- 28 Jun, 2006 40 commits
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Linus Torvalds authored
* master.kernel.org:/pub/scm/linux/kernel/git/mchehab/v4l-dvb: (26 commits) V4L/DVB (4263): Fix warning when compiling on 64 bit machines V4L/DVB (4261): Included required header for in-kernel compilation V4L/DVB (4260): Stradis.c: make 2 functions static V4L/DVB (4259): Pass an explicit log prefix to cx2341x_log_status V4L/DVB (4257): Fix 64-bit compile warnings. V4L/DVB (4255): Tda9887 default TOP value is 0x10 V4L/DVB (4254): Remove obsoleted tuner_debug option. V4L/DVB (4253): IVTV VBI format description too long. V4L/DVB (4252): Remove duplicate 'tda9887' in info messages. V4L/DVB (4245): Reduce the amount of pvrusb2-sourced noise going into the system log V4L/DVB (4244): Implement use of cx2341x module in pvrusb2 driver V4L/DVB (4243): Exploit new V4L control features in pvrusb2 V4L/DVB (4242): Don't suspend encoder when changing its attributes (in pvrusb2) V4L/DVB (4241): Fix faulty encoder error recovery in pvrusb2 V4L/DVB (4240): Various V4L control enhancements in pvrusb2 V4L/DVB (4239): Handle boolean controls in pvrusb2 V4L/DVB (4238): Make sure flags field is initialized when quering a control in pvrusb2 V4L/DVB (4237): Move LOG_STATUS bracketing to a different part of the pvrusb2 driver V4L/DVB (4236): Rearrange things in pvrusb2 driver in preparation for using cx2341x module V4L/DVB (4235): Increase the maximum number of controls that pvrusb2-sysfs.c can handle. ...
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Karsten Keil authored
DLE masking was non-functional since the new tty handling. Found by Peter Evertz <leo2@pec.homeip.net> Signed-off-by:
Karsten Keil <kkeil@suse.de> Signed-off-by:
Linus Torvalds <torvalds@osdl.org>
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Andrew Morton authored
arch/i386/kernel/irq.c: In function 'do_IRQ': arch/i386/kernel/irq.c:104: warning: suggest parentheses around arithmetic in operand of | Signed-off-by:
Andrew Morton <akpm@osdl.org> Signed-off-by:
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Adrian Bunk authored
It's available in a header as a static inline - there's no need to export it. Signed-off-by:
Adrian Bunk <bunk@stusta.de> Cc: Markus Lidel <Markus.Lidel@shadowconnect.com> Signed-off-by:
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Adrian Bunk authored
Make struct proc_ipmi_root static. Besides this, tremove removes an unused #ifdef CONFIG_PROC_FS from include/linux/ipmi.h. Acked-by:
Corey Minyard <minyard@acm.org> Signed-off-by:
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Jesper Juhl authored
Remove redundant NULL chck before kfree + tiny CodingStyle cleanup for drivers/ Signed-off-by:
Jesper Juhl <jesper.juhl@gmail.com> Signed-off-by:
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Jesper Juhl authored
Remove redundant kfree NULL checks from kernel/ Signed-off-by:
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Jesper Juhl authored
Remove redundant NULL checks before kfree for fs/ Signed-off-by:
Mark Fasheh <mark.fasheh@oracle.com> Signed-off-by:
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Sebastien Dugue authored
In futex_requeue(), when the 2 futexes keys hash to the same bucket, there is no need to move the futex_q to the end of the bucket list. Signed-off-by:
Sebastien Dugue <sebastien.dugue@bull.net> Cc: Ingo Molnar <mingo@elte.hu> Cc: Thomas Gleixner <tglx@linutronix.de> Signed-off-by:
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Thomas Gleixner authored
When the priority of a task, which is blocked on a lock, changes we must propagate this change into the PI lock chain. Therefor the chain walk code is changed to get rid of the references to current to avoid false positives in the deadlock detector, as setscheduler might be called by a task which holds the lock on which the task whose priority is changed is blocked. Also add some comments about the get/put_task_struct usage to avoid confusion. Signed-off-by:
Thomas Gleixner <tglx@linutronix.de> Signed-off-by:
Ingo Molnar <mingo@elte.hu> Cc: Steven Rostedt <rostedt@goodmis.org> Signed-off-by:
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Thomas Gleixner authored
Make test suite setscheduler calls asynchronously. Remove the waits in the test cases and add a new testcase to verify the correctness of the setscheduler priority propagation. Signed-off-by:
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Thomas Gleixner authored
