- 20 May, 2016 40 commits
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Joonsoo Kim authored
There is a system thats node's pfns are overlapped as follows: -----pfn--------> N0 N1 N2 N0 N1 N2 Therefore, we need to care this overlapping when iterating pfn range. There are one place in page_owner.c that iterates pfn range and it doesn't consider this overlapping. Add it. Without this patch, above system could over count early allocated page number before page_owner is activated. Signed-off-by:
Joonsoo Kim <iamjoonsoo.kim@lge.com> Acked-by:
Vlastimil Babka <vbabka@suse.cz> Cc: Rik van Riel <riel@redhat.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Mel Gorman <mgorman@techsingularity.net> Cc: Laura Abbott <lauraa@codeaurora.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Marek Szyprowski <m.szyprowski@samsung.com> Cc: Michal Nazarewicz <mina86@mina86.com> Cc: "Aneesh Kumar K.V" <aneesh.kumar@linux.vnet.ibm.com> Cc: "Rafael J. Wysocki" <rjw@rjwysocki.net> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Michael Ellerman <mpe@ellerman.id.au> Signed-off-by:
Andrew Morton <akpm@linux-foundation.org> Signed-off-by:
Linus Torvalds <torvalds@linux-foundation.org>
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Joonsoo Kim authored
There is a system thats node's pfns are overlapped as follows: -----pfn--------> N0 N1 N2 N0 N1 N2 Therefore, we need to care this overlapping when iterating pfn range. There are two places in vmstat.c that iterates pfn range and they don't consider this overlapping. Add it. Without this patch, above system could over count pageblock number on a zone. Signed-off-by:
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Joonsoo Kim authored
__offline_isolated_pages() and test_pages_isolated() are used by memory hotplug. These functions require that range is in a single zone but there is no code to do this because memory hotplug checks it before calling these functions. To avoid confusing future user of these functions, this patch adds comments to them. Signed-off-by:
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Joonsoo Kim authored
This patchset deals with some problematic sites that iterate pfn ranges. There is a system thats node's pfns are overlapped as follows: -----pfn--------> N0 N1 N2 N0 N1 N2 Therefore, we need to take care of this overlapping when iterating pfn range. I audit many iterating sites that uses pfn_valid(), pfn_valid_within(), zone_start_pfn and etc. and others looks safe to me. This is a preparation step for a new CMA implementation, ZONE_CMA (https://lkml.org/lkml/2015/2/12/95), because it would be easily overlapped with other zones. But, zone overlap check is also needed for the general case so I send it separately. This patch (of 5): alloc_gigantic_page() uses alloc_contig_range() and this requires that the requested range is in a single zone. To satisfy this requirement, add this check to pfn_range_valid_gigantic(). Signed-off-by:
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Andrew Morton authored
It's huge. Uninlining it saves 206 bytes per callsite. Shaves 4924 bytes from the x86_64 allmodconfig vmlinux. [akpm@linux-foundation.org: coding-style fixes] Cc: Steve Capper <steve.capper@arm.com> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Signed-off-by:
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Chanho Min authored
is_highmem() can be simplified by use of is_highmem_idx(). This patch removes redundant code and will make it easier to maintain if the zone policy is changed or a new zone is added. (akpm: saves me 25 bytes of text per is_highmem() callsite) Signed-off-by:
Chanho Min <chanho.min@lge.com> Reviewed-by:
Dan Williams <dan.j.williams@intel.com> Cc: Michal Hocko <mhocko@kernel.org> Signed-off-by:
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Vlastimil Babka authored
