Currently swsusp saves the contents of highmem pages by copying them to the
normal zone which is quite inefficient (eg. it requires two normal pages
to be used for saving one highmem page). This may be improved by using
highmem for saving the contents of saveable highmem pages.
Namely, during the suspend phase of the suspend-resume cycle we try to
allocate as many free highmem pages as there are saveable highmem pages.
If there are not enough highmem image pages to store the contents of all of
the saveable highmem pages, some of them will be stored in the "normal"
memory. Next, we allocate as many free "normal" pages as needed to store
the (remaining) image data. We use a memory bitmap to mark the allocated
free pages (ie. highmem as well as "normal" image pages).
Now, we use another memory bitmap to mark all of the saveable pages
(highmem as well as "normal") and the contents of the saveable pages are
copied into the image pages. Then, the second bitmap is used to save the
pfns corresponding to the saveable pages and the first one is used to save
their data.
During the resume phase the pfns of the pages that were saveable during the
suspend are loaded from the image and used to mark the "unsafe" page
frames. Next, we try to allocate as many free highmem page frames as to
load all of the image data that had been in the highmem before the suspend
and we allocate so many free "normal" page frames that the total number of
allocated free pages (highmem and "normal") is equal to the size of the
image. While doing this we have to make sure that there will be some extra
free "normal" and "safe" page frames for two lists of PBEs constructed
later.
Now, the image data are loaded, if possible, into their "original" page
frames. The image data that cannot be written into their "original" page
frames are loaded into "safe" page frames and their "original" kernel
virtual addresses, as well as the addresses of the "safe" pages containing
their copies, are stored in one of two lists of PBEs.
One list of PBEs is for the copies of "normal" suspend pages (ie. "normal"
pages that were saveable during the suspend) and it is used in the same way
as previously (ie. by the architecture-dependent parts of swsusp). The
other list of PBEs is for the copies of highmem suspend pages. The pages
in this list are restored (in a reversible way) right before the
arch-dependent code is called.
Signed-off-by: Rafael J. Wysocki <rjw@sisk.pl>
Cc: Pavel Machek <pavel@ucw.cz>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Make swsusp use block device offsets instead of swap offsets to identify swap
locations and make it use the same code paths for writing as well as for
reading data.
This allows us to use the same code for handling swap files and swap
partitions and to simplify the code, eg. by dropping rw_swap_page_sync().
Signed-off-by: Rafael J. Wysocki <rjw@sisk.pl>
Cc: Pavel Machek <pavel@ucw.cz>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
The Linux kernel handles swap files almost in the same way as it handles swap
partitions and there are only two differences between these two types of swap
areas:
(1) swap files need not be contiguous,
(2) the header of a swap file is not in the first block of the partition
that holds it. From the swsusp's point of view (1) is not a problem,
because it is already taken care of by the swap-handling code, but (2) has
to be taken into consideration.
In principle the location of a swap file's header may be determined with the
help of appropriate filesystem driver. Unfortunately, however, it requires
the filesystem holding the swap file to be mounted, and if this filesystem is
journaled, it cannot be mounted during a resume from disk. For this reason we
need some other means by which swap areas can be identified.
For example, to identify a swap area we can use the partition that holds the
area and the offset from the beginning of this partition at which the swap
header is located.
The following patch allows swsusp to identify swap areas this way. It changes
swap_type_of() so that it takes an additional argument representing an offset
of the swap header within the partition represented by its first argument.
Signed-off-by: Rafael J. Wysocki <rjw@sisk.pl>
Acked-by: Pavel Machek <pavel@ucw.cz>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Make radix tree lookups safe to be performed without locks. Readers are
protected against nodes being deleted by using RCU based freeing. Readers
are protected against new node insertion by using memory barriers to ensure
the node itself will be properly written before it is visible in the radix
tree.
Each radix tree node keeps a record of their height (above leaf nodes).
This height does not change after insertion -- when the radix tree is
extended, higher nodes are only inserted in the top. So a lookup can take
the pointer to what is *now* the root node, and traverse down it even if
the tree is concurrently extended and this node becomes a subtree of a new
root.
"Direct" pointers (tree height of 0, where root->rnode points directly to
the data item) are handled by using the low bit of the pointer to signal
whether rnode is a direct pointer or a pointer to a radix tree node.
