- 10 Oct, 2014 40 commits
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Oleg Nesterov authored
Trivial cleanup. alloc_pages_vma() can use mpol_cond_put(). Signed-off-by:
Oleg Nesterov <oleg@redhat.com> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: David Rientjes <rientjes@google.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Cyrill Gorcunov <gorcunov@openvz.org> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: "Kirill A. Shutemov" <kirill.shutemov@linux.intel.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Hugh Dickins <hughd@google.com> Cc: Andi Kleen <andi@firstfloor.org> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by:
Andrew Morton <akpm@linux-foundation.org> Signed-off-by:
Linus Torvalds <torvalds@linux-foundation.org>
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Johannes Weiner authored
The deprecation warnings for the scan_unevictable interface triggers by scripts doing `sysctl -a | grep something else'. This is annoying and not helpful. The interface has been defunct since 264e56d8 ("mm: disable user interface to manually rescue unevictable pages"), which was in 2011, and there haven't been any reports of usecases for it, only reports that the deprecation warnings are annying. It's unlikely that anybody is using this interface specifically at this point, so remove it. Signed-off-by:
Johannes Weiner <hannes@cmpxchg.org> Signed-off-by:
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Cyrill Gorcunov authored
During development of c/r we've noticed that in case if we need to support user namespaces we face a problem with capabilities in prctl(PR_SET_MM, ...) call, in particular once new user namespace is created capable(CAP_SYS_RESOURCE) no longer passes. A approach is to eliminate CAP_SYS_RESOURCE check but pass all new values in one bundle, which would allow the kernel to make more intensive test for sanity of values and same time allow us to support checkpoint/restore of user namespaces. Thus a new command PR_SET_MM_MAP introduced. It takes a pointer of prctl_mm_map structure which carries all the members to be updated. prctl(PR_SET_MM, PR_SET_MM_MAP, struct prctl_mm_map *, size) struct prctl_mm_map { __u64 start_code; __u64 end_code; __u64 start_data; __u64 end_data; __u64 start_brk; __u64 brk; __u64 start_stack; __u64 arg_start; __u64 arg_end; __u64 env_start; __u64 env_end; __u64 *auxv; __u32 auxv_size; __u32 exe_fd; }; All members except @exe_fd correspond ones of struct mm_struct. To figure out which available values these members may take here are meanings of the members. - start_code, end_code: represent bounds of executable code area - start_data, end_data: represent bounds of data area - start_brk, brk: used to calculate bounds for brk() syscall - start_stack: used when accounting space needed for command line arguments, environment and shmat() syscall - arg_start, arg_end, env_start, env_end: represent memory area supplied for command line arguments and environment variables - auxv, auxv_size: carries auxiliary vector, Elf format specifics - exe_fd: file descriptor number for executable link (/proc/self/exe) Thus we apply the following requirements to the values 1) Any member except @auxv, @auxv_size, @exe_fd is rather an address in user space thus it must be laying inside [mmap_min_addr, mmap_max_addr) interval. 2) While @[start|end]_code and @[start|end]_data may point to an nonexisting VMAs (say a program maps own new .text and .data segments during execution) the rest of members should belong to VMA which must exist. 3) Addresses must be ordered, ie @start_ member must not be greater or equal to appropriate @end_ member. 4) As in regular Elf loading procedure we require that @start_brk and @brk be greater than @end_data. 5) If RLIMIT_DATA rlimit is set to non-infinity new values should not exceed existing limit. Same applies to RLIMIT_STACK. 6) Auxiliary vector size must not exceed existing one (which is predefined as AT_VECTOR_SIZE and depends on architecture). 7) File descriptor passed in @exe_file should be pointing to executable file (because we use existing prctl_set_mm_exe_file_locked helper it ensures that the file we are going to use as exe link has all required permission granted). Now about where these members are involved inside kernel code: - @start_code and @end_code are used in /proc/$pid/[stat|statm] output; - @start_data and @end_data are used in /proc/$pid/[stat|statm] output, also they are considered if there enough space for brk() syscall result if RLIMIT_DATA is set; - @start_brk shown in /proc/$pid/stat output and accounted in brk() syscall if RLIMIT_DATA is set; also this member is tested to find a symbolic name of mmap event for perf system (we choose if event is generated for "heap" area); one more aplication is selinux -- we test if a process has PROCESS__EXECHEAP permission if trying to make heap area being executable with mprotect() syscall; - @brk is a current value for brk() syscall which lays inside heap area, it's shown in /proc/$pid/stat. When syscall brk() succesfully provides new memory area to a user space upon brk() completion the mm::brk is updated to carry new value; Both @start_brk and @brk are actively used in /proc/$pid/maps and /proc/$pid/smaps output to find a symbolic name "heap" for VMA being scanned; - @start_stack is printed out in /proc/$pid/stat and used to find a symbolic name "stack" for task and threads in /proc/$pid/maps and /proc/$pid/smaps output, and as the same as with @start_brk -- perf system uses it for event naming. Also kernel treat this member as a start address of where to map vDSO pages and to check if there is enough space for shmat() syscall; - @arg_start, @arg_end, @env_start and @env_end are printed out in /proc/$pid/stat. Another access to the data these members represent is to read /proc/$pid/environ or /proc/$pid/cmdline. Any attempt to read these areas kernel tests with access_process_vm helper so a user must have enough rights for this action; - @auxv and @auxv_size may be read from /proc/$pid/auxv. Strictly speaking kernel doesn't care much about which exactly data is sitting there because it is solely for userspace; - @exe_fd is referred from /proc/$pid/exe and when generating coredump. We uses prctl_set_mm_exe_file_locked helper to update this member, so exe-file link modification remains one-shot action. Still note that updating exe-file link now doesn't require sys-resource capability anymore, after all there is no much profit in preventing setup own file link (there are a number of ways to execute own code -- ptrace, ld-preload, so that the only reliable way to find which exactly code is executed is to inspect running program memory). Still we require the caller to be at least user-namespace root user. I believe the old interface should be deprecated and ripped off in a couple of kernel releases if no one against. To test if new interface is implemented in the kernel one can pass PR_SET_MM_MAP_SIZE opcode and the kernel returns the size of currently supported struct prctl_mm_map. [akpm@linux-foundation.org: fix 80-col wordwrap in macro definitions] Signed-off-by:Cyrill Gorcunov <gorcunov@openvz.org> Cc: Kees Cook <keescook@chromium.org> Cc: Tejun Heo <tj@kernel.org> Acked-by:
Andrew Vagin <avagin@openvz.org> Tested-by:
Serge Hallyn <serge.hallyn@canonical.com> Cc: Pavel Emelyanov <xemul@parallels.com> Cc: Vasiliy Kulikov <segoon@openwall.com> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Julien Tinnes <jln@google.com> Signed-off-by:
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Cyrill Gorcunov authored
Instead of taking mm->mmap_sem inside prctl_set_mm_exe_file() move it out and rename the helper to prctl_set_mm_exe_file_locked(). This will allow to reuse this function in a next patch. Signed-off-by:
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Cyrill Gorcunov authored
Signed-off-by:
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Cyrill Gorcunov authored
To eliminate code duplication lets introduce check_data_rlimit helper which we will use in brk() and prctl() syscalls. Signed-off-by:
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David Rientjes authored
struct compact_control currently converts the gfp mask to a migratetype, but we need the entire gfp mask in a follow-up patch. Pass the entire gfp mask as part of struct compact_control. Signed-off-by:
David Rientjes <rientjes@google.com> Signed-off-by:
Vlastimil Babka <vbabka@suse.cz> Reviewed-by:
Zhang Yanfei <zhangyanfei@cn.fujitsu.com> Acked-by:
Minchan Kim <minchan@kernel.org> Acked-by:
Mel Gorman <mgorman@suse.de> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Michal Nazarewicz <mina86@mina86.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: Christoph Lameter <cl@linux.com> Cc: Rik van Riel <riel@redhat.com> Signed-off-by:
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David Rientjes authored
The page allocator has gfp flags (like __GFP_WAIT) and alloc flags (like ALLOC_CPUSET) that have separate semantics. The function allocflags_to_migratetype() actually takes gfp flags, not alloc flags, and returns a migratetype. Rename it to gfpflags_to_migratetype(). Signed-off-by:
Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Acked-by:
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Vlastimil Babka authored
The migration scanner skips PageBuddy pages, but does not consider their order as checking page_order() is generally unsafe without holding the zone->lock, and acquiring the lock just for the check wouldn't be a good tradeoff. Still, this could avoid some iterations over the rest of the buddy page, and if we are careful, the race window between PageBuddy() check and page_order() is small, and the worst thing that can happen is that we skip too much and miss some isolation candidates. This is not that bad, as compaction can already fail for many other reasons like parallel allocations, and those have much larger race window. This patch therefore makes the migration scanner obtain the buddy page order and use it to skip the whole buddy page, if the order appears to be in the valid range. It's important that the page_order() is read only once, so that the value used in the checks and in the pfn calculation is the same. But in theory the compiler can replace the local variable by multiple inlines of page_order(). Therefore, the patch introduces page_order_unsafe() that uses ACCESS_ONCE to prevent this. Testing with stress-highalloc from mmtests shows a 15% reduction in number of pages scanned by migration scanner. The reduction is >60% with __GFP_NO_KSWAPD allocations, along with success rates better by few percent. Signed-off-by:
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Vlastimil Babka authored
