Adding upstream version 4.19.249.upstream/4.19.249upstream

Signed-off-by: Daniel Baumann <daniel.baumann@progress-linux.org>

11 files changed, 2250 insertions, 0 deletions

diff --git a/Documentation/scheduler/00-INDEX b/Documentation/scheduler/00-INDEX
new file mode 100644
index 000000000..eccf7ad2e
--- /dev/null
+++ b/Documentation/scheduler/00-INDEX
@@ -0,0 +1,18 @@
+00-INDEX
+	- this file.
+sched-arch.txt
+	- CPU Scheduler implementation hints for architecture specific code.
+sched-bwc.txt
+	- CFS bandwidth control overview.
+sched-design-CFS.txt
+	- goals, design and implementation of the Completely Fair Scheduler.
+sched-domains.txt
+	- information on scheduling domains.
+sched-nice-design.txt
+	- How and why the scheduler's nice levels are implemented.
+sched-rt-group.txt
+	- real-time group scheduling.
+sched-deadline.txt
+	- deadline scheduling.
+sched-stats.txt
+	- information on schedstats (Linux Scheduler Statistics).
diff --git a/Documentation/scheduler/completion.txt b/Documentation/scheduler/completion.txt
new file mode 100644
index 000000000..656cf803c
--- /dev/null
+++ b/Documentation/scheduler/completion.txt
@@ -0,0 +1,247 @@
+completions - wait for completion handling
+==========================================
+
+This document was originally written based on 3.18.0 (linux-next)
+
+Introduction:
+-------------
+
+If you have one or more threads of execution that must wait for some process
+to have reached a point or a specific state, completions can provide a
+race-free solution to this problem. Semantically they are somewhat like a
+pthread_barrier and have similar use-cases.
+
+Completions are a code synchronization mechanism which is preferable to any
+misuse of locks. Any time you think of using yield() or some quirky
+msleep(1) loop to allow something else to proceed, you probably want to
+look into using one of the wait_for_completion*() calls instead. The
+advantage of using completions is clear intent of the code, but also more
+efficient code as both threads can continue until the result is actually
+needed.
+
+Completions are built on top of the generic event infrastructure in Linux,
+with the event reduced to a simple flag (appropriately called "done") in
+struct completion that tells the waiting threads of execution if they
+can continue safely.
+
+As completions are scheduling related, the code is found in
+kernel/sched/completion.c.
+
+
+Usage:
+------
+
+There are three parts to using completions, the initialization of the
+struct completion, the waiting part through a call to one of the variants of
+wait_for_completion() and the signaling side through a call to complete()
+or complete_all(). Further there are some helper functions for checking the
+state of completions.
+
+To use completions one needs to include <linux/completion.h> and
+create a variable of type struct completion. The structure used for
+handling of completions is:
+
+	struct completion {
+		unsigned int done;
+		wait_queue_head_t wait;
+	};
+
+providing the wait queue to place tasks on for waiting and the flag for
+indicating the state of affairs.
+
+Completions should be named to convey the intent of the waiter. A good
+example is:
+
+	wait_for_completion(&early_console_added);
+
+	complete(&early_console_added);
+
+Good naming (as always) helps code readability.
+
+
+Initializing completions:
+-------------------------
+
+Initialization of dynamically allocated completions, often embedded in
+other structures, is done with:
+
+	void init_completion(&done);
+
+Initialization is accomplished by initializing the wait queue and setting
+the default state to "not available", that is, "done" is set to 0.
+
+The re-initialization function, reinit_completion(), simply resets the
+done element to "not available", thus again to 0, without touching the
+wait queue. Calling init_completion() twice on the same completion object is
+most likely a bug as it re-initializes the queue to an empty queue and
+enqueued tasks could get "lost" - use reinit_completion() in that case.
+
+For static declaration and initialization, macros are available. These are:
+
+	static DECLARE_COMPLETION(setup_done)
+
+used for static declarations in file scope. Within functions the static
+initialization should always use:
+
+	DECLARE_COMPLETION_ONSTACK(setup_done)
+
+suitable for automatic/local variables on the stack and will make lockdep
+happy. Note also that one needs to make *sure* the completion passed to
+work threads remains in-scope, and no references remain to on-stack data
+when the initiating function returns.