The goal of direct compaction is to quickly make a high-order page available for the pending allocation. Within an aligned block of pages of desired order, a single allocated page that cannot be isolated for migration means that the block cannot fully merge to a buddy page that would satisfy the allocation request. Therefore we can reduce the allocation stall by skipping the rest of the block immediately on isolation failure. For async compaction, this also means a higher chance of succeeding until it detects contention. We however shouldn't completely sacrifice the second objective of compaction, which is to reduce overal long-term memory fragmentation. As a compromise, perform the eager skipping only in direct async compaction, while sync compaction (including kcompactd) remains thorough. Testing was done using stress-highalloc from mmtests, configured for order-4 GFP_KERNEL allocations: 4.6-rc1 4.6-rc1 before after Success 1 Min 24.00 ( 0.00%) 27.00 (-12.50%) Success 1 Mean 30.20 ( 0.00%) 31.60 ( -4.64%) Success 1 Max 37.00 ( 0.00%) 35.00 ( 5.41%) Success 2 Min 42.00 ( 0.00%) 32.00 ( 23.81%) Success 2 Mean 44.00 ( 0.00%) 44.80 ( -1.82%) Success 2 Max 48.00 ( 0.00%) 52.00 ( -8.33%) Success 3 Min 91.00 ( 0.00%) 92.00 ( -1.10%) Success 3 Mean 92.20 ( 0.00%) 92.80 ( -0.65%) Success 3 Max 94.00 ( 0.00%) 93.00 ( 1.06%) We can see that success rates are unaffected by the skipping. 4.6-rc1 4.6-rc1 before after User 2587.42 2566.53 System 482.89 471.20 Elapsed 1395.68 1382.00 Times are not so useful metric for this benchmark as main portion is the interfering kernel builds, but results do hint at reduced system times. 4.6-rc1 4.6-rc1 before after Direct pages scanned 163614 159608 Kswapd pages scanned 2070139 2078790 Kswapd pages reclaimed 2061707 2069757 Direct pages reclaimed 163354 159505 Reduced direct reclaim was unintended, but could be explained by more successful first attempt at (async) direct compaction, which is attempted before the first reclaim attempt in __alloc_pages_slowpath(). Compaction stalls 33052 39853 Compaction success 12121 19773 Compaction failures 20931 20079 Compaction is indeed more successful, and thus less likely to get deferred, so there are also more direct compaction stalls. Page migrate success 3781876 3326819 Page migrate failure 45817 41774 Compaction pages isolated 7868232 6941457 Compaction migrate scanned 168160492 127269354 Compaction migrate prescanned 0 0 Compaction free scanned 2522142582 2326342620 Compaction free direct alloc 0 0 Compaction free dir. all. miss 0 0 Compaction cost 5252 4476 The patch reduces migration scanned pages by 25% thanks to the eager skipping. [hughd@google.com: prevent nr_isolated_* from going negative] Signed-off-by:Hugh Dickins <hughd@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Mel Gorman <mgorman@techsingularity.net> Cc: Rik van Riel <riel@redhat.com> Cc: David Rientjes <rientjes@google.com> Cc: Minchan Kim <minchan@kernel.org> Cc: Michal Hocko <mhocko@suse.com> Signed-off-by:
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Vlastimil Babka authored
Compaction drains the local pcplists each time migration scanner moves away from a cc->order aligned block where it isolated pages for migration, so that the pages freed by migrations can merge into higher orders. The detection is currently coarser than it could be. The cc->last_migrated_pfn variable should track the lowest pfn that was isolated for migration. But it is set to the pfn where isolate_migratepages_block() starts scanning, which is typically the first pfn of the pageblock. There, the scanner might fail to isolate several order-aligned blocks, and then isolate COMPACT_CLUSTER_MAX in another block. This would cause the pcplists drain to be performed, although the scanner didn't yet finish the block where it isolated from. This patch thus makes cc->last_migrated_pfn handling more accurate by setting it to the pfn of an actually isolated page in isolate_migratepages_block(). Although practical effects of this patch are likely low, it arguably makes the intent of the code more obvious. Also the next patch will make async direct compaction skip blocks more aggressively, and draining pcplists due to skipped blocks is wasteful. Signed-off-by:
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Vlastimil Babka authored
