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  1. 21 10月, 2007 1 次提交
    • A
      [PATCH] audit: watching subtrees · 74c3cbe3
      由 Al Viro 提交于 
      New kind of audit rule predicates: "object is visible in given subtree".
The part that can be sanely implemented, that is.  Limitations:
	* if you have hardlink from outside of tree, you'd better watch
it too (or just watch the object itself, obviously)
	* if you mount something under a watched tree, tell audit
that new chunk should be added to watched subtrees
	* if you umount something in a watched tree and it's still mounted
elsewhere, you will get matches on events happening there.  New command
tells audit to recalculate the trees, trimming such sources of false
positives.

Note that it's _not_ about path - if something mounted in several places
(multiple mount, bindings, different namespaces, etc.), the match does
_not_ depend on which one we are using for access.
Signed-off-by: NAl Viro <viro@zeniv.linux.org.uk>
      74c3cbe3
  3. 20 10月, 2007 6 次提交
    • S
      Hook up group scheduler with control groups · 68318b8e
      由 Srivatsa Vaddagiri 提交于 
      Enable "cgroup" (formerly containers) based fair group scheduling.  This
will let administrator create arbitrary groups of tasks (using "cgroup"
pseudo filesystem) and control their cpu bandwidth usage.

[akpm@linux-foundation.org: fix cpp condition]
Signed-off-by: NSrivatsa Vaddagiri <vatsa@linux.vnet.ibm.com>
Signed-off-by: NDhaval Giani <dhaval@linux.vnet.ibm.com>
Cc: Randy Dunlap <randy.dunlap@oracle.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Paul Menage <menage@google.com>
Cc: Ingo Molnar <mingo@elte.hu>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      68318b8e
    • S
      cgroups: implement namespace tracking subsystem · 858d72ea
      由 Serge E. Hallyn 提交于 
      When a task enters a new namespace via a clone() or unshare(), a new cgroup
is created and the task moves into it.

This version names cgroups which are automatically created using
cgroup_clone() as "node_<pid>" where pid is the pid of the unsharing or
cloned process.  (Thanks Pavel for the idea) This is safe because if the
process unshares again, it will create

	/cgroups/(...)/node_<pid>/node_<pid>

The only possibilities (AFAICT) for a -EEXIST on unshare are

	1. pid wraparound
	2. a process fails an unshare, then tries again.

Case 1 is unlikely enough that I ignore it (at least for now).  In case 2, the
node_<pid> will be empty and can be rmdir'ed to make the subsequent unshare()
succeed.

Changelog:
	Name cloned cgroups as "node_<pid>".

[clg@fr.ibm.com: fix order of cgroup subsystems in init/Kconfig]
Signed-off-by: NSerge E. Hallyn <serue@us.ibm.com>
Cc: Paul Menage <menage@google.com>
Signed-off-by: NCedric Le Goater <clg@fr.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      858d72ea
    • P
      Task Control Groups: simple task cgroup debug info subsystem · 006cb992
      由 Paul Menage 提交于 
      This example subsystem exports debugging information as an aid to diagnosing
refcount leaks, etc, in the cgroup framework.
Signed-off-by: NPaul Menage <menage@google.com>
Cc: Serge E. Hallyn <serue@us.ibm.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Paul Jackson <pj@sgi.com>
Cc: Kirill Korotaev <dev@openvz.org>
Cc: Herbert Poetzl <herbert@13thfloor.at>
Cc: Srivatsa Vaddagiri <vatsa@in.ibm.com>
Cc: Cedric Le Goater <clg@fr.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      006cb992
    • P
      Task Control Groups: example CPU accounting subsystem · 62d0df64
      由 Paul Menage 提交于 
      This example demonstrates how to use the generic cgroup subsystem for a
simple resource tracker that counts, for the processes in a cgroup, the
total CPU time used and the %CPU used in the last complete 10 second interval.