+
+Using on-stack completions for code that calls any of the _timeout or
+_interruptible/_killable variants is not advisable as they will require
+additional synchronization to prevent the on-stack completion object in
+the timeout/signal cases from going out of scope. Consider using dynamically
+allocated completions when intending to use the _interruptible/_killable
+or _timeout variants of wait_for_completion().
+
+
+Waiting for completions:
+------------------------
+
+For a thread of execution to wait for some concurrent work to finish, it
+calls wait_for_completion() on the initialized completion structure.
+A typical usage scenario is:
+
+	struct completion setup_done;
+	init_completion(&setup_done);
+	initialize_work(...,&setup_done,...)
+
+	/* run non-dependent code */              /* do setup */
+
+	wait_for_completion(&setup_done);         complete(setup_done)
+
+This is not implying any temporal order on wait_for_completion() and the
+call to complete() - if the call to complete() happened before the call
+to wait_for_completion() then the waiting side simply will continue
+immediately as all dependencies are satisfied if not it will block until
+completion is signaled by complete().
+
+Note that wait_for_completion() is calling spin_lock_irq()/spin_unlock_irq(),
+so it can only be called safely when you know that interrupts are enabled.
+Calling it from hard-irq or irqs-off atomic contexts will result in
+hard-to-detect spurious enabling of interrupts.
+
+wait_for_completion():
+
+	void wait_for_completion(struct completion *done):
+
+The default behavior is to wait without a timeout and to mark the task as
+uninterruptible. wait_for_completion() and its variants are only safe
+in process context (as they can sleep) but not in atomic context,
+interrupt context, with disabled irqs. or preemption is disabled - see also
+try_wait_for_completion() below for handling completion in atomic/interrupt
+context.
+
+As all variants of wait_for_completion() can (obviously) block for a long
+time, you probably don't want to call this with held mutexes.
+
+
+Variants available:
+-------------------
+
+The below variants all return status and this status should be checked in
+most(/all) cases - in cases where the status is deliberately not checked you
+probably want to make a note explaining this (e.g. see
+arch/arm/kernel/smp.c:__cpu_up()).
+
+A common problem that occurs is to have unclean assignment of return types,
+so care should be taken with assigning return-values to variables of proper
+type. Checking for the specific meaning of return values also has been found
+to be quite inaccurate e.g. constructs like
+if (!wait_for_completion_interruptible_timeout(...)) would execute the same
+code path for successful completion and for the interrupted case - which is
+probably not what you want.
+
+	int wait_for_completion_interruptible(struct completion *done)
+
+This function marks the task TASK_INTERRUPTIBLE. If a signal was received
+while waiting it will return -ERESTARTSYS; 0 otherwise.
+
+	unsigned long wait_for_completion_timeout(struct completion *done,
+		unsigned long timeout)
+
+The task is marked as TASK_UNINTERRUPTIBLE and will wait at most 'timeout'
+(in jiffies). If timeout occurs it returns 0 else the remaining time in
+jiffies (but at least 1). Timeouts are preferably calculated with
+msecs_to_jiffies() or usecs_to_jiffies(). If the returned timeout value is
+deliberately ignored a comment should probably explain why (e.g. see
+drivers/mfd/wm8350-core.c wm8350_read_auxadc())
+
+	long wait_for_completion_interruptible_timeout(
+		struct completion *done, unsigned long timeout)
+
+This function passes a timeout in jiffies and marks the task as
+TASK_INTERRUPTIBLE. If a signal was received it will return -ERESTARTSYS;
+otherwise it returns 0 if the completion timed out or the remaining time in
+jiffies if completion occurred.
+
+Further variants include _killable which uses TASK_KILLABLE as the
+designated tasks state and will return -ERESTARTSYS if it is interrupted or
+else 0 if completion was achieved.  There is a _timeout variant as well:
+
+	long wait_for_completion_killable(struct completion *done)
+	long wait_for_completion_killable_timeout(struct completion *done,
+		unsigned long timeout)
+
+The _io variants wait_for_completion_io() behave the same as the non-_io
+variants, except for accounting waiting time as waiting on IO, which has
+an impact on how the task is accounted in scheduling stats.