When a reader wants to traverse the next branch, they will take a copy of
the pointer. This pointer will be either NULL (and the branch is empty) or
non-NULL (and will point to a valid node).
[akpm@osdl.org: cleanups]
[Lee.Schermerhorn@hp.com: bugfixes, comments, simplifications]
[clameter@sgi.com: build fix]
Signed-off-by: Nick Piggin <npiggin@suse.de>
Cc: "Paul E. McKenney" <paulmck@us.ibm.com>
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Currently we we use the lru head link of the second page of a compound page
to hold its destructor. This was ok when it was purely an internal
implmentation detail. However, hugetlbfs overrides this destructor
violating the layering. Abstract this out as explicit calls, also
introduce a type for the callback function allowing them to be type
checked. For each callback we pre-declare the function, causing a type
error on definition rather than on use elsewhere.
[akpm@osdl.org: cleanups]
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Replace all uses of kmem_cache_t with struct kmem_cache.
The patch was generated using the following script:
#!/bin/sh
#
# Replace one string by another in all the kernel sources.
#
set -e
for file in `find * -name "*.c" -o -name "*.h"|xargs grep -l $1`; do
quilt add $file
sed -e "1,\$s/$1/$2/g" $file >/tmp/$$
mv /tmp/$$ $file
quilt refresh
done
The script was run like this
sh replace kmem_cache_t "struct kmem_cache"
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
SLAB_DMA is an alias of GFP_DMA. This is the last one so we
remove the leftover comment too.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
SLAB_KERNEL is an alias of GFP_KERNEL.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
SLAB_ATOMIC is an alias of GFP_ATOMIC
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
SLAB_USER is an alias of GFP_USER
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
SLAB_NOFS is an alias of GFP_NOFS.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
SLAB_NOIO is an alias of GFP_NOIO with a single instance of use.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
SLAB_LEVEL_MASK is only used internally to the slab and is
and alias of GFP_LEVEL_MASK.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
It is only used internally in the slab.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
x86 NUMA systems only define bootmem for node 0. alloc_bootmem_node() and
friends therefore ignore the passed pgdat and use NODE_DATA(0) in all
cases. This leads to the following warnings as we are not using the passed
parameter:
.../mm/page_alloc.c: In function 'zone_wait_table_init':
.../mm/page_alloc.c:2259: warning: unused variable 'pgdat'
One option would be to define all variables used with these macros
__attribute__ ((unused)), but this would leave us exposed should these
become genuinely unused.
The key here is that we _are_ using the value, we ignore it but that is a
deliberate action. This patch adds a nested local variable within the
alloc_bootmem_node helper to which the pgdat parameter is assigned making
it 'used'. The nested local is marked __attribute__ ((unused)) to silence
this same warning for it.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: Andi Kleen <ak@suse.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
NUMA node ids are passed as either int or unsigned int almost exclusivly
page_to_nid and zone_to_nid both return unsigned long. This is a throw
back to when page_to_nid was a #define and was thus exposing the real type
of the page flags field.
In addition to fixing up the definitions of page_to_nid and zone_to_nid I
audited the users of these functions identifying the following incorrect
uses:
1) mm/page_alloc.c show_node() -- printk dumping the node id,
2) include/asm-ia64/pgalloc.h pgtable_quicklist_free() -- comparison
against numa_node_id() which returns an int from cpu_to_node(), and
3) mm/mpolicy.c check_pte_range -- used as an index in node_isset which
uses bit_set which in generic code takes an int.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: "Luck, Tony" <tony.luck@intel.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Remove all uses of kmem_cache_t (the most were left in slab.h). The
typedef for kmem_cache_t is then only necessary for other kernel
subsystems. Add a comment to that effect.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
The names_cachep is used for getname() and putname(). So lets put it into
fs.h near those two definitions.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
fs_cachep is only used in kernel/exit.c and in kernel/fork.c.
It is used to store fs_struct items so it should be placed in linux/fs_struct.h
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
filp_cachep is only used in fs/file_table.c and in fs/dcache.c where
it is defined.
Move it to related definitions in linux/file.h.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Proper place is in file.h since files_cachep uses are rated to file I/O.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
vm_area_cachep is used to store vm_area_structs. So move to mm.h.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Move sighand_cachep definitioni to linux/signal.h
The sighand cache is only used in fs/exec.c and kernel/fork.c. It is defined
in kernel/fork.c but only used in fs/exec.c.