Portions contributed by Balbir Singh <balbir@in.ibm.com>
Signed-off-by: NPaul Menage <menage@google.com>
Cc: Serge E. Hallyn <serue@us.ibm.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Paul Jackson <pj@sgi.com>
Cc: Kirill Korotaev <dev@openvz.org>
Cc: Herbert Poetzl <herbert@13thfloor.at>
Cc: Srivatsa Vaddagiri <vatsa@in.ibm.com>
Cc: Cedric Le Goater <clg@fr.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      62d0df64
    • P
      Task Control Groups: make cpusets a client of cgroups · 8793d854
      由 Paul Menage 提交于 
      Remove the filesystem support logic from the cpusets system and makes cpusets
a cgroup subsystem

The "cpuset" filesystem becomes a dummy filesystem; attempts to mount it get
passed through to the cgroup filesystem with the appropriate options to
emulate the old cpuset filesystem behaviour.
Signed-off-by: NPaul Menage <menage@google.com>
Cc: Serge E. Hallyn <serue@us.ibm.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Paul Jackson <pj@sgi.com>
Cc: Kirill Korotaev <dev@openvz.org>
Cc: Herbert Poetzl <herbert@13thfloor.at>
Cc: Srivatsa Vaddagiri <vatsa@in.ibm.com>
Cc: Cedric Le Goater <clg@fr.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      8793d854
    • P
      Task Control Groups: basic task cgroup framework · ddbcc7e8
      由 Paul Menage 提交于 
      Generic Process Control Groups
--------------------------

There have recently been various proposals floating around for
resource management/accounting and other task grouping subsystems in
the kernel, including ResGroups, User BeanCounters, NSProxy
cgroups, and others.  These all need the basic abstraction of being
able to group together multiple processes in an aggregate, in order to
track/limit the resources permitted to those processes, or control
other behaviour of the processes, and all implement this grouping in
different ways.

This patchset provides a framework for tracking and grouping processes
into arbitrary "cgroups" and assigning arbitrary state to those
groupings, in order to control the behaviour of the cgroup as an
aggregate.

The intention is that the various resource management and
virtualization/cgroup efforts can also become task cgroup
clients, with the result that:

- the userspace APIs are (somewhat) normalised

- it's easier to test e.g. the ResGroups CPU controller in
 conjunction with the BeanCounters memory controller, or use either of
them as the resource-control portion of a virtual server system.

- the additional kernel footprint of any of the competing resource
 management systems is substantially reduced, since it doesn't need
 to provide process grouping/containment, hence improving their
 chances of getting into the kernel

This patch:

Add the main task cgroups framework - the cgroup filesystem, and the
basic structures for tracking membership and associating subsystem state
objects to tasks.
Signed-off-by: NPaul Menage <menage@google.com>
Cc: Serge E. Hallyn <serue@us.ibm.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Paul Jackson <pj@sgi.com>
Cc: Kirill Korotaev <dev@openvz.org>
Cc: Herbert Poetzl <herbert@13thfloor.at>
Cc: Srivatsa Vaddagiri <vatsa@in.ibm.com>
Cc: Cedric Le Goater <clg@fr.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      ddbcc7e8
  5. 17 10月, 2007 3 次提交
  7. 15 10月, 2007 6 次提交
  9. 20 9月, 2007 1 次提交
  11. 01 8月, 2007 2 次提交
  13. 26 7月, 2007 1 次提交
  15. 18 7月, 2007 1 次提交
  17. 17 7月, 2007 4 次提交
  19. 10 7月, 2007 1 次提交
  21. 17 5月, 2007 2 次提交
  23. 11 5月, 2007 5 次提交
    • D
      signal/timer/event: eventfd core · e1ad7468
      由 Davide Libenzi 提交于 
      This is a very simple and light file descriptor, that can be used as event
wait/dispatch by userspace (both wait and dispatch) and by the kernel
(dispatch only).  It can be used instead of pipe(2) in all cases where those
would simply be used to signal events.  Their kernel overhead is much lower
than pipes, and they do not consume two fds.  When used in the kernel, it can
offer an fd-bridge to enable, for example, functionalities like KAIO or
syslets/threadlets to signal to an fd the completion of certain operations.
But more in general, an eventfd can be used by the kernel to signal readiness,
in a POSIX poll/select way, of interfaces that would otherwise be incompatible
with it.  The API is:

int eventfd(unsigned int count);

The eventfd API accepts an initial "count" parameter, and returns an eventfd
fd.  It supports poll(2) (POLLIN, POLLOUT, POLLERR), read(2) and write(2).

