- 16 Oct, 2007 40 commits
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Christoph Lameter authored
The kmem_cache_cpu structures introduced are currently an array placed in the kmem_cache struct. Meaning the kmem_cache_cpu structures are overwhelmingly on the wrong node for systems with a higher amount of nodes. These are performance critical structures since the per node information has to be touched for every alloc and free in a slab. In order to place the kmem_cache_cpu structure optimally we put an array of pointers to kmem_cache_cpu structs in kmem_cache (similar to SLAB). However, the kmem_cache_cpu structures can now be allocated in a more intelligent way. We would like to put per cpu structures for the same cpu but different slab caches in cachelines together to save space and decrease the cache footprint. However, the slab allocators itself control only allocations per node. We set up a simple per cpu array for every processor with 100 per cpu structures which is usually enough to get them all set up right. If we run out then we fall back to kmalloc_node. This also solves the bootstrap problem since we do not have to use slab allocator functions early in boot to get memory for the small per cpu structures. Pro: - NUMA aware placement improves memory performance - All global structures in struct kmem_cache become readonly - Dense packing of per cpu structures reduces cacheline footprint in SMP and NUMA. - Potential avoidance of exclusive cacheline fetches on the free and alloc hotpath since multiple kmem_cache_cpu structures are in one cacheline. This is particularly important for the kmalloc array. Cons: - Additional reference to one read only cacheline (per cpu array of pointers to kmem_cache_cpu) in both slab_alloc() and slab_free(). [akinobu.mita@gmail.com: fix cpu hotplug offline/online path] Signed-off-by:
Christoph Lameter <clameter@sgi.com> Cc: "Pekka Enberg" <penberg@cs.helsinki.fi> Cc: Akinobu Mita <akinobu.mita@gmail.com> Signed-off-by:
Andrew Morton <akpm@linux-foundation.org> Signed-off-by:
Linus Torvalds <torvalds@linux-foundation.org>
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Christoph Lameter authored
Set c->node to -1 if we allocate from a debug slab instead for SlabDebug which requires access the page struct cacheline. Signed-off-by:
Alexey Dobriyan <adobriyan@sw.ru> Signed-off-by:
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Christoph Lameter authored
We need the offset from the page struct during slab_alloc and slab_free. In both cases we also reference the cacheline of the kmem_cache_cpu structure. We can therefore move the offset field into the kmem_cache_cpu structure freeing up 16 bits in the page struct. Moving the offset allows an allocation from slab_alloc() without touching the page struct in the hot path. The only thing left in slab_free() that touches the page struct cacheline for per cpu freeing is the checking of SlabDebug(page). The next patch deals with that. Use the available 16 bits to broaden page->inuse. More than 64k objects per slab become possible and we can get rid of the checks for that limitation. No need anymore to shrink the order of slabs if we boot with 2M sized slabs (slub_min_order=9). No need anymore to switch off the offset calculation for very large slabs since the field in the kmem_cache_cpu structure is 32 bits and so the offset field can now handle slab sizes of up to 8GB. Signed-off-by:
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Christoph Lameter authored
After moving the lockless_freelist to kmem_cache_cpu we no longer need page->lockless_freelist. Restructure the use of the struct page fields in such a way that we never touch the mapping field. This is turn allows us to remove the special casing of SLUB when determining the mapping of a page (needed for corner cases of virtual caches machines that need to flush caches of processors mapping a page). Signed-off-by:
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Christoph Lameter authored
A remote free may access the same page struct that also contains the lockless freelist for the cpu slab. If objects have a short lifetime and are freed by a different processor then remote frees back to the slab from which we are currently allocating are frequent. The cacheline with the page struct needs to be repeately acquired in exclusive mode by both the allocating thread and the freeing thread. If this is frequent enough then performance will suffer because of cacheline bouncing. This patchset puts the lockless_freelist pointer in its own cacheline. In order to make that happen we introduce a per cpu structure called kmem_cache_cpu. Instead of keeping an array of pointers to page structs we now keep an array to a per cpu structure that--among other things--contains the pointer to the lockless freelist. The freeing thread can then keep possession of exclusive access to the page struct cacheline while the allocating thread keeps its exclusive access to the cacheline containing the per cpu structure. This works as long as the allocating cpu is able to service its request from the lockless freelist. If the lockless freelist runs empty then the allocating thread needs to acquire exclusive access to the cacheline with the page struct lock the slab. The allocating thread will then check if new objects were freed to the per cpu slab. If so it will keep the slab as the cpu slab and continue with the recently remote freed objects. So the allocating thread can take a series of just freed remote pages and dish them out again. Ideally allocations could be just recycling objects in the same slab