Compaction code has accumulated numerous instances of manual calculations of the first (inclusive) and last (exclusive) pfn of a pageblock (or a smaller block of given order), given a pfn within the pageblock. Wrap these calculations by introducing pageblock_start_pfn(pfn) and pageblock_end_pfn(pfn) macros. [vbabka@suse.cz: fix crash in get_pfnblock_flags_mask() from isolate_freepages():] Signed-off-by:
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Konstantin Khlebnikov authored
This check effectively catches anon vma hierarchy inconsistence and some vma corruptions. It was effective for catching corner cases in anon vma reusing logic. For now this code seems stable so check could be hidden under CONFIG_DEBUG_VM and replaced with WARN because it's not so fatal. Signed-off-by:
Konstantin Khlebnikov <khlebnikov@yandex-team.ru> Suggested-by:
Vasily Averin <vvs@virtuozzo.com> Acked-by:
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Andrew Morton authored
This code was pretty obscure and was relying upon obscure side-effects of next_node(-1, ...) and was relying upon NUMA_NO_NODE being equal to -1. Clean that all up and document the function's intent. Acked-by:
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Andrew Morton authored
Instead of open-coding it. Signed-off-by:
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Li Zhang authored
__free_pages_boot_core has parameter pfn which is not used at all. Remove it. Signed-off-by:
Li Zhang <zhlcindy@linux.vnet.ibm.com> Reviewed-by:
Pan Xinhui <xinhui.pan@linux.vnet.ibm.com> Signed-off-by:
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Michal Hocko authored
> The comment seems to have not much to do with the code? I guess the comment tries to say that the code path is triggered when we charge the page which happens _before_ it is added to the LRU list and so last_scanned_node might contain the stale data. Cc: Johannes Weiner <hannes@cmpxchg.org> Signed-off-by:
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Yaowei Bai authored
Make vma_migratable() return bool due to this particular function only using either one or zero as its return value. Signed-off-by:
Yaowei Bai <baiyaowei@cmss.chinamobile.com> Signed-off-by:
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Yaowei Bai authored
Make is_vmalloc_addr() return bool to improve readability due to this particular function only using either one or zero as its return value. Signed-off-by:
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Yaowei Bai authored
Make is_mem_section_removable() return bool to improve readability due to this particular function only using either one or zero as its return value. Signed-off-by:
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Yaowei Bai authored
Make is_vm_hugetlb_page() return bool to improve readability due to this particular function only using either one or zero as its return value. Signed-off-by:
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Vaishali Thakkar authored
Update setup_hugepagesz() to call hugetlb_bad_size() when unsupported hugepage size is found. Signed-off-by:
Vaishali Thakkar <vaishali.thakkar@oracle.com> Reviewed-by:
Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Reviewed-by:
Mike Kravetz <mike.kravetz@oracle.com> Acked-by:
Michal Hocko <mhocko@suse.com> Cc: Hillf Danton <hillf.zj@alibaba-inc.com> Cc: Yaowei Bai <baiyaowei@cmss.chinamobile.com> Cc: Dominik Dingel <dingel@linux.vnet.ibm.com> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Paul Gortmaker <paul.gortmaker@windriver.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Signed-off-by:
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Vaishali Thakkar authored
Update setup_hugepagesz() to call hugetlb_bad_size() when unsupported hugepage size is found. Signed-off-by:
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Vaishali Thakkar authored
Update setup_hugepagesz() to call hugetlb_bad_size() when unsupported hugepage size is found. Signed-off-by:
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Vaishali Thakkar authored
Update setup_hugepagesz() to call hugetlb_bad_size() when unsupported hugepage size is found. Signed-off-by:
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Vaishali Thakkar authored
Update setup_hugepagesz() to call hugetlb_bad_size() when unsupported hugepage size is found. Signed-off-by:
