linux-stable/kernel/sys_ni.c

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License cleanup: add SPDX GPL-2.0 license identifier to files with no license Many source files in the tree are missing licensing information, which makes it harder for compliance tools to determine the correct license. By default all files without license information are under the default license of the kernel, which is GPL version 2. Update the files which contain no license information with the 'GPL-2.0' SPDX license identifier. The SPDX identifier is a legally binding shorthand, which can be used instead of the full boiler plate text. This patch is based on work done by Thomas Gleixner and Kate Stewart and Philippe Ombredanne. How this work was done: Patches were generated and checked against linux-4.14-rc6 for a subset of the use cases: - file had no licensing information it it. - file was a */uapi/* one with no licensing information in it, - file was a */uapi/* one with existing licensing information, Further patches will be generated in subsequent months to fix up cases where non-standard license headers were used, and references to license had to be inferred by heuristics based on keywords. The analysis to determine which SPDX License Identifier to be applied to a file was done in a spreadsheet of side by side results from of the output of two independent scanners (ScanCode & Windriver) producing SPDX tag:value files created by Philippe Ombredanne. Philippe prepared the base worksheet, and did an initial spot review of a few 1000 files. The 4.13 kernel was the starting point of the analysis with 60,537 files assessed. Kate Stewart did a file by file comparison of the scanner results in the spreadsheet to determine which SPDX license identifier(s) to be applied to the file. She confirmed any determination that was not immediately clear with lawyers working with the Linux Foundation. Criteria used to select files for SPDX license identifier tagging was: - Files considered eligible had to be source code files. - Make and config files were included as candidates if they contained >5 lines of source - File already had some variant of a license header in it (even if <5 lines). All documentation files were explicitly excluded. The following heuristics were used to determine which SPDX license identifiers to apply. - when both scanners couldn't find any license traces, file was considered to have no license information in it, and the top level COPYING file license applied. For non */uapi/* files that summary was: SPDX license identifier # files ---------------------------------------------------|------- GPL-2.0 11139 and resulted in the first patch in this series. If that file was a */uapi/* path one, it was "GPL-2.0 WITH Linux-syscall-note" otherwise it was "GPL-2.0". Results of that was: SPDX license identifier # files ---------------------------------------------------|------- GPL-2.0 WITH Linux-syscall-note 930 and resulted in the second patch in this series. - if a file had some form of licensing information in it, and was one of the */uapi/* ones, it was denoted with the Linux-syscall-note if any GPL family license was found in the file or had no licensing in it (per prior point). Results summary: SPDX license identifier # files ---------------------------------------------------|------ GPL-2.0 WITH Linux-syscall-note 270 GPL-2.0+ WITH Linux-syscall-note 169 ((GPL-2.0 WITH Linux-syscall-note) OR BSD-2-Clause) 21 ((GPL-2.0 WITH Linux-syscall-note) OR BSD-3-Clause) 17 LGPL-2.1+ WITH Linux-syscall-note 15 GPL-1.0+ WITH Linux-syscall-note 14 ((GPL-2.0+ WITH Linux-syscall-note) OR BSD-3-Clause) 5 LGPL-2.0+ WITH Linux-syscall-note 4 LGPL-2.1 WITH Linux-syscall-note 3 ((GPL-2.0 WITH Linux-syscall-note) OR MIT) 3 ((GPL-2.0 WITH Linux-syscall-note) AND MIT) 1 and that resulted in the third patch in this series. - when the two scanners agreed on the detected license(s), that became the concluded license(s). - when there was disagreement between the two scanners (one detected a license but the other didn't, or they both detected different licenses) a manual inspection of the file occurred. - In most cases a manual inspection of the information in the file resulted in a clear resolution of the license that should apply (and which scanner probably needed to revisit its heuristics). - When it was not immediately clear, the license identifier was confirmed with lawyers working with the Linux Foundation. - If there was any question as to the appropriate license identifier, the file was flagged for further research and to be revisited later in time. In total, over 70 hours of logged manual review was done on the spreadsheet to determine the SPDX license identifiers to apply to the source files by Kate, Philippe, Thomas and, in some cases, confirmation by lawyers working with the Linux Foundation. Kate also obtained a third independent scan of the 4.13 code base from FOSSology, and compared selected files where the other two scanners disagreed against that SPDX file, to see if there was new insights. The Windriver scanner is based on an older version of FOSSology in part, so they are related. Thomas did random spot checks in about 500 files from the spreadsheets for the uapi headers and agreed with SPDX license identifier in the files he inspected. For the non-uapi files Thomas did random spot checks in about 15000 files. In initial set of patches against 4.14-rc6, 3 files were found to have copy/paste license identifier errors, and have been fixed to reflect the correct identifier. Additionally Philippe spent 10 hours this week doing a detailed manual inspection and review of the 12,461 patched files from the initial patch version early this week with: - a full scancode scan run, collecting the matched texts, detected license ids and scores - reviewing anything where there was a license detected (about 500+ files) to ensure that the applied SPDX license was correct - reviewing anything where there was no detection but the patch license was not GPL-2.0 WITH Linux-syscall-note to ensure that the applied SPDX license was correct This produced a worksheet with 20 files needing minor correction. This worksheet was then exported into 3 different .csv files for the different types of files to be modified. These .csv files were then reviewed by Greg. Thomas wrote a script to parse the csv files and add the proper SPDX tag to the file, in the format that the file expected. This script was further refined by Greg based on the output to detect more types of files automatically and to distinguish between header and source .c files (which need different comment types.) Finally Greg ran the script using the .csv files to generate the patches. Reviewed-by: Kate Stewart <kstewart@linuxfoundation.org> Reviewed-by: Philippe Ombredanne <pombredanne@nexb.com> Reviewed-by: Thomas Gleixner <tglx@linutronix.de> Signed-off-by: Greg Kroah-Hartman <gregkh@linuxfoundation.org>
2017-11-01 14:07:57 +00:00
// SPDX-License-Identifier: GPL-2.0
#include <linux/linkage.h>
#include <linux/errno.h>
#include <asm/unistd.h>
syscalls/core: Prepare CONFIG_ARCH_HAS_SYSCALL_WRAPPER=y for compat syscalls It may be useful for an architecture to override the definitions of the COMPAT_SYSCALL_DEFINE0() and __COMPAT_SYSCALL_DEFINEx() macros in <linux/compat.h>, in particular to use a different calling convention for syscalls. This patch provides a mechanism to do so, based on the previously introduced CONFIG_ARCH_HAS_SYSCALL_WRAPPER. If it is enabled, <asm/sycall_wrapper.h> is included in <linux/compat.h> and may be used to define the macros mentioned above. Moreover, as the syscall calling convention may be different if CONFIG_ARCH_HAS_SYSCALL_WRAPPER is set, the compat syscall function prototypes in <linux/compat.h> are #ifndef'd out in that case. As some of the syscalls and/or compat syscalls may not be present, the COND_SYSCALL() and COND_SYSCALL_COMPAT() macros in kernel/sys_ni.c as well as the SYS_NI() and COMPAT_SYS_NI() macros in kernel/time/posix-stubs.c can be re-defined in <asm/syscall_wrapper.h> iff CONFIG_ARCH_HAS_SYSCALL_WRAPPER is enabled. Signed-off-by: Dominik Brodowski <linux@dominikbrodowski.net> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Andy Lutomirski <luto@kernel.org> Cc: Borislav Petkov <bp@alien8.de> Cc: Brian Gerst <brgerst@gmail.com> Cc: Denys Vlasenko <dvlasenk@redhat.com> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Josh Poimboeuf <jpoimboe@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Link: http://lkml.kernel.org/r/20180405095307.3730-5-linux@dominikbrodowski.net Signed-off-by: Ingo Molnar <mingo@kernel.org>
2018-04-05 09:53:03 +00:00
#ifdef CONFIG_ARCH_HAS_SYSCALL_WRAPPER
/* Architectures may override COND_SYSCALL and COND_SYSCALL_COMPAT */
#include <asm/syscall_wrapper.h>
#endif /* CONFIG_ARCH_HAS_SYSCALL_WRAPPER */
/* we can't #include <linux/syscalls.h> here,
but tell gcc to not warn with -Wmissing-prototypes */
asmlinkage long sys_ni_syscall(void);
/*
* Non-implemented system calls get redirected here.
*/
asmlinkage long sys_ni_syscall(void)
{
return -ENOSYS;
}
syscalls/core: Prepare CONFIG_ARCH_HAS_SYSCALL_WRAPPER=y for compat syscalls It may be useful for an architecture to override the definitions of the COMPAT_SYSCALL_DEFINE0() and __COMPAT_SYSCALL_DEFINEx() macros in <linux/compat.h>, in particular to use a different calling convention for syscalls. This patch provides a mechanism to do so, based on the previously introduced CONFIG_ARCH_HAS_SYSCALL_WRAPPER. If it is enabled, <asm/sycall_wrapper.h> is included in <linux/compat.h> and may be used to define the macros mentioned above. Moreover, as the syscall calling convention may be different if CONFIG_ARCH_HAS_SYSCALL_WRAPPER is set, the compat syscall function prototypes in <linux/compat.h> are #ifndef'd out in that case. As some of the syscalls and/or compat syscalls may not be present, the COND_SYSCALL() and COND_SYSCALL_COMPAT() macros in kernel/sys_ni.c as well as the SYS_NI() and COMPAT_SYS_NI() macros in kernel/time/posix-stubs.c can be re-defined in <asm/syscall_wrapper.h> iff CONFIG_ARCH_HAS_SYSCALL_WRAPPER is enabled. Signed-off-by: Dominik Brodowski <linux@dominikbrodowski.net> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Andy Lutomirski <luto@kernel.org> Cc: Borislav Petkov <bp@alien8.de> Cc: Brian Gerst <brgerst@gmail.com> Cc: Denys Vlasenko <dvlasenk@redhat.com> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Josh Poimboeuf <jpoimboe@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Link: http://lkml.kernel.org/r/20180405095307.3730-5-linux@dominikbrodowski.net Signed-off-by: Ingo Molnar <mingo@kernel.org>
2018-04-05 09:53:03 +00:00
#ifndef COND_SYSCALL
#define COND_SYSCALL(name) cond_syscall(sys_##name)
syscalls/core: Prepare CONFIG_ARCH_HAS_SYSCALL_WRAPPER=y for compat syscalls It may be useful for an architecture to override the definitions of the COMPAT_SYSCALL_DEFINE0() and __COMPAT_SYSCALL_DEFINEx() macros in <linux/compat.h>, in particular to use a different calling convention for syscalls. This patch provides a mechanism to do so, based on the previously introduced CONFIG_ARCH_HAS_SYSCALL_WRAPPER. If it is enabled, <asm/sycall_wrapper.h> is included in <linux/compat.h> and may be used to define the macros mentioned above. Moreover, as the syscall calling convention may be different if CONFIG_ARCH_HAS_SYSCALL_WRAPPER is set, the compat syscall function prototypes in <linux/compat.h> are #ifndef'd out in that case. As some of the syscalls and/or compat syscalls may not be present, the COND_SYSCALL() and COND_SYSCALL_COMPAT() macros in kernel/sys_ni.c as well as the SYS_NI() and COMPAT_SYS_NI() macros in kernel/time/posix-stubs.c can be re-defined in <asm/syscall_wrapper.h> iff CONFIG_ARCH_HAS_SYSCALL_WRAPPER is enabled. Signed-off-by: Dominik Brodowski <linux@dominikbrodowski.net> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Andy Lutomirski <luto@kernel.org> Cc: Borislav Petkov <bp@alien8.de> Cc: Brian Gerst <brgerst@gmail.com> Cc: Denys Vlasenko <dvlasenk@redhat.com> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Josh Poimboeuf <jpoimboe@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Link: http://lkml.kernel.org/r/20180405095307.3730-5-linux@dominikbrodowski.net Signed-off-by: Ingo Molnar <mingo@kernel.org>
2018-04-05 09:53:03 +00:00
#endif /* COND_SYSCALL */
#ifndef COND_SYSCALL_COMPAT
#define COND_SYSCALL_COMPAT(name) cond_syscall(compat_sys_##name)
syscalls/core: Prepare CONFIG_ARCH_HAS_SYSCALL_WRAPPER=y for compat syscalls It may be useful for an architecture to override the definitions of the COMPAT_SYSCALL_DEFINE0() and __COMPAT_SYSCALL_DEFINEx() macros in <linux/compat.h>, in particular to use a different calling convention for syscalls. This patch provides a mechanism to do so, based on the previously introduced CONFIG_ARCH_HAS_SYSCALL_WRAPPER. If it is enabled, <asm/sycall_wrapper.h> is included in <linux/compat.h> and may be used to define the macros mentioned above. Moreover, as the syscall calling convention may be different if CONFIG_ARCH_HAS_SYSCALL_WRAPPER is set, the compat syscall function prototypes in <linux/compat.h> are #ifndef'd out in that case. As some of the syscalls and/or compat syscalls may not be present, the COND_SYSCALL() and COND_SYSCALL_COMPAT() macros in kernel/sys_ni.c as well as the SYS_NI() and COMPAT_SYS_NI() macros in kernel/time/posix-stubs.c can be re-defined in <asm/syscall_wrapper.h> iff CONFIG_ARCH_HAS_SYSCALL_WRAPPER is enabled. Signed-off-by: Dominik Brodowski <linux@dominikbrodowski.net> Acked-by: Linus Torvalds <torvalds@linux-foundation.org> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Andy Lutomirski <luto@kernel.org> Cc: Borislav Petkov <bp@alien8.de> Cc: Brian Gerst <brgerst@gmail.com> Cc: Denys Vlasenko <dvlasenk@redhat.com> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Josh Poimboeuf <jpoimboe@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Thomas Gleixner <tglx@linutronix.de> Link: http://lkml.kernel.org/r/20180405095307.3730-5-linux@dominikbrodowski.net Signed-off-by: Ingo Molnar <mingo@kernel.org>
2018-04-05 09:53:03 +00:00
#endif /* COND_SYSCALL_COMPAT */
/*
* This list is kept in the same order as include/uapi/asm-generic/unistd.h.