this way which will lead to an ideal allocation / remote free pattern. The number of objects that can be handled in this way is limited by the capacity of one slab. Increasing slab size via slub_min_objects/ slub_max_order may increase the number of objects and therefore performance. If the allocating thread runs out of objects and finds that no objects were put back by the remote processor then it will retrieve a new slab (from the partial lists or from the page allocator) and start with a whole new set of objects while the remote thread may still be freeing objects to the old cpu slab. This may then repeat until the new slab is also exhausted. If remote freeing has freed objects in the earlier slab then that earlier slab will now be on the partial freelist and the allocating thread will pick that slab next for allocation. So the loop is extended. However, both threads need to take the list_lock to make the swizzling via the partial list happen. It is likely that this kind of scheme will keep the objects being passed around to a small set that can be kept in the cpu caches leading to increased performance. More code cleanups become possible: - Instead of passing a cpu we can now pass a kmem_cache_cpu structure around. Allows reducing the number of parameters to various functions. - Can define a new node_match() function for NUMA to encapsulate locality checks. Effect on allocations: Cachelines touched before this patch: Write: page cache struct and first cacheline of object Cachelines touched after this patch: Write: kmem_cache_cpu cacheline and first cacheline of object Read: page cache struct (but see later patch that avoids touching that cacheline) The handling when the lockless alloc list runs empty gets to be a bit more complicated since another cacheline has now to be written to. But that is halfway out of the hot path. Effect on freeing: Cachelines touched before this patch: Write: page_struct and first cacheline of object Cachelines touched after this patch depending on how we free: Write(to cpu_slab): kmem_cache_cpu struct and first cacheline of object Write(to other): page struct and first cacheline of object Read(to cpu_slab): page struct to id slab etc. (but see later patch that avoids touching the page struct on free) Read(to other): cpu local kmem_cache_cpu struct to verify its not the cpu slab. Summary: Pro: - Distinct cachelines so that concurrent remote frees and local allocs on a cpuslab can occur without cacheline bouncing. - Avoids potential bouncing cachelines because of neighboring per cpu pointer updates in kmem_cache's cpu_slab structure since it now grows to a cacheline (Therefore remove the comment that talks about that concern). Cons: - Freeing objects now requires the reading of one additional cacheline. That can be mitigated for some cases by the following patches but its not possible to completely eliminate these references. - Memory usage grows slightly. The size of each per cpu object is blown up from one word (pointing to the page_struct) to one cacheline with various data. So this is NR_CPUS*NR_SLABS*L1_BYTES more memory use. Lets say NR_SLABS is 100 and a cache line size of 128 then we have just increased SLAB metadata requirements by 12.8k per cpu. (Another later patch reduces these requirements) Signed-off-by:
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Adrian Bunk authored
This patch makes needlessly global code static. Signed-off-by:
Adrian Bunk <bunk@stusta.de> Signed-off-by:
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Mel Gorman authored
This patch provides fragmentation avoidance statistics via /proc/pagetypeinfo. The information is collected only on request so there is no runtime overhead. The statistics are in three parts: The first part prints information on the size of blocks that pages are being grouped on and looks like Page block order: 10 Pages per block: 1024 The second part is a more detailed version of /proc/buddyinfo and looks like Free pages count per migrate type at order 0 1 2 3 4 5 6 7 8 9 10 Node 0, zone DMA, type Unmovable 0 0 0 0 0 0 0 0 0 0 0 Node 0, zone DMA, type Reclaimable 1 0 0 0 0 0 0 0 0 0 0 Node 0, zone DMA, type Movable 0 0 0 0 0 0 0 0 0 0 0 Node 0, zone DMA, type Reserve 0 4 4 0 0 0 0 1 0 1 0 Node 0, zone Normal, type Unmovable 111 8 4 4 2 3 1 0 0 0 0 Node 0, zone Normal, type Reclaimable 293 89 8 0 0 0 0 0 0 0 0 Node 0, zone Normal, type Movable 1 6 13 9 7 6 3 0 0 0 0 Node 0, zone Normal, type Reserve 0 0 0 0 0 0 0 0 0 0 4 The third part looks like Number of blocks type Unmovable Reclaimable Movable Reserve Node 0, zone DMA 0 1 2 1 Node 0, zone Normal 3 17 94 4 To walk the zones within a node with interrupts disabled, walk_zones_in_node() is introduced and shared between /proc/buddyinfo, /proc/zoneinfo and /proc/pagetypeinfo to reduce code duplication. It seems specific to what vmstat.c requires but could be broken out as a general utility function in mmzone.c if there were other other potential users. Signed-off-by:
Mel Gorman <mel@csn.ul.ie> Acked-by:
Andy Whitcroft <apw@shadowen.org> Acked-by:
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Mel Gorman authored