The POLLIN flag is raised when the internal counter is greater than zero.

The POLLOUT flag is raised when at least a value of "1" can be written to the
internal counter.

The POLLERR flag is raised when an overflow in the counter value is detected.

The write(2) operation can never overflow the counter, since it blocks (unless
O_NONBLOCK is set, in which case -EAGAIN is returned).

But the eventfd_signal() function can do it, since it's supposed to not sleep
during its operation.

The read(2) function reads the __u64 counter value, and reset the internal
value to zero.  If the value read is equal to (__u64) -1, an overflow happened
on the internal counter (due to 2^64 eventfd_signal() posts that has never
been retired - unlickely, but possible).

The write(2) call writes an __u64 count value, and adds it to the current
counter.  The eventfd fd supports O_NONBLOCK also.

On the kernel side, we have:

struct file *eventfd_fget(int fd);
int eventfd_signal(struct file *file, unsigned int n);

The eventfd_fget() should be called to get a struct file* from an eventfd fd
(this is an fget() + check of f_op being an eventfd fops pointer).

The kernel can then call eventfd_signal() every time it wants to post an event
to userspace.  The eventfd_signal() function can be called from any context.
An eventfd() simple test and bench is available here:

http://www.xmailserver.org/eventfd-bench.c

This is the eventfd-based version of pipetest-4 (pipe(2) based):

http://www.xmailserver.org/pipetest-4.c

Not that performance matters much in the eventfd case, but eventfd-bench
shows almost as double as performance than pipetest-4.

[akpm@linux-foundation.org: fix i386 build]
[akpm@linux-foundation.org: add sys_eventfd to sys_ni.c]
Signed-off-by: NDavide Libenzi <davidel@xmailserver.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      e1ad7468
    • D
      signal/timer/event: timerfd core · b215e283
      由 Davide Libenzi 提交于 
      This patch introduces a new system call for timers events delivered though
file descriptors.  This allows timer event to be used with standard POSIX
poll(2), select(2) and read(2).  As a consequence of supporting the Linux
f_op->poll subsystem, they can be used with epoll(2) too.

The system call is defined as:

int timerfd(int ufd, int clockid, int flags, const struct itimerspec *utmr);

The "ufd" parameter allows for re-use (re-programming) of an existing timerfd
w/out going through the close/open cycle (same as signalfd).  If "ufd" is -1,
s new file descriptor will be created, otherwise the existing "ufd" will be
re-programmed.

The "clockid" parameter is either CLOCK_MONOTONIC or CLOCK_REALTIME.  The time
specified in the "utmr->it_value" parameter is the expiry time for the timer.

If the TFD_TIMER_ABSTIME flag is set in "flags", this is an absolute time,
otherwise it's a relative time.

If the time specified in the "utmr->it_interval" is not zero (.tv_sec == 0,
tv_nsec == 0), this is the period at which the following ticks should be
generated.

The "utmr->it_interval" should be set to zero if only one tick is requested.
Setting the "utmr->it_value" to zero will disable the timer, or will create a
timerfd without the timer enabled.

The function returns the new (or same, in case "ufd" is a valid timerfd
descriptor) file, or -1 in case of error.

As stated before, the timerfd file descriptor supports poll(2), select(2) and
epoll(2).  When a timer event happened on the timerfd, a POLLIN mask will be
returned.

The read(2) call can be used, and it will return a u32 variable holding the
number of "ticks" that happened on the interface since the last call to
read(2).  The read(2) call supportes the O_NONBLOCK flag too, and EAGAIN will
be returned if no ticks happened.