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Vaishali Thakkar authored
When any unsupported hugepage size is specified, 'hugepagesz=' and 'hugepages=' should be ignored during command line parsing until any supported hugepage size is found. But currently incorrect number of hugepages are allocated when unsupported size is specified as it fails to ignore the 'hugepages=' command. Test case: Note that this is specific to x86 architecture. Boot the kernel with command line option 'hugepagesz=256M hugepages=X'. After boot, dmesg output shows that X number of hugepages of the size 2M is pre-allocated instead of 0. So, to handle such command line options, introduce new routine hugetlb_bad_size. The routine hugetlb_bad_size sets the global variable parsed_valid_hugepagesz. We are using parsed_valid_hugepagesz to save the state when unsupported hugepagesize is found so that we can ignore the 'hugepages=' parameters after that and then reset the variable when supported hugepage size is found. The routine hugetlb_bad_size can be called while setting 'hugepagesz=' parameter in an architecture specific code. Signed-off-by:
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Mike Kravetz authored
It was observed that minimum size accounting associated with the hugetlbfs min_size mount option may not perform optimally and as expected. As huge pages/reservations are released from the filesystem and given back to the global pools, they are reserved for subsequent filesystem use as long as the subpool reserved count is less than subpool minimum size. It does not take into account used pages within the filesystem. The filesystem size limits are not exceeded and this is technically not a bug. However, better behavior would be to wait for the number of used pages/reservations associated with the filesystem to drop below the minimum size before taking reservations to satisfy minimum size. An optimization is also made to the hugepage_subpool_get_pages() routine which is called when pages/reservations are allocated. This does not change behavior, but simply avoids the accounting if all reservations have already been taken (subpool reserved count == 0). Signed-off-by:
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Andrew Morton authored
Lots of code does node = next_node(node, XXX); if (node == MAX_NUMNODES) node = first_node(XXX); so create next_node_in() to do this and use it in various places. [mhocko@suse.com: use next_node_in() helper] Acked-by:
Michal Hocko <mhocko@kernel.org> Signed-off-by:
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Rasmus Villemoes authored
Attach the malloc attribute to a few allocation functions. This helps gcc generate better code by telling it that the return value doesn't alias any existing pointers (which is even more valuable given the pessimizations implied by -fno-strict-aliasing). A simple example of what this allows gcc to do can be seen by looking at the last part of drm_atomic_helper_plane_reset: plane->state = kzalloc(sizeof(*plane->state), GFP_KERNEL); if (plane->state) { plane->state->plane = plane; plane->state->rotation = BIT(DRM_ROTATE_0); } which compiles to e8 99 bf d6 ff callq ffffffff8116d540 <kmem_cache_alloc_trace> 48 85 c0 test %rax,%rax 48 89 83 40 02 00 00 mov %rax,0x240(%rbx) 74 11 je ffffffff814015c4 <drm_atomic_helper_plane_reset+0x64> 48 89 18 mov %rbx,(%rax) 48 8b 83 40 02 00 00 mov 0x240(%rbx),%rax [*] c7 40 40 01 00 00 00 movl $0x1,0x40(%rax) With this patch applied, the instruction at [*] is elided, since the store to plane->state->plane is known to not alter the value of plane->state. [akpm@linux-foundation.org: coding-style fixes] Signed-off-by:Rasmus Villemoes <linux@rasmusvillemoes.dk> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Andi Kleen <ak@linux.intel.com> Signed-off-by:
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Rasmus Villemoes authored