Currently mobility grouping works at the MAX_ORDER_NR_PAGES level. This makes sense for the majority of users where this is also the huge page size. However, on platforms like ia64 where the huge page size is runtime configurable it is desirable to group at a lower order. On x86_64 and occasionally on x86, the hugepage size may not always be MAX_ORDER_NR_PAGES. This patch groups pages together based on the value of HUGETLB_PAGE_ORDER. It uses a compile-time constant if possible and a variable where the huge page size is runtime configurable. It is assumed that grouping should be done at the lowest sensible order and that the user would not want to override this. If this is not true, page_block order could be forced to a variable initialised via a boot-time kernel parameter. One potential issue with this patch is that IA64 now parses hugepagesz with early_param() instead of __setup(). __setup() is called after the memory allocator has been initialised and the pageblock bitmaps already setup. In tests on one IA64 there did not seem to be any problem with using early_param() and in fact may be more correct as it guarantees the parameter is handled before the parsing of hugepages=. Signed-off-by:
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Mel Gorman authored
move_freepages_block() returns the number of blocks moved. This value is used to determine if a block of pages should be stolen for the exclusive use of a migrate type or not. However, the value returned is being used correctly. This patch fixes the calculation to return the number of base pages that have been moved. This should be considered a fix to the patch move-free-pages-between-lists-on-steal.patch Credit to Andy Whitcroft for spotting the problem. Signed-off-by:
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Mel Gorman authored
Grouping high-order atomic allocations together was intended to allow bursty users of atomic allocations to work such as e1000 in situations where their preallocated buffers were depleted. This did not work in at least one case with a wireless network adapter needing order-1 allocations frequently. To resolve that, the free pages used for min_free_kbytes were moved to separate contiguous blocks with the patch bias-the-location-of-pages-freed-for-min_free_kbytes-in-the-same-max_order_nr_pages-blocks. It is felt that keeping the free pages in the same contiguous blocks should be sufficient for bursty short-lived high-order atomic allocations to succeed, maybe even with the e1000. Even if there is a failure, increasing the value of min_free_kbytes will free pages as contiguous bloks in contrast to the standard buddy allocator which makes no attempt to keep the minimum number of free pages contiguous. This patch backs out grouping high order atomic allocations together to determine if it is really needed or not. If a new report comes in about high-order atomic allocations failing, the feature can be reintroduced to determine if it fixes the problem or not. As a side-effect, this patch reduces by 1 the number of bits required to track the mobility type of pages within a MAX_ORDER_NR_PAGES block. Signed-off-by:
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Mel Gorman authored
Grouping pages by mobility can be disabled at compile-time. This was considered undesirable by a number of people. However, in the current stack of patches, it is not a simple case of just dropping the configurable patch as it would cause merge conflicts. This patch backs out the configuration option. Signed-off-by:
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Mel Gorman authored
The standard buddy allocator always favours the smallest block of pages. The effect of this is that the pages free to satisfy min_free_kbytes tends to be preserved since boot time at the same location of memory ffor a very long time and as a contiguous block. When an administrator sets the reserve at 16384 at boot time, it tends to be the same MAX_ORDER blocks that remain free. This allows the occasional high atomic allocation to succeed up until the point the blocks are split. In practice, it is difficult to split these blocks but when they do split, the benefit of having min_free_kbytes for contiguous blocks disappears. Additionally, increasing min_free_kbytes once the system has been running for some time has no guarantee of creating contiguous blocks. On the other hand, CONFIG_PAGE_GROUP_BY_MOBILITY favours splitting large blocks when there are no free pages of the appropriate type available. A side-effect of this is that all blocks in memory tends to be used up and the contiguous free blocks from boot time are not preserved like in the vanilla allocator. This can cause a problem if a new caller is unwilling to reclaim or does not reclaim for long enough. A failure scenario was found for a wireless network device allocating order-1 atomic allocations but the allocations were not intense or frequent enough for a whole block of pages to be preserved for MIGRATE_HIGHALLOC. This was reproduced on a desktop by booting with mem=256mb, forcing the driver to allocate at order-1, running a bittorrent client (downloading a debian ISO) and building a kernel with -j2. This patch addresses the problem on the desktop machine booted with mem=256mb. It works by setting aside a reserve of MAX_ORDER_NR_PAGES blocks, the number of which depends on the value of min_free_kbytes. These blocks are only fallen back to when there is no other free pages. Then the smallest possible page is used just like the normal buddy allocator instead of the largest possible page to preserve contiguous pages The pages in free lists in the reserve blocks are never taken for another migrate type. The results is that even if min_free_kbytes is set to a low value, contiguous blocks will be preserved in the MIGRATE_RESERVE blocks. This works better than the vanilla allocator because if min_free_kbytes is increased, a new reserve block will be chosen based on the location of reclaimable pages and the block will free up as contiguous pages. In the vanilla allocator, no effort is made