Unlike the migration scanner, the free scanner remembers the beginning of the last scanned pageblock in cc->free_pfn. It might be therefore rescanning pages uselessly when called several times during single compaction. This might have been useful when pages were returned to the buddy allocator after a failed migration, but this is no longer the case. This patch changes the meaning of cc->free_pfn so that if it points to a middle of a pageblock, that pageblock is scanned only from cc->free_pfn to the end. isolate_freepages_block() will record the pfn of the last page it looked at, which is then used to update cc->free_pfn. In the mmtests stress-highalloc benchmark, this has resulted in lowering the ratio between pages scanned by both scanners, from 2.5 free pages per migrate page, to 2.25 free pages per migrate page, without affecting success rates. With __GFP_NO_KSWAPD allocations, this appears to result in a worse ratio (2.1 instead of 1.8), but page migration successes increased by 10%, so this could mean that more useful work can be done until need_resched() aborts this kind of compaction. Signed-off-by:
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Vlastimil Babka authored
Compaction scanners try to lock zone locks as late as possible by checking many page or pageblock properties opportunistically without lock and skipping them if not unsuitable. For pages that pass the initial checks, some properties have to be checked again safely under lock. However, if the lock was already held from a previous iteration in the initial checks, the rechecks are unnecessary. This patch therefore skips the rechecks when the lock was already held. This is now possible to do, since we don't (potentially) drop and reacquire the lock between the initial checks and the safe rechecks anymore. Signed-off-by:
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Vlastimil Babka authored
Compaction scanners regularly check for lock contention and need_resched() through the compact_checklock_irqsave() function. However, if there is no contention, the lock can be held and IRQ disabled for potentially long time. This has been addressed by commit b2eef8c0 ("mm: compaction: minimise the time IRQs are disabled while isolating pages for migration") for the migration scanner. However, the refactoring done by commit 2a1402aa ("mm: compaction: acquire the zone->lru_lock as late as possible") has changed the conditions so that the lock is dropped only when there's contention on the lock or need_resched() is true. Also, need_resched() is checked only when the lock is already held. The comment "give a chance to irqs before checking need_resched" is therefore misleading, as IRQs remain disabled when the check is done. This patch restores the behavior intended by commit b2eef8c0 and also tries to better balance and make more deterministic the time spent by checking for contention vs the time the scanners might run between the checks. It also avoids situations where checking has not been done often enough before. The result should be avoiding both too frequent and too infrequent contention checking, and especially the potentially long-running scans with IRQs disabled and no checking of need_resched() or for fatal signal pending, which can happen when many consecutive pages or pageblocks fail the preliminary tests and do not reach the later call site to compact_checklock_irqsave(), as explained below. Before the patch: In the migration scanner, compact_checklock_irqsave() was called each loop, if reached. If not reached, some lower-frequency checking could still be done if the lock was already held, but this would not result in aborting contended async compaction until reaching compact_checklock_irqsave() or end of pageblock. In the free scanner, it was similar but completely without the periodical checking, so lock can be potentially held until reaching the end of pageblock. After the patch, in both scanners: The periodical check is done as the first thing in the loop on each SWAP_CLUSTER_MAX aligned pfn, using the new compact_unlock_should_abort() function, which always unlocks the lock (if locked) and aborts async compaction if scheduling is needed. It also aborts any type of compaction when a fatal signal is pending. The compact_checklock_irqsave() function is replaced with a slightly different compact_trylock_irqsave(). The biggest difference is that the function is not called at all if the lock is already held. The periodical need_resched() checking is left solely to compact_unlock_should_abort(). The lock contention avoidance for async compaction is achieved by the periodical unlock by compact_unlock_should_abort() and by using trylock in compact_trylock_irqsave() and aborting when trylock fails. Sync compaction does not use trylock. Signed-off-by:
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Vlastimil Babka authored