The sighand_cachep is related to signal processing. So add the definition to
signal.h.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Remove bio_cachep from slab.h - it no longer exists.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Node-aware allocation of skbs for the receive path.
Details:
- __alloc_skb gets a new node argument and cals the node-aware
slab functions with it.
- netdev_alloc_skb passed the node number it gets from dev_to_node
to it, everyone else passes -1 (any node)
Signed-off-by: Christoph Hellwig <hch@lst.de>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: "David S. Miller" <davem@davemloft.net>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
For node-aware skb allocations we need information about the node in struct
net_device or struct device. Davem suggested to put it into struct device
which this patch does.
In particular:
- struct device gets a new int numa_node member if CONFIG_NUMA is set
- there are two new helpers, dev_to_node and set_dev_node to
transparently deal with the non-numa case
- for pci devices the node-info is set to the value we get from
pcibus_to_node.
Note that for some architectures pcibus_to_node doesn't work yet at the time
we call it currently. This is harmless and will just mean skb allocations
aren't node-local on this architectures until the implementation of
pcibus_to_node on these architectures have been updated (There are patches for
x86 and x86_64 floating around)
[akpm@osdl.org: cleanup]
Signed-off-by: Christoph Hellwig <hch@lst.de>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: Greg KH <greg@kroah.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
We have variants of kmalloc and kmem_cache_alloc that leave leak tracking to
the caller. This is used for subsystem-specific allocators like skb_alloc.
To make skb_alloc node-aware we need similar routines for the node-aware slab
allocator, which this patch adds.
Note that the code is rather ugly, but it mirrors the non-node-aware code 1:1:
[akpm@osdl.org: add module export]
Signed-off-by: Christoph Hellwig <hch@lst.de>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Make kmap_atomic/kunmap_atomic denote a pagefault disabled scope. All non
trivial implementations already do this anyway.
Signed-off-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Acked-by: Nick Piggin <npiggin@suse.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Introduce pagefault_{disable,enable}() and use these where previously we did
manual preempt increments/decrements to make the pagefault handler do the
atomic thing.
Currently they still rely on the increased preempt count, but do not rely on
the disabled preemption, this might go away in the future.
(NOTE: the extra barrier() in pagefault_disable might fix some holes on
machines which have too many registers for their own good)
[heiko.carstens@de.ibm.com: s390 fix]
Signed-off-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Acked-by: Nick Piggin <npiggin@suse.de>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Signed-off-by: Heiko Carstens <heiko.carstens@de.ibm.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Following up with the work on shared page table done by Dave McCracken. This
set of patch target shared page table for hugetlb memory only.
The shared page table is particular useful in the situation of large number of
independent processes sharing large shared memory segments. In the normal
page case, the amount of memory saved from process' page table is quite
significant. For hugetlb, the saving on page table memory is not the primary
objective (as hugetlb itself already cuts down page table overhead
significantly), instead, the purpose of using shared page table on hugetlb is
to allow faster TLB refill and smaller cache pollution upon TLB miss.
With PT sharing, pte entries are shared among hundreds of processes, the cache
consumption used by all the page table is smaller and in return, application
gets much higher cache hit ratio. One other effect is that cache hit ratio
with hardware page walker hitting on pte in cache will be higher and this
helps to reduce tlb miss latency. These two effects contribute to higher
application performance.
Signed-off-by: Ken Chen <kenneth.w.chen@intel.com>
Acked-by: Hugh Dickins <hugh@veritas.com>
Cc: Dave McCracken <dmccr@us.ibm.com>
Cc: William Lee Irwin III <wli@holomorphy.com>
Cc: "Luck, Tony" <tony.luck@intel.com>
Cc: Paul Mackerras <paulus@samba.org>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: David Gibson <david@gibson.dropbear.id.au>
Cc: Adam Litke <agl@us.ibm.com>
Cc: Paul Mundt <lethal@linux-sh.org>
Cc: "David S. Miller" <davem@davemloft.net>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Add an arch_alloc_page to match arch_free_page.