to target a block of pages to free as contiguous pages and min_free_kbytes pages are scattered randomly. This effect has been observed on the test machine. min_free_kbytes was set initially low but it was kept as a contiguous free block within MIGRATE_RESERVE. min_free_kbytes was then set to a higher value and over a period of time, the free blocks were within the reserve and coalescing. How long it takes to free up depends on how quickly LRU is rotating. Amusingly, this means that more activity will free the blocks faster. This mechanism potentially replaces MIGRATE_HIGHALLOC as it may be more effective than grouping contiguous free pages together. It all depends on whether the number of active atomic high allocations exceeds min_free_kbytes or not. If the number of active allocations exceeds min_free_kbytes, it's worth it but maybe in that situation, min_free_kbytes should be set higher. Once there are no more reports of allocation failures, a patch will be submitted that backs out MIGRATE_HIGHALLOC and see if the reports stay missing. Credit to Mariusz Kozlowski for discovering the problem, describing the failure scenario and testing patches and scenarios. [akpm@linux-foundation.org: cleanups] Signed-off-by:
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Mel Gorman authored
There are problems in the use of SPARSEMEM and pageblock flags that causes problems on ia64. The first part of the problem is that units are incorrect in SECTION_BLOCKFLAGS_BITS computation. This results in a map_section's section_mem_map being treated as part of a bitmap which isn't good. This was evident with an invalid virtual address when mem_init attempted to free bootmem pages while relinquishing control from the bootmem allocator. The second part of the problem occurs because the pageblock flags bitmap is be located with the mem_section. The SECTIONS_PER_ROOT computation using sizeof (mem_section) may not be a power of 2 depending on the size of the bitmap. This renders masks and other such things not power of 2 base. This issue was seen with SPARSEMEM_EXTREME on ia64. This patch moves the bitmap outside of mem_section and uses a pointer instead in the mem_section. The bitmaps are allocated when the section is being initialised. Note that sparse_early_usemap_alloc() does not use alloc_remap() like sparse_early_mem_map_alloc(). The allocation required for the bitmap on x86, the only architecture that uses alloc_remap is typically smaller than a cache line. alloc_remap() pads out allocations to the cache size which would be a needless waste. Credit to Bob Picco for identifying the original problem and effecting a fix for the SECTION_BLOCKFLAGS_BITS calculation. Credit to Andy Whitcroft for devising the best way of allocating the bitmaps only when required for the section. [wli@holomorphy.com: warning fix] Signed-off-by:
Bob Picco <bob.picco@hp.com> Signed-off-by:
William Irwin <bill.irwin@oracle.com> Signed-off-by:
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Mel Gorman authored
MIGRATE_RECLAIMABLE allocations tend to be very bursty in nature like when updatedb starts. It is likely this will occur in situations where MAX_ORDER blocks of pages are not free. This means that updatedb can scatter MIGRATE_RECLAIMABLE pages throughout the address space. This patch is more agressive about stealing blocks of pages for MIGRATE_RECLAIMABLE. Signed-off-by:
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Mel Gorman authored
This patch chooses blocks with lower PFNs when placing kernel allocations. This is particularly important during fallback in low memory situations to stop unmovable pages being placed throughout the entire address space. Signed-off-by:
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Mel Gorman authored
Grouping pages by mobility can only successfully operate when there are more MAX_ORDER_NR_PAGES areas than mobility types. When there are insufficient areas, fallbacks cannot be avoided. This has noticeable performance impacts on machines with small amounts of memory in comparison to MAX_ORDER_NR_PAGES. For example, on IA64 with a configuration including huge pages spans 1GiB with MAX_ORDER_NR_PAGES so would need at least 4GiB of RAM before grouping pages by mobility would be useful. In comparison, an x86 would need 16MB. This patch checks the size of vm_total_pages in build_all_zonelists(). If there are not enough areas, mobility is effectivly disabled by considering all allocations as the same type (UNMOVABLE). This is achived via a __read_mostly flag. With this patch, performance is comparable to disabling grouping pages by mobility at compile-time on a test machine with insufficient memory. With this patch, it is reasonable to get rid of grouping pages by mobility a compile-time option. Signed-off-by:
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Mel Gorman authored
In rare cases, the kernel needs to allocate a high-order block of pages without sleeping. For example, this is the case with e1000 cards configured to use jumbo frames. Migrating or reclaiming pages in this situation is not an option. This patch groups these allocations together as much as possible by adding a new MIGRATE_TYPE. The MIGRATE_HIGHATOMIC type are exactly what they sound like. Care is taken that pages of other migrate types do not use the same blocks as high-order atomic allocations. Signed-off-by:
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Mel Gorman authored