gcc as far back as at least 3.04 documents the function attribute __malloc__. Add a shorthand for attaching that to a function declaration. This was also suggested by Andi Kleen way back in 2002 [1], but didn't get applied, perhaps because gcc at that time generated the exact same code with and without this attribute. This attribute tells the compiler that the return value (if non-NULL) can be assumed not to alias any other valid pointers at the time of the call. Please note that the documentation for a range of gcc versions (starting from around 4.7) contained a somewhat confusing and self-contradicting text: The malloc attribute is used to tell the compiler that a function may be treated as if any non-NULL pointer it returns cannot alias any other pointer valid when the function returns and *that the memory has undefined content*. [...] Standard functions with this property include malloc and *calloc*. (emphasis mine). The intended meaning has later been clarified [2]: This tells the compiler that a function is malloc-like, i.e., that the pointer P returned by the function cannot alias any other pointer valid when the function returns, and moreover no pointers to valid objects occur in any storage addressed by P. What this means is that we can apply the attribute to kmalloc and friends, and it is ok for the returned memory to have well-defined contents (__GFP_ZERO). But it is not ok to apply it to kmemdup(), nor to other functions which both allocate and possibly initialize the memory with existing pointers. So unless someone is doing something pretty perverted kstrdup() should also be a fine candidate. [1] http://thread.gmane.org/gmane.linux.kernel/57172 [2] https://gcc.gnu.org/bugzilla/show_bug.cgi?id=56955Signed-off-by:
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Joonsoo Kim authored
Many developers already know that field for reference count of the struct page is _count and atomic type. They would try to handle it directly and this could break the purpose of page reference count tracepoint. To prevent direct _count modification, this patch rename it to _refcount and add warning message on the code. After that, developer who need to handle reference count will find that field should not be accessed directly. [akpm@linux-foundation.org: fix comments, per Vlastimil] [akpm@linux-foundation.org: Documentation/vm/transhuge.txt too] [sfr@canb.auug.org.au: sync ethernet driver changes] Signed-off-by:
Stephen Rothwell <sfr@canb.auug.org.au> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Hugh Dickins <hughd@google.com> Cc: Johannes Berg <johannes@sipsolutions.net> Cc: "David S. Miller" <davem@davemloft.net> Cc: Sunil Goutham <sgoutham@cavium.com> Cc: Chris Metcalf <cmetcalf@mellanox.com> Cc: Manish Chopra <manish.chopra@qlogic.com> Cc: Yuval Mintz <yuval.mintz@qlogic.com> Cc: Tariq Toukan <tariqt@mellanox.com> Cc: Saeed Mahameed <saeedm@mellanox.com> Signed-off-by:
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Joonsoo Kim authored
page_reference manipulation functions are introduced to track down reference count change of the page. Use it instead of direct modification of _count. Signed-off-by:
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Li Peng authored
/sys/kernel/slab/xx/defrag_ratio should be remote_node_defrag_ratio. Link: http://lkml.kernel.org/r/1463449242-5366-1-git-send-email-lip@dtdream.comSigned-off-by:
Li Peng <lip@dtdream.com> Acked-by:
Christoph Lameter <cl@linux.com> Signed-off-by:
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Yang Shi authored
Now we have IS_ENABLED helper to check if a Kconfig option is enabled or not, so ZONE_DMA_FLAG sounds no longer useful. And, the use of ZONE_DMA_FLAG in slab looks pointless according to the comment [1] from Johannes Weiner, so remove them and ORing passed in flags with the cache gfp flags has been done in kmem_getpages(). [1] https://lkml.org/lkml/2014/9/25/553 Link: http://lkml.kernel.org/r/1462381297-11009-1-git-send-email-yang.shi@linaro.orgSigned-off-by:
Yang Shi <yang.shi@linaro.org> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Signed-off-by:
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Thomas Garnier authored