Signed-off-by: Nick Piggin <npiggin@suse.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
The new swap token patches replace the current token traversal algo. The old
algo had a crude timeout parameter that was used to handover the token from
one task to another. This algo, transfers the token to the tasks that are in
need of the token. The urgency for the token is based on the number of times
a task is required to swap-in pages. Accordingly, the priority of a task is
incremented if it has been badly affected due to swap-outs. To ensure that
the token doesnt bounce around rapidly, the token holders are given a priority
boost. The priority of tasks is also decremented, if their rate of swap-in's
keeps reducing. This way, the condition to check whether to pre-empt the swap
token, is a matter of comparing two task's priority fields.
[akpm@osdl.org: cleanups]
Signed-off-by: Ashwin Chaugule <ashwin.chaugule@celunite.com>
Cc: Rik van Riel <riel@redhat.com>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Rearrange the struct members in the 'struct zonelist_cache' structure, so
as to put the readonly (once initialized) z_to_n[] array first, where it
will come right after the zones[] array in struct zonelist.
This pretty much eliminates the chance that the two frequently written
elements of 'struct zonelist_cache', the fullzones bitmap and last_full_zap
times, will end up on the same cache line as the performance sensitive,
frequently read, never (after init) written zones[] array.
Keeping frequently written data off frequently read cache lines is good for
performance.
Thanks to Rohit Seth for the suggestion.
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Optimize the critical zonelist scanning for free pages in the kernel memory
allocator by caching the zones that were found to be full recently, and
skipping them.
Remembers the zones in a zonelist that were short of free memory in the
last second. And it stashes a zone-to-node table in the zonelist struct,
to optimize that conversion (minimize its cache footprint.)
Recent changes:
This differs in a significant way from a similar patch that I
posted a week ago. Now, instead of having a nodemask_t of
recently full nodes, I have a bitmask of recently full zones.
This solves a problem that last weeks patch had, which on
systems with multiple zones per node (such as DMA zone) would
take seeing any of these zones full as meaning that all zones
on that node were full.
Also I changed names - from "zonelist faster" to "zonelist cache",
as that seemed to better convey what we're doing here - caching
some of the key zonelist state (for faster access.)
See below for some performance benchmark results. After all that
discussion with David on why I didn't need them, I went and got
some ;). I wanted to verify that I had not hurt the normal case
of memory allocation noticeably. At least for my one little
microbenchmark, I found (1) the normal case wasn't affected, and
(2) workloads that forced scanning across multiple nodes for
memory improved up to 10% fewer System CPU cycles and lower
elapsed clock time ('sys' and 'real'). Good. See details, below.
I didn't have the logic in get_page_from_freelist() for various
full nodes and zone reclaim failures correct. That should be
fixed up now - notice the new goto labels zonelist_scan,
this_zone_full, and try_next_zone, in get_page_from_freelist().
There are two reasons I persued this alternative, over some earlier
proposals that would have focused on optimizing the fake numa
emulation case by caching the last useful zone:
1) Contrary to what I said before, we (SGI, on large ia64 sn2 systems)
have seen real customer loads where the cost to scan the zonelist
was a problem, due to many nodes being full of memory before
we got to a node we could use. Or at least, I think we have.
This was related to me by another engineer, based on experiences
from some time past. So this is not guaranteed. Most likely, though.
The following approach should help such real numa systems just as
much as it helps fake numa systems, or any combination thereof.
2) The effort to distinguish fake from real numa, using node_distance,
so that we could cache a fake numa node and optimize choosing
it over equivalent distance fake nodes, while continuing to
properly scan all real nodes in distance order, was going to
require a nasty blob of zonelist and node distance munging.
The following approach has no new dependency on node distances or
zone sorting.
See comment in the patch below for a description of what it actually does.
Technical details of note (or controversy):
- See the use of "zlc_active" and "did_zlc_setup" below, to delay
adding any work for this new mechanism until we've looked at the
first zone in zonelist. I figured the odds of the first zone
having the memory we needed were high enough that we should just
look there, first, then get fancy only if we need to keep looking.
- Some odd hackery was needed to add items to struct zonelist, while
not tripping up the custom zonelists built by the mm/mempolicy.c
code for MPOL_BIND. My usual wordy comments below explain this.
Search for "MPOL_BIND".
- Some per-node data in the struct zonelist is now modified frequently,
with no locking. Multiple CPU cores on a node could hit and mangle
this data. The theory is that this is just performance hint data,
and the memory allocator will work just fine despite any such mangling.
