and sysfs metadata, as well as
excess page cache shadow nodes, to the new zero-seek setting.
Link: http://lkml.kernel.org/r/20181009184732.762-5-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Reported-by: Domas Mituzas <dmituzas@fb.com>
Reviewed-by: Andrew Morton <akpm@linux-foundation.org>
Reviewed-by: Rik van Riel <riel@surriel.com>
Acked-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Make it easier to catch bugs in the shadow node shrinker by adding a
counter for the shadow nodes in circulation.
[akpm@linux-foundation.org: assert that irqs are disabled, for __inc_lruvec_page_state()]
[akpm@linux-foundation.org: s/WARN_ON_ONCE/VM_WARN_ON_ONCE/, per Johannes]
Link: http://lkml.kernel.org/r/20181009184732.762-4-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Reviewed-by: Andrew Morton <akpm@linux-foundation.org>
Acked-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Rik van Riel <riel@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
No need to use the preemption-safe lruvec state function inside the
reclaim region that has irqs disabled.
Link: http://lkml.kernel.org/r/20181009184732.762-3-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Reviewed-by: Andrew Morton <akpm@linux-foundation.org>
Reviewed-by: Rik van Riel <riel@surriel.com>
Acked-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Refaults happen during transitions between workingsets as well as in-place
thrashing. Knowing the difference between the two has a range of
applications, including measuring the impact of memory shortage on the
system performance, as well as the ability to smarter balance pressure
between the filesystem cache and the swap-backed workingset.
During workingset transitions, inactive cache refaults and pushes out
established active cache. When that active cache isn't stale, however,
and also ends up refaulting, that's bonafide thrashing.
Introduce a new page flag that tells on eviction whether the page has been
active or not in its lifetime. This bit is then stored in the shadow
entry, to classify refaults as transitioning or thrashing.
How many page->flags does this leave us with on 32-bit?
20 bits are always page flags
21 if you have an MMU
23 with the zone bits for DMA, Normal, HighMem, Movable
29 with the sparsemem section bits
30 if PAE is enabled
31 with this patch.
So on 32-bit PAE, that leaves 1 bit for distinguishing two NUMA nodes. If
that's not enough, the system can switch to discontigmem and re-gain the 6
or 7 sparsemem section bits.
Link: http://lkml.kernel.org/r/20180828172258.3185-3-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Tested-by: Daniel Drake <drake@endlessm.com>
Tested-by: Suren Baghdasaryan <surenb@google.com>
Cc: Christopher Lameter <cl@linux.com>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: Johannes Weiner <jweiner@fb.com>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Enderborg <peter.enderborg@sony.com>
Cc: Randy Dunlap <rdunlap@infradead.org>
Cc: Shakeel Butt <shakeelb@google.com>
Cc: Tejun Heo <tj@kernel.org>
Cc: Vinayak Menon <vinmenon@codeaurora.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Patch series "psi: pressure stall information for CPU, memory, and IO", v4.
Overview
PSI reports the overall wallclock time in which the tasks in a system (or
cgroup) wait for (contended) hardware resources.
This helps users understand the resource pressure their workloads are
under, which allows them to rootcause and fix throughput and latency
problems caused by overcommitting, underprovisioning, suboptimal job
placement in a grid; as well as anticipate major disruptions like OOM.
Real-world applications
We're using the data collected by PSI (and its previous incarnation,
memdelay) quite extensively at Facebook, and with several success stories.
One usecase is avoiding OOM hangs/livelocks. The reason these happen is
because the OOM killer is triggered by reclaim not being able to free
pages, but with fast flash devices there is *always* some clean and
uptodate cache to reclaim; the OOM killer never kicks in, even as tasks
spend 90% of the time thrashing the cache pages of their own executables.
There is no situation where this ever makes sense in practice. We wrote a
<100 line POC python script to monitor memory pressure and kill stuff way
before such pathological thrashing leads to full system losses that would
require forcible hard resets.
We've since extended and deployed this code into other places to guarantee
latency and throughput SLAs, since they're usually violated way before the
kernel OOM killer would ever kick in.
It is available here: https://github.com/facebookincubator/oomd
Eventually we probably want to trigger the in-kernel OOM killer based on
extreme sustained pressure as well, so that Linux can avoid memory
livelocks - which technically aren't deadlocks, but to the user
indistinguishable from them - out of the box. We'd continue using OOMD as
the first line of defense to ensure workload health and implement complex
kill policies that are beyond the scope of the kernel.
We also use PSI memory pressure for loadshedding. Our batch job
infrastructure used to use heuristics based on various VM stats to
anticipate OOM situations, with lackluster success. We switched it to PSI
and managed to anticipate and avoid OOM kills and lockups fairly reliably.