This patch marks a number of allocations that are either short-lived such as network buffers or are reclaimable such as inode allocations. When something like updatedb is called, long-lived and unmovable kernel allocations tend to be spread throughout the address space which increases fragmentation. This patch groups these allocations together as much as possible by adding a new MIGRATE_TYPE. The MIGRATE_RECLAIMABLE type is for allocations that can be reclaimed on demand, but not moved. i.e. they can be migrated by deleting them and re-reading the information from elsewhere. Signed-off-by:
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Mel Gorman authored
When a fallback occurs, there will be free pages for one allocation type stored on the list for another. When a large steal occurs, this patch will move all the free pages within one list to the other. [y-goto@jp.fujitsu.com: fix BUG_ON check at move_freepages()] [apw@shadowen.org: Move to using pfn_valid_within()] Signed-off-by:
Yasunori Goto <y-goto@jp.fujitsu.com> Cc: Bjorn Helgaas <bjorn.helgaas@hp.com> Signed-off-by:
Andy Whitcroft <andyw@uk.ibm.com> Cc: Bob Picco <bob.picco@hp.com> Signed-off-by:
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Mel Gorman authored
Per-cpu pages can accidentally cause fragmentation because they are free, but pinned pages in an otherwise contiguous block. When this patch is applied, the per-cpu caches are drained after the direct-reclaim is entered if the requested order is greater than 0. It simply reuses the code used by suspend and hotplug. Signed-off-by:
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Mel Gorman authored
The grouping mechanism has some memory overhead and a more complex allocation path. This patch allows the strategy to be disabled for small memory systems or if it is known the workload is suffering because of the strategy. It also acts to show where the page groupings strategy interacts with the standard buddy allocator. Signed-off-by:
Joel Schopp <jschopp@austin.ibm.com> Cc: Andy Whitcroft <apw@shadowen.org> Signed-off-by:
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Mel Gorman authored
The freelists for each migrate type can slowly become polluted due to the per-cpu list. Consider what happens when the following happens 1. A 2^(MAX_ORDER-1) list is reserved for __GFP_MOVABLE pages 2. An order-0 page is allocated from the newly reserved block 3. The page is freed and placed on the per-cpu list 4. alloc_page() is called with GFP_KERNEL as the gfp_mask 5. The per-cpu list is used to satisfy the allocation This results in a kernel page is in the middle of a migratable region. This patch prevents this leak occuring by storing the MIGRATE_ type of the page in page->private. On allocate, a page will only be returned of the desired type, else more pages will be allocated. This may temporarily allow a per-cpu list to go over the pcp->high limit but it'll be corrected on the next free. Care is taken to preserve the hotness of pages recently freed. The additional code is not measurably slower for the workloads we've tested. Signed-off-by:
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Mel Gorman authored
This patch adds the core of the fragmentation reduction strategy. It works by grouping pages together based on their ability to migrate or be reclaimed. Basically, it works by breaking the list in zone->free_area list into MIGRATE_TYPES number of lists. Signed-off-by:
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Mel Gorman authored
Here is the latest revision of the anti-fragmentation patches. Of particular note in this version is special treatment of high-order atomic allocations. Care is taken to group them together and avoid grouping pages of other types near them. Artifical tests imply that it works. I'm trying to get the hardware together that would allow setting up of a "real" test. If anyone already has a setup and test that can trigger the atomic-allocation problem, I'd appreciate a test of these patches and a report. The second major change is that these patches will apply cleanly with patches that implement anti-fragmentation through zones. kernbench shows effectively no performance difference varying between -0.2% and +2% on a variety of test machines. Success rates for huge page allocation are dramatically increased. For example, on a ppc64 machine, the vanilla kernel was only able to allocate 1% of memory as a hugepage and this was due to a single hugepage reserved as min_free_kbytes. With these patches applied, 17% was allocatable as superpages. With reclaim-related fixes from Andy Whitcroft, it was 40% and further reclaim-related improvements should increase this further. Changelog Since V28 o Group high-order atomic allocations together o It is no longer required to set min_free_kbytes to 10% of memory. A value of 16384 in most cases will be sufficient o Now applied with zone-based anti-fragmentation o Fix incorrect VM_BUG_ON within buffered_rmqueue() o Reorder the stack so later patches do not back out work from earlier patches o Fix bug were journal pages were being treated as movable o Bias placement of non-movable pages to lower PFNs o More agressive clustering of reclaimable pages in reactions to workloads like updatedb that flood the size of inode caches Changelog Since V27 o Renamed anti-fragmentation to Page Clustering. Anti-fragmentation was giving the mistaken impression that it was the 100% solution for high order allocations. Instead, it greatly increases the chances high-order allocations will succeed and lays the foundation for defragmentation and memory hot-remove to work properly o Redefine page groupings based on ability to migrate or reclaim instead of basing on reclaimability alone o Get rid of spurious inits o Per-cpu