There is no need to hold tasklist_lock across the setscheduler call, when we pin the task structure with get_task_struct(). Interrupts are disabled in setscheduler anyway and the permission checks do not need interrupts disabled. Signed-off-by:
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Roman Zippel authored
Fix defaults and dependencies. Signed-off-by:
Roman Zippel <zippel@linux-m68k.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@elte.hu> Signed-off-by:
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Ingo Molnar authored
This adds the actual pi-futex implementation, based on rt-mutexes. [dino@in.ibm.com: fix an oops-causing race] Signed-off-by:
Arjan van de Ven <arjan@linux.intel.com> Signed-off-by:
Dinakar Guniguntala <dino@in.ibm.com> Signed-off-by:
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Ingo Molnar authored
Add proxy-locking rt-mutex functionality needed by pi-futexes. Signed-off-by:
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Thomas Gleixner authored
RT-mutex tester: scriptable tester for rt mutexes, which allows userspace scripting of mutex unit-tests (and dynamic tests as well), using the actual rt-mutex implementation of the kernel. [akpm@osdl.org: fixlet] Signed-off-by:
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Ingo Molnar authored
Runtime debugging functionality for rt-mutexes. Signed-off-by:
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Steven Rostedt authored
Add rt-mutex documentation. [rostedt@goodmis.org: Update rt-mutex-design.txt as per Randy Dunlap suggestions] Signed-off-by:
Steven Rostedt <rostedt@goodmis.org> Signed-off-by:
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Ingo Molnar authored
Core functions for the rt-mutex subsystem. Signed-off-by:
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Ingo Molnar authored
Add framework to boost/unboost the priority of RT tasks. This consists of: - caching the 'normal' priority in ->normal_prio - providing a functions to set/get the priority of the task - make sched_setscheduler() aware of boosting The effective_prio() cleanups also fix a priority-calculation bug pointed out by Andrey Gelman, in set_user_nice(). has_rt_policy() fix: Peter Williams <pwil3058@bigpond.net.au> Signed-off-by:
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Ingo Molnar authored
Add the priority-sorted list (plist) implementation. Signed-off-by:
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Ingo Molnar authored
Introduce a new WARN_ON variant: WARN_ON_SMP(cond). Signed-off-by:
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Ingo Molnar authored
Add debug_check_no_locks_freed(), as a central inline to add bad-lock-free-debugging functionality to. Signed-off-by:
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Ingo Molnar authored
Fix typo in Documentation/robust-futexes.txt. Signed-off-by:
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Ingo Molnar authored
We are pleased to announce "lightweight userspace priority inheritance" (PI) support for futexes. The following patchset and glibc patch implements it, ontop of the robust-futexes patchset which is included in 2.6.16-mm1. We are calling it lightweight for 3 reasons: - in the user-space fastpath a PI-enabled futex involves no kernel work (or any other PI complexity) at all. No registration, no extra kernel calls - just pure fast atomic ops in userspace. - in the slowpath (in the lock-contention case), the system call and scheduling pattern is in fact better than that of normal futexes, due to the 'integrated' nature of FUTEX_LOCK_PI. [more about that further down] - the in-kernel PI implementation is streamlined around the mutex abstraction, with strict rules that keep the implementation relatively simple: only a single owner may own a lock (i.e. no read-write lock support), only the owner may unlock a lock, no recursive locking, etc. Priority Inheritance - why, oh why??? ------------------------------------- Many of you heard the horror stories about the evil PI code circling Linux for years, which makes no real sense at all and is only used by buggy applications and which has horrible overhead. Some of you have dreaded this very moment, when someone actually submits working PI code ;-) So why would we like to see PI support for futexes? We'd like to see it done purely for technological reasons. We dont think it's a buggy concept, we think it's useful functionality to offer to applications, which functionality cannot be achieved in other ways. We also think it's the right thing to do, and we think we've got the right arguments and the right numbers to prove that. We also believe that we can address all the counter-arguments