Async compaction aborts when it detects zone lock contention or need_resched() is true. David Rientjes has reported that in practice, most direct async compactions for THP allocation abort due to need_resched(). This means that a second direct compaction is never attempted, which might be OK for a page fault, but khugepaged is intended to attempt a sync compaction in such case and in these cases it won't. This patch replaces "bool contended" in compact_control with an int that distinguishes between aborting due to need_resched() and aborting due to lock contention. This allows propagating the abort through all compaction functions as before, but passing the abort reason up to __alloc_pages_slowpath() which decides when to continue with direct reclaim and another compaction attempt. Another problem is that try_to_compact_pages() did not act upon the reported contention (both need_resched() or lock contention) immediately and would proceed with another zone from the zonelist. When need_resched() is true, that means initializing another zone compaction, only to check again need_resched() in isolate_migratepages() and aborting. For zone lock contention, the unintended consequence is that the lock contended status reported back to the allocator is detrmined from the last zone where compaction was attempted, which is rather arbitrary. This patch fixes the problem in the following way: - async compaction of a zone aborting due to need_resched() or fatal signal pending means that further zones should not be tried. We report COMPACT_CONTENDED_SCHED to the allocator. - aborting zone compaction due to lock contention means we can still try another zone, since it has different set of locks. We report back COMPACT_CONTENDED_LOCK only if *all* zones where compaction was attempted, it was aborted due to lock contention. As a result of these fixes, khugepaged will proceed with second sync compaction as intended, when the preceding async compaction aborted due to need_resched(). Page fault compactions aborting due to need_resched() will spare some cycles previously wasted by initializing another zone compaction only to abort again. Lock contention will be reported only when compaction in all zones aborted due to lock contention, and therefore it's not a good idea to try again after reclaim. In stress-highalloc from mmtests configured to use __GFP_NO_KSWAPD, this has improved number of THP collapse allocations by 10%, which shows positive effect on khugepaged. The benchmark's success rates are unchanged as it is not recognized as khugepaged. Numbers of compact_stall and compact_fail events have however decreased by 20%, with compact_success still a bit improved, which is good. With benchmark configured not to use __GFP_NO_KSWAPD, there is 6% improvement in THP collapse allocations, and only slight improvement in stalls and failures. [akpm@linux-foundation.org: fix warnings] Reported-by:
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Vlastimil Babka authored
The unification of the migrate and free scanner families of function has highlighted a difference in how the scanners ensure they only isolate pages of the intended zone. This is important for taking zone lock or lru lock of the correct zone. Due to nodes overlapping, it is however possible to encounter a different zone within the range of the zone being compacted. The free scanner, since its inception by commit 748446bb ("mm: compaction: memory compaction core"), has been checking the zone of the first valid page in a pageblock, and skipping the whole pageblock if the zone does not match. This checking was completely missing from the migration scanner at first, and later added by commit dc908600 ("mm: compaction: check for overlapping nodes during isolation for migration") in a reaction to a bug report. But the zone comparison in migration scanner is done once per a single scanned page, which is more defensive and thus more costly than a check per pageblock. This patch unifies the checking done in both scanners to once per pageblock, through a new pageblock_pfn_to_page() function, which also includes pfn_valid() checks. It is more defensive than the current free scanner checks, as it checks both the first and last page of the pageblock, but less defensive by the migration scanner per-page checks. It assumes that node overlapping may result (on some architecture) in a boundary between two nodes falling into the middle of a pageblock, but that there cannot be a node0 node1 node0 interleaving within a single pageblock. The result is more code being shared and a bit less per-page CPU cost in the migration scanner. Signed-off-by:
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Vlastimil Babka authored
isolate_migratepages_range() is the main function of the compaction scanner, called either on a single pageblock by isolate_migratepages() during regular compaction, or on an arbitrary range by CMA's __alloc_contig_migrate_range(). It currently perfoms two pageblock-wide compaction suitability checks, and because of the CMA callpath, it tracks if it crossed a pageblock boundary in order to repeat those checks. However, closer inspection shows that those checks are always true for CMA: - isolation_suitable() is true because CMA sets cc->ignore_skip_hint to true - migrate_async_suitable() check is skipped because CMA uses sync compaction We can therefore move the compaction-specific checks to isolate_migratepages() and simplify isolate_migratepages_range(). Furthermore, we can mimic the freepage scanner family of functions, which has isolate_freepages_block() function called both by compaction from isolate_freepages() and by CMA from isolate_freepages_range(), where each use-case adds own specific glue code. This allows further code simplification. Thus, we rename isolate_migratepages_range() to isolate_migratepages_block() and limit its functionality to a single pageblock (or its subset). For CMA, a new different isolate_migratepages_range() is created as a CMA-specific wrapper for the _block() function. The checks specific to compaction are moved to isolate_migratepages(). As part of the unification of these two families of functions, we remove the redundant zone parameter where applicable, since zone pointer is already passed in cc->zone. Furthermore, going back to compact_zone() and compact_finished() when pageblock is found unsuitable (now by isolate_migratepages()) is wasteful - the checks are meant to skip pageblocks quickly. The patch therefore also introduces a simple loop into isolate_migratepages() so that it does not return immediately on failed pageblock checks, but keeps going until isolate_migratepages_range() gets called once. Similarily to isolate_freepages(), the function periodically checks if it needs to reschedule or abort async compaction. [iamjoonsoo.kim@lge.com: fix isolated page counting bug in compaction] Signed-off-by:
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Vlastimil Babka authored
isolate_freepages_block() rechecks if the pageblock is suitable to be a target for migration after it has taken the zone->lock. However, the check has been optimized to occur only once per pageblock, and compact_checklock_irqsave() might be dropping and reacquiring lock, which means somebody else might have changed the pageblock's migratetype meanwhile. Furthermore, nothing prevents the migratetype to change right after isolate_freepages_block() has finished isolating. Given how imperfect this is, it's simpler to just rely on the check done in isolate_freepages() without lock, and not pretend that the recheck under lock guarantees anything. It is just a heuristic after all. Signed-off-by:
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Vlastimil Babka authored
The compact_stall vmstat counter counts the number of allocations stalled by direct compaction. It does not count when all attempted zones had deferred compaction, but it does count when all zones skipped compaction. The skipping is decided based on very early check of compaction_suitable(), based on watermarks and memory fragmentation. Therefore it makes sense not to count skipped compactions as stalls. Moreover, compact_success or compact_fail is also already not being counted when compaction was skipped, so this patch changes the compact_stall counting to match the other two. Additionally, restructure __alloc_pages_direct_compact() code for better readability. Signed-off-by:
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Vlastimil Babka authored
When direct sync compaction is often unsuccessful, it may become deferred for some time to avoid further useless attempts, both sync and async. Successful high-order allocations un-defer compaction, while further unsuccessful compaction attempts prolong the compaction deferred period. Currently the checking and setting deferred status is performed only on the preferred zone of the allocation that invoked direct compaction. But compaction itself is attempted on all eligible zones in the zonelist, so the behavior is suboptimal and may lead both to scenarios where 1) compaction is attempted uselessly, or 2) where it's not attempted despite good chances of succeeding, as shown on the examples below: 1) A direct compaction with Normal preferred zone failed and set deferred compaction for the Normal zone. Another unrelated direct compaction with DMA32 as preferred zone will attempt to compact DMA32 zone even though the first compaction attempt also included DMA32 zone. In another scenario, compaction with Normal preferred zone failed to compact Normal zone, but succeeded in the DMA32 zone, so it will not defer compaction. In the next attempt, it will try Normal zone which will fail again, instead of skipping Normal zone and trying DMA32 directly. 2) Kswapd will balance DMA32 zone and reset defer status based on watermarks looking good. A direct compaction with preferred Normal zone will skip compaction of all zones including DMA32 because Normal was still deferred. The allocation might have succeeded in DMA32, but won't. This patch makes compaction deferring work on individual zone basis instead of preferred zone. For each zone, it checks compaction_deferred() to decide if the zone should be skipped. If watermarks fail after compacting the zone, defer_compaction() is called. The zone where watermarks passed can still be deferred when the allocation attempt is unsuccessful. When allocation is successful, compaction_defer_reset() is called for the zone containing the allocated page. This approach should approximate calling defer_compaction() only on zones where compaction was attempted and did not yield allocated page. There might be corner cases but that is inevitable as long as the decision to stop compacting dues not guarantee that a page will be allocated. Due to a new COMPACT_DEFERRED return value, some functions relying implicitly on COMPACT_SKIPPED = 0 had to be updated, with comments made more accurate. The did_some_progress output parameter of __alloc_pages_direct_compact() is removed completely, as the caller actually does not use it after compaction sets it - it is only considered when direct reclaim sets it. During testing on a two-node machine with a single very small Normal zone on node 1, this patch has improved success rates in stress-highalloc mmtests benchmark. The success here were previously made worse by commit 3a025760 ("mm: page_alloc: spill to remote nodes before waking kswapd") as kswapd was no longer resetting often enough the deferred compaction for the Normal zone, and DMA32 zones on both nodes were thus not considered for compaction. On different machine, success rates were improved with __GFP_NO_KSWAPD allocations. [akpm@linux-foundation.org: fix CONFIG_COMPACTION=n build] Signed-off-by:
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Vlastimil Babka authored
When allocating huge page for collapsing, khugepaged currently holds mmap_sem for reading on the mm where collapsing occurs. Afterwards the read lock is dropped before write lock is taken on the same mmap_sem. Holding mmap_sem during whole huge page allocation is therefore useless, the vma needs to be rechecked after taking the write lock anyway. Furthemore, huge page allocation might involve a rather long sync compaction, and thus block any mmap_sem writers and i.e. affect workloads that perform frequent m(un)map or mprotect oterations. This patch simply releases the read lock before allocating a huge page. It also deletes an outdated comment that assumed vma must be stable, as it was using alloc_hugepage_vma(). This is no longer true since commit 9f1b868a ("mm: thp: khugepaged: add policy for finding target node"). Signed-off-by:
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Akinobu Mita authored
Sequential read from a block device is expected to be equal or faster than from the file on a filesystem. But it is not correct due to the lack of effective readpages() in the address space operations for block device. This implements readpages() operation for block device by using mpage_readpages() which can create multipage BIOs instead of BIOs for each page and reduce system CPU time consumption. Install 1GB of RAM disk storage: # modprobe scsi_debug dev_size_mb=1024 delay=0 Sequential read from file on a filesystem: # mkfs.ext4 /dev/$DEV # mount /dev/$DEV /mnt # fio --name=t --size=512m --rw=read --filename=/mnt/file ... read : io=524288KB, bw=2133.4MB/s, iops=546133, runt= 240msec Sequential read from a block device: # fio --name=t --size=512m --rw=read --filename=/dev/$DEV ... (Without this commit) read : io=524288KB, bw=1700.2MB/s, iops=435455, runt= 301msec (With this commit) read : io=524288KB, bw=2160.4MB/s, iops=553046, runt= 237msec Signed-off-by:
Akinobu Mita <akinobu.mita@gmail.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Jeff Moyer <jmoyer@redhat.com> Signed-off-by:
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Akinobu Mita authored
Add guard_bio_eod() check for mpage code in order to allow us to do IO even on the odd last sectors of a device, even if the block size is some multiple of the physical sector size. Using mpage_readpages() for block device requires this guard check. Signed-off-by:
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Akinobu Mita authored
This patchset implements readpages() operation for block device by using mpage_readpages() which can create multipage BIOs instead of BIOs for each page and reduce system CPU time consumption. This patch (of 3): guard_bh_eod() is used in submit_bh() to allow us to do IO even on the odd last sectors of a device, even if the block size is some multiple of the physical sector size. This makes guard_bh_eod() more generic and renames it guard_bio_eod() so that we can use it without struct buffer_head argument. The reason for this change is that using mpage_readpages() for block device requires to add this guard check in mpage code. Signed-off-by:
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Vlastimil Babka authored
The check for ALLOC_CMA in __alloc_pages_nodemask() derives migratetype from gfp_mask in each retry pass, although the migratetype variable already has the value determined and it does not change. Use the variable and perform the check only once. Also convert #ifdef CONFIG_CMA to IS_ENABLED. Signed-off-by:
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Marek Szyprowski authored
DMA-mapping supports CMA regions places either in low or high memory, so there is no longer needed to limit default CMA regions only to low memory. The real limit is still defined by architecture specific DMA limit. Signed-off-by:
Marek Szyprowski <m.szyprowski@samsung.com> Reported-by:
Russell King - ARM Linux <linux@arm.linux.org.uk> Acked-by:
Michal Nazarewicz <mina86@mina86.com> Cc: Daniel Drake <drake@endlessm.com> Cc: Minchan Kim <minchan@kernel.org> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Signed-off-by:
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Marek Szyprowski authored
Russell King recently noticed that limiting default CMA region only to low memory on ARM architecture causes serious memory management issues with machines having a lot of memory (which is mainly available as high memory). More information can be found the following thread: http://thread.gmane.org/gmane.linux.ports.arm.kernel/348441/ Those two patches removes this limit letting kernel to put default CMA region into high memory when this is possible (there is enough high memory available and architecture specific DMA limit fits). This should solve strange OOM issues on systems with lots of RAM (i.e. >1GiB) and large (>256M) CMA area. This patch (of 2): Automatically allocated regions should not cross low/high memory boundary, because such regions cannot be later correctly initialized due to spanning across two memory zones. This patch adds a check for this case and a simple code for moving region to low memory if automatically selected address might not fit completely into high memory. Signed-off-by:
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Laura Abbott authored
Neither CMA nor noncoherent allocations support atomic allocations. Add a dedicated atomic pool to support this. Reviewed-by:
Catalin Marinas <catalin.marinas@arm.com> Signed-off-by:
Laura Abbott <lauraa@codeaurora.org> Cc: Arnd Bergmann <arnd@arndb.de> Cc: David Riley <davidriley@chromium.org> Cc: Olof Johansson <olof@lixom.net> Cc: Ritesh Harjain <ritesh.harjani@gmail.com> Cc: Russell King <linux@arm.linux.org.uk> Cc: Thierry Reding <thierry.reding@gmail.com> Cc: Will Deacon <will.deacon@arm.com> Signed-off-by:
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Laura Abbott authored
ARM currently uses a bitmap for tracking atomic allocations. genalloc already handles this type of memory pool allocation so switch to using that instead. Signed-off-by:
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Laura Abbott authored
For architectures without coherent DMA, memory for DMA may need to be remapped with coherent attributes. Factor out the the remapping code from arm and put it in a common location to reduce code duplication. As part of this, the arm APIs are now migrated away from ioremap_page_range to the common APIs which use map_vm_area for remapping. This should be an equivalent change and using map_vm_area is more correct as ioremap_page_range is intended to bring in io addresses into the cpu space and not regular kernel managed memory. Signed-off-by:
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Laura Abbott authored
After allocating an address from a particular genpool, there is no good way to verify if that address actually belongs to a genpool. Introduce addr_in_gen_pool which will return if an address plus size falls completely within the genpool range. Signed-off-by:
Will Deacon <will.deacon@arm.com> Reviewed-by:
Olof Johansson <olof@lixom.net> Reviewed-by:
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Laura Abbott authored
One of the more common algorithms used for allocation is to align the start address of the allocation to the order of size requested. Add this as an algorithm option for genalloc. Signed-off-by:
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Mel Gorman authored
ARCH_USES_NUMA_PROT_NONE was defined for architectures that implemented _PAGE_NUMA using _PROT_NONE. This saved using an additional PTE bit and relied on the fact that PROT_NONE vmas were skipped by the NUMA hinting fault scanner. This was found to be conceptually confusing with a lot of implicit assumptions and it was asked that an alternative be found. Commit c46a7c81 "x86: define _PAGE_NUMA by reusing software bits on the PMD and PTE levels" redefined _PAGE_NUMA on x86 to be one of the swap PTE bits and shrunk the maximum possible swap size but it did not go far enough. There are no architectures that reuse _PROT_NONE as _PROT_NUMA but the relics still exist. This patch removes ARCH_USES_NUMA_PROT_NONE and removes some unnecessary duplication in powerpc vs the generic implementation by defining the types the core NUMA helpers expected to exist from x86 with their ppc64 equivalent. This necessitated that a PTE bit mask be created that identified the bits that distinguish present from NUMA pte entries but it is expected this will only differ between arches based on _PAGE_PROTNONE. The naming for the generic helpers was taken from x86 originally but ppc64 has types that are equivalent for the purposes of the helper so they are mapped instead of duplicating code. Signed-off-by:
Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Signed-off-by:
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Zhang Zhen authored
Currently memory-hotplug has two limits: 1. If the memory block is in ZONE_NORMAL, you can change it to ZONE_MOVABLE, but this memory block must be adjacent to ZONE_MOVABLE. 2. If the memory block is in ZONE_MOVABLE, you can change it to ZONE_NORMAL, but this memory block must be adjacent to ZONE_NORMAL. With this patch, we can easy to know a memory block can be onlined to which zone, and don't need to know the above two limits. Updated the related Documentation. [akpm@linux-foundation.org: use conventional comment layout] [akpm@linux-foundation.org: fix build with CONFIG_MEMORY_HOTREMOVE=n] [akpm@linux-foundation.org: remove unused local zone_prev] Signed-off-by:
Zhang Zhen <zhenzhang.zhang@huawei.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: David Rientjes <rientjes@google.com> Cc: Toshi Kani <toshi.kani@hp.com> Cc: Yasuaki Ishimatsu <isimatu.yasuaki@jp.fujitsu.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: Wang Nan <wangnan0@huawei.com> Signed-off-by:
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vishnu.ps authored
Signed-off-by:
vishnu.ps <vishnu.ps@samsung.com> Signed-off-by:
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Joonsoo Kim authored
Because of chicken and egg problem, initialization of SLAB is really complicated. We need to allocate cpu cache through SLAB to make the kmem_cache work, but before initialization of kmem_cache, allocation through SLAB is impossible. On the other hand, SLUB does initialization in a more simple way. It uses percpu allocator to allocate cpu cache so there is no chicken and egg problem. So, this patch try to use percpu allocator in SLAB. This simplifies the initialization step in SLAB so that we could maintain SLAB code more easily. In my testing there is no performance difference. This implementation relies on percpu allocator. Because percpu allocator uses vmalloc address space, vmalloc address space could be exhausted by this change on many cpu system with *32 bit* kernel. This implementation can cover 1024 cpus in worst case by following calculation. Worst: 1024 cpus * 4 bytes for pointer * 300 kmem_caches * 120 objects per cpu_cache = 140 MB Normal: 1024 cpus * 4 bytes for pointer * 150 kmem_caches(slab merge) * 80 objects per cpu_cache = 46 MB Signed-off-by:
Joonsoo Kim <iamjoonsoo.kim@lge.com> Acked-by:
Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Jeremiah Mahler <jmmahler@gmail.com> Signed-off-by:
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Joonsoo Kim authored
Slab merge is good feature to reduce fragmentation. If new creating slab have similar size and property with exsitent slab, this feature reuse it rather than creating new one. As a result, objects are packed into fewer slabs so that fragmentation is reduced. Below is result of my testing. * After boot, sleep 20; cat /proc/meminfo | grep Slab <Before> Slab: 25136 kB <After> Slab: 24364 kB We can save 3% memory used by slab. For supporting this feature in SLAB, we need to implement SLAB specific kmem_cache_flag() and __kmem_cache_alias(), because SLUB implements some SLUB specific processing related to debug flag and object size change on these functions. Signed-off-by:
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Joonsoo Kim authored
Slab merge is good feature to reduce fragmentation. Now, it is only applied to SLUB, but, it would be good to apply it to SLAB. This patch is preparation step to apply slab merge to SLAB by commonizing slab merge logic. Signed-off-by:
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Mikulas Patocka authored
Fix a bug (discovered with kmemcheck) in for_each_kmem_cache_node(). The for loop reads the array "node" before verifying that the index is within the range. This results in kmemcheck warning. Signed-off-by:
Mikulas Patocka <mpatocka@redhat.com> Reviewed-by:
Pekka Enberg <penberg@kernel.org> Acked-by:
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Nishanth Aravamudan authored
kernel/kthread.c: partial revert of 81c98869 ("kthread: ensure locality of task_struct allocations") After discussions with Tejun, we don't want to spread the use of cpu_to_mem() (and thus knowledge of allocators/NUMA topology details) into callers, but would rather ensure the callees correctly handle memoryless nodes. With the previous patches ("topology: add support for node_to_mem_node() to determine the fallback node" and "slub: fallback to node_to_mem_node() node if allocating on memoryless node") adding and using node_to_mem_node(), we can safely undo part of the change to the kthread logic from 81c98869. Signed-off-by:
Nishanth Aravamudan <nacc@linux.vnet.ibm.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: David Rientjes <rientjes@google.com> Cc: Han Pingtian <hanpt@linux.vnet.ibm.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Anton Blanchard <anton@samba.org> Cc: Christoph Lameter <cl@linux.com> Cc: Wanpeng Li <liwanp@linux.vnet.ibm.com> Cc: Tejun Heo <tj@kernel.org> Signed-off-by:
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Joonsoo Kim authored
Update the SLUB code to search for partial slabs on the nearest node with memory in the presence of memoryless nodes. Additionally, do not consider it to be an ALLOC_NODE_MISMATCH (and deactivate the slab) when a memoryless-node specified allocation goes off-node. Signed-off-by:
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Joonsoo Kim authored
Anton noticed (http://www.spinics.net/lists/linux-mm/msg67489.html) that on ppc LPARs with memoryless nodes, a large amount of memory was consumed by slabs and was marked unreclaimable. He tracked it down to slab deactivations in the SLUB core when we allocate remotely, leading to poor efficiency always when memoryless nodes are present. After much discussion, Joonsoo provided a few patches that help significantly. They don't resolve the problem altogether: - memory hotplug still needs testing, that is when a memoryless node becomes memory-ful, we want to dtrt - there are other reasons for going off-node than memoryless nodes, e.g., fully exhausted local nodes Neither case is resolved with this series, but I don't think that should block their acceptance, as they can be explored/resolved with follow-on patches. The series consists of: [1/3] topology: add support for node_to_mem_node() to determine the fallback node [2/3] slub: fallback to node_to_mem_node() node if allocating on memoryless node - Joonsoo's patches to cache the nearest node with memory for each NUMA node [3/3] Partial revert of 81c98869 (""kthread: ensure locality of task_struct allocations") - At Tejun's request, keep the knowledge of memoryless node fallback to the allocator core. This patch (of 3): We need to determine the fallback node in slub allocator if the allocation target node is memoryless node. Without it, the SLUB wrongly select the node which has no memory and can't use a partial slab, because of node mismatch. Introduced function, node_to_mem_node(X), will return a node Y with memory that has the nearest distance. If X is memoryless node, it will return nearest distance node, but, if X is normal node, it will return itself. We will use this function in following patch to determine the fallback node. Signed-off-by:
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