* Architecture specific entries go below, followed by deprecated or obsolete
* system calls.
*/
COND_SYSCALL(io_setup);
COND_SYSCALL_COMPAT(io_setup);
COND_SYSCALL(io_destroy);
COND_SYSCALL(io_submit);
COND_SYSCALL_COMPAT(io_submit);
COND_SYSCALL(io_cancel);
COND_SYSCALL(io_getevents_time32);
COND_SYSCALL(io_getevents);
COND_SYSCALL(io_pgetevents_time32);
COND_SYSCALL(io_pgetevents);
COND_SYSCALL_COMPAT(io_pgetevents_time32);
COND_SYSCALL_COMPAT(io_pgetevents);
Add io_uring IO interface The submission queue (SQ) and completion queue (CQ) rings are shared between the application and the kernel. This eliminates the need to copy data back and forth to submit and complete IO. IO submissions use the io_uring_sqe data structure, and completions are generated in the form of io_uring_cqe data structures. The SQ ring is an index into the io_uring_sqe array, which makes it possible to submit a batch of IOs without them being contiguous in the ring. The CQ ring is always contiguous, as completion events are inherently unordered, and hence any io_uring_cqe entry can point back to an arbitrary submission. Two new system calls are added for this: io_uring_setup(entries, params) Sets up an io_uring instance for doing async IO. On success, returns a file descriptor that the application can mmap to gain access to the SQ ring, CQ ring, and io_uring_sqes. io_uring_enter(fd, to_submit, min_complete, flags, sigset, sigsetsize) Initiates IO against the rings mapped to this fd, or waits for them to complete, or both. The behavior is controlled by the parameters passed in. If 'to_submit' is non-zero, then we'll try and submit new IO. If IORING_ENTER_GETEVENTS is set, the kernel will wait for 'min_complete' events, if they aren't already available. It's valid to set IORING_ENTER_GETEVENTS and 'min_complete' == 0 at the same time, this allows the kernel to return already completed events without waiting for them. This is useful only for polling, as for IRQ driven IO, the application can just check the CQ ring without entering the kernel. With this setup, it's possible to do async IO with a single system call. Future developments will enable polled IO with this interface, and polled submission as well. The latter will enable an application to do IO without doing ANY system calls at all. For IRQ driven IO, an application only needs to enter the kernel for completions if it wants to wait for them to occur. Each io_uring is backed by a workqueue, to support buffered async IO as well. We will only punt to an async context if the command would need to wait for IO on the device side. Any data that can be accessed directly in the page cache is done inline. This avoids the slowness issue of usual threadpools, since cached data is accessed as quickly as a sync interface. Sample application: http://git.kernel.dk/cgit/fio/plain/t/io_uring.c Reviewed-by: Hannes Reinecke <hare@suse.com> Signed-off-by: Jens Axboe <axboe@kernel.dk>
2019-01-07 17:46:33 +00:00
COND_SYSCALL(io_uring_setup);
COND_SYSCALL(io_uring_enter);
io_uring: add support for pre-mapped user IO buffers If we have fixed user buffers, we can map them into the kernel when we setup the io_uring. That avoids the need to do get_user_pages() for each and every IO. To utilize this feature, the application must call io_uring_register() after having setup an io_uring instance, passing in IORING_REGISTER_BUFFERS as the opcode. The argument must be a pointer to an iovec array, and the nr_args should contain how many iovecs the application wishes to map. If successful, these buffers are now mapped into the kernel, eligible for IO. To use these fixed buffers, the application must use the IORING_OP_READ_FIXED and IORING_OP_WRITE_FIXED opcodes, and then set sqe->index to the desired buffer index. sqe->addr..sqe->addr+seq->len must point to somewhere inside the indexed buffer. The application may register buffers throughout the lifetime of the io_uring instance. It can call io_uring_register() with IORING_UNREGISTER_BUFFERS as the opcode to unregister the current set of buffers, and then register a new set. The application need not unregister buffers explicitly before shutting down the io_uring instance. It's perfectly valid to setup a larger buffer, and then sometimes only use parts of it for an IO. As long as the range is within the originally mapped region, it will work just fine. For now, buffers must not be file backed. If file backed buffers are passed in, the registration will fail with -1/EOPNOTSUPP. This restriction may be relaxed in the future. RLIMIT_MEMLOCK is used to check how much memory we can pin. A somewhat arbitrary 1G per buffer size is also imposed. Reviewed-by: Hannes Reinecke <hare@suse.com> Signed-off-by: Jens Axboe <axboe@kernel.dk>
2019-01-09 16:16:05 +00:00
COND_SYSCALL(io_uring_register);
COND_SYSCALL(lookup_dcookie);
COND_SYSCALL_COMPAT(lookup_dcookie);
COND_SYSCALL(eventfd2);
COND_SYSCALL(epoll_create1);
COND_SYSCALL(epoll_ctl);
COND_SYSCALL(epoll_pwait);
COND_SYSCALL_COMPAT(epoll_pwait);
COND_SYSCALL(epoll_pwait2);
COND_SYSCALL_COMPAT(epoll_pwait2);
COND_SYSCALL(inotify_init1);
COND_SYSCALL(inotify_add_watch);
COND_SYSCALL(inotify_rm_watch);
COND_SYSCALL(ioprio_set);
COND_SYSCALL(ioprio_get);
COND_SYSCALL(flock);
COND_SYSCALL(quotactl);
COND_SYSCALL(quotactl_fd);
COND_SYSCALL(signalfd4);
COND_SYSCALL_COMPAT(signalfd4);
COND_SYSCALL(timerfd_create);
COND_SYSCALL(timerfd_settime);
COND_SYSCALL(timerfd_settime32);
COND_SYSCALL(timerfd_gettime);
COND_SYSCALL(timerfd_gettime32);
COND_SYSCALL(acct);
COND_SYSCALL(capget);
COND_SYSCALL(capset);
2019-06-20 23:26:35 +00:00
/* __ARCH_WANT_SYS_CLONE3 */
COND_SYSCALL(clone3);
COND_SYSCALL(futex);
COND_SYSCALL(futex_time32);
COND_SYSCALL(set_robust_list);
COND_SYSCALL_COMPAT(set_robust_list);
COND_SYSCALL(get_robust_list);
COND_SYSCALL_COMPAT(get_robust_list);
COND_SYSCALL(futex_waitv);
COND_SYSCALL(kexec_load);
COND_SYSCALL_COMPAT(kexec_load);
COND_SYSCALL(init_module);
COND_SYSCALL(delete_module);
COND_SYSCALL(syslog);
COND_SYSCALL(setregid);
COND_SYSCALL(setgid);
COND_SYSCALL(setreuid);
COND_SYSCALL(setuid);
COND_SYSCALL(setresuid);
COND_SYSCALL(getresuid);
COND_SYSCALL(setresgid);
COND_SYSCALL(getresgid);
COND_SYSCALL(setfsuid);
COND_SYSCALL(setfsgid);
COND_SYSCALL(setgroups);
COND_SYSCALL(getgroups);
COND_SYSCALL(mq_open);
COND_SYSCALL_COMPAT(mq_open);
COND_SYSCALL(mq_unlink);
COND_SYSCALL(mq_timedsend);
COND_SYSCALL(mq_timedsend_time32);
COND_SYSCALL(mq_timedreceive);
COND_SYSCALL(mq_timedreceive_time32);
COND_SYSCALL(mq_notify);
COND_SYSCALL_COMPAT(mq_notify);
COND_SYSCALL(mq_getsetattr);
COND_SYSCALL_COMPAT(mq_getsetattr);
COND_SYSCALL(msgget);
ipc: rename old-style shmctl/semctl/msgctl syscalls The behavior of these system calls is slightly different between architectures, as determined by the CONFIG_ARCH_WANT_IPC_PARSE_VERSION symbol. Most architectures that implement the split IPC syscalls don't set that symbol and only get the modern version, but alpha, arm, microblaze, mips-n32, mips-n64 and xtensa expect the caller to pass the IPC_64 flag. For the architectures that so far only implement sys_ipc(), i.e. m68k, mips-o32, powerpc, s390, sh, sparc, and x86-32, we want the new behavior when adding the split syscalls, so we need to distinguish between the two groups of architectures. The method I picked for this distinction is to have a separate system call entry point: sys_old_*ctl() now uses ipc_parse_version, while sys_*ctl() does not. The system call tables of the five architectures are changed accordingly. As an additional benefit, we no longer need the configuration specific definition for ipc_parse_version(), it always does the same thing now, but simply won't get called on architectures with the modern interface. A small downside is that on architectures that do set ARCH_WANT_IPC_PARSE_VERSION, we now have an extra set of entry points that are never called. They only add a few bytes of bloat, so it seems better to keep them compared to adding yet another Kconfig symbol. I considered adding new syscall numbers for the IPC_64 variants for consistency, but decided against that for now. Signed-off-by: Arnd Bergmann <arnd@arndb.de>
2018-12-31 21:22:40 +00:00
COND_SYSCALL(old_msgctl);
COND_SYSCALL(msgctl);
COND_SYSCALL_COMPAT(msgctl);
COND_SYSCALL_COMPAT(old_msgctl);
COND_SYSCALL(msgrcv);
COND_SYSCALL_COMPAT(msgrcv);
COND_SYSCALL(msgsnd);
COND_SYSCALL_COMPAT(msgsnd);
COND_SYSCALL(semget);
ipc: rename old-style shmctl/semctl/msgctl syscalls The behavior of these system calls is slightly different between architectures, as determined by the CONFIG_ARCH_WANT_IPC_PARSE_VERSION symbol. Most architectures that implement the split IPC syscalls don't set that symbol and only get the modern version, but alpha, arm, microblaze, mips-n32, mips-n64 and xtensa expect the caller to pass the IPC_64 flag. For the architectures that so far only implement sys_ipc(), i.e. m68k, mips-o32, powerpc, s390, sh, sparc, and x86-32, we want the new behavior when adding the split syscalls, so we need to distinguish between the two groups of architectures. The method I picked for this distinction is to have a separate system call entry point: sys_old_*ctl() now uses ipc_parse_version, while sys_*ctl() does not. The system call tables of the five architectures are changed accordingly. As an additional benefit, we no longer need the configuration specific definition for ipc_parse_version(), it always does the same thing now, but simply won't get called on architectures with the modern interface. A small downside is that on architectures that do set ARCH_WANT_IPC_PARSE_VERSION, we now have an extra set of entry points that are never called. They only add a few bytes of bloat, so it seems better to keep them compared to adding yet another Kconfig symbol. I considered adding new syscall numbers for the IPC_64 variants for consistency, but decided against that for now. Signed-off-by: Arnd Bergmann <arnd@arndb.de>
2018-12-31 21:22:40 +00:00
COND_SYSCALL(old_semctl);
COND_SYSCALL(semctl);
COND_SYSCALL_COMPAT(semctl);
COND_SYSCALL_COMPAT(old_semctl);
COND_SYSCALL(semtimedop);
COND_SYSCALL(semtimedop_time32);
COND_SYSCALL(semop);
COND_SYSCALL(shmget);
ipc: rename old-style shmctl/semctl/msgctl syscalls The behavior of these system calls is slightly different between architectures, as determined by the CONFIG_ARCH_WANT_IPC_PARSE_VERSION symbol. Most architectures that implement the split IPC syscalls don't set that symbol and only get the modern version, but alpha, arm, microblaze, mips-n32, mips-n64 and xtensa expect the caller to pass the IPC_64 flag. For the architectures that so far only implement sys_ipc(), i.e. m68k, mips-o32, powerpc, s390, sh, sparc, and x86-32, we want the new behavior when adding the split syscalls, so we need to distinguish between the two groups of architectures. The method I picked for this distinction is to have a separate system call entry point: sys_old_*ctl() now uses ipc_parse_version, while sys_*ctl() does not. The system call tables of the five architectures are changed accordingly. As an additional benefit, we no longer need the configuration specific definition for ipc_parse_version(), it always does the same thing now, but simply won't get called on architectures with the modern interface. A small downside is that on architectures that do set ARCH_WANT_IPC_PARSE_VERSION, we now have an extra set of entry points that are never called. They only add a few bytes of bloat, so it seems better to keep them compared to adding yet another Kconfig symbol. I considered adding new syscall numbers for the IPC_64 variants for consistency, but decided against that for now. Signed-off-by: Arnd Bergmann <arnd@arndb.de>