lists are no longer split up per-type. Instead the per-cpu list is searched for a page of the appropriate type o Added more explanation commentary o Fix up bug in pageblock code where bitmap was used before being initalised Changelog Since V26 o Fix double init of lists in setup_pageset Changelog Since V25 o Fix loop order of for_each_rclmtype_order so that order of loop matches args o gfpflags_to_rclmtype uses gfp_t instead of unsigned long o Rename get_pageblock_type() to get_page_rclmtype() o Fix alignment problem in move_freepages() o Add mechanism for assigning flags to blocks of pages instead of page->flags o On fallback, do not examine the preferred list of free pages a second time The purpose of these patches is to reduce external fragmentation by grouping pages of related types together. When pages are migrated (or reclaimed under memory pressure), large contiguous pages will be freed. This patch works by categorising allocations by their ability to migrate; Movable - The pages may be moved with the page migration mechanism. These are generally userspace pages. Reclaimable - These are allocations for some kernel caches that are reclaimable or allocations that are known to be very short-lived. Unmovable - These are pages that are allocated by the kernel that are not trivially reclaimed. For example, the memory allocated for a loaded module would be in this category. By default, allocations are considered to be of this type HighAtomic - These are high-order allocations belonging to callers that cannot sleep or perform any IO. In practice, this is restricted to jumbo frame allocation for network receive. It is assumed that the allocations are short-lived Instead of having one MAX_ORDER-sized array of free lists in struct free_area, there is one for each type of reclaimability. Once a 2^MAX_ORDER block of pages is split for a type of allocation, it is added to the free-lists for that type, in effect reserving it. Hence, over time, pages of the different types can be clustered together. When the preferred freelists are expired, the largest possible block is taken from an alternative list. Buddies that are split from that large block are placed on the preferred allocation-type freelists to mitigate fragmentation. This implementation gives best-effort for low fragmentation in all zones. Ideally, min_free_kbytes needs to be set to a value equal to 4 * (1 << (MAX_ORDER-1)) pages in most cases. This would be 16384 on x86 and x86_64 for example. Our tests show that about 60-70% of physical memory can be allocated on a desktop after a few days uptime. In benchmarks and stress tests, we are finding that 80% of memory is available as contiguous blocks at the end of the test. To compare, a standard kernel was getting < 1% of memory as large pages on a desktop and about 8-12% of memory as large pages at the end of stress tests. Following this email are 12 patches that implement thie page grouping feature. The first patch introduces a mechanism for storing flags related to a whole block of pages. Then allocations are split between movable and all other allocations. Following that are patches to deal with per-cpu pages and make the mechanism configurable. The next patch moves free pages between lists when partially allocated blocks are used for pages of another migrate type. The second last patch groups reclaimable kernel allocations such as inode caches together. The final patch related to groupings keeps high-order atomic allocations. The last two patches are more concerned with control of fragmentation. The second last patch biases placement of non-movable allocations towards the start of memory. This is with a view of supporting memory hot-remove of DIMMs with higher PFNs in the future. The biasing could be enforced a lot heavier but it would cost. The last patch agressively clusters reclaimable pages like inode caches together. The fragmentation reduction strategy needs to track if pages within a block can be moved or reclaimed so that pages are freed to the appropriate list. This patch adds a bitmap for flags affecting a whole a MAX_ORDER block of pages. In non-SPARSEMEM configurations, the bitmap is stored in the struct zone and allocated during initialisation. SPARSEMEM statically allocates the bitmap in a struct mem_section so that bitmaps do not have to be resized during memory hotadd. This wastes a small amount of memory per unused section (usually sizeof(unsigned long)) but the complexity of dynamically allocating the memory is quite high. Additional credit to Andy Whitcroft who reviewed up an earlier implementation of the mechanism an suggested how to make it a *lot* cleaner. Signed-off-by:
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KAMEZAWA Hiroyuki authored
Current ia64 kernel flushes icache by lazy_mmu_prot_update() *after* set_pte(). This is too late. This patch removes lazy_mmu_prot_update and add modfied set_pte() for flushing if necessary. This patch flush icache of a page when new pte has exec bit. && new pte has present bit && new pte is user's page. && (old *ptep is not present || new pte's pfn is not same to old *ptep's ptn) && new pte's page has no Pg_arch_1 bit. Pg_arch_1 is set when a page is cache consistent. I think this condition checks are much easier to understand than considering "Where sync_icache_dcache() should be inserted ?". pte_user() for ia64 was removed by http://lkml.org/lkml/2007/6/12/67 as clean-up. So, I added it again. Signed-off-by:KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: "Luck, Tony" <tony.luck@intel.com> Cc: Christoph Lameter <clameter@sgi.com> Cc: Hugh Dickins <hugh@veritas.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Acked-by:
David S. Miller <davem@davemloft.net> Signed-off-by:
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KAMEZAWA Hiroyuki authored
In migration, a new page should be cache flushed before set_pte() in some archs which have virtually-tagged cache. Signed-off-by:
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Andrea Arcangeli authored
Swappiness isn't a safe sysctl. Setting it to 0 for example can hang a system. That's a corner case but even setting it to 10 or lower can waste enormous amounts of cpu without making much progress. We've customers who wants to use swappiness but they can't because of the current implementation (if you change it so the system stops swapping it really stops swapping and nothing works sane anymore if you really had to swap something to make progress). This patch from Kurt Garloff makes swappiness safer to use (no more huge cpu usage or hangs with low swappiness values). I think the prev_priority can also be nuked since it wastes 4 bytes per zone (that would be an incremental patch but I wait the nr_scan_[in]active to be nuked first for similar reasons). Clearly somebody at some point noticed how broken that thing was and they had to add min(priority, prev_priority) to give it some reliability, but they didn't go the last mile to nuke prev_priority too. Calculating distress only in function of not-racy priority is correct and sure more than enough without having to add randomness into the equation. Patch is tested on older kernels but it compiles and it's quite simple so... Overall I'm not very satisified by the swappiness tweak, since it doesn't rally do anything with the dirty pagecache that may be inactive. We need another kind of tweak that controls the inactive scan and tunes the can_writepage feature (not yet in mainline despite having submitted it a few times), not only the active one. That new tweak will tell the kernel how hard to scan the inactive list for pure clean pagecache (something the mainline kernel isn't capable of yet). We already have that feature working in all our enterprise kernels with the default reasonable tune, or they can't even run a readonly backup with tar without triggering huge write I/O. I think it should be available also in mainline later. Cc: Nick Piggin <npiggin@suse.de> Signed-off-by:
Kurt Garloff <garloff@suse.de> Signed-off-by:
Andrea Arcangeli <andrea@suse.de> Signed-off-by:
Fengguang Wu <wfg@mail.ustc.edu.cn> Signed-off-by:
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Christoph Lameter authored
The function of GFP_LEVEL_MASK seems to be unclear. In order to clear up the mystery we get rid of it and replace GFP_LEVEL_MASK with 3 sets of GFP flags: GFP_RECLAIM_MASK Flags used to control page allocator reclaim behavior. GFP_CONSTRAINT_MASK Flags used to limit where allocations can occur. GFP_SLAB_BUG_MASK Flags that the slab allocator BUG()s on. These replace the uses of GFP_LEVEL mask in the slab allocators and in vmalloc.c. The use of the flags not included in these sets may occur as a result of a slab allocation standing in for a page allocation when constructing scatter gather lists. Extraneous flags are cleared and not passed through to the page allocator. __GFP_MOVABLE/RECLAIMABLE, __GFP_COLD and __GFP_COMP will now be ignored if passed to a slab allocator. Change the allocation of allocator meta data in SLAB and vmalloc to not pass through flags listed in GFP_CONSTRAINT_MASK. SLAB already removes the __GFP_THISNODE flag for such allocations. Generalize that to also cover vmalloc. The use of GFP_CONSTRAINT_MASK also includes __GFP_HARDWALL. The impact of allocator metadata placement on access latency to the cachelines of the object itself is minimal since metadata is only referenced on alloc and free. The attempt is still made to place the meta data optimally but we consistently allow fallback both in SLAB and vmalloc (SLUB does not need to allocate metadata like that). Allocator metadata may serve multiple in kernel users and thus should not be subject to the limitations arising from a single allocation context. [akpm@linux-foundation.org: fix fallback_alloc()] Signed-off-by:
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Yasunori Goto authored
When a cpu is onlined on memory-less-node box, kernel panics due to touch NULL pointer of pgdat->kswapd. Current kswapd runs only nodes which have memory. So, calling of set_cpus_allowed() is not necessary for memory-less node. This is fix for it. Signed-off-by:
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Lee Schermerhorn authored
Here's a cut at fixing up uses of the online node map in generic code. mm/shmem.c:shmem_parse_mpol() Ensure nodelist is subset of nodes with memory. Use node_states[N_HIGH_MEMORY] as default for missing nodelist for interleave policy. mm/shmem.c:shmem_fill_super() initialize policy_nodes to node_states[N_HIGH_MEMORY] mm/page-writeback.c:highmem_dirtyable_memory() sum over nodes with memory mm/page_alloc.c:zlc_setup() allowednodes - use nodes with memory. mm/page_alloc.c:default_zonelist_order() average over nodes with memory. mm/page_alloc.c:find_next_best_node() skip nodes w/o memory. N_HIGH_MEMORY state mask may not be initialized at this time, unless we want to depend on early_calculate_totalpages() [see below]. Will ZONE_MOVABLE ever be configurable? mm/page_alloc.c:find_zone_movable_pfns_for_nodes() spread kernelcore over nodes with memory. This required calling early_calculate_totalpages() unconditionally, and populating N_HIGH_MEMORY node state therein from nodes in the early_node_map[]. If we can depend on this, we can eliminate the population of N_HIGH_MEMORY mask from __build_all_zonelists() and use the N_HIGH_MEMORY mask in find_next_best_node(). mm/mempolicy.c:mpol_check_policy() Ensure nodes specified for policy are subset of nodes with memory. [akpm@linux-foundation.org: fix warnings] Signed-off-by:
Lee Schermerhorn <lee.schermerhorn@hp.com> Acked-by:
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Christoph Lameter authored
cpusets try to ensure that any node added to a cpuset's mems_allowed is on-line and contains memory. The assumption was that online nodes contained memory. Thus, it is possible to add memoryless nodes to a cpuset and then add tasks to this cpuset. This results in continuous series of oom-kill and apparent system hang. Change cpusets to use node_states[N_HIGH_MEMORY] [a.k.a. node_memory_map] in place of node_online_map when vetting memories. Return error if admin attempts to write a non-empty mems_allowed node mask containing only memoryless-nodes. Signed-off-by:
Nishanth Aravamudan <nacc@us.ibm.com> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Mel Gorman <mel@skynet.ie> Signed-off-by:
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Christoph Lameter authored
GFP_THISNODE checks that the zone selected is within the pgdat (node) of the first zone of a nodelist. That only works if the node has memory. A memoryless node will have its first node on another pgdat (node). GFP_THISNODE currently will return simply memory on the first pgdat. Thus it is returning memory on other nodes. GFP_THISNODE should fail if there is no local memory on a node. Add a new set of zonelists for each node that only contain the nodes that belong to the zones itself so that no fallback is possible. Then modify gfp_type to pickup the right zone based on the presence of __GFP_THISNODE. Drop the existing GFP_THISNODE checks from the page_allocators hot path. Signed-off-by:
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Christoph Lameter authored
get_pfn_range_for_nid() is called multiple times for each node at boot time. Each time, it will warn about nodes with no memory, resulting in boot messages like: Node 0 active with no memory Node 0 active with no memory Node 0 active with no memory Node 0 active with no memory Node 0 active with no memory Node 0 active with no memory On node 0 totalpages: 0 Node 0 active with no memory Node 0 active with no memory DMA zone: 0 pages used for memmap Node 0 active with no memory Node 0 active with no memory Normal zone: 0 pages used for memmap Node 0 active with no memory Node 0 active with no memory Movable zone: 0 pages used for memmap and so on for each memoryless node. We already have the "On node N totalpages: ..." and other related messages, so drop the "Node N active with no memory" warnings. Signed-off-by: -
Christoph Lameter authored
We need the check for a node with cpu in zone reclaim. Zone reclaim will not allow remote zone reclaim if a node has a cpu. [Lee.Schermerhorn@hp.com: Move setup of N_CPU node state mask] Signed-off-by:
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Christoph Lameter authored
Online nodes now may have no memory. The checks and initialization must therefore be changed to no longer use the online functions. This will correctly initialize the interleave on bootup to only target nodes with memory and will make sys_move_pages return an error when a page is to be moved to a memoryless node. Similarly we will get an error if MPOL_BIND and MPOL_INTERLEAVE is used on a memoryless node. These are somewhat new semantics. So far one could specify memoryless nodes and we would maybe do the right thing and just ignore the node (or we'd do something strange like with MPOL_INTERLEAVE). If we want to allow the specification of memoryless nodes via memory policies then we need to keep checking for online nodes. Signed-off-by:
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Christoph Lameter authored
Processors on memoryless nodes must be able to fall back to remote nodes in order to get a profiling buffer. This may lead to excessive NUMA traffic but I think we should allow this rather than failing. Signed-off-by:
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Christoph Lameter authored
The checks for node_online in the uncached allocator are made to make sure that memory is available on these nodes. Thus switch all the checks to use N_HIGH_MEMORY and to N_ONLINE. Signed-off-by:
Jes Sorensen <jes@sgi.com> Acked-by:
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Christoph Lameter authored
Simply switch all for_each_online_node to for_each_node_state(NORMAL_MEMORY). That way SLUB only operates on nodes with regular memory. Any allocation attempt on a memoryless node or a node with just highmem will fall whereupon SLUB will fetch memory from a nearby node (depending on how memory policies and cpuset describe fallback). Signed-off-by:
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Christoph Lameter authored
Slab should not allocate control structures for nodes without memory. This may seem to work right now but its unreliable since not all allocations can fall back due to the use of GFP_THISNODE. Switching a few for_each_online_node's to N_NORMAL_MEMORY will allow us to only allocate for nodes that have regular memory. Signed-off-by:
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Christoph Lameter authored
A node without memory does not need a kswapd. So use the memory map instead of the online map when starting kswapd. Signed-off-by:
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