as well. For these reasons (and the reasons outlined below) we are submitting this patch-set for upstream kernel inclusion. What are the benefits of PI? The short reply: ---------------- User-space PI helps achieving/improving determinism for user-space applications. In the best-case, it can help achieve determinism and well-bound latencies. Even in the worst-case, PI will improve the statistical distribution of locking related application delays. The longer reply: ----------------- Firstly, sharing locks between multiple tasks is a common programming technique that often cannot be replaced with lockless algorithms. As we can see it in the kernel [which is a quite complex program in itself], lockless structures are rather the exception than the norm - the current ratio of lockless vs. locky code for shared data structures is somewhere between 1:10 and 1:100. Lockless is hard, and the complexity of lockless algorithms often endangers to ability to do robust reviews of said code. I.e. critical RT apps often choose lock structures to protect critical data structures, instead of lockless algorithms. Furthermore, there are cases (like shared hardware, or other resource limits) where lockless access is mathematically impossible. Media players (such as Jack) are an example of reasonable application design with multiple tasks (with multiple priority levels) sharing short-held locks: for example, a highprio audio playback thread is combined with medium-prio construct-audio-data threads and low-prio display-colory-stuff threads. Add video and decoding to the mix and we've got even more priority levels. So once we accept that synchronization objects (locks) are an unavoidable fact of life, and once we accept that multi-task userspace apps have a very fair expectation of being able to use locks, we've got to think about how to offer the option of a deterministic locking implementation to user-space. Most of the technical counter-arguments against doing priority inheritance only apply to kernel-space locks. But user-space locks are different, there we cannot disable interrupts or make the task non-preemptible in a critical section, so the 'use spinlocks' argument does not apply (user-space spinlocks have the same priority inversion problems as other user-space locking constructs). Fact is, pretty much the only technique that currently enables good determinism for userspace locks (such as futex-based pthread mutexes) is priority inheritance: Currently (without PI), if a high-prio and a low-prio task shares a lock [this is a quite common scenario for most non-trivial RT applications], even if all critical sections are coded carefully to be deterministic (i.e. all critical sections are short in duration and only execute a limited number of instructions), the kernel cannot guarantee any deterministic execution of the high-prio task: any medium-priority task could preempt the low-prio task while it holds the shared lock and executes the critical section, and could delay it indefinitely. Implementation: --------------- As mentioned before, the userspace fastpath of PI-enabled pthread mutexes involves no kernel work at all - they behave quite similarly to normal futex-based locks: a 0 value means unlocked, and a value==TID means locked. (This is the same method as used by list-based robust futexes.) Userspace uses atomic ops to lock/unlock these mutexes without entering the kernel. To handle the slowpath, we have added two new futex ops: FUTEX_LOCK_PI FUTEX_UNLOCK_PI If the lock-acquire fastpath fails, [i.e. an atomic transition from 0 to TID fails], then FUTEX_LOCK_PI is called. The kernel does all the remaining work: if there is no futex-queue attached to the futex address yet then the code looks up the task that owns the futex [it has put its own TID into the futex value], and attaches a 'PI state' structure to the futex-queue. The pi_state includes an rt-mutex, which is a PI-aware, kernel-based synchronization object. The 'other' task is made the owner of the rt-mutex, and the FUTEX_WAITERS bit is atomically set in the futex value. Then this task tries to lock the rt-mutex, on which it blocks. Once it returns, it has the mutex acquired, and it sets the futex value to its own TID and returns. Userspace has no other work to perform - it now owns the lock, and futex value contains FUTEX_WAITERS|TID. If the unlock side fastpath succeeds, [i.e. userspace manages to do a TID -> 0 atomic transition of the futex value], then no kernel work is triggered. If the unlock fastpath fails (because the FUTEX_WAITERS bit is set), then