2018-12-31 21:22:40 +00:00
COND_SYSCALL(old_shmctl);
COND_SYSCALL(shmctl);
COND_SYSCALL_COMPAT(shmctl);
COND_SYSCALL_COMPAT(old_shmctl);
COND_SYSCALL(shmat);
COND_SYSCALL_COMPAT(shmat);
COND_SYSCALL(shmdt);
COND_SYSCALL(socket);
COND_SYSCALL(socketpair);
COND_SYSCALL(bind);
COND_SYSCALL(listen);
COND_SYSCALL(accept);
COND_SYSCALL(connect);
COND_SYSCALL(getsockname);
COND_SYSCALL(getpeername);
COND_SYSCALL(setsockopt);
COND_SYSCALL_COMPAT(setsockopt);
COND_SYSCALL(getsockopt);
COND_SYSCALL_COMPAT(getsockopt);
COND_SYSCALL(sendto);
COND_SYSCALL(shutdown);
COND_SYSCALL(recvfrom);
COND_SYSCALL_COMPAT(recvfrom);
COND_SYSCALL(sendmsg);
COND_SYSCALL_COMPAT(sendmsg);
COND_SYSCALL(recvmsg);
COND_SYSCALL_COMPAT(recvmsg);
COND_SYSCALL(mremap);
COND_SYSCALL(add_key);
COND_SYSCALL(request_key);
COND_SYSCALL(keyctl);
COND_SYSCALL_COMPAT(keyctl);
landlock: Add syscall implementations These 3 system calls are designed to be used by unprivileged processes to sandbox themselves: * landlock_create_ruleset(2): Creates a ruleset and returns its file descriptor. * landlock_add_rule(2): Adds a rule (e.g. file hierarchy access) to a ruleset, identified by the dedicated file descriptor. * landlock_restrict_self(2): Enforces a ruleset on the calling thread and its future children (similar to seccomp). This syscall has the same usage restrictions as seccomp(2): the caller must have the no_new_privs attribute set or have CAP_SYS_ADMIN in the current user namespace. All these syscalls have a "flags" argument (not currently used) to enable extensibility. Here are the motivations for these new syscalls: * A sandboxed process may not have access to file systems, including /dev, /sys or /proc, but it should still be able to add more restrictions to itself. * Neither prctl(2) nor seccomp(2) (which was used in a previous version) fit well with the current definition of a Landlock security policy. All passed structs (attributes) are checked at build time to ensure that they don't contain holes and that they are aligned the same way for each architecture. See the user and kernel documentation for more details (provided by a following commit): * Documentation/userspace-api/landlock.rst * Documentation/security/landlock.rst Cc: Arnd Bergmann <arnd@arndb.de> Cc: James Morris <jmorris@namei.org> Cc: Jann Horn <jannh@google.com> Cc: Kees Cook <keescook@chromium.org> Signed-off-by: Mickaël Salaün <mic@linux.microsoft.com> Acked-by: Serge Hallyn <serge@hallyn.com> Link: https://lore.kernel.org/r/20210422154123.13086-9-mic@digikod.net Signed-off-by: James Morris <jamorris@linux.microsoft.com>
2021-04-22 15:41:18 +00:00
COND_SYSCALL(landlock_create_ruleset);
COND_SYSCALL(landlock_add_rule);
COND_SYSCALL(landlock_restrict_self);
COND_SYSCALL(fadvise64_64);
COND_SYSCALL_COMPAT(fadvise64_64);
/* CONFIG_MMU only */
COND_SYSCALL(swapon);
COND_SYSCALL(swapoff);
COND_SYSCALL(mprotect);
COND_SYSCALL(msync);
COND_SYSCALL(mlock);
COND_SYSCALL(munlock);
COND_SYSCALL(mlockall);
COND_SYSCALL(munlockall);
COND_SYSCALL(mincore);
COND_SYSCALL(madvise);
mm/madvise: introduce process_madvise() syscall: an external memory hinting API There is usecase that System Management Software(SMS) want to give a memory hint like MADV_[COLD|PAGEEOUT] to other processes and in the case of Android, it is the ActivityManagerService. The information required to make the reclaim decision is not known to the app. Instead, it is known to the centralized userspace daemon(ActivityManagerService), and that daemon must be able to initiate reclaim on its own without any app involvement. To solve the issue, this patch introduces a new syscall process_madvise(2). It uses pidfd of an external process to give the hint. It also supports vector address range because Android app has thousands of vmas due to zygote so it's totally waste of CPU and power if we should call the syscall one by one for each vma.(With testing 2000-vma syscall vs 1-vector syscall, it showed 15% performance improvement. I think it would be bigger in real practice because the testing ran very cache friendly environment). Another potential use case for the vector range is to amortize the cost ofTLB shootdowns for multiple ranges when using MADV_DONTNEED; this could benefit users like TCP receive zerocopy and malloc implementations. In future, we could find more usecases for other advises so let's make it happens as API since we introduce a new syscall at this moment. With that, existing madvise(2) user could replace it with process_madvise(2) with their own pid if they want to have batch address ranges support feature. ince it could affect other process's address range, only privileged process(PTRACE_MODE_ATTACH_FSCREDS) or something else(e.g., being the same UID) gives it the right to ptrace the process could use it successfully. The flag argument is reserved for future use if we need to extend the API. I think supporting all hints madvise has/will supported/support to process_madvise is rather risky. Because we are not sure all hints make sense from external process and implementation for the hint may rely on the caller being in the current context so it could be error-prone. Thus, I just limited hints as MADV_[COLD|PAGEOUT] in this patch. If someone want to add other hints, we could hear the usecase and review it for each hint. It's safer for maintenance rather than introducing a buggy syscall but hard to fix it later. So finally, the API is as follows, ssize_t process_madvise(int pidfd, const struct iovec *iovec, unsigned long vlen, int advice, unsigned int flags); DESCRIPTION The process_madvise() system call is used to give advice or directions to the kernel about the address ranges from external process as well as local process. It provides the advice to address ranges of process described by iovec and vlen. The goal of such advice is to improve system or application performance. The pidfd selects the process referred to by the PID file descriptor specified in pidfd. (See pidofd_open(2) for further information) The pointer iovec points to an array of iovec structures, defined in <sys/uio.h> as: struct iovec { void *iov_base; /* starting address */ size_t iov_len; /* number of bytes to be advised */ }; The iovec describes address ranges beginning at address(iov_base) and with size length of bytes(iov_len). The vlen represents the number of elements in iovec. The advice is indicated in the advice argument, which is one of the following at this moment if the target process specified by pidfd is external. MADV_COLD MADV_PAGEOUT Permission to provide a hint to external process is governed by a ptrace access mode PTRACE_MODE_ATTACH_FSCREDS check; see ptrace(2). The process_madvise supports every advice madvise(2) has if target process is in same thread group with calling process so user could use process_madvise(2) to extend existing madvise(2) to support vector address ranges. RETURN VALUE On success, process_madvise() returns the number of bytes advised. This return value may be less than the total number of requested bytes, if an error occurred. The caller should check return value to determine whether a partial advice occurred. FAQ: Q.1 - Why does any external entity have better knowledge? Quote from Sandeep "For Android, every application (including the special SystemServer) are forked from Zygote. The reason of course is to share as many libraries and classes between the two as possible to benefit from the preloading during boot. After applications start, (almost) all of the APIs end up calling into this SystemServer process over IPC (binder) and back to the application. In a fully running system, the SystemServer monitors every single process periodically to calculate their PSS / RSS and also decides which process is "important" to the user for interactivity. So, because of how these processes start _and_ the fact that the SystemServer is looping to monitor each process, it does tend to *know* which address range of the application is not used / useful. Besides, we can never rely on applications to clean things up themselves. We've had the "hey app1, the system is low on memory, please trim your memory usage down" notifications for a long time[1]. They rely on applications honoring the broadcasts and very few do. So, if we want to avoid the inevitable killing of the application and restarting it, some way to be able to tell the OS about unimportant memory in these applications will be useful. - ssp Q.2 - How to guarantee the race(i.e., object validation) between when giving a hint from an external process and get the hint from the target process? process_madvise operates on the target process's address space as it exists at the instant that process_madvise is called. If the space target process can run between the time the process_madvise process inspects the target process address space and the time that process_madvise is actually called, process_madvise may operate on memory regions that the calling process does not expect. It's the responsibility of the process calling process_madvise to close this race condition. For example, the calling process can suspend the target process with ptrace, SIGSTOP, or the freezer cgroup so that it doesn't have an opportunity to change its own address space before process_madvise is called. Another option is to operate on memory regions that the caller knows a priori will be unchanged in the target process. Yet another option is to accept the race for certain process_madvise calls after reasoning that mistargeting will do no harm. The suggested API itself does not provide synchronization. It also apply other APIs like move_pages, process_vm_write. The race isn't really a problem though. Why is it so wrong to require that callers do their own synchronization in some manner? Nobody objects to write(2) merely because it's possible for two processes to open the same file and clobber each other's writes --- instead, we tell people to use flock or something. Think about mmap. It never guarantees newly allocated address space is still valid when the user tries to access it because other threads could unmap the memory right before. That's where we need synchronization by using other API or design from userside. It shouldn't be part of API itself. If someone needs more fine-grained synchronization rather than process level, there were two ideas suggested - cookie[2] and anon-fd[3]. Both are applicable via using last reserved argument of the API but I don't think it's necessary right now since we have already ways to prevent the race so don't want to add additional complexity with more fine-grained optimization model. To make the API extend, it reserved