A quick test program, shows timerfd working correctly on my amd64 box:

http://www.xmailserver.org/timerfd-test.c

[akpm@linux-foundation.org: add sys_timerfd to sys_ni.c]
Signed-off-by: NDavide Libenzi <davidel@xmailserver.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      b215e283
    • D
      signal/timer/event: signalfd core · fba2afaa
      由 Davide Libenzi 提交于 
      This patch series implements the new signalfd() system call.

I took part of the original Linus code (and you know how badly it can be
broken :), and I added even more breakage ;) Signals are fetched from the same
signal queue used by the process, so signalfd will compete with standard
kernel delivery in dequeue_signal().  If you want to reliably fetch signals on
the signalfd file, you need to block them with sigprocmask(SIG_BLOCK).  This
seems to be working fine on my Dual Opteron machine.  I made a quick test
program for it:

http://www.xmailserver.org/signafd-test.c

The signalfd() system call implements signal delivery into a file descriptor
receiver.  The signalfd file descriptor if created with the following API:

int signalfd(int ufd, const sigset_t *mask, size_t masksize);

The "ufd" parameter allows to change an existing signalfd sigmask, w/out going
to close/create cycle (Linus idea).  Use "ufd" == -1 if you want a brand new
signalfd file.

The "mask" allows to specify the signal mask of signals that we are interested
in.  The "masksize" parameter is the size of "mask".

The signalfd fd supports the poll(2) and read(2) system calls.  The poll(2)
will return POLLIN when signals are available to be dequeued.  As a direct
consequence of supporting the Linux poll subsystem, the signalfd fd can use
used together with epoll(2) too.

The read(2) system call will return a "struct signalfd_siginfo" structure in
the userspace supplied buffer.  The return value is the number of bytes copied
in the supplied buffer, or -1 in case of error.  The read(2) call can also
return 0, in case the sighand structure to which the signalfd was attached,
has been orphaned.  The O_NONBLOCK flag is also supported, and read(2) will
return -EAGAIN in case no signal is available.

If the size of the buffer passed to read(2) is lower than sizeof(struct
signalfd_siginfo), -EINVAL is returned.  A read from the signalfd can also
return -ERESTARTSYS in case a signal hits the process.  The format of the
struct signalfd_siginfo is, and the valid fields depends of the (->code &
__SI_MASK) value, in the same way a struct siginfo would:

struct signalfd_siginfo {
	__u32 signo;	/* si_signo */
	__s32 err;	/* si_errno */
	__s32 code;	/* si_code */
	__u32 pid;	/* si_pid */
	__u32 uid;	/* si_uid */
	__s32 fd;	/* si_fd */
	__u32 tid;	/* si_fd */
	__u32 band;	/* si_band */
	__u32 overrun;	/* si_overrun */
	__u32 trapno;	/* si_trapno */
	__s32 status;	/* si_status */
	__s32 svint;	/* si_int */
	__u64 svptr;	/* si_ptr */
	__u64 utime;	/* si_utime */
	__u64 stime;	/* si_stime */
	__u64 addr;	/* si_addr */
};

[akpm@linux-foundation.org: fix signalfd_copyinfo() on i386]
Signed-off-by: NDavide Libenzi <davidel@xmailserver.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      fba2afaa
    • D
      signal/timer/event fds: anonymous inode source · 5dc8bf81
      由 Davide Libenzi 提交于 
      This patch add an anonymous inode source, to be used for files that need
and inode only in order to create a file*. We do not care of having an
inode for each file, and we do not even care of having different names in
the associated dentries (dentry names will be same for classes of file*).
This allow code reuse, and will be used by epoll, signalfd and timerfd
(and whatever else there'll be).
Signed-off-by: NDavide Libenzi <davidel@xmailserver.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      5dc8bf81
    • C
      SLUB: SLUB_DEBUG must depend on SLUB · d4751a27
      由 Christoph Lameter 提交于 
      Otherwise people get asked about SLUB_DEBUG even if they have another
slab allocator enabled.
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      d4751a27
  25. 10 5月, 2007 2 次提交
  27. 09 5月, 2007 3 次提交
  29. 08 5月, 2007 2 次提交
    • B
      blackfin architecture · 1394f032
      由 Bryan Wu 提交于 
      This adds support for the Analog Devices Blackfin processor architecture, and
currently supports the BF533, BF532, BF531, BF537, BF536, BF534, and BF561
(Dual Core) devices, with a variety of development platforms including those
avaliable from Analog Devices (BF533-EZKit, BF533-STAMP, BF537-STAMP,
BF561-EZKIT), and Bluetechnix!  Tinyboards.