FUTEX_UNLOCK_PI is called, and the kernel unlocks the futex on the behalf of userspace - and it also unlocks the attached pi_state->rt_mutex and thus wakes up any potential waiters. Note that under this approach, contrary to other PI-futex approaches, there is no prior 'registration' of a PI-futex. [which is not quite possible anyway, due to existing ABI properties of pthread mutexes.] Also, under this scheme, 'robustness' and 'PI' are two orthogonal properties of futexes, and all four combinations are possible: futex, robust-futex, PI-futex, robust+PI-futex. glibc support: -------------- Ulrich Drepper and Jakub Jelinek have written glibc support for PI-futexes (and robust futexes), enabling robust and PI (PTHREAD_PRIO_INHERIT) POSIX mutexes. (PTHREAD_PRIO_PROTECT support will be added later on too, no additional kernel changes are needed for that). [NOTE: The glibc patch is obviously inofficial and unsupported without matching upstream kernel functionality.] the patch-queue and the glibc patch can also be downloaded from: http://redhat.com/~mingo/PI-futex-patches/ Many thanks go to the people who helped us create this kernel feature: Steven Rostedt, Esben Nielsen, Benedikt Spranger, Daniel Walker, John Cooper, Arjan van de Ven, Oleg Nesterov and others. Credits for related prior projects goes to Dirk Grambow, Inaky Perez-Gonzalez, Bill Huey and many others. Clean up the futex code, before adding more features to it: - use u32 as the futex field type - that's the ABI - use __user and pointers to u32 instead of unsigned long - code style / comment style cleanups - rename hash-bucket name from 'bh' to 'hb'. I checked the pre and post futex.o object files to make sure this patch has no code effects. Signed-off-by:
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Steven Rostedt authored
Thomas Gleixner is adding the call to a rtmutex function in setscheduler. This call grabs a spin_lock that is not always protected by interrupts disabled. So this means that setscheduler cant be called from interrupt context. To prevent this from happening in the future, this patch adds a BUG_ON(in_interrupt()) in that function. (Thanks to akpm <aka. Andrew Morton> for this suggestion). Signed-off-by:
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Oleg Nesterov authored
Saves 543 bytes from sched.o (gcc 3.3.3). Signed-off-by:
Oleg Nesterov <oleg@tv-sign.ru> Cc: Ingo Molnar <mingo@elte.hu> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Cc: Con Kolivas <kernel@kolivas.org> Cc: Peter Williams <pwil3058@bigpond.net.au> Signed-off-by:
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Siddha, Suresh B authored
sysfs entries 'sched_mc_power_savings' and 'sched_smt_power_savings' in /sys/devices/system/cpu/ control the MC/SMT power savings policy for the scheduler. Based on the values (1-enable, 0-disable) for these controls, sched groups cpu power will be determined for different domains. When power savings policy is enabled and under light load conditions, scheduler will minimize the physical packages/cpu cores carrying the load and thus conserving power(with a perf impact based on the workload characteristics... see OLS 2005 CMP kernel scheduler paper for more details..) Signed-off-by:
Suresh Siddha <suresh.b.siddha@intel.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Cc: Con Kolivas <kernel@kolivas.org> Cc: "Chen, Kenneth W" <kenneth.w.chen@intel.com> Cc: "David S. Miller" <davem@davemloft.net> Signed-off-by:
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Srivatsa Vaddagiri authored
As explained here: http://marc.theaimsgroup.com/?l=linux-kernel&m=114327539012323&w=2 there is a problem with sharing sched_group structures between two separate sched_group structures for different sched_domains. The patch has been tested and found to avoid the kernel lockup problem described in above URL. Signed-off-by:
Srivatsa Vaddagiri <vatsa@in.ibm.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Cc: Ingo Molnar <mingo@elte.hu> Cc: "Siddha, Suresh B" <suresh.b.siddha@intel.com> Signed-off-by:
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Srivatsa Vaddagiri authored
The sched group structures used to represent various nodes need to be allocated from respective nodes (as suggested here also: http://uwsg.ucs.indiana.edu/hypermail/linux/kernel/0603.3/0051.html) Signed-off-by:
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Srivatsa Vaddagiri authored
Replace GFP_ATOMIC allocation for sched_group_nodes with GFP_KERNEL based allocation. Signed-off-by: Srivatsa Vaddagiri <vatsa@in.ibm.com Cc: Nick Piggin <nickpiggin@yahoo.com.au> Cc: Ingo Molnar <mingo@elte.hu> Cc: "Siddha, Suresh B" <suresh.b.siddha@intel.com> Signed-off-by:
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Srivatsa Vaddagiri authored
Try to handle mem allocation failures in build_sched_domains by bailing out and cleaning up thus-far allocated memory. The patch has a direct consequence that we disable load balancing completely (even at sibling level) upon *any* memory allocation failure. [Lee.Schermerhorn@hp.com: bugfix] Signed-off-by:
Lee Schermerhorn <lee.schermerhorn@hp.com> Signed-off-by:
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Peter Williams authored
Problem: To help distribute high priority tasks evenly across the available CPUs move_tasks() does not, under some circumstances, skip tasks whose load weight is bigger than the designated amount. Because the highest priority task on the busiest queue may be on the expired array it may be moved as a result of this mechanism. Apart from not being the most desirable way to redistribute the high priority tasks (we'd rather move the second highest priority task), there is a risk that this could set up a loop with this task bouncing backwards and forwards between the two queues. (This latter possibility can be demonstrated by running a nice==-20 CPU bound task on an otherwise quiet 2 CPU system.) Solution: Modify the mechanism so that it does not override skip for the highest priority task on the CPU. Of course, if there are more than one tasks at the highest priority then it will allow the override for one of them as this is a desirable redistribution of high priority tasks. Signed-off-by:
Peter Williams <pwil3058@bigpond.com.au> Cc: Ingo Molnar <mingo@elte.hu> Cc: "Siddha, Suresh B" <suresh.b.siddha@intel.com> Signed-off-by:
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Peter Williams authored
Problem: The move_tasks() function is designed to move UP TO the amount of load it is asked to move and in doing this it skips over tasks looking for ones whose load weights are less than or equal to the remaining load to be moved. This is (in general) a good thing but it has the unfortunate result of breaking one of the original load balancer's good points: namely, that (within the limits imposed by the active/expired array model and the fact the expired is processed first) it moves high priority tasks before low priority ones and this means there's a good chance (see active/expired problem for why it's only a chance) that the highest priority task on the queue but not actually on the CPU will be moved to the other CPU where (as a high priority task) it may preempt the current task. Solution: Modify move_tasks() so that high priority tasks are not skipped when moving them will make them the highest priority task on their new run queue. Signed-off-by:
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Peter Williams authored
Problem: The introduction of separate run queues per CPU has brought with it "nice" enforcement problems that are best described by a simple example. For the sake of argument suppose that on a single CPU machine with a nice==19 hard spinner and a nice==0 hard spinner running that the nice==0 task gets 95% of the CPU and the nice==19 task gets 5% of the CPU. Now suppose that there is a system with 2 CPUs and 2 nice==19 hard spinners and 2 nice==0 hard spinners running. The user of this system would be entitled to expect that the nice==0 tasks each get 95% of a CPU and the nice==19 tasks only get 5% each. However, whether this expectation is met is pretty much down to luck as there are four equally likely distributions of the tasks to the CPUs that the load balancing code will consider to be balanced with loads of 2.0 for each CPU. Two of these distributions involve one nice==0 and one nice==19 task per CPU and in these circumstances the users expectations will be met. The other two distributions both involve both nice==0 tasks being on one CPU and both nice==19 being on the other CPU and each task will get 50% of a CPU and the user's expectations will not be met. Solution: The solution to this problem that is implemented in the attached patch is to use weighted loads when determining if the system is balanced and, when an imbalance is detected, to move an amount of weighted load between run queues (as opposed to a number of tasks) to restore the balance. Once again, the easiest way to explain why both of these measures are necessary is to use a simple example. Suppose that (in a slight variation of the above example) that we have a two CPU system with 4 nice==0 and 4 nice=19 hard