an unsigned long as last argument so we could support it in future if someone really needs it. Q.3 - Why doesn't ptrace work? Injecting an madvise in the target process using ptrace would not work for us because such injected madvise would have to be executed by the target process, which means that process would have to be runnable and that creates the risk of the abovementioned race and hinting a wrong VMA. Furthermore, we want to act the hint in caller's context, not the callee's, because the callee is usually limited in cpuset/cgroups or even freezed state so they can't act by themselves quick enough, which causes more thrashing/kill. It doesn't work if the target process are ptraced(e.g., strace, debugger, minidump) because a process can have at most one ptracer. [1] https://developer.android.com/topic/performance/memory" [2] process_getinfo for getting the cookie which is updated whenever vma of process address layout are changed - Daniel Colascione - https://lore.kernel.org/lkml/20190520035254.57579-1-minchan@kernel.org/T/#m7694416fd179b2066a2c62b5b139b14e3894e224 [3] anonymous fd which is used for the object(i.e., address range) validation - Michal Hocko - https://lore.kernel.org/lkml/20200120112722.GY18451@dhcp22.suse.cz/ [minchan@kernel.org: fix process_madvise build break for arm64] Link: http://lkml.kernel.org/r/20200303145756.GA219683@google.com [minchan@kernel.org: fix build error for mips of process_madvise] Link: http://lkml.kernel.org/r/20200508052517.GA197378@google.com [akpm@linux-foundation.org: fix patch ordering issue] [akpm@linux-foundation.org: fix arm64 whoops] [minchan@kernel.org: make process_madvise() vlen arg have type size_t, per Florian] [akpm@linux-foundation.org: fix i386 build] [sfr@canb.auug.org.au: fix syscall numbering] Link: https://lkml.kernel.org/r/20200905142639.49fc3f1a@canb.auug.org.au [sfr@canb.auug.org.au: madvise.c needs compat.h] Link: https://lkml.kernel.org/r/20200908204547.285646b4@canb.auug.org.au [minchan@kernel.org: fix mips build] Link: https://lkml.kernel.org/r/20200909173655.GC2435453@google.com [yuehaibing@huawei.com: remove duplicate header which is included twice] Link: https://lkml.kernel.org/r/20200915121550.30584-1-yuehaibing@huawei.com [minchan@kernel.org: do not use helper functions for process_madvise] Link: https://lkml.kernel.org/r/20200921175539.GB387368@google.com [akpm@linux-foundation.org: pidfd_get_pid() gained an argument] [sfr@canb.auug.org.au: fix up for "iov_iter: transparently handle compat iovecs in import_iovec"] Link: https://lkml.kernel.org/r/20200928212542.468e1fef@canb.auug.org.au Signed-off-by: Minchan Kim <minchan@kernel.org> Signed-off-by: YueHaibing <yuehaibing@huawei.com> Signed-off-by: Stephen Rothwell <sfr@canb.auug.org.au> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Reviewed-by: Suren Baghdasaryan <surenb@google.com> Reviewed-by: Vlastimil Babka <vbabka@suse.cz> Acked-by: David Rientjes <rientjes@google.com> Cc: Alexander Duyck <alexander.h.duyck@linux.intel.com> Cc: Brian Geffon <bgeffon@google.com> Cc: Christian Brauner <christian@brauner.io> Cc: Daniel Colascione <dancol@google.com> Cc: Jann Horn <jannh@google.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Joel Fernandes <joel@joelfernandes.org> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: John Dias <joaodias@google.com> Cc: Kirill Tkhai <ktkhai@virtuozzo.com> Cc: Michal Hocko <mhocko@suse.com> Cc: Oleksandr Natalenko <oleksandr@redhat.com> Cc: Sandeep Patil <sspatil@google.com> Cc: SeongJae Park <sj38.park@gmail.com> Cc: SeongJae Park <sjpark@amazon.de> Cc: Shakeel Butt <shakeelb@google.com> Cc: Sonny Rao <sonnyrao@google.com> Cc: Tim Murray <timmurray@google.com> Cc: Christian Brauner <christian.brauner@ubuntu.com> Cc: Florian Weimer <fw@deneb.enyo.de> Cc: <linux-man@vger.kernel.org> Link: http://lkml.kernel.org/r/20200302193630.68771-3-minchan@kernel.org Link: http://lkml.kernel.org/r/20200508183320.GA125527@google.com Link: http://lkml.kernel.org/r/20200622192900.22757-4-minchan@kernel.org Link: https://lkml.kernel.org/r/20200901000633.1920247-4-minchan@kernel.org Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-10-17 23:14:59 +00:00
COND_SYSCALL(process_madvise);
COND_SYSCALL(process_mrelease);
COND_SYSCALL(remap_file_pages);
COND_SYSCALL(mbind);
COND_SYSCALL(get_mempolicy);
COND_SYSCALL(set_mempolicy);
COND_SYSCALL(migrate_pages);
COND_SYSCALL(move_pages);
COND_SYSCALL(set_mempolicy_home_node);
cachestat: implement cachestat syscall There is currently no good way to query the page cache state of large file sets and directory trees. There is mincore(), but it scales poorly: the kernel writes out a lot of bitmap data that userspace has to aggregate, when the user really doesn not care about per-page information in that case. The user also needs to mmap and unmap each file as it goes along, which can be quite slow as well. Some use cases where this information could come in handy: * Allowing database to decide whether to perform an index scan or direct table queries based on the in-memory cache state of the index. * Visibility into the writeback algorithm, for performance issues diagnostic. * Workload-aware writeback pacing: estimating IO fulfilled by page cache (and IO to be done) within a range of a file, allowing for more frequent syncing when and where there is IO capacity, and batching when there is not. * Computing memory usage of large files/directory trees, analogous to the du tool for disk usage. More information about these use cases could be found in the following thread: https://lore.kernel.org/lkml/20230315170934.GA97793@cmpxchg.org/ This patch implements a new syscall that queries cache state of a file and summarizes the number of cached pages, number of dirty pages, number of pages marked for writeback, number of (recently) evicted pages, etc. in a given range. Currently, the syscall is only wired in for x86 architecture. NAME cachestat - query the page cache statistics of a file. SYNOPSIS #include <sys/mman.h> struct cachestat_range { __u64 off; __u64 len; }; struct cachestat { __u64 nr_cache; __u64 nr_dirty; __u64 nr_writeback; __u64 nr_evicted; __u64 nr_recently_evicted; }; int cachestat(unsigned int fd, struct cachestat_range *cstat_range, struct cachestat *cstat, unsigned int flags); DESCRIPTION cachestat() queries the number of cached pages, number of dirty pages, number of pages marked for writeback, number of evicted pages, number of recently evicted pages, in the bytes range given by `off` and `len`. An evicted page is a page that is previously in the page cache but has been evicted since. A page is recently evicted if its last eviction was recent enough that its reentry to the cache would indicate that it is actively being used by the system, and that there is memory pressure on the system. These values are returned in a cachestat struct, whose address is given by the `cstat` argument. The `off` and `len` arguments must be non-negative integers. If `len` > 0, the queried range is [`off`, `off` + `len`]. If `len` == 0, we will query in the range from `off` to the end of the file. The `flags` argument is unused for now, but is included for future extensibility. User should pass 0 (i.e no flag specified). Currently, hugetlbfs is not supported. Because the status of a page can change after cachestat() checks it but before it returns to the application, the returned values may contain stale information. RETURN VALUE On success, cachestat returns 0. On error, -1 is returned, and errno is set to indicate the error. ERRORS EFAULT cstat or cstat_args points to an invalid address. EINVAL invalid flags. EBADF invalid file descriptor. EOPNOTSUPP file descriptor is of a hugetlbfs file [nphamcs@gmail.com: replace rounddown logic with the existing helper] Link: https://lkml.kernel.org/r/20230504022044.3675469-1-nphamcs@gmail.com Link: https://lkml.kernel.org/r/20230503013608.2431726-3-nphamcs@gmail.com Signed-off-by: Nhat Pham <nphamcs@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Cc: Brian Foster <bfoster@redhat.com> Cc: Matthew Wilcox (Oracle) <willy@infradead.org> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2023-05-03 01:36:07 +00:00
COND_SYSCALL(cachestat);
COND_SYSCALL(perf_event_open);
COND_SYSCALL(accept4);
COND_SYSCALL(recvmmsg);
y2038: socket: Add compat_sys_recvmmsg_time64 recvmmsg() takes two arguments to pointers of structures that differ between 32-bit and 64-bit architectures: mmsghdr and timespec. For y2038 compatbility, we are changing the native system call from timespec to __kernel_timespec with a 64-bit time_t (in another patch), and use the existing compat system call on both 32-bit and 64-bit architectures for compatibility with traditional 32-bit user space. As we now have two variants of recvmmsg() for 32-bit tasks that are both different from the variant that we use on 64-bit tasks, this means we also require two compat system calls! The solution I picked is to flip things around: The existing compat_sys_recvmmsg() call gets moved from net/compat.c into net/socket.c and now handles the case for old user space on all architectures that have set CONFIG_COMPAT_32BIT_TIME. A new compat_sys_recvmmsg_time64() call gets added in the old place for 64-bit architectures only, this one handles the case of a compat mmsghdr structure combined with __kernel_timespec. In the indirect sys_socketcall(), we now need to call either do_sys_recvmmsg() or __compat_sys_recvmmsg(), depending on what kind of architecture we are on. For compat_sys_socketcall(), no such change is needed, we always call __compat_sys_recvmmsg(). I decided to not add a new SYS_RECVMMSG_TIME64 socketcall: Any libc implementation for 64-bit time_t will need significant changes including an updated asm/unistd.h, and it seems better to consistently use the separate syscalls that configuration, leaving the socketcall only for backward compatibility with 32-bit time_t based libc. The naming is asymmetric for the moment, so both existing syscalls entry points keep their names, while the new ones are recvmmsg_time32 and compat_recvmmsg_time64 respectively. I expect that we will rename the compat syscalls later as we start using generated syscall tables everywhere and add these entry points. Signed-off-by: Arnd Bergmann <arnd@arndb.de>
2018-04-18 11:43:52 +00:00
COND_SYSCALL(recvmmsg_time32);
COND_SYSCALL_COMPAT(recvmmsg_time32);