The Blackfin architecture was jointly developed by Intel and Analog Devices
Inc.  (ADI) as the Micro Signal Architecture (MSA) core and introduced it in
December of 2000.  Since then ADI has put this core into its Blackfin
processor family of devices.  The Blackfin core has the advantages of a clean,
orthogonal,RISC-like microprocessor instruction set.  It combines a dual-MAC
(Multiply/Accumulate), state-of-the-art signal processing engine and
single-instruction, multiple-data (SIMD) multimedia capabilities into a single
instruction-set architecture.

The Blackfin architecture, including the instruction set, is described by the
ADSP-BF53x/BF56x Blackfin Processor Programming Reference
http://blackfin.uclinux.org/gf/download/frsrelease/29/2549/Blackfin_PRM.pdf

The Blackfin processor is already supported by major releases of gcc, and
there are binary and source rpms/tarballs for many architectures at:
http://blackfin.uclinux.org/gf/project/toolchain/frs There is complete
documentation, including "getting started" guides available at:
http://docs.blackfin.uclinux.org/ which provides links to the sources and
patches you will need in order to set up a cross-compiling environment for
bfin-linux-uclibc

This patch, as well as the other patches (toolchain, distribution,
uClibc) are actively supported by Analog Devices Inc, at:
http://blackfin.uclinux.org/

We have tested this on LTP, and our test plan (including pass/fails) can
be found at:
http://docs.blackfin.uclinux.org/doku.php?id=testing_the_linux_kernel

[m.kozlowski@tuxland.pl: balance parenthesis in blackfin header files]
Signed-off-by: NBryan Wu <bryan.wu@analog.com>
Signed-off-by: NMariusz Kozlowski <m.kozlowski@tuxland.pl>
Signed-off-by: NAubrey Li <aubrey.li@analog.com>
Signed-off-by: NJie Zhang <jie.zhang@analog.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      1394f032
    • C
      SLUB core · 81819f0f
      由 Christoph Lameter 提交于 
      This is a new slab allocator which was motivated by the complexity of the
existing code in mm/slab.c. It attempts to address a variety of concerns
with the existing implementation.

A. Management of object queues

   A particular concern was the complex management of the numerous object
   queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for
   each allocating CPU and use objects from a slab directly instead of
   queueing them up.

B. Storage overhead of object queues

   SLAB Object queues exist per node, per CPU. The alien cache queue even
   has a queue array that contain a queue for each processor on each
   node. For very large systems the number of queues and the number of
   objects that may be caught in those queues grows exponentially. On our
   systems with 1k nodes / processors we have several gigabytes just tied up
   for storing references to objects for those queues  This does not include
   the objects that could be on those queues. One fears that the whole
   memory of the machine could one day be consumed by those queues.

C. SLAB meta data overhead

   SLAB has overhead at the beginning of each slab. This means that data
   cannot be naturally aligned at the beginning of a slab block. SLUB keeps
   all meta data in the corresponding page_struct. Objects can be naturally
   aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte
   boundaries and can fit tightly into a 4k page with no bytes left over.
   SLAB cannot do this.

D. SLAB has a complex cache reaper

   SLUB does not need a cache reaper for UP systems. On SMP systems
   the per CPU slab may be pushed back into partial list but that
   operation is simple and does not require an iteration over a list
   of objects. SLAB expires per CPU, shared and alien object queues
   during cache reaping which may cause strange hold offs.