+
+	void wait_for_completion_io(struct completion *done)
+	unsigned long wait_for_completion_io_timeout(struct completion *done
+		unsigned long timeout)
+
+
+Signaling completions:
+----------------------
+
+A thread that wants to signal that the conditions for continuation have been
+achieved calls complete() to signal exactly one of the waiters that it can
+continue.
+
+	void complete(struct completion *done)
+
+or calls complete_all() to signal all current and future waiters.
+
+	void complete_all(struct completion *done)
+
+The signaling will work as expected even if completions are signaled before
+a thread starts waiting. This is achieved by the waiter "consuming"
+(decrementing) the done element of struct completion. Waiting threads
+wakeup order is the same in which they were enqueued (FIFO order).
+
+If complete() is called multiple times then this will allow for that number
+of waiters to continue - each call to complete() will simply increment the
+done element. Calling complete_all() multiple times is a bug though. Both
+complete() and complete_all() can be called in hard-irq/atomic context safely.
+
+There only can be one thread calling complete() or complete_all() on a
+particular struct completion at any time - serialized through the wait
+queue spinlock. Any such concurrent calls to complete() or complete_all()
+probably are a design bug.
+
+Signaling completion from hard-irq context is fine as it will appropriately
+lock with spin_lock_irqsave/spin_unlock_irqrestore and it will never sleep.
+
+
+try_wait_for_completion()/completion_done():
+--------------------------------------------
+
+The try_wait_for_completion() function will not put the thread on the wait
+queue but rather returns false if it would need to enqueue (block) the thread,
+else it consumes one posted completion and returns true.
+
+	bool try_wait_for_completion(struct completion *done)
+
+Finally, to check the state of a completion without changing it in any way, 
+call completion_done(), which returns false if there are no posted
+completions that were not yet consumed by waiters (implying that there are
+waiters) and true otherwise;
+
+	bool completion_done(struct completion *done)
+
+Both try_wait_for_completion() and completion_done() are safe to be called in
+hard-irq or atomic context.
diff --git a/Documentation/scheduler/sched-arch.txt b/Documentation/scheduler/sched-arch.txt
new file mode 100644
index 000000000..a2f27bbf2
--- /dev/null
+++ b/Documentation/scheduler/sched-arch.txt
@@ -0,0 +1,74 @@
+	CPU Scheduler implementation hints for architecture specific code
+
+	Nick Piggin, 2005
+
+Context switch
+==============
+1. Runqueue locking
+By default, the switch_to arch function is called with the runqueue
+locked. This is usually not a problem unless switch_to may need to
+take the runqueue lock. This is usually due to a wake up operation in
+the context switch. See arch/ia64/include/asm/switch_to.h for an example.
+
+To request the scheduler call switch_to with the runqueue unlocked,
+you must `#define __ARCH_WANT_UNLOCKED_CTXSW` in a header file
+(typically the one where switch_to is defined).
+
+Unlocked context switches introduce only a very minor performance
+penalty to the core scheduler implementation in the CONFIG_SMP case.
+
+CPU idle
+========
+Your cpu_idle routines need to obey the following rules:
+
+1. Preempt should now disabled over idle routines. Should only
+   be enabled to call schedule() then disabled again.
+
+2. need_resched/TIF_NEED_RESCHED is only ever set, and will never
+   be cleared until the running task has called schedule(). Idle
+   threads need only ever query need_resched, and may never set or
+   clear it.
+
+3. When cpu_idle finds (need_resched() == 'true'), it should call
+   schedule(). It should not call schedule() otherwise.
+
+4. The only time interrupts need to be disabled when checking
+   need_resched is if we are about to sleep the processor until
+   the next interrupt (this doesn't provide any protection of
+   need_resched, it prevents losing an interrupt).
+
+	4a. Common problem with this type of sleep appears to be:
+	        local_irq_disable();
+	        if (!need_resched()) {
+	                local_irq_enable();
+	                *** resched interrupt arrives here ***
+	                __asm__("sleep until next interrupt");
+	        }
+
+5. TIF_POLLING_NRFLAG can be set by idle routines that do not
+   need an interrupt to wake them up when need_resched goes high.