y2038: socket: Add compat_sys_recvmmsg_time64 recvmmsg() takes two arguments to pointers of structures that differ between 32-bit and 64-bit architectures: mmsghdr and timespec. For y2038 compatbility, we are changing the native system call from timespec to __kernel_timespec with a 64-bit time_t (in another patch), and use the existing compat system call on both 32-bit and 64-bit architectures for compatibility with traditional 32-bit user space. As we now have two variants of recvmmsg() for 32-bit tasks that are both different from the variant that we use on 64-bit tasks, this means we also require two compat system calls! The solution I picked is to flip things around: The existing compat_sys_recvmmsg() call gets moved from net/compat.c into net/socket.c and now handles the case for old user space on all architectures that have set CONFIG_COMPAT_32BIT_TIME. A new compat_sys_recvmmsg_time64() call gets added in the old place for 64-bit architectures only, this one handles the case of a compat mmsghdr structure combined with __kernel_timespec. In the indirect sys_socketcall(), we now need to call either do_sys_recvmmsg() or __compat_sys_recvmmsg(), depending on what kind of architecture we are on. For compat_sys_socketcall(), no such change is needed, we always call __compat_sys_recvmmsg(). I decided to not add a new SYS_RECVMMSG_TIME64 socketcall: Any libc implementation for 64-bit time_t will need significant changes including an updated asm/unistd.h, and it seems better to consistently use the separate syscalls that configuration, leaving the socketcall only for backward compatibility with 32-bit time_t based libc. The naming is asymmetric for the moment, so both existing syscalls entry points keep their names, while the new ones are recvmmsg_time32 and compat_recvmmsg_time64 respectively. I expect that we will rename the compat syscalls later as we start using generated syscall tables everywhere and add these entry points. Signed-off-by: Arnd Bergmann <arnd@arndb.de>
2018-04-18 11:43:52 +00:00
COND_SYSCALL_COMPAT(recvmmsg_time64);
/*
* Architecture specific syscalls: see further below
*/
/* fanotify */
COND_SYSCALL(fanotify_init);
COND_SYSCALL(fanotify_mark);
/* open by handle */
COND_SYSCALL(name_to_handle_at);
COND_SYSCALL(open_by_handle_at);
COND_SYSCALL_COMPAT(open_by_handle_at);
syscalls, x86: add __NR_kcmp syscall While doing the checkpoint-restore in the user space one need to determine whether various kernel objects (like mm_struct-s of file_struct-s) are shared between tasks and restore this state. The 2nd step can be solved by using appropriate CLONE_ flags and the unshare syscall, while there's currently no ways for solving the 1st one. One of the ways for checking whether two tasks share e.g. mm_struct is to provide some mm_struct ID of a task to its proc file, but showing such info considered to be not that good for security reasons. Thus after some debates we end up in conclusion that using that named 'comparison' syscall might be the best candidate. So here is it -- __NR_kcmp. It takes up to 5 arguments - the pids of the two tasks (which characteristics should be compared), the comparison type and (in case of comparison of files) two file descriptors. Lookups for pids are done in the caller's PID namespace only. At moment only x86 is supported and tested. [akpm@linux-foundation.org: fix up selftests, warnings] [akpm@linux-foundation.org: include errno.h] [akpm@linux-foundation.org: tweak comment text] Signed-off-by: Cyrill Gorcunov <gorcunov@openvz.org> Acked-by: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Pavel Emelyanov <xemul@parallels.com> Cc: Andrey Vagin <avagin@openvz.org> Cc: KOSAKI Motohiro <kosaki.motohiro@gmail.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Glauber Costa <glommer@parallels.com> Cc: Andi Kleen <andi@firstfloor.org> Cc: Tejun Heo <tj@kernel.org> Cc: Matt Helsley <matthltc@us.ibm.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: Eric Dumazet <eric.dumazet@gmail.com> Cc: Vasiliy Kulikov <segoon@openwall.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Valdis.Kletnieks@vt.edu Cc: Michal Marek <mmarek@suse.cz> Cc: Frederic Weisbecker <fweisbec@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-05-31 23:26:44 +00:00
COND_SYSCALL(sendmmsg);
COND_SYSCALL_COMPAT(sendmmsg);
COND_SYSCALL(process_vm_readv);
COND_SYSCALL_COMPAT(process_vm_readv);
COND_SYSCALL(process_vm_writev);
COND_SYSCALL_COMPAT(process_vm_writev);
syscalls, x86: add __NR_kcmp syscall While doing the checkpoint-restore in the user space one need to determine whether various kernel objects (like mm_struct-s of file_struct-s) are shared between tasks and restore this state. The 2nd step can be solved by using appropriate CLONE_ flags and the unshare syscall, while there's currently no ways for solving the 1st one. One of the ways for checking whether two tasks share e.g. mm_struct is to provide some mm_struct ID of a task to its proc file, but showing such info considered to be not that good for security reasons. Thus after some debates we end up in conclusion that using that named 'comparison' syscall might be the best candidate. So here is it -- __NR_kcmp. It takes up to 5 arguments - the pids of the two tasks (which characteristics should be compared), the comparison type and (in case of comparison of files) two file descriptors. Lookups for pids are done in the caller's PID namespace only. At moment only x86 is supported and tested. [akpm@linux-foundation.org: fix up selftests, warnings] [akpm@linux-foundation.org: include errno.h] [akpm@linux-foundation.org: tweak comment text] Signed-off-by: Cyrill Gorcunov <gorcunov@openvz.org> Acked-by: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Pavel Emelyanov <xemul@parallels.com> Cc: Andrey Vagin <avagin@openvz.org> Cc: KOSAKI Motohiro <kosaki.motohiro@gmail.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Glauber Costa <glommer@parallels.com> Cc: Andi Kleen <andi@firstfloor.org> Cc: Tejun Heo <tj@kernel.org> Cc: Matt Helsley <matthltc@us.ibm.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: Eric Dumazet <eric.dumazet@gmail.com> Cc: Vasiliy Kulikov <segoon@openwall.com> Cc: Alexey Dobriyan <adobriyan@gmail.com> Cc: Valdis.Kletnieks@vt.edu Cc: Michal Marek <mmarek@suse.cz> Cc: Frederic Weisbecker <fweisbec@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-05-31 23:26:44 +00:00
/* compare kernel pointers */
COND_SYSCALL(kcmp);
COND_SYSCALL(finit_module);
/* operate on Secure Computing state */
COND_SYSCALL(seccomp);
COND_SYSCALL(memfd_create);
/* access BPF programs and maps */
COND_SYSCALL(bpf);
syscalls: implement execveat() system call This patchset adds execveat(2) for x86, and is derived from Meredydd Luff's patch from Sept 2012 (https://lkml.org/lkml/2012/9/11/528). The primary aim of adding an execveat syscall is to allow an implementation of fexecve(3) that does not rely on the /proc filesystem, at least for executables (rather than scripts). The current glibc version of fexecve(3) is implemented via /proc, which causes problems in sandboxed or otherwise restricted environments. Given the desire for a /proc-free fexecve() implementation, HPA suggested (https://lkml.org/lkml/2006/7/11/556) that an execveat(2) syscall would be an appropriate generalization. Also, having a new syscall means that it can take a flags argument without back-compatibility concerns. The current implementation just defines the AT_EMPTY_PATH and AT_SYMLINK_NOFOLLOW flags, but other flags could be added in future -- for example, flags for new namespaces (as suggested at https://lkml.org/lkml/2006/7/11/474). Related history: - https://lkml.org/lkml/2006/12/27/123 is an example of someone realizing that fexecve() is likely to fail in a chroot environment. - http://bugs.debian.org/cgi-bin/bugreport.cgi?bug=514043 covered documenting the /proc requirement of fexecve(3) in its manpage, to "prevent other people from wasting their time". - https://bugzilla.redhat.com/show_bug.cgi?id=241609 described a problem where a process that did setuid() could not fexecve() because it no longer had access to /proc/self/fd; this has since been fixed. This patch (of 4): Add a new execveat(2) system call. execveat() is to execve() as openat() is to open(): it takes a file descriptor that refers to a directory, and resolves the filename relative to that. In addition, if the filename is empty and AT_EMPTY_PATH is specified, execveat() executes the file to which the file descriptor refers. This replicates the functionality of fexecve(), which is a system call in other UNIXen, but in Linux glibc it depends on opening "/proc/self/fd/<fd>" (and so relies on /proc being mounted). The filename fed to the executed program as argv[0] (or the name of the script fed to a script interpreter) will be of the form "/dev/fd/<fd>" (for an empty filename) or "/dev/fd/<fd>/<filename>", effectively reflecting how the executable was found. This does however mean that execution of a script in a /proc-less environment won't work; also, script execution via an O_CLOEXEC file descriptor fails (as the file will not be accessible after exec). Based on patches by Meredydd Luff. Signed-off-by: David Drysdale <drysdale@google.com> Cc: Meredydd Luff <meredydd@senatehouse.org> Cc: Shuah Khan <shuah.kh@samsung.com> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Andy Lutomirski <luto@amacapital.net> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Kees Cook <keescook@chromium.org> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Rich Felker <dalias@aerifal.cx> Cc: Christoph Hellwig <hch@infradead.org> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-12-13 00:57:29 +00:00
/* execveat */
COND_SYSCALL(execveat);
sys_membarrier(): system-wide memory barrier (generic, x86) Here is an implementation of a new system call, sys_membarrier(), which executes a memory barrier on all threads running on the system. It is implemented by calling synchronize_sched(). It can be used to distribute the cost of user-space memory barriers asymmetrically by transforming pairs of memory barriers into pairs consisting of sys_membarrier() and a compiler barrier. For synchronization primitives that distinguish between read-side and write-side (e.g. userspace RCU [1], rwlocks), the read-side can be accelerated significantly by moving the bulk of the memory barrier overhead to the write-side. The existing applications of which I am aware that would be improved by this system call are as follows: * Through Userspace RCU library (http://urcu.so) - DNS server (Knot DNS) https://www.knot-dns.cz/ - Network sniffer (http://netsniff-ng.org/) - Distributed object storage (https://sheepdog.github.io/sheepdog/) - User-space tracing (http://lttng.org) - Network storage system (https://www.gluster.org/) - Virtual routers (https://events.linuxfoundation.org/sites/events/files/slides/DPDK_RCU_0MQ.pdf) - Financial software (https://lkml.org/lkml/2015/3/23/189) Those projects use RCU in userspace to increase read-side speed and scalability compared to locking. Especially in the case of RCU used by libraries, sys_membarrier can speed up the read-side by moving the bulk of the memory barrier cost to synchronize_rcu(). * Direct users of sys_membarrier - core dotnet garbage collector (https://github.com/dotnet/coreclr/issues/198) Microsoft core dotnet GC developers are planning to use the mprotect() side-effect of issuing memory barriers through IPIs as a way to implement Windows FlushProcessWriteBuffers() on Linux. They are referring to sys_membarrier in their github thread, specifically stating that sys_membarrier() is what they are looking for. To explain the benefit of this scheme, let's introduce two example threads: Thread A (non-frequent, e.g. executing liburcu synchronize_rcu()) Thread B (frequent, e.g. executing liburcu rcu_read_lock()/rcu_read_unlock()) In a scheme where all smp_mb() in thread A are ordering memory accesses with respect to smp_mb() present in Thread B, we can change each smp_mb() within Thread A into calls to sys_membarrier() and each smp_mb() within Thread B into compiler barriers "barrier()". Before the change, we had, for each smp_mb() pairs: Thread A Thread B previous mem accesses previous mem accesses smp_mb() smp_mb() following mem accesses following mem accesses After the change, these pairs become: Thread A Thread B prev mem accesses prev mem accesses sys_membarrier() barrier() follow mem accesses follow mem accesses As we can see, there are two possible scenarios: either Thread B memory accesses do not happen concurrently with Thread A accesses (1), or they do (2). 