The fields at risk are the struct 'zonelist_cache' fields 'fullzones'
(a bitmask) and 'last_full_zap' (unsigned long jiffies). It should
all be self correcting after at most a one second delay.
- This still does a linear scan of the same lengths as before. All
I've optimized is making the scan faster, not algorithmically
shorter. It is now able to scan a compact array of 'unsigned
short' in the case of many full nodes, so one cache line should
cover quite a few nodes, rather than each node hitting another
one or two new and distinct cache lines.
- If both Andi and Nick don't find this too complicated, I will be
(pleasantly) flabbergasted.
- I removed the comment claiming we only use one cachline's worth of
zonelist. We seem, at least in the fake numa case, to have put the
lie to that claim.
- I pay no attention to the various watermarks and such in this performance
hint. A node could be marked full for one watermark, and then skipped
over when searching for a page using a different watermark. I think
that's actually quite ok, as it will tend to slightly increase the
spreading of memory over other nodes, away from a memory stressed node.
===============
Performance - some benchmark results and analysis:
This benchmark runs a memory hog program that uses multiple
threads to touch alot of memory as quickly as it can.
Multiple runs were made, touching 12, 38, 64 or 90 GBytes out of
the total 96 GBytes on the system, and using 1, 19, 37, or 55
threads (on a 56 CPU system.) System, user and real (elapsed)
timings were recorded for each run, shown in units of seconds,
in the table below.
Two kernels were tested - 2.6.18-mm3 and the same kernel with
this zonelist caching patch added. The table also shows the
percentage improvement the zonelist caching sys time is over
(lower than) the stock *-mm kernel.
number 2.6.18-mm3 zonelist-cache delta (< 0 good) percent
GBs N ------------ -------------- ---------------- systime
mem threads sys user real sys user real sys user real better
12 1 153 24 177 151 24 176 -2 0 -1 1%
12 19 99 22 8 99 22 8 0 0 0 0%
12 37 111 25 6 112 25 6 1 0 0 -0%
12 55 115 25 5 110 23 5 -5 -2 0 4%
38 1 502 74 576 497 73 570 -5 -1 -6 0%
38 19 426 78 48 373 76 39 -53 -2 -9 12%
38 37 544 83 36 547 82 36 3 -1 0 -0%
38 55 501 77 23 511 80 24 10 3 1 -1%
64 1 917 125 1042 890 124 1014 -27 -1 -28 2%
64 19 1118 138 119 965 141 103 -153 3 -16 13%
64 37 1202 151 94 1136 150 81 -66 -1 -13 5%
64 55 1118 141 61 1072 140 58 -46 -1 -3 4%
90 1 1342 177 1519 1275 174 1450 -67 -3 -69 4%
90 19 2392 199 192 2116 189 176 -276 -10 -16 11%
90 37 3313 238 175 2972 225 145 -341 -13 -30 10%
90 55 1948 210 104 1843 213 100 -105 3 -4 5%
Notes:
1) This test ran a memory hog program that started a specified number N of
threads, and had each thread allocate and touch 1/N'th of
the total memory to be used in the test run in a single loop,
writing a constant word to memory, one store every 4096 bytes.
Watching this test during some earlier trial runs, I would see
each of these threads sit down on one CPU and stay there, for
the remainder of the pass, a different CPU for each thread.
2) The 'real' column is not comparable to the 'sys' or 'user' columns.
The 'real' column is seconds wall clock time elapsed, from beginning
to end of that test pass. The 'sys' and 'user' columns are total
CPU seconds spent on that test pass. For a 19 thread test run,
for example, the sum of 'sys' and 'user' could be up to 19 times the
number of 'real' elapsed wall clock seconds.
3) Tests were run on a fresh, single-user boot, to minimize the amount
of memory already in use at the start of the test, and to minimize
the amount of background activity that might interfere.
4) Tests were done on a 56 CPU, 28 Node system with 96 GBytes of RAM.
5) Notice that the 'real' time gets large for the single thread runs, even
though the measured 'sys' and 'user' times are modest. I'm not sure what
that means - probably something to do with it being slow for one thread to
be accessing memory along ways away. Perhaps the fake numa system, running
ostensibly the same workload, would not show this substantial degradation
of 'real' time for one thread on many nodes -- lets hope not.