+   In other words, they must be periodically polling need_resched,
+   although it may be reasonable to do some background work or enter
+   a low CPU priority.
+
+   	5a. If TIF_POLLING_NRFLAG is set, and we do decide to enter
+	    an interrupt sleep, it needs to be cleared then a memory
+	    barrier issued (followed by a test of need_resched with
+	    interrupts disabled, as explained in 3).
+
+arch/x86/kernel/process.c has examples of both polling and
+sleeping idle functions.
+
+
+Possible arch/ problems
+=======================
+
+Possible arch problems I found (and either tried to fix or didn't):
+
+ia64 - is safe_halt call racy vs interrupts? (does it sleep?) (See #4a)
+
+sh64 - Is sleeping racy vs interrupts? (See #4a)
+
+sparc - IRQs on at this point(?), change local_irq_save to _disable.
+      - TODO: needs secondary CPUs to disable preempt (See #1)
+
diff --git a/Documentation/scheduler/sched-bwc.txt b/Documentation/scheduler/sched-bwc.txt
new file mode 100644
index 000000000..de583fbbf
--- /dev/null
+++ b/Documentation/scheduler/sched-bwc.txt
@@ -0,0 +1,167 @@
+CFS Bandwidth Control
+=====================
+
+[ This document only discusses CPU bandwidth control for SCHED_NORMAL.
+  The SCHED_RT case is covered in Documentation/scheduler/sched-rt-group.txt ]
+
+CFS bandwidth control is a CONFIG_FAIR_GROUP_SCHED extension which allows the
+specification of the maximum CPU bandwidth available to a group or hierarchy.
+
+The bandwidth allowed for a group is specified using a quota and period. Within
+each given "period" (microseconds), a group is allowed to consume only up to
+"quota" microseconds of CPU time.  When the CPU bandwidth consumption of a
+group exceeds this limit (for that period), the tasks belonging to its
+hierarchy will be throttled and are not allowed to run again until the next
+period.
+
+A group's unused runtime is globally tracked, being refreshed with quota units
+above at each period boundary.  As threads consume this bandwidth it is
+transferred to cpu-local "silos" on a demand basis.  The amount transferred
+within each of these updates is tunable and described as the "slice".
+
+Management
+----------
+Quota and period are managed within the cpu subsystem via cgroupfs.
+
+cpu.cfs_quota_us: the total available run-time within a period (in microseconds)
+cpu.cfs_period_us: the length of a period (in microseconds)
+cpu.stat: exports throttling statistics [explained further below]
+
+The default values are:
+	cpu.cfs_period_us=100ms
+	cpu.cfs_quota=-1
+
+A value of -1 for cpu.cfs_quota_us indicates that the group does not have any
+bandwidth restriction in place, such a group is described as an unconstrained
+bandwidth group.  This represents the traditional work-conserving behavior for
+CFS.
+
+Writing any (valid) positive value(s) will enact the specified bandwidth limit.
+The minimum quota allowed for the quota or period is 1ms.  There is also an
+upper bound on the period length of 1s.  Additional restrictions exist when
+bandwidth limits are used in a hierarchical fashion, these are explained in
+more detail below.
+
+Writing any negative value to cpu.cfs_quota_us will remove the bandwidth limit
+and return the group to an unconstrained state once more.
+
+Any updates to a group's bandwidth specification will result in it becoming
+unthrottled if it is in a constrained state.
+
+System wide settings
+--------------------
+For efficiency run-time is transferred between the global pool and CPU local
+"silos" in a batch fashion.  This greatly reduces global accounting pressure
+on large systems.  The amount transferred each time such an update is required
+is described as the "slice".
+
+This is tunable via procfs:
+	/proc/sys/kernel/sched_cfs_bandwidth_slice_us (default=5ms)
+
+Larger slice values will reduce transfer overheads, while smaller values allow
+for more fine-grained consumption.
+
+Statistics
+----------
+A group's bandwidth statistics are exported via 3 fields in cpu.stat.