1) Non-concurrent Thread A vs Thread B accesses: Thread A Thread B prev mem accesses sys_membarrier() follow mem accesses prev mem accesses barrier() follow mem accesses In this case, thread B accesses will be weakly ordered. This is OK, because at that point, thread A is not particularly interested in ordering them with respect to its own accesses. 2) Concurrent Thread A vs Thread B accesses Thread A Thread B prev mem accesses prev mem accesses sys_membarrier() barrier() follow mem accesses follow mem accesses In this case, thread B accesses, which are ensured to be in program order thanks to the compiler barrier, will be "upgraded" to full smp_mb() by synchronize_sched(). * Benchmarks On Intel Xeon E5405 (8 cores) (one thread is calling sys_membarrier, the other 7 threads are busy looping) 1000 non-expedited sys_membarrier calls in 33s =3D 33 milliseconds/call. * User-space user of this system call: Userspace RCU library Both the signal-based and the sys_membarrier userspace RCU schemes permit us to remove the memory barrier from the userspace RCU rcu_read_lock() and rcu_read_unlock() primitives, thus significantly accelerating them. These memory barriers are replaced by compiler barriers on the read-side, and all matching memory barriers on the write-side are turned into an invocation of a memory barrier on all active threads in the process. By letting the kernel perform this synchronization rather than dumbly sending a signal to every process threads (as we currently do), we diminish the number of unnecessary wake ups and only issue the memory barriers on active threads. Non-running threads do not need to execute such barrier anyway, because these are implied by the scheduler context switches. Results in liburcu: Operations in 10s, 6 readers, 2 writers: memory barriers in reader: 1701557485 reads, 2202847 writes signal-based scheme: 9830061167 reads, 6700 writes sys_membarrier: 9952759104 reads, 425 writes sys_membarrier (dyn. check): 7970328887 reads, 425 writes The dynamic sys_membarrier availability check adds some overhead to the read-side compared to the signal-based scheme, but besides that, sys_membarrier slightly outperforms the signal-based scheme. However, this non-expedited sys_membarrier implementation has a much slower grace period than signal and memory barrier schemes. Besides diminishing the number of wake-ups, one major advantage of the membarrier system call over the signal-based scheme is that it does not need to reserve a signal. This plays much more nicely with libraries, and with processes injected into for tracing purposes, for which we cannot expect that signals will be unused by the application. An expedited version of this system call can be added later on to speed up the grace period. Its implementation will likely depend on reading the cpu_curr()->mm without holding each CPU's rq lock. This patch adds the system call to x86 and to asm-generic. [1] http://urcu.so membarrier(2) man page: MEMBARRIER(2) Linux Programmer's Manual MEMBARRIER(2) NAME membarrier - issue memory barriers on a set of threads SYNOPSIS #include <linux/membarrier.h> int membarrier(int cmd, int flags); DESCRIPTION The cmd argument is one of the following: MEMBARRIER_CMD_QUERY Query the set of supported commands. It returns a bitmask of supported commands. MEMBARRIER_CMD_SHARED Execute a memory barrier on all threads running on the system. Upon return from system call, the caller thread is ensured that all running threads have passed through a state where all memory accesses to user-space addresses match program order between entry to and return from the system call (non-running threads are de facto in such a state). This covers threads from all pro=E2=80=90 cesses running on the system. This command returns 0. The flags argument needs to be 0. For future extensions. All memory accesses performed in program order from each targeted thread is guaranteed to be ordered with respect to sys_membarrier(). If we use the semantic "barrier()" to represent a compiler barrier forcing memory accesses to be performed in program order across the barrier, and smp_mb() to represent explicit memory barriers forcing full memory ordering across the barrier, we have the following ordering table for each pair of barrier(), sys_membarrier() and smp_mb(): The pair ordering is detailed as (O: ordered, X: not ordered): barrier() smp_mb() sys_membarrier() barrier() X X O smp_mb() X O O sys_membarrier() O O O RETURN VALUE On success, these system calls return zero. On error, -1 is returned, and errno is set appropriately. For a given command, with flags argument set to 0, this system call is guaranteed to always return the same value until reboot. ERRORS ENOSYS System call is not implemented. EINVAL Invalid arguments. Linux 2015-04-15 MEMBARRIER(2) Signed-off-by: Mathieu Desnoyers <mathieu.desnoyers@efficios.com> Reviewed-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Reviewed-by: Josh Triplett <josh@joshtriplett.org> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Nicholas Miell <nmiell@comcast.net> Cc: Ingo Molnar <mingo@redhat.com> Cc: Alan Cox <gnomes@lxorguk.ukuu.org.uk> Cc: Lai Jiangshan <laijs@cn.fujitsu.com> Cc: Stephen Hemminger <stephen@networkplumber.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Peter Zijlstra <peterz@infradead.org> Cc: David Howells <dhowells@redhat.com> Cc: Pranith Kumar <bobby.prani@gmail.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Shuah Khan <shuahkh@osg.samsung.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-09-11 20:07:39 +00:00
COND_SYSCALL(userfaultfd);
sys_membarrier(): system-wide memory barrier (generic, x86) Here is an implementation of a new system call, sys_membarrier(), which executes a memory barrier on all threads running on the system. It is implemented by calling synchronize_sched(). It can be used to distribute the cost of user-space memory barriers asymmetrically by transforming pairs of memory barriers into pairs consisting of sys_membarrier() and a compiler barrier. For synchronization primitives that distinguish between read-side and write-side (e.g. userspace RCU [1], rwlocks), the read-side can be accelerated significantly by moving the bulk of the memory barrier overhead to the write-side. The existing applications of which I am aware that would be improved by this system call are as follows: * Through Userspace RCU library (http://urcu.so) - DNS server (Knot DNS) https://www.knot-dns.cz/ - Network sniffer (http://netsniff-ng.org/) - Distributed object storage (https://sheepdog.github.io/sheepdog/) - User-space tracing (http://lttng.org) - Network storage system (https://www.gluster.org/) - Virtual routers (https://events.linuxfoundation.org/sites/events/files/slides/DPDK_RCU_0MQ.pdf) - Financial software (https://lkml.org/lkml/2015/3/23/189) Those projects use RCU in userspace to increase read-side speed and scalability compared to locking. Especially in the case of RCU used by libraries, sys_membarrier can speed up the read-side by moving the bulk of the memory barrier cost to synchronize_rcu(). * Direct users of sys_membarrier - core dotnet garbage collector (https://github.com/dotnet/coreclr/issues/198) Microsoft core dotnet GC developers are planning to use the mprotect() side-effect of issuing memory barriers through IPIs as a way to implement Windows FlushProcessWriteBuffers() on Linux. They are referring to sys_membarrier in their github thread, specifically stating that sys_membarrier() is what they are looking for. To explain the benefit of this scheme, let's introduce two example threads: Thread A (non-frequent, e.g. executing liburcu synchronize_rcu()) Thread B (frequent, e.g. executing liburcu rcu_read_lock()/rcu_read_unlock()) In a scheme where all smp_mb() in thread A are ordering memory accesses with respect to smp_mb() present in Thread B, we can change each smp_mb() within Thread A into calls to sys_membarrier() and each smp_mb() within Thread B into compiler barriers "barrier()". Before the change, we had, for each smp_mb() pairs: Thread A Thread B previous mem accesses previous mem accesses smp_mb() smp_mb() following mem accesses following mem accesses After the change, these pairs become: Thread A Thread B prev mem accesses prev mem accesses sys_membarrier() barrier() follow mem accesses follow mem accesses As we can see, there are two possible scenarios: either Thread B memory accesses do not happen concurrently with Thread A accesses (1), or they do (2). 1) Non-concurrent Thread A vs Thread B accesses: Thread A Thread B prev mem accesses sys_membarrier() follow mem accesses prev mem accesses barrier() follow mem accesses In this case, thread B accesses will be weakly ordered. This is OK, because at that point, thread A is not particularly interested in ordering them with respect to its own accesses. 2) Concurrent Thread A vs Thread B accesses Thread A Thread B prev mem accesses prev mem accesses sys_membarrier() barrier() follow mem accesses follow mem accesses In this case, thread B accesses, which are ensured to be in program order thanks to the compiler barrier, will be "upgraded" to full smp_mb() by synchronize_sched(). * Benchmarks On Intel Xeon E5405 (8 cores) (one thread is calling sys_membarrier, the other 7 threads are busy looping) 1000 non-expedited sys_membarrier calls in 33s =3D 33 milliseconds/call. * User-space user of this system call: Userspace RCU library Both the signal-based and the sys_membarrier userspace RCU schemes permit us to remove the memory barrier from the userspace RCU rcu_read_lock() and rcu_read_unlock() primitives, thus significantly accelerating them. These memory barriers are replaced by compiler barriers on the read-side, and all matching memory barriers on the write-side are turned into an invocation of a memory barrier on all active threads in the process. By letting the kernel perform this synchronization rather than dumbly sending a signal to every process threads (as we currently do), we diminish the number of unnecessary wake ups and only issue the memory barriers on active threads. Non-running threads do not need to execute such barrier anyway, because these are implied by the scheduler context switches. Results in liburcu: Operations in 10s, 6 readers, 2 writers: memory barriers in reader: 1701557485 reads, 2202847 writes signal-based scheme: 9830061167 reads, 6700 writes sys_membarrier: 9952759104 reads, 425 writes sys_membarrier (dyn. check): 7970328887 reads, 425 writes The dynamic sys_membarrier availability check adds some