E. SLAB has complex NUMA policy layer support

   SLUB pushes NUMA policy handling into the page allocator. This means that
   allocation is coarser (SLUB does interleave on a page level) but that
   situation was also present before 2.6.13. SLABs application of
   policies to individual slab objects allocated in SLAB is
   certainly a performance concern due to the frequent references to
   memory policies which may lead a sequence of objects to come from
   one node after another. SLUB will get a slab full of objects
   from one node and then will switch to the next.

F. Reduction of the size of partial slab lists

   SLAB has per node partial lists. This means that over time a large
   number of partial slabs may accumulate on those lists. These can
   only be reused if allocator occur on specific nodes. SLUB has a global
   pool of partial slabs and will consume slabs from that pool to
   decrease fragmentation.

G. Tunables

   SLAB has sophisticated tuning abilities for each slab cache. One can
   manipulate the queue sizes in detail. However, filling the queues still
   requires the uses of the spin lock to check out slabs. SLUB has a global
   parameter (min_slab_order) for tuning. Increasing the minimum slab
   order can decrease the locking overhead. The bigger the slab order the
   less motions of pages between per CPU and partial lists occur and the
   better SLUB will be scaling.

G. Slab merging

   We often have slab caches with similar parameters. SLUB detects those
   on boot up and merges them into the corresponding general caches. This
   leads to more effective memory use. About 50% of all caches can
   be eliminated through slab merging. This will also decrease
   slab fragmentation because partial allocated slabs can be filled
   up again. Slab merging can be switched off by specifying
   slub_nomerge on boot up.

   Note that merging can expose heretofore unknown bugs in the kernel
   because corrupted objects may now be placed differently and corrupt
   differing neighboring objects. Enable sanity checks to find those.

H. Diagnostics

   The current slab diagnostics are difficult to use and require a
   recompilation of the kernel. SLUB contains debugging code that
   is always available (but is kept out of the hot code paths).
   SLUB diagnostics can be enabled via the "slab_debug" option.
   Parameters can be specified to select a single or a group of
   slab caches for diagnostics. This means that the system is running
   with the usual performance and it is much more likely that
   race conditions can be reproduced.

I. Resiliency

   If basic sanity checks are on then SLUB is capable of detecting
   common error conditions and recover as best as possible to allow the
   system to continue.

J. Tracing

   Tracing can be enabled via the slab_debug=T,<slabcache> option
   during boot. SLUB will then protocol all actions on that slabcache
   and dump the object contents on free.

K. On demand DMA cache creation.

   Generally DMA caches are not needed. If a kmalloc is used with
   __GFP_DMA then just create this single slabcache that is needed.
   For systems that have no ZONE_DMA requirement the support is
   completely eliminated.

L. Performance increase

   Some benchmarks have shown speed improvements on kernbench in the
   range of 5-10%. The locking overhead of slub is based on the
   underlying base allocation size. If we can reliably allocate
   larger order pages then it is possible to increase slub
   performance much further. The anti-fragmentation patches may
   enable further performance increases.

Tested on:
i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator

SLUB Boot options

slub_nomerge		Disable merging of slabs
slub_min_order=x	Require a minimum order for slab caches. This
			increases the managed chunk size and therefore
			reduces meta data and locking overhead.
slub_min_objects=x	Mininum objects per slab. Default is 8.
slub_max_order=x	Avoid generating slabs larger than order specified.
slub_debug		Enable all diagnostics for all caches
slub_debug=<options>	Enable selective options for all caches
slub_debug=<o>,<cache>	Enable selective options for a certain set of
			caches

Available Debug options
F		Double Free checking, sanity and resiliency
R		Red zoning
P		Object / padding poisoning
U		Track last free / alloc
T		Trace all allocs / frees (only use for individual slabs).

To use SLUB: Apply this patch and then select SLUB as the default slab
allocator.

[hugh@veritas.com: fix an oops-causing locking error]
[akpm@linux-foundation.org: various stupid cleanups and small fixes]
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NHugh Dickins <hugh@veritas.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      81819f0f