+
+cpu.stat:
+- nr_periods: Number of enforcement intervals that have elapsed.
+- nr_throttled: Number of times the group has been throttled/limited.
+- throttled_time: The total time duration (in nanoseconds) for which entities
+  of the group have been throttled.
+
+This interface is read-only.
+
+Hierarchical considerations
+---------------------------
+The interface enforces that an individual entity's bandwidth is always
+attainable, that is: max(c_i) <= C. However, over-subscription in the
+aggregate case is explicitly allowed to enable work-conserving semantics
+within a hierarchy.
+  e.g. \Sum (c_i) may exceed C
+[ Where C is the parent's bandwidth, and c_i its children ]
+
+
+There are two ways in which a group may become throttled:
+	a. it fully consumes its own quota within a period
+	b. a parent's quota is fully consumed within its period
+
+In case b) above, even though the child may have runtime remaining it will not
+be allowed to until the parent's runtime is refreshed.
+
+CFS Bandwidth Quota Caveats
+---------------------------
+Once a slice is assigned to a cpu it does not expire.  However all but 1ms of
+the slice may be returned to the global pool if all threads on that cpu become
+unrunnable. This is configured at compile time by the min_cfs_rq_runtime
+variable. This is a performance tweak that helps prevent added contention on
+the global lock.
+
+The fact that cpu-local slices do not expire results in some interesting corner
+cases that should be understood.
+
+For cgroup cpu constrained applications that are cpu limited this is a
+relatively moot point because they will naturally consume the entirety of their
+quota as well as the entirety of each cpu-local slice in each period. As a
+result it is expected that nr_periods roughly equal nr_throttled, and that
+cpuacct.usage will increase roughly equal to cfs_quota_us in each period.
+
+For highly-threaded, non-cpu bound applications this non-expiration nuance
+allows applications to briefly burst past their quota limits by the amount of
+unused slice on each cpu that the task group is running on (typically at most
+1ms per cpu or as defined by min_cfs_rq_runtime).  This slight burst only
+applies if quota had been assigned to a cpu and then not fully used or returned
+in previous periods. This burst amount will not be transferred between cores.
+As a result, this mechanism still strictly limits the task group to quota
+average usage, albeit over a longer time window than a single period.  This
+also limits the burst ability to no more than 1ms per cpu.  This provides
+better more predictable user experience for highly threaded applications with
+small quota limits on high core count machines. It also eliminates the
+propensity to throttle these applications while simultanously using less than
+quota amounts of cpu. Another way to say this, is that by allowing the unused
+portion of a slice to remain valid across periods we have decreased the
+possibility of wastefully expiring quota on cpu-local silos that don't need a
+full slice's amount of cpu time.
+
+The interaction between cpu-bound and non-cpu-bound-interactive applications
+should also be considered, especially when single core usage hits 100%. If you
+gave each of these applications half of a cpu-core and they both got scheduled
+on the same CPU it is theoretically possible that the non-cpu bound application
+will use up to 1ms additional quota in some periods, thereby preventing the
+cpu-bound application from fully using its quota by that same amount. In these
+instances it will be up to the CFS algorithm (see sched-design-CFS.rst) to
+decide which application is chosen to run, as they will both be runnable and
+have remaining quota. This runtime discrepancy will be made up in the following
+periods when the interactive application idles.
+
+Examples
+--------
+1. Limit a group to 1 CPU worth of runtime.
+
+	If period is 250ms and quota is also 250ms, the group will get
+	1 CPU worth of runtime every 250ms.
+
+	# echo 250000 > cpu.cfs_quota_us /* quota = 250ms */
+	# echo 250000 > cpu.cfs_period_us /* period = 250ms */
+
+2. Limit a group to 2 CPUs worth of runtime on a multi-CPU machine.
+
+	With 500ms period and 1000ms quota, the group can get 2 CPUs worth of
+	runtime every 500ms.
+
+	# echo 1000000 > cpu.cfs_quota_us /* quota = 1000ms */
+	# echo 500000 > cpu.cfs_period_us /* period = 500ms */
+
+	The larger period here allows for increased burst capacity.