Provides an optional config (CONFIG_SLAB_FREELIST_RANDOM) to randomize the SLAB freelist. The list is randomized during initialization of a new set of pages. The order on different freelist sizes is pre-computed at boot for performance. Each kmem_cache has its own randomized freelist. Before pre-computed lists are available freelists are generated dynamically. This security feature reduces the predictability of the kernel SLAB allocator against heap overflows rendering attacks much less stable. For example this attack against SLUB (also applicable against SLAB) would be affected: https://jon.oberheide.org/blog/2010/09/10/linux-kernel-can-slub-overflow/ Also, since v4.6 the freelist was moved at the end of the SLAB. It means a controllable heap is opened to new attacks not yet publicly discussed. A kernel heap overflow can be transformed to multiple use-after-free. This feature makes this type of attack harder too. To generate entropy, we use get_random_bytes_arch because 0 bits of entropy is available in the boot stage. In the worse case this function will fallback to the get_random_bytes sub API. We also generate a shift random number to shift pre-computed freelist for each new set of pages. The config option name is not specific to the SLAB as this approach will be extended to other allocators like SLUB. Performance results highlighted no major changes: Hackbench (running 90 10 times): Before average: 0.0698 After average: 0.0663 (-5.01%) slab_test 1 run on boot. Difference only seen on the 2048 size test being the worse case scenario covered by freelist randomization. New slab pages are constantly being created on the 10000 allocations. Variance should be mainly due to getting new pages every few allocations. Before: Single thread testing ===================== 1. Kmalloc: Repeatedly allocate then free test 10000 times kmalloc(8) -> 99 cycles kfree -> 112 cycles 10000 times kmalloc(16) -> 109 cycles kfree -> 140 cycles 10000 times kmalloc(32) -> 129 cycles kfree -> 137 cycles 10000 times kmalloc(64) -> 141 cycles kfree -> 141 cycles 10000 times kmalloc(128) -> 152 cycles kfree -> 148 cycles 10000 times kmalloc(256) -> 195 cycles kfree -> 167 cycles 10000 times kmalloc(512) -> 257 cycles kfree -> 199 cycles 10000 times kmalloc(1024) -> 393 cycles kfree -> 251 cycles 10000 times kmalloc(2048) -> 649 cycles kfree -> 228 cycles 10000 times kmalloc(4096) -> 806 cycles kfree -> 370 cycles 10000 times kmalloc(8192) -> 814 cycles kfree -> 411 cycles 10000 times kmalloc(16384) -> 892 cycles kfree -> 455 cycles 2. Kmalloc: alloc/free test 10000 times kmalloc(8)/kfree -> 121 cycles 10000 times kmalloc(16)/kfree -> 121 cycles 10000 times kmalloc(32)/kfree -> 121 cycles 10000 times kmalloc(64)/kfree -> 121 cycles 10000 times kmalloc(128)/kfree -> 121 cycles 10000 times kmalloc(256)/kfree -> 119 cycles 10000 times kmalloc(512)/kfree -> 119 cycles 10000 times kmalloc(1024)/kfree -> 119 cycles 10000 times kmalloc(2048)/kfree -> 119 cycles 10000 times kmalloc(4096)/kfree -> 121 cycles 10000 times kmalloc(8192)/kfree -> 119 cycles 10000 times kmalloc(16384)/kfree -> 119 cycles After: Single thread testing ===================== 1. Kmalloc: Repeatedly allocate then free test 10000 times kmalloc(8) -> 130 cycles kfree -> 86 cycles 10000 times kmalloc(16) -> 118 cycles kfree -> 86 cycles 10000 times kmalloc(32) -> 121 cycles kfree -> 85 cycles 10000 times kmalloc(64) -> 176 cycles kfree -> 102 cycles 10000 times kmalloc(128) -> 178 cycles kfree -> 100 cycles 10000 times kmalloc(256) -> 205 cycles kfree -> 109 cycles 10000 times kmalloc(512) -> 262 cycles kfree -> 136 cycles 10000 times kmalloc(1024) -> 342 cycles kfree -> 157 cycles 10000 times kmalloc(2048) -> 701 cycles kfree -> 238 cycles 10000 times kmalloc(4096) -> 803 cycles kfree -> 364 cycles 10000 times kmalloc(8192) -> 835 cycles kfree -> 404 cycles 10000 times kmalloc(16384) -> 896 cycles kfree -> 441 cycles 2. Kmalloc: alloc/free test 10000 times kmalloc(8)/kfree -> 121 cycles 10000 times kmalloc(16)/kfree -> 121 cycles 10000 times kmalloc(32)/kfree -> 123 cycles 10000 times kmalloc(64)/kfree -> 142 cycles 10000 times kmalloc(128)/kfree -> 121 cycles 10000 times kmalloc(256)/kfree -> 119 cycles 10000 times kmalloc(512)/kfree -> 119 cycles 10000 times kmalloc(1024)/kfree -> 119 cycles 10000 times kmalloc(2048)/kfree -> 119 cycles 10000 times kmalloc(4096)/kfree -> 119 cycles 10000 times kmalloc(8192)/kfree -> 119 cycles 10000 times kmalloc(16384)/kfree -> 119 cycles [akpm@linux-foundation.org: propagate gfp_t into cache_random_seq_create()] Signed-off-by:
Thomas Garnier <thgarnie@google.com> Acked-by:
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Vladimir Davydov authored
When we call __kmem_cache_shrink on memory cgroup removal, we need to synchronize kmem_cache->cpu_partial update with put_cpu_partial that might be running on other cpus. Currently, we achieve that by using kick_all_cpus_sync, which works as a system wide memory barrier. Though fast it is, this method has a flaw - it issues a lot of IPIs, which might hurt high performance or real-time workloads. To fix this, let's replace kick_all_cpus_sync with synchronize_sched. Although the latter one may take much longer to finish, it shouldn't be a problem in this particular case, because memory cgroups are destroyed asynchronously from a workqueue so that no user visible effects should be introduced. OTOH, it will save us from excessive IPIs when someone removes a cgroup. Anyway, even if using synchronize_sched turns out to take too long, we can always introduce a kind of __kmem_cache_shrink batching so that this method would only be called once per one cgroup destruction (not per each per memcg kmem cache as it is now). Signed-off-by:
Vladimir Davydov <vdavydov@virtuozzo.com> Reported-by:
Peter Zijlstra <peterz@infradead.org> Suggested-by:
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Joonsoo Kim authored
To check whether free objects exist or not precisely, we need to grab a lock. But, accuracy isn't that important because race window would be even small and if there is too much free object, cache reaper would reap it. So, this patch makes the check for free object exisistence not to hold a lock. This will reduce lock contention in heavily allocation case. Note that until now, n->shared can be freed during the processing by writing slabinfo, but, with some trick in this patch, we can access it freely within interrupt disabled period. Below is the result of concurrent allocation/free in slab allocation benchmark made by Christoph a long time ago. I make the output simpler. The number shows cycle count during alloc/free respectively so less is better. * Before Kmalloc N*alloc N*free(32): Average=248/966 Kmalloc N*alloc N*free(64): Average=261/949 Kmalloc N*alloc N*free(128): Average=314/1016 Kmalloc N*alloc N*free(256): Average=741/1061 Kmalloc N*alloc N*free(512): Average=1246/1152 Kmalloc N*alloc N*free(1024): Average=2437/1259 Kmalloc N*alloc N*free(2048): Average=4980/1800 Kmalloc N*alloc N*free(4096): Average=9000/2078 * After Kmalloc N*alloc N*free(32): Average=344/792 Kmalloc N*alloc N*free(64): Average=347/882 Kmalloc N*alloc N*free(128): Average=390/959 Kmalloc N*alloc N*free(256): Average=393/1067 Kmalloc N*alloc N*free(512): Average=683/1229 Kmalloc N*alloc N*free(1024): Average=1295/1325 Kmalloc N*alloc N*free(2048): Average=2513/1664 Kmalloc N*alloc N*free(4096): Average=4742/2172 It shows that allocation performance decreases for the object size up to 128 and it may be due to extra checks in cache_alloc_refill(). But, with considering improvement of free performance, net result looks the same. Result for other size class looks very promising, roughly, 50% performance improvement. Signed-off-by:
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Joonsoo Kim authored
Until now, cache growing makes a free slab on node's slab list and then we can allocate free objects from it. This necessarily requires to hold a node lock which is very contended. If we refill cpu cache before attaching it to node's slab list, we can avoid holding a node lock as much as possible because this newly allocated slab is only visible to the current task. This will reduce lock contention. Below is the result of concurrent allocation/free in slab allocation benchmark made by Christoph a long time ago. I make the output simpler. The number shows cycle count during alloc/free respectively so less is better. * Before Kmalloc N*alloc N*free(32): Average=355/750 Kmalloc N*alloc N*free(64): Average=452/812 Kmalloc N*alloc N*free(128): Average=559/1070 Kmalloc N*alloc