spinning tasks running and that the 4 nice==0 tasks are on one CPU and the 4 nice==19 tasks are on the other CPU. The weighted loads for the two CPUs would be 4.0 and 0.2 respectively and the load balancing code would move 2 tasks resulting in one CPU with a load of 2.0 and the other with load of 2.2. If this was considered to be a big enough imbalance to justify moving a task and that task was moved using the current move_tasks() then it would move the highest priority task that it found and this would result in one CPU with a load of 3.0 and the other with a load of 1.2 which would result in the movement of a task in the opposite direction and so on -- infinite loop. If, on the other hand, an amount of load to be moved is calculated from the imbalance (in this case 0.1) and move_tasks() skips tasks until it find ones whose contributions to the weighted load are less than this amount it would move two of the nice==19 tasks resulting in a system with 2 nice==0 and 2 nice=19 on each CPU with loads of 2.1 for each CPU. One of the advantages of this mechanism is that on a system where all tasks have nice==0 the load balancing calculations would be mathematically identical to the current load balancing code. Notes: struct task_struct: has a new field load_weight which (in a trade off of space for speed) stores the contribution that this task makes to a CPU's weighted load when it is runnable. struct runqueue: has a new field raw_weighted_load which is the sum of the load_weight values for the currently runnable tasks on this run queue. This field always needs to be updated when nr_running is updated so two new inline functions inc_nr_running() and dec_nr_running() have been created to make sure that this happens. This also offers a convenient way to optimize away this part of the smpnice mechanism when CONFIG_SMP is not defined. int try_to_wake_up(): in this function the value SCHED_LOAD_BALANCE is used to represent the load contribution of a single task in various calculations in the code that decides which CPU to put the waking task on. While this would be a valid on a system where the nice values for the runnable tasks were distributed evenly around zero it will lead to anomalous load balancing if the distribution is skewed in either direction. To overcome this problem SCHED_LOAD_SCALE has been replaced by the load_weight for the relevant task or by the average load_weight per task for the queue in question (as appropriate). int move_tasks(): The modifications to this function were complicated by the fact that active_load_balance() uses it to move exactly one task without checking whether an imbalance actually exists. This precluded the simple overloading of max_nr_move with max_load_move and necessitated the addition of the latter as an extra argument to the function. The internal implementation is then modified to move up to max_nr_move tasks and max_load_move of weighted load. This slightly complicates the code where move_tasks() is called and if ever active_load_balance() is changed to not use move_tasks() the implementation of move_tasks() should be simplified accordingly. struct sched_group *find_busiest_group(): Similar to try_to_wake_up(), there are places in this function where SCHED_LOAD_SCALE is used to represent the load contribution of a single task and the same issues are created. A similar solution is adopted except that it is now the average per task contribution to a group's load (as opposed to a run queue) that is required. As this value is not directly available from the group it is calculated on the fly as the queues in the groups are visited when determining the busiest group. A key change to this function is that it is no longer to scale down *imbalance on exit as move_tasks() uses the load in its scaled form. void set_user_nice(): has been modified to update the task's load_weight field when it's nice value and also to ensure that its run queue's raw_weighted_load field is updated if it was runnable. From: "Siddha, Suresh B" <suresh.b.siddha@intel.com> With smpnice, sched groups with highest priority tasks can mask the imbalance between the other sched groups with in the same domain. This patch fixes some of the listed down scenarios by not considering the sched groups which are lightly loaded. a) on a simple 4-way MP system, if we have one high priority and 4 normal priority tasks, with smpnice we would like to see the high priority task scheduled on one cpu, two other cpus getting one normal task each and the fourth cpu getting the remaining two normal tasks. but with current smpnice extra