6) The high thread count passes (one thread per CPU - on 55 of 56 CPUs)
ran quite efficiently, as one might expect. Each pair of threads needed
to allocate and touch the memory on the node the two threads shared, a
pleasantly parallizable workload.
7) The intermediate thread count passes, when asking for alot of memory forcing
them to go to a few neighboring nodes, improved the most with this zonelist
caching patch.
Conclusions:
* This zonelist cache patch probably makes little difference one way or the
other for most workloads on real numa hardware, if those workloads avoid
heavy off node allocations.
* For memory intensive workloads requiring substantial off-node allocations
on real numa hardware, this patch improves both kernel and elapsed timings
up to ten per-cent.
* For fake numa systems, I'm optimistic, but will have to leave that up to
Rohit Seth to actually test (once I get him a 2.6.18 backport.)
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: David Rientjes <rientjes@cs.washington.edu>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
The zone table is mostly not needed. If we have a node in the page flags
then we can get to the zone via NODE_DATA() which is much more likely to be
already in the cpu cache.
In case of SMP and UP NODE_DATA() is a constant pointer which allows us to
access an exact replica of zonetable in the node_zones field. In all of
the above cases there will be no need at all for the zone table.
The only remaining case is if in a NUMA system the node numbers do not fit
into the page flags. In that case we make sparse generate a table that
maps sections to nodes and use that table to to figure out the node number.
This table is sized to fit in a single cache line for the known 32 bit
NUMA platform which makes it very likely that the information can be
obtained without a cache miss.
For sparsemem the zone table seems to be have been fairly large based on
the maximum possible number of sections and the number of zones per node.
There is some memory saving by removing zone_table. The main benefit is to
reduce the cache foootprint of the VM from the frequent lookups of zones.
Plus it simplifies the page allocator.
[akpm@osdl.org: build fix]
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
With CONFIG_SMP=n:
drivers/input/ff-memless.c:384: warning: implicit declaration of function 'local_bh_disable'
drivers/input/ff-memless.c:393: warning: implicit declaration of function 'local_bh_enable'
Really linux/spinlock.h should include linux/interrupt.h. But interrupt.h
includes sched.h which will need spinlock.h.
So the patch breaks the _bh declarations out into a separate header and
includes it in both interrupt.h and spinlock.h.
Cc: "Randy.Dunlap" <rdunlap@xenotime.net>
Cc: Andi Kleen <ak@suse.de>
Cc: <stable@kernel.org>
Cc: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Enhanced resolution for time measurement functions.
Signed-off-by: Pavel Pisa <pisa@cmp.felk.cvut.cz>
Signed-off-by: Sascha Hauer <s.hauer@pengutronix.de>
Signed-off-by: Russell King <rmk+kernel@arm.linux.org.uk>
Add support to suspend and resume, using the
H1940's bootloader
Signed-off-by: Ben Dooks <ben-linux@fluff.org>
Signed-off-by: Arnaud Patard <arnaud.patard@rtp-net.org>
Signed-off-by: Russell King <rmk+kernel@arm.linux.org.uk>
This patch fixes an error in the numbering of the disk LEDs on the
Linksys NSLU2. The error crept in because the physical location
of the LEDs has the Disk 2 LED *above* the Disk 1 LED.
Thanks to Gordon Farquharson for reporting this.
Signed-off-by: Rod Whitby <rod@whitby.id.au>
Signed-off-by: Russell King <rmk+kernel@arm.linux.org.uk>
This is done in a completely lockless fashion. Bits 0 to 31 of the count
are provided by the hardware while bits 32 to 62 are stored in memory.
The top bit in memory is used to synchronize with the hardware count
half-period. When the top bit of both counters (hardware and in memory)
differ then the memory is updated with a new value, incrementing it when
the hardware counter wraps around. Because a word store in memory is
atomic then the incremented value will always be in synch with the top
bit indicating to any potential concurrent reader if the value in memory
is up to date or not wrt the needed increment. And any race in updating
the value in memory is harmless as the same value would be stored more
than once.
Signed-off-by: Nicolas Pitre <nico@cam.org>
Signed-off-by: Russell King <rmk+kernel@arm.linux.org.uk>
On ARM all divisions have to be performed "manually". For 64-bit
divisions that may take more than a hundred cycles in many cases.