The reduction of OOM outages in the worker pool raised the pool's
aggregate productivity, and we were able to switch that service to smaller
machines.
Lastly, we use cgroups to isolate a machine's main workload from
maintenance crap like package upgrades, logging, configuration, as well as
to prevent multiple workloads on a machine from stepping on each others'
toes. We were not able to configure this properly without the pressure
metrics; we would see latency or bandwidth drops, but it would often be
hard to impossible to rootcause it post-mortem.
We now log and graph pressure for the containers in our fleet and can
trivially link latency spikes and throughput drops to shortages of
specific resources after the fact, and fix the job config/scheduling.
PSI has also received testing, feedback, and feature requests from Android
and EndlessOS for the purpose of low-latency OOM killing, to intervene in
pressure situations before the UI starts hanging.
How do you use this feature?
A kernel with CONFIG_PSI=y will create a /proc/pressure directory with 3
files: cpu, memory, and io. If using cgroup2, cgroups will also have
cpu.pressure, memory.pressure and io.pressure files, which simply
aggregate task stalls at the cgroup level instead of system-wide.
The cpu file contains one line:
some avg10=2.04 avg60=0.75 avg300=0.40 total=157656722
The averages give the percentage of walltime in which one or more tasks
are delayed on the runqueue while another task has the CPU. They're
recent averages over 10s, 1m, 5m windows, so you can tell short term
trends from long term ones, similarly to the load average.
The total= value gives the absolute stall time in microseconds. This
allows detecting latency spikes that might be too short to sway the
running averages. It also allows custom time averaging in case the
10s/1m/5m windows aren't adequate for the usecase (or are too coarse with
future hardware).
What to make of this "some" metric? If CPU utilization is at 100% and CPU
pressure is 0, it means the system is perfectly utilized, with one
runnable thread per CPU and nobody waiting. At two or more runnable tasks
per CPU, the system is 100% overcommitted and the pressure average will
indicate as much. From a utilization perspective this is a great state of
course: no CPU cycles are being wasted, even when 50% of the threads were
to go idle (as most workloads do vary). From the perspective of the
individual job it's not great, however, and they would do better with more
resources. Depending on what your priority and options are, raised "some"
numbers may or may not require action.
The memory file contains two lines:
some avg10=70.24 avg60=68.52 avg300=69.91 total=3559632828
full avg10=57.59 avg60=58.06 avg300=60.38 total=3300487258
The some line is the same as for cpu, the time in which at least one task
is stalled on the resource. In the case of memory, this includes waiting
on swap-in, page cache refaults and page reclaim.
The full line, however, indicates time in which *nobody* is using the CPU
productively due to pressure: all non-idle tasks are waiting for memory in
one form or another. Significant time spent in there is a good trigger
for killing things, moving jobs to other machines, or dropping incoming
requests, since neither the jobs nor the machine overall are making too
much headway.
The io file is similar to memory. Because the block layer doesn't have a
concept of hardware contention right now (how much longer is my IO request
taking due to other tasks?), it reports CPU potential lost on all IO
delays, not just the potential lost due to competition.
FAQ
Q: How is PSI's CPU component different from the load average?
A: There are several quirks in the load average that make it hard to
impossible to tell how overcommitted the CPU really is.
1. The load average is reported as a raw number of active tasks.
You need to know how many CPUs there are in the system, how many
CPUs the workload is allowed to use, then think about what the
proportion between load and the number of CPUs mean for the
tasks trying to run.
PSI reports the percentage of wallclock time in which tasks are
waiting for a CPU to run on. It doesn't matter how many CPUs are
present or usable. The number always tells the quality of life
of tasks in the system or in a particular cgroup.
2. The shortest averaging window is 1m, which is extremely coarse,
and it's sampled in 5s intervals. A *lot* can happen on a CPU in
5 seconds. This *may* be able to identify persistent long-term
trends and very clear and obvious overloads, but it's unusable
for latency spikes and more subtle overutilization.
PSI's shortest window is 10s. It also exports the cumulative
stall times (in microseconds) of synchronously recorded events.
3. On Linux, the load average for historical reasons includes all
TASK_UNINTERRUPTIBLE tasks. This gives a broader sense of how
busy the system is, but on the flipside it doesn't distinguish
whether tasks are likely to contend over the CPU or IO - which
obviously requires very different interventions from a sys admin
or a job scheduler.
PSI reports independent metrics for CPU and IO. You can tell
which resource is making the tasks wait, but in conjunction
still see how overloaded the system is overall.
Q: What's the cost / performance impact of this feature?
A: PSI's primary cost is in the scheduler, in particular task wakeups
and sleeps.