normal priority task keeps jumping from one cpu to another cpu having the normal priority task. This is because of the busiest_has_loaded_cpus, nr_loaded_cpus logic.. We are not including the cpu with high priority task in max_load calculations but including that in total and avg_load calcuations.. leading to max_load < avg_load and load balance between cpus running normal priority tasks(2 Vs 1) will always show imbalanace as one normal priority and the extra normal priority task will keep moving from one cpu to another cpu having normal priority task.. b) 4-way system with HT (8 logical processors). Package-P0 T0 has a highest priority task, T1 is idle. Package-P1 Both T0 and T1 have 1 normal priority task each.. P2 and P3 are idle. With this patch, one of the normal priority tasks on P1 will be moved to P2 or P3.. c) With the current weighted smp nice calculations, it doesn't always make sense to look at the highest weighted runqueue in the busy group.. Consider a load balance scenario on a DP with HT system, with Package-0 containing one high priority and one low priority, Package-1 containing one low priority(with other thread being idle).. Package-1 thinks that it need to take the low priority thread from Package-0. And find_busiest_queue() returns the cpu thread with highest priority task.. And ultimately(with help of active load balance) we move high priority task to Package-1. And same continues with Package-0 now, moving high priority task from package-1 to package-0.. Even without the presence of active load balance, load balance will fail to balance the above scenario.. Fix find_busiest_queue to use "imbalance" when it is lightly loaded. [kernel@kolivas.org: sched: store weighted load on up] [kernel@kolivas.org: sched: add discrete weighted cpu load function] [suresh.b.siddha@intel.com: sched: remove dead code] Signed-off-by:
Con Kolivas <kernel@kolivas.org> Cc: John Hawkes <hawkes@sgi.com> Signed-off-by:
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Kirill Korotaev authored
There is a race between set_cpus_allowed() and move_task_off_dead_cpu(). __migrate_task() doesn't report any err code, so task can be left on its runqueue if its cpus_allowed mask changed so that dest_cpu is not longer a possible target. Also, chaning cpus_allowed mask requires rq->lock being held. Signed-off-by:
Kirill Korotaev <dev@openvz.org> Acked-By:
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Steven Rostedt authored
Unless we expect to have more than 2G CPUs, there's no reason to have 'i' as a long long here. Signed-off-by:
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Con Kolivas authored
The relationship between INTERACTIVE_SLEEP and the ceiling is not perfect and not explicit enough. The sleep boost is not supposed to be any larger than without this code and the comment is not clear enough about what exactly it does, just the reason it does it. Fix it. There is a ceiling to the priority beyond which tasks that only ever sleep for very long periods cannot surpass. Fix it. Prevent the on-runqueue bonus logic from defeating the idle sleep logic. Opportunity to micro-optimise. Signed-off-by:
Mike Galbraith <efault@gmx.de> Acked-by:
Ken Chen <kenneth.w.chen@intel.com> Signed-off-by:
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Steven Rostedt authored
Signed-off-by:
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Chen, Kenneth W authored
Initial report and lock contention fix from Chris Mason: Recent benchmarks showed some performance regressions between 2.6.16 and 2.6.5. We tracked down one of the regressions to lock contention in schedule heavy workloads (~70,000 context switches per second) kernel/sched.c:dependent_sleeper() was responsible for most of the lock contention, hammering on the run queue locks. The patch below is more of a discussion point than a suggested fix (although it does reduce lock contention significantly). The dependent_sleeper code looks very expensive to me, especially for using a spinlock to bounce control between two different siblings in the same cpu. It is further optimized: * perform dependent_sleeper check after next task is determined * convert wake_sleeping_dependent to use trylock * skip smt runqueue check if trylock fails * optimize double_rq_lock now that smt nice is converted to trylock * early exit in searching first SD_SHARE_CPUPOWER domain * speedup fast path of dependent_sleeper [akpm@osdl.org: cleanup] Signed-off-by:
Nick Piggin <npiggin@suse.de> Acked-by:
Chris Mason <mason@suse.com> Signed-off-by:
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