overhead to the read-side compared to the signal-based scheme, but besides that, sys_membarrier slightly outperforms the signal-based scheme. However, this non-expedited sys_membarrier implementation has a much slower grace period than signal and memory barrier schemes. Besides diminishing the number of wake-ups, one major advantage of the membarrier system call over the signal-based scheme is that it does not need to reserve a signal. This plays much more nicely with libraries, and with processes injected into for tracing purposes, for which we cannot expect that signals will be unused by the application. An expedited version of this system call can be added later on to speed up the grace period. Its implementation will likely depend on reading the cpu_curr()->mm without holding each CPU's rq lock. This patch adds the system call to x86 and to asm-generic. [1] http://urcu.so membarrier(2) man page: MEMBARRIER(2) Linux Programmer's Manual MEMBARRIER(2) NAME membarrier - issue memory barriers on a set of threads SYNOPSIS #include <linux/membarrier.h> int membarrier(int cmd, int flags); DESCRIPTION The cmd argument is one of the following: MEMBARRIER_CMD_QUERY Query the set of supported commands. It returns a bitmask of supported commands. MEMBARRIER_CMD_SHARED Execute a memory barrier on all threads running on the system. Upon return from system call, the caller thread is ensured that all running threads have passed through a state where all memory accesses to user-space addresses match program order between entry to and return from the system call (non-running threads are de facto in such a state). This covers threads from all pro=E2=80=90 cesses running on the system. This command returns 0. The flags argument needs to be 0. For future extensions. All memory accesses performed in program order from each targeted thread is guaranteed to be ordered with respect to sys_membarrier(). If we use the semantic "barrier()" to represent a compiler barrier forcing memory accesses to be performed in program order across the barrier, and smp_mb() to represent explicit memory barriers forcing full memory ordering across the barrier, we have the following ordering table for each pair of barrier(), sys_membarrier() and smp_mb(): The pair ordering is detailed as (O: ordered, X: not ordered): barrier() smp_mb() sys_membarrier() barrier() X X O smp_mb() X O O sys_membarrier() O O O RETURN VALUE On success, these system calls return zero. On error, -1 is returned, and errno is set appropriately. For a given command, with flags argument set to 0, this system call is guaranteed to always return the same value until reboot. ERRORS ENOSYS System call is not implemented. EINVAL Invalid arguments. Linux 2015-04-15 MEMBARRIER(2) Signed-off-by: Mathieu Desnoyers <mathieu.desnoyers@efficios.com> Reviewed-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Reviewed-by: Josh Triplett <josh@joshtriplett.org> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Nicholas Miell <nmiell@comcast.net> Cc: Ingo Molnar <mingo@redhat.com> Cc: Alan Cox <gnomes@lxorguk.ukuu.org.uk> Cc: Lai Jiangshan <laijs@cn.fujitsu.com> Cc: Stephen Hemminger <stephen@networkplumber.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Peter Zijlstra <peterz@infradead.org> Cc: David Howells <dhowells@redhat.com> Cc: Pranith Kumar <bobby.prani@gmail.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Shuah Khan <shuahkh@osg.samsung.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-09-11 20:07:39 +00:00
/* membarrier */
COND_SYSCALL(membarrier);
COND_SYSCALL(mlock2);
COND_SYSCALL(copy_file_range);
/* memory protection keys */
COND_SYSCALL(pkey_mprotect);
COND_SYSCALL(pkey_alloc);
COND_SYSCALL(pkey_free);
mm: introduce memfd_secret system call to create "secret" memory areas Introduce "memfd_secret" system call with the ability to create memory areas visible only in the context of the owning process and not mapped not only to other processes but in the kernel page tables as well. The secretmem feature is off by default and the user must explicitly enable it at the boot time. Once secretmem is enabled, the user will be able to create a file descriptor using the memfd_secret() system call. The memory areas created by mmap() calls from this file descriptor will be unmapped from the kernel direct map and they will be only mapped in the page table of the processes that have access to the file descriptor. Secretmem is designed to provide the following protections: * Enhanced protection (in conjunction with all the other in-kernel attack prevention systems) against ROP attacks. Seceretmem makes "simple" ROP insufficient to perform exfiltration, which increases the required complexity of the attack. Along with other protections like the kernel stack size limit and address space layout randomization which make finding gadgets is really hard, absence of any in-kernel primitive for accessing secret memory means the one gadget ROP attack can't work. Since the only way to access secret memory is to reconstruct the missing mapping entry, the attacker has to recover the physical page and insert a PTE pointing to it in the kernel and then retrieve the contents. That takes at least three gadgets which is a level of difficulty beyond most standard attacks. * Prevent cross-process secret userspace memory exposures. Once the secret memory is allocated, the user can't accidentally pass it into the kernel to be transmitted somewhere. The secreremem pages cannot be accessed via the direct map and they are disallowed in GUP. * Harden against exploited kernel flaws. In order to access secretmem, a kernel-side attack would need to either walk the page tables and create new ones, or spawn a new privileged uiserspace process to perform secrets exfiltration using ptrace. The file descriptor based memory has several advantages over the "traditional" mm interfaces, such as mlock(), mprotect(), madvise(). File descriptor approach allows explicit and controlled sharing of the memory areas, it allows to seal the operations. Besides, file descriptor based memory paves the way for VMMs to remove the secret memory range from the userspace hipervisor process, for instance QEMU. Andy Lutomirski says: "Getting fd-backed memory into a guest will take some possibly major work in the kernel, but getting vma-backed memory into a guest without mapping it in the host user address space seems much, much worse." memfd_secret() is made a dedicated system call rather than an extension to memfd_create() because it's purpose is to allow the user to create more secure memory mappings rather than to simply allow file based access to the memory. Nowadays a new system call cost is negligible while it is way simpler for userspace to deal with a clear-cut system calls than with a multiplexer or an overloaded syscall. Moreover, the initial implementation of memfd_secret() is completely distinct from memfd_create() so there is no much sense in overloading memfd_create() to begin with. If there will be a need for code sharing between these implementation it can be easily achieved without a need to adjust user visible APIs. The secret memory remains accessible in the process context using uaccess primitives, but it is not exposed to the kernel otherwise; secret memory areas are removed from the direct map and functions in the follow_page()/get_user_page() family will refuse to return a page that belongs to the secret memory area. Once there will be a use case that will require exposing secretmem to the kernel it will be an opt-in request in the system call flags so that user would have to decide what data can be exposed to the kernel. Removing of the pages from the direct map may cause its fragmentation on architectures that use large pages to map the physical memory which affects the system performance. However, the original Kconfig text for CONFIG_DIRECT_GBPAGES said that gigabyte pages in the direct map "... can improve the kernel's performance a tiny bit ..." (commit 00d1c5e05736 ("x86: add gbpages switches")) and the recent report [1] showed that "... although 1G mappings are a good default choice, there is no compelling evidence that it must be the only choice". Hence, it is sufficient to have secretmem disabled by default with the ability of a system administrator to enable it at boot time. Pages in the secretmem regions are unevictable and unmovable to avoid accidental exposure of the sensitive data via swap or during page migration. Since the secretmem mappings are locked in memory they cannot exceed RLIMIT_MEMLOCK. Since these mappings are already locked independently from mlock(), an attempt to mlock()/munlock() secretmem range would fail and mlockall()/munlockall() will ignore secretmem mappings. However, unlike mlock()ed memory, secretmem currently behaves more like long-term GUP: secretmem mappings are unmovable mappings directly consumed by user space. With default limits, there is no excessive use of secretmem and it poses no real problem in combination with ZONE_MOVABLE/CMA, but in the future this should be addressed to allow balanced use of large amounts of secretmem along with ZONE_MOVABLE/CMA. A page that was a part of the secret memory area is cleared when it is freed to ensure the data is not exposed to the next user of that page. The following example demonstrates creation of a secret mapping (error handling is omitted): fd = memfd_secret(0); ftruncate(fd, MAP_SIZE); ptr = mmap(NULL, MAP_SIZE, PROT_READ | PROT_WRITE, MAP_SHARED, fd, 0); [1] https://lore.kernel.org/linux-mm/213b4567-46ce-f116-9cdf-bbd0c884eb3c@linux.intel.com/ [akpm@linux-foundation.org: suppress Kconfig whine] Link: https://lkml.kernel.org/r/20210518072034.31572-5-rppt@kernel.org Signed-off-by: Mike Rapoport <rppt@linux.ibm.com> Acked-by: Hagen Paul Pfeifer <hagen@jauu.net> Acked-by: James Bottomley <James.Bottomley@HansenPartnership.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Andy Lutomirski <luto@kernel.org> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Borislav Petkov <bp@alien8.de> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Christopher Lameter <cl@linux.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Elena Reshetova <elena.reshetova@intel.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: James Bottomley <jejb@linux.ibm.com> Cc: "Kirill A. Shutemov" <kirill@shutemov.name> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mark Rutland <mark.rutland@arm.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Palmer Dabbelt <palmer@dabbelt.com> Cc: Palmer Dabbelt <palmerdabbelt@google.com> Cc: Paul Walmsley <paul.walmsley@sifive.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rick Edgecombe <rick.p.edgecombe@intel.com> Cc: Roman Gushchin <guro@fb.com> Cc: Shakeel Butt <shakeelb@google.com> Cc: Shuah Khan <shuah@kernel.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Tycho Andersen <tycho@tycho.ws> Cc: Will Deacon <will@kernel.org> Cc: David Hildenbrand <david@redhat.com> Cc: kernel test robot <lkp@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-07-08 01:08:03 +00:00
/* memfd_secret */
COND_SYSCALL(memfd_secret);
/*
* Architecture specific weak syscall entries.