With 32-bit divisions gcc already use the recyprocal of constant
divisors to perform a multiplication, but not with 64-bit divisions.
Since the kernel is increasingly relying upon 64-bit divisions it is
worth optimizing at least those cases where the divisor is a constant.
This is what this patch does using plain C code that gets optimized away
at compile time.
For example, despite the amount of added C code, do_div(x, 10000) now
produces the following assembly code (where x is assigned to r0-r1):
adr r4, .L0
ldmia r4, {r4-r5}
umull r2, r3, r4, r0
mov r2, #0
umlal r3, r2, r5, r0
umlal r3, r2, r4, r1
mov r3, #0
umlal r2, r3, r5, r1
mov r0, r2, lsr #11
orr r0, r0, r3, lsl #21
mov r1, r3, lsr #11
...
.L0:
.word 948328779
.word 879609302
which is the fastest that can be done for any value of x in that case,
many times faster than the __do_div64 code (except for the small x value
space for which the result ends up being zero or a single bit).
The fact that this code is generated inline produces a tiny increase in
.text size, but not significant compared to the needed code around each
__do_div64 call site this code is replacing.
The algorithm used has been validated on a 16-bit scale for all possible
values, and then recodified for 64-bit values. Furthermore I've been
running it with the final BUG_ON() uncommented for over two months now
with no problem.
Note that this new code is compiled with gcc versions 4.0 or later.
Earlier gcc versions proved themselves too problematic and only the
original code is used with them.
Signed-off-by: Nicolas Pitre <nico@cam.org>
Signed-off-by: Russell King <rmk+kernel@arm.linux.org.uk>
Add Tg3_FLG2_IS_NIC flag to unambiguously determine whether the
device is NIC or onboard. Previously, the EEPROM_WRITE_PROT flag was
overloaded to also mean onboard. With the separation, we can
support some devices that are onboard but do not use eeprom write
protect.
Signed-off-by: Michael Chan <mchan@broadcom.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
Disables auditing in ipsec when CONFIG_AUDITSYSCALL is
disabled in the kernel.
Also includes a bug fix for xfrm_state.c as a result of
original ipsec audit patch.
Signed-off-by: Joy Latten <latten@austin.ibm.com>
Signed-off-by: James Morris <jmorris@namei.org>
Signed-off-by: David S. Miller <davem@davemloft.net>
An audit message occurs when an ipsec SA
or ipsec policy is created/deleted.
Signed-off-by: Joy Latten <latten@austin.ibm.com>
Signed-off-by: James Morris <jmorris@namei.org>
Signed-off-by: David S. Miller <davem@davemloft.net>
include/net/irda/irlan_filter.h:31: warning: 'struct seq_file' declared inside parameter list
include/net/irda/irlan_filter.h:31: warning: its scope is only this definition or declaration, which is probably not what you want
Signed-off-by: Randy Dunlap <randy.dunlap@oracle.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
Signed-off-by: Yasuyuki Kozakai <yasuyuki.kozakai@toshiba.co.jp>
Signed-off-by: Patrick McHardy <kaber@trash.net>
Signed-off-by: David S. Miller <davem@davemloft.net>
A lot of cypher modes need multiplications in GF(2^128). LRW, ABL, GCM...
I use functions from this library in my LRW implementation and I will
also use them in my ABL (Arbitrary Block Length, an unencumbered (correct
me if I am wrong, wide block cipher mode).
Elements of GF(2^128) must be presented as u128 *, it encourages automatic
and proper alignment.
The library contains support for two different representations of GF(2^128),
see the comment in gf128mul.h. There different levels of optimization
(memory/speed tradeoff).
The code is based on work by Dr Brian Gladman. Notable changes:
- deletion of two optimization modes
- change from u32 to u64 for faster handling on 64bit machines
- support for 'bbe' representation in addition to the, already implemented,
'lle' representation.
- move 'inline void' functions from header to 'static void' in the
source file
- update to use the linux coding style conventions
The original can be found at:
http://fp.gladman.plus.com/AES/modes.vc8.19-06-06.zip
The copyright (and GPL statement) of the original author is preserved.
Signed-off-by: Rik Snel <rsnel@cube.dyndns.org>
Signed-off-by: Herbert Xu <herbert@gondor.apana.org.au>
Rafael J. Wysocki