N*free(256): Average=1176/980 Kmalloc N*alloc N*free(512): Average=1939/1189 Kmalloc N*alloc N*free(1024): Average=3521/1278 Kmalloc N*alloc N*free(2048): Average=7152/1838 Kmalloc N*alloc N*free(4096): Average=13438/2013 * After Kmalloc N*alloc N*free(32): Average=248/966 Kmalloc N*alloc N*free(64): Average=261/949 Kmalloc N*alloc N*free(128): Average=314/1016 Kmalloc N*alloc N*free(256): Average=741/1061 Kmalloc N*alloc N*free(512): Average=1246/1152 Kmalloc N*alloc N*free(1024): Average=2437/1259 Kmalloc N*alloc N*free(2048): Average=4980/1800 Kmalloc N*alloc N*free(4096): Average=9000/2078 It shows that contention is reduced for all the object sizes and performance increases by 30 ~ 40%. Signed-off-by:
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Joonsoo Kim authored
This is a preparation step to implement lockless allocation path when there is no free objects in kmem_cache. What we'd like to do here is to refill cpu cache without holding a node lock. To accomplish this purpose, refill should be done after new slab allocation but before attaching the slab to the management list. So, this patch separates cache_grow() to two parts, allocation and attaching to the list in order to add some code inbetween them in the following patch. Signed-off-by:
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Joonsoo Kim authored
Currently, cache_grow() assumes that allocated page's nodeid would be same with parameter nodeid which is used for allocation request. If we discard this assumption, we can handle fallback_alloc() case gracefully. So, this patch makes cache_grow() handle the page allocated on arbitrary node and clean-up relevant code. Signed-off-by:
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Joonsoo Kim authored
Slab color isn't needed to be changed strictly. Because locking for changing slab color could cause more lock contention so this patch implements racy access/modify the slab color. This is a preparation step to implement lockless allocation path when there is no free objects in the kmem_cache. Below is the result of concurrent allocation/free in slab allocation benchmark made by Christoph a long time ago. I make the output simpler. The number shows cycle count during alloc/free respectively so less is better. * Before Kmalloc N*alloc N*free(32): Average=365/806 Kmalloc N*alloc N*free(64): Average=452/690 Kmalloc N*alloc N*free(128): Average=736/886 Kmalloc N*alloc N*free(256): Average=1167/985 Kmalloc N*alloc N*free(512): Average=2088/1125 Kmalloc N*alloc N*free(1024): Average=4115/1184 Kmalloc N*alloc N*free(2048): Average=8451/1748 Kmalloc N*alloc N*free(4096): Average=16024/2048 * After Kmalloc N*alloc N*free(32): Average=355/750 Kmalloc N*alloc N*free(64): Average=452/812 Kmalloc N*alloc N*free(128): Average=559/1070 Kmalloc N*alloc N*free(256): Average=1176/980 Kmalloc N*alloc N*free(512): Average=1939/1189 Kmalloc N*alloc N*free(1024): Average=3521/1278 Kmalloc N*alloc N*free(2048): Average=7152/1838 Kmalloc N*alloc N*free(4096): Average=13438/2013 It shows that contention is reduced for object size >= 1024 and performance increases by roughly 15%. Signed-off-by:
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Joonsoo Kim authored
Currently, determination to free a slab is done whenever each freed object is put into the slab. This has a following problem. Assume free_limit = 10 and nr_free = 9. Free happens as following sequence and nr_free changes as following. free(become a free slab) free(not become a free slab) nr_free: 9 -> 10 (at first free) -> 11 (at second free) If we try to check if we can free current slab or not on each object free, we can't free any slab in this situation because current slab isn't a free slab when nr_free exceed free_limit (at second free) even if there is a free slab. However, if we check it lastly, we can free 1 free slab. This problem would cause to keep too much memory in the slab subsystem. This patch try to fix it by checking number of free object after all free work is done. If there is free slab at that time, we can free slab as much as possible so we keep free slab as minimal. Signed-off-by:
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