*/
/* pciconfig: alpha, arm, arm64, ia64, sparc */
COND_SYSCALL(pciconfig_read);
COND_SYSCALL(pciconfig_write);
COND_SYSCALL(pciconfig_iobase);
/* sys_socketcall: arm, mips, x86, ... */
COND_SYSCALL(socketcall);
COND_SYSCALL_COMPAT(socketcall);
/* compat syscalls for arm64, x86, ... */
COND_SYSCALL_COMPAT(fanotify_mark);
/* x86 */
COND_SYSCALL(vm86old);
COND_SYSCALL(modify_ldt);
COND_SYSCALL(vm86);
COND_SYSCALL(kexec_file_load);
x86/shstk: Introduce map_shadow_stack syscall When operating with shadow stacks enabled, the kernel will automatically allocate shadow stacks for new threads, however in some cases userspace will need additional shadow stacks. The main example of this is the ucontext family of functions, which require userspace allocating and pivoting to userspace managed stacks. Unlike most other user memory permissions, shadow stacks need to be provisioned with special data in order to be useful. They need to be setup with a restore token so that userspace can pivot to them via the RSTORSSP instruction. But, the security design of shadow stacks is that they should not be written to except in limited circumstances. This presents a problem for userspace, as to how userspace can provision this special data, without allowing for the shadow stack to be generally writable. Previously, a new PROT_SHADOW_STACK was attempted, which could be mprotect()ed from RW permissions after the data was provisioned. This was found to not be secure enough, as other threads could write to the shadow stack during the writable window. The kernel can use a special instruction, WRUSS, to write directly to userspace shadow stacks. So the solution can be that memory can be mapped as shadow stack permissions from the beginning (never generally writable in userspace), and the kernel itself can write the restore token. First, a new madvise() flag was explored, which could operate on the PROT_SHADOW_STACK memory. This had a couple of downsides: 1. Extra checks were needed in mprotect() to prevent writable memory from ever becoming PROT_SHADOW_STACK. 2. Extra checks/vma state were needed in the new madvise() to prevent restore tokens being written into the middle of pre-used shadow stacks. It is ideal to prevent restore tokens being added at arbitrary locations, so the check was to make sure the shadow stack had never been written to. 3. It stood out from the rest of the madvise flags, as more of direct action than a hint at future desired behavior. So rather than repurpose two existing syscalls (mmap, madvise) that don't quite fit, just implement a new map_shadow_stack syscall to allow userspace to map and setup new shadow stacks in one step. While ucontext is the primary motivator, userspace may have other unforeseen reasons to setup its own shadow stacks using the WRSS instruction. Towards this provide a flag so that stacks can be optionally setup securely for the common case of ucontext without enabling WRSS. Or potentially have the kernel set up the shadow stack in some new way. The following example demonstrates how to create a new shadow stack with map_shadow_stack: void *shstk = map_shadow_stack(addr, stack_size, SHADOW_STACK_SET_TOKEN); Signed-off-by: Rick Edgecombe <rick.p.edgecombe@intel.com> Signed-off-by: Dave Hansen <dave.hansen@linux.intel.com> Reviewed-by: Borislav Petkov (AMD) <bp@alien8.de> Reviewed-by: Kees Cook <keescook@chromium.org> Acked-by: Mike Rapoport (IBM) <rppt@kernel.org> Tested-by: Pengfei Xu <pengfei.xu@intel.com> Tested-by: John Allen <john.allen@amd.com> Tested-by: Kees Cook <keescook@chromium.org> Link: https://lore.kernel.org/all/20230613001108.3040476-35-rick.p.edgecombe%40intel.com
2023-06-13 00:11:00 +00:00
COND_SYSCALL(map_shadow_stack);
/* s390 */
COND_SYSCALL(s390_pci_mmio_read);
COND_SYSCALL(s390_pci_mmio_write);
COND_SYSCALL(s390_ipc);
COND_SYSCALL_COMPAT(s390_ipc);
/* powerpc */
COND_SYSCALL(rtas);
COND_SYSCALL(spu_run);
COND_SYSCALL(spu_create);
COND_SYSCALL(subpage_prot);
/*
* Deprecated system calls which are still defined in
* include/uapi/asm-generic/unistd.h and wanted by >= 1 arch
*/
/* __ARCH_WANT_SYSCALL_NO_FLAGS */
COND_SYSCALL(epoll_create);
COND_SYSCALL(inotify_init);
COND_SYSCALL(eventfd);
COND_SYSCALL(signalfd);
COND_SYSCALL_COMPAT(signalfd);
/* __ARCH_WANT_SYSCALL_OFF_T */
COND_SYSCALL(fadvise64);
/* __ARCH_WANT_SYSCALL_DEPRECATED */
COND_SYSCALL(epoll_wait);
COND_SYSCALL(recv);
COND_SYSCALL_COMPAT(recv);
COND_SYSCALL(send);
COND_SYSCALL(uselib);
/* optional: time32 */
COND_SYSCALL(time32);
COND_SYSCALL(stime32);
COND_SYSCALL(utime32);
COND_SYSCALL(adjtimex_time32);
COND_SYSCALL(sched_rr_get_interval_time32);
COND_SYSCALL(nanosleep_time32);
COND_SYSCALL(rt_sigtimedwait_time32);
COND_SYSCALL_COMPAT(rt_sigtimedwait_time32);
COND_SYSCALL(timer_settime32);
COND_SYSCALL(timer_gettime32);
COND_SYSCALL(clock_settime32);
COND_SYSCALL(clock_gettime32);
COND_SYSCALL(clock_getres_time32);
COND_SYSCALL(clock_nanosleep_time32);
COND_SYSCALL(utimes_time32);
COND_SYSCALL(futimesat_time32);
COND_SYSCALL(pselect6_time32);
COND_SYSCALL_COMPAT(pselect6_time32);
COND_SYSCALL(ppoll_time32);
COND_SYSCALL_COMPAT(ppoll_time32);
COND_SYSCALL(utimensat_time32);
COND_SYSCALL(clock_adjtime32);
/*
* The syscalls below are not found in include/uapi/asm-generic/unistd.h
*/
/* obsolete: SGETMASK_SYSCALL */
COND_SYSCALL(sgetmask);
COND_SYSCALL(ssetmask);
/* obsolete: SYSFS_SYSCALL */
COND_SYSCALL(sysfs);
/* obsolete: __ARCH_WANT_SYS_IPC */
COND_SYSCALL(ipc);
COND_SYSCALL_COMPAT(ipc);
/* obsolete: UID16 */
COND_SYSCALL(chown16);
COND_SYSCALL(fchown16);
COND_SYSCALL(getegid16);
COND_SYSCALL(geteuid16);
COND_SYSCALL(getgid16);
COND_SYSCALL(getgroups16);
COND_SYSCALL(getresgid16);
COND_SYSCALL(getresuid16);
COND_SYSCALL(getuid16);
COND_SYSCALL(lchown16);
COND_SYSCALL(setfsgid16);
COND_SYSCALL(setfsuid16);
COND_SYSCALL(setgid16);
COND_SYSCALL(setgroups16);
COND_SYSCALL(setregid16);
COND_SYSCALL(setresgid16);
COND_SYSCALL(setresuid16);
COND_SYSCALL(setreuid16);
COND_SYSCALL(setuid16);
rseq: Introduce restartable sequences system call Expose a new system call allowing each thread to register one userspace memory area to be used as an ABI between kernel and user-space for two purposes: user-space restartable sequences and quick access to read the current CPU number value from user-space. * Restartable sequences (per-cpu atomics) Restartables sequences allow user-space to perform update operations on per-cpu data without requiring heavy-weight atomic operations. The restartable critical sections (percpu atomics) work has been started by Paul Turner and Andrew Hunter. It lets the kernel handle restart of critical sections. [1] [2] The re-implementation proposed here brings a few simplifications to the ABI which facilitates porting to other architectures and speeds up the user-space fast path. Here are benchmarks of various rseq use-cases. Test hardware: arm32: ARMv7 Processor rev 4 (v7l) "Cubietruck", 2-core x86-64: Intel E5-2630 v3@2.40GHz, 16-core, hyperthreading The following benchmarks were all performed on a single thread. * Per-CPU statistic counter increment getcpu+atomic (ns/op) rseq (ns/op) speedup arm32: 344.0 31.4 11.0 x86-64: 15.3 2.0 7.7 * LTTng-UST: write event 32-bit header, 32-bit payload into tracer per-cpu buffer getcpu+atomic (ns/op) rseq (ns/op) speedup arm32: 2502.0 2250.0 1.1 x86-64: 117.4 98.0 1.2 * liburcu percpu: lock-unlock pair, dereference, read/compare word getcpu+atomic (ns/op) rseq (ns/op) speedup arm32: 751.0 128.5 5.8 x86-64: 53.4 28.6 1.9 * jemalloc memory allocator adapted to use rseq Using rseq with per-cpu memory pools in jemalloc at Facebook (based on rseq 2016 implementation): The production workload response-time has 1-2% gain avg. latency, and the P99 overall latency drops by 2-3%. * Reading the current CPU number Speeding up reading the current CPU number on which the caller thread is running is done by keeping the current CPU number up do date within the cpu_id field of the memory area registered by the thread. This is done by making scheduler preemption set the TIF_NOTIFY_RESUME flag on the current thread. Upon return to user-space, a notify-resume handler updates the current CPU value within the registered user-space memory area. User-space can then read the current CPU number directly from memory. Keeping the current cpu id in a memory area shared between kernel and user-space is an improvement over current mechanisms available to read the current CPU number, which has the following benefits over alternative approaches: - 35x speedup on ARM vs system call through glibc - 20x speedup on x86 compared to calling glibc, which calls vdso executing a "lsl" instruction, - 14x speedup on x86 compared to inlined "lsl" instruction, - Unlike vdso approaches, this cpu_id value can be read from an inline assembly, which makes it a useful building block for restartable sequences. - The approach of reading the cpu id through memory mapping shared between kernel and user-space is portable (e.g. ARM), which is not the case for the lsl-based x86 vdso. On x86, yet another possible approach would be to use the gs segment selector to point to user-space per-cpu data. This approach performs similarly to the cpu id cache, but it has two disadvantages: it is not portable, and it is incompatible with existing applications already using the gs segment selector for other purposes. Benchmarking various approaches for reading the current CPU number: ARMv7 Processor rev 4 (v7l) Machine model: Cubietruck - Baseline (empty loop): 8.4 ns - Read CPU from rseq cpu_id: 16.7 ns - Read CPU from rseq cpu_id (lazy register): 19.8 ns - glibc 2.19-0ubuntu6.6 getcpu: 301.8 ns - getcpu system call: 234.9 ns x86-64 Intel(R) Xeon(R) CPU E5-2630 v3 @ 2.40GHz: - Baseline (empty loop): 0.8 ns - Read CPU from rseq cpu_id: 0.8 ns - Read CPU from rseq cpu_id (lazy register): 0.8 ns - Read using gs segment selector: 0.8 ns - "lsl" inline assembly: 13.0 ns - glibc 2.19-0ubuntu6 getcpu: 16.6 ns - getcpu system call: 53.9 ns - Speed (benchmark taken on v8 of patchset) Running 10 runs of hackbench -l 100000 seems to indicate, contrary to expectations, that enabling CONFIG_RSEQ slightly accelerates the scheduler: Configuration: 2 sockets * 8-core Intel(R) Xeon(R) CPU E5-2630 v3 @ 2.40GHz (directly on hardware, hyperthreading disabled in BIOS, energy saving disabled in BIOS, turboboost disabled in BIOS, cpuidle.off=1 kernel parameter), with a Linux v4.6 defconfig+localyesconfig, restartable sequences series applied. * CONFIG_RSEQ=n avg.: 41.37 s std.dev.: 0.36 s * CONFIG_RSEQ=y avg.: 40.46 s std.dev.: 0.33 s - Size On x86-64, between CONFIG_RSEQ=n/y, the text size increase of vmlinux is 567 bytes, and the data size increase of vmlinux is 5696 bytes. [1] https://lwn.net/Articles/650333/ [2] http://www.linuxplumbersconf.org/2013/ocw/system/presentations/1695/original/LPC%20-%20PerCpu%20Atomics.pdf Signed-off-by: Mathieu Desnoyers <mathieu.desnoyers@efficios.com> Signed-off-by: Thomas Gleixner <tglx@linutronix.de> Acked-by: Peter Zijlstra (Intel) <peterz@infradead.org> Cc: Joel Fernandes <joelaf@google.com> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Dave Watson <davejwatson@fb.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Andi Kleen <andi@firstfloor.org> Cc: "H . Peter Anvin" <hpa@zytor.com> Cc: Chris Lameter <cl@linux.com> Cc: Russell King <linux@arm.linux.org.uk> Cc: Andrew Hunter <ahh@google.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: "Paul E . McKenney" <paulmck@linux.vnet.ibm.com> Cc: Paul Turner <pjt@google.com> Cc: Boqun Feng <boqun.feng@gmail.com> Cc: Josh Triplett <josh@joshtriplett.org> Cc: Steven Rostedt <rostedt@goodmis.org> Cc: Ben Maurer <bmaurer@fb.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: linux-api@vger.kernel.org Cc: Andy Lutomirski <luto@amacapital.net> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Linus Torvalds <torvalds@linux-foundation.org> Link: http://lkml.kernel.org/r/20151027235635.16059.11630.stgit@pjt-glaptop.roam.corp.google.com Link: http://lkml.kernel.org/r/20150624222609.6116.86035.stgit@kitami.mtv.corp.google.com Link: https://lkml.kernel.org/r/20180602124408.8430-3-mathieu.desnoyers@efficios.com
2018-06-02 12:43:54 +00:00
/* restartable sequence */
COND_SYSCALL(rseq);