+
+3. Limit a group to 20% of 1 CPU.
+
+	With 50ms period, 10ms quota will be equivalent to 20% of 1 CPU.
+
+	# echo 10000 > cpu.cfs_quota_us /* quota = 10ms */
+	# echo 50000 > cpu.cfs_period_us /* period = 50ms */
+
+	By using a small period here we are ensuring a consistent latency
+	response at the expense of burst capacity.
+
diff --git a/Documentation/scheduler/sched-deadline.txt b/Documentation/scheduler/sched-deadline.txt
new file mode 100644
index 000000000..b14e03ff3
--- /dev/null
+++ b/Documentation/scheduler/sched-deadline.txt
@@ -0,0 +1,871 @@
+			  Deadline Task Scheduling
+			  ------------------------
+
+CONTENTS
+========
+
+ 0. WARNING
+ 1. Overview
+ 2. Scheduling algorithm
+   2.1 Main algorithm
+   2.2 Bandwidth reclaiming
+ 3. Scheduling Real-Time Tasks
+   3.1 Definitions
+   3.2 Schedulability Analysis for Uniprocessor Systems
+   3.3 Schedulability Analysis for Multiprocessor Systems
+   3.4 Relationship with SCHED_DEADLINE Parameters
+ 4. Bandwidth management
+   4.1 System-wide settings
+   4.2 Task interface
+   4.3 Default behavior
+   4.4 Behavior of sched_yield()
+ 5. Tasks CPU affinity
+   5.1 SCHED_DEADLINE and cpusets HOWTO
+ 6. Future plans
+ A. Test suite
+ B. Minimal main()
+
+
+0. WARNING
+==========
+
+ Fiddling with these settings can result in an unpredictable or even unstable
+ system behavior. As for -rt (group) scheduling, it is assumed that root users
+ know what they're doing.
+
+
+1. Overview
+===========
+
+ The SCHED_DEADLINE policy contained inside the sched_dl scheduling class is
+ basically an implementation of the Earliest Deadline First (EDF) scheduling
+ algorithm, augmented with a mechanism (called Constant Bandwidth Server, CBS)
+ that makes it possible to isolate the behavior of tasks between each other.
+
+
+2. Scheduling algorithm
+==================
+
+2.1 Main algorithm
+------------------
+
+ SCHED_DEADLINE [18] uses three parameters, named "runtime", "period", and
+ "deadline", to schedule tasks. A SCHED_DEADLINE task should receive
+ "runtime" microseconds of execution time every "period" microseconds, and
+ these "runtime" microseconds are available within "deadline" microseconds
+ from the beginning of the period.  In order to implement this behavior,
+ every time the task wakes up, the scheduler computes a "scheduling deadline"
+ consistent with the guarantee (using the CBS[2,3] algorithm). Tasks are then
+ scheduled using EDF[1] on these scheduling deadlines (the task with the
+ earliest scheduling deadline is selected for execution). Notice that the
+ task actually receives "runtime" time units within "deadline" if a proper
+ "admission control" strategy (see Section "4. Bandwidth management") is used
+ (clearly, if the system is overloaded this guarantee cannot be respected).
+
+ Summing up, the CBS[2,3] algorithm assigns scheduling deadlines to tasks so
+ that each task runs for at most its runtime every period, avoiding any
+ interference between different tasks (bandwidth isolation), while the EDF[1]
+ algorithm selects the task with the earliest scheduling deadline as the one
+ to be executed next. Thanks to this feature, tasks that do not strictly comply
+ with the "traditional" real-time task model (see Section 3) can effectively
+ use the new policy.