I benchmarked this code using Facebook's two most scheduling
sensitive workloads: memcache and webserver. They handle a ton of
small requests - lots of wakeups and sleeps with little actual work
in between - so they tend to be canaries for scheduler regressions.
In the tests, the boxes were handling live traffic over the course
of several hours. Half the machines, the control, ran with
CONFIG_PSI=n.
For memcache I used eight machines total. They're 2-socket, 14
core, 56 thread boxes. The test runs for half the test period,
flips the test and control kernels on the hardware to rule out HW
factors, DC location etc., then runs the other half of the test.
For the webservers, I used 32 machines total. They're single
socket, 16 core, 32 thread machines.
During the memcache test, CPU load was nopsi=78.05% psi=78.98% in
the first half and nopsi=77.52% psi=78.25%, so PSI added between
0.7 and 0.9 percentage points to the CPU load, a difference of
about 1%.
UPDATE: I re-ran this test with the v3 version of this patch set
and the CPU utilization was equivalent between test and control.
UPDATE: v4 is on par with v3.
As far as end-to-end request latency from the client perspective
goes, we don't sample those finely enough to capture the requests
going to those particular machines during the test, but we know the
p50 turnaround time in this workload is 54us, and perf bench sched
pipe on those machines show nopsi=5.232666 us/op and psi=5.587347
us/op, so this doesn't add much here either.
The profile for the pipe benchmark shows:
0.87% sched-pipe [kernel.vmlinux] [k] psi_group_change
0.83% perf.real [kernel.vmlinux] [k] psi_group_change
0.82% perf.real [kernel.vmlinux] [k] psi_task_change
0.58% sched-pipe [kernel.vmlinux] [k] psi_task_change
The webserver load is running inside 4 nested cgroup levels. The
CPU load with both nopsi and psi kernels was indistinguishable at
81%.
For comparison, we had to disable the cgroup cpu controller on the
webservers because it added 4 percentage points to the CPU% during
this same exact test.
Versions of this accounting code now run on 80% of our fleet. None
of our workloads have reported regressions during the rollout.
Daniel Drake said:
: I just retested the latest version at
: http://git.cmpxchg.org/cgit.cgi/linux-psi.git (Linux 4.18) and the results
: are great.
:
: Test setup:
: Endless OS
: GeminiLake N4200 low end laptop
: 2GB RAM
: swap (and zram swap) disabled
:
: Baseline test: open a handful of large-ish apps and several website
: tabs in Google Chrome.
:
: Results: after a couple of minutes, system is excessively thrashing, mouse
: cursor can barely be moved, UI is not responding to mouse clicks, so it's
: impractical to recover from this situation as an ordinary user
:
: Add my simple killer:
: https://gist.github.com/dsd/a8988bf0b81a6163475988120fe8d9cd
:
: Results: when the thrashing causes the UI to become sluggish, the killer
: steps in and kills something (usually a chrome tab), and the system
: remains usable. I repeatedly opened more apps and more websites over a 15
: minute period but I wasn't able to get the system to a point of UI
: unresponsiveness.
Suren said:
: Backported to 4.9 and retested on ARMv8 8 code system running Android.
: Signals behave as expected reacting to memory pressure, no jumps in
: "total" counters that would indicate an overflow/underflow issues. Nicely
: done!
This patch (of 9):
If we keep just enough refault information to match the *current* page
cache during reclaim time, we could lose a lot of events when there is
only a temporary spike in non-cache memory consumption that pushes out all
the cache. Once cache comes back, we won't see those refaults. They
might not be actionable for LRU aging, but we want to know about them for
measuring memory pressure.
[hannes@cmpxchg.org: switch to NUMA-aware lru and slab counters]
Link: http://lkml.kernel.org/r/20181009184732.762-2-hannes@cmpxchg.org
Link: http://lkml.kernel.org/r/20180828172258.3185-2-hannes@cmpxchg.org
Signed-off-by: Johannes Weiner <jweiner@fb.com>
Acked-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Reviewed-by: Rik van Riel <riel@surriel.com>
Tested-by: Daniel Drake <drake@endlessm.com>
Tested-by: Suren Baghdasaryan <surenb@google.com>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: Tejun Heo <tj@kernel.org>
Cc: Vinayak Menon <vinmenon@codeaurora.org>
Cc: Christopher Lameter <cl@linux.com>
Cc: Peter Enderborg <peter.enderborg@sony.com>
Cc: Shakeel Butt <shakeelb@google.com>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Randy Dunlap <rdunlap@infradead.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
We construct an XA_STATE and use it to delete the node with
xas_store() rather than adding a special function for this unique
use case. Includes a test that simulates this usage for the
test suite.
Signed-off-by: Matthew Wilcox <willy@infradead.org>
Yu Zhao