+
+ In more details, the CBS algorithm assigns scheduling deadlines to
+ tasks in the following way:
+
+  - Each SCHED_DEADLINE task is characterized by the "runtime",
+    "deadline", and "period" parameters;
+
+  - The state of the task is described by a "scheduling deadline", and
+    a "remaining runtime". These two parameters are initially set to 0;
+
+  - When a SCHED_DEADLINE task wakes up (becomes ready for execution),
+    the scheduler checks if
+
+                 remaining runtime                  runtime
+        ----------------------------------    >    ---------
+        scheduling deadline - current time           period
+
+    then, if the scheduling deadline is smaller than the current time, or
+    this condition is verified, the scheduling deadline and the
+    remaining runtime are re-initialized as
+
+         scheduling deadline = current time + deadline
+         remaining runtime = runtime
+
+    otherwise, the scheduling deadline and the remaining runtime are
+    left unchanged;
+
+  - When a SCHED_DEADLINE task executes for an amount of time t, its
+    remaining runtime is decreased as
+
+         remaining runtime = remaining runtime - t
+
+    (technically, the runtime is decreased at every tick, or when the
+    task is descheduled / preempted);
+
+  - When the remaining runtime becomes less or equal than 0, the task is
+    said to be "throttled" (also known as "depleted" in real-time literature)
+    and cannot be scheduled until its scheduling deadline. The "replenishment
+    time" for this task (see next item) is set to be equal to the current
+    value of the scheduling deadline;
+
+  - When the current time is equal to the replenishment time of a
+    throttled task, the scheduling deadline and the remaining runtime are
+    updated as
+
+         scheduling deadline = scheduling deadline + period
+         remaining runtime = remaining runtime + runtime
+
+ The SCHED_FLAG_DL_OVERRUN flag in sched_attr's sched_flags field allows a task
+ to get informed about runtime overruns through the delivery of SIGXCPU
+ signals.
+
+
+2.2 Bandwidth reclaiming
+------------------------
+
+ Bandwidth reclaiming for deadline tasks is based on the GRUB (Greedy
+ Reclamation of Unused Bandwidth) algorithm [15, 16, 17] and it is enabled
+ when flag SCHED_FLAG_RECLAIM is set.
+
+ The following diagram illustrates the state names for tasks handled by GRUB:
+
+                             ------------
+                 (d)        |   Active   |
+              ------------->|            |
+              |             | Contending |
+              |              ------------
+              |                A      |
+          ----------           |      |
+         |          |          |      |
+         | Inactive |          |(b)   | (a)
+         |          |          |      |
+          ----------           |      |
+              A                |      V
+              |              ------------
+              |             |   Active   |
+              --------------|     Non    |
+                 (c)        | Contending |
+                             ------------
+
+ A task can be in one of the following states:
+
+  - ActiveContending: if it is ready for execution (or executing);
+
+  - ActiveNonContending: if it just blocked and has not yet surpassed the 0-lag
+    time;
+
+  - Inactive: if it is blocked and has surpassed the 0-lag time.
+
+ State transitions:
+
+  (a) When a task blocks, it does not become immediately inactive since its
+      bandwidth cannot be immediately reclaimed without breaking the
+      real-time guarantees. It therefore enters a transitional state called
+      ActiveNonContending. The scheduler arms the "inactive timer" to fire at
+      the 0-lag time, when the task's bandwidth can be reclaimed without
+      breaking the real-time guarantees.
+
+      The 0-lag time for a task entering the ActiveNonContending state is
+      computed as
+
+                        (runtime * dl_period)
+             deadline - ---------------------
+                             dl_runtime
+
+      where runtime is the remaining runtime, while dl_runtime and dl_period
+      are the reservation parameters.
+
+  (b) If the task wakes up before the inactive timer fires, the task re-enters
+      the ActiveContending state and the "inactive timer" is canceled.
+      In addition, if the task wakes up on a different runqueue, then
+      the task's utilization must be removed from the previous runqueue's active
+      utilization and must be added to the new runqueue's active utilization.
+      In order to avoid races between a task waking up on a runqueue while the
+       "inactive timer" is running on a different CPU, the "dl_non_contending"
+      flag is used to indicate that a task is not on a runqueue but is active
+      (so, the flag is set when the task blocks and is cleared when the
+      "inactive timer" fires or when the task  wakes up).
+
+  (c) When the "inactive timer" fires, the task enters the Inactive state and
+      its utilization is removed from the runqueue's active utilization.
+
+  (d) When an inactive task wakes up, it enters the ActiveContending state and
+      its utilization is added to the active utilization of the runqueue where
+      it has been enqueued.
+
+ For each runqueue, the algorithm GRUB keeps track of two different bandwidths:
+
+  - Active bandwidth (running_bw): this is the sum of the bandwidths of all
+    tasks in active state (i.e., ActiveContending or ActiveNonContending);
+
+  - Total bandwidth (this_bw): this is the sum of all tasks "belonging" to the
+    runqueue, including the tasks in Inactive state.
+
+
+ The algorithm reclaims the bandwidth of the tasks in Inactive state.
+ It does so by decrementing the runtime of the executing task Ti at a pace equal
+ to
+
+           dq = -max{ Ui / Umax, (1 - Uinact - Uextra) } dt
+
+ where:
+
+  - Ui is the bandwidth of task Ti;
+  - Umax is the maximum reclaimable utilization (subjected to RT throttling
+    limits);
+  - Uinact is the (per runqueue) inactive utilization, computed as
+    (this_bq - running_bw);
+  - Uextra is the (per runqueue) extra reclaimable utilization
+    (subjected to RT throttling limits).
+
+
+ Let's now see a trivial example of two deadline tasks with runtime equal
+ to 4 and period equal to 8 (i.e., bandwidth equal to 0.5):
+
+     A            Task T1
+     |
+     |                               |
+     |                               |
+     |--------                       |----
+     |       |                       V
+     |---|---|---|---|---|---|---|---|--------->t
+     0   1   2   3   4   5   6   7   8
+
+
+     A            Task T2
+     |
+     |                               |
+     |                               |
+     |       ------------------------|
+     |       |                       V
+     |---|---|---|---|---|---|---|---|--------->t
+     0   1   2   3   4   5   6   7   8
+
+
+     A            running_bw
+     |
+   1 -----------------               ------
+     |               |               |
+  0.5-               -----------------
+     |                               |
+     |---|---|---|---|---|---|---|---|--------->t
+     0   1   2   3   4   5   6   7   8
+
+
+  - Time t = 0:
+
+    Both tasks are ready for execution and therefore in ActiveContending state.
+    Suppose Task T1 is the first task to start execution.
+    Since there are no inactive tasks, its runtime is decreased as dq = -1 dt.
+
+  - Time t = 2:
+
+    Suppose that task T1 blocks
+    Task T1 therefore enters the ActiveNonContending state. Since its remaining
+    runtime is equal to 2, its 0-lag time is equal to t = 4.
+    Task T2 start execution, with runtime still decreased as dq = -1 dt since
+    there are no inactive tasks.
+
+  - Time t = 4:
+
+    This is the 0-lag time for Task T1. Since it didn't woken up in the
+    meantime, it enters the Inactive state. Its bandwidth is removed from
+    running_bw.
+    Task T2 continues its execution. However, its runtime is now decreased as
+    dq = - 0.5 dt because Uinact = 0.5.
+    Task T2 therefore reclaims the bandwidth unused by Task T1.
+
+  - Time t = 8:
+
+    Task T1 wakes up. It enters the ActiveContending state again, and the
+    running_bw is incremented.
+
+
+2.3 Energy-aware scheduling
+------------------------
+
+ When cpufreq's schedutil governor is selected, SCHED_DEADLINE implements the
+ GRUB-PA [19] algorithm, reducing the CPU operating frequency to the minimum
+ value that still allows to meet the deadlines. This behavior is currently
+ implemented only for ARM architectures.
+
+ A particular care must be taken in case the time needed for changing frequency
+ is of the same order of magnitude of the reservation period. In such cases,
+ setting a fixed CPU frequency results in a lower amount of deadline misses.
+
+
+3. Scheduling Real-Time Tasks
+=============================
+
+ * BIG FAT WARNING ******************************************************
+ *
+ * This section contains a (not-thorough) summary on classical deadline
+ * scheduling theory, and how it applies to SCHED_DEADLINE.
+ * The reader can "safely" skip to Section 4 if only interested in seeing
+ * how the scheduling policy can be used. Anyway, we strongly recommend
+ * to come back here and continue reading (once the urge for testing is
+ * satisfied :P) to be sure of fully understanding all technical details.
+ ************************************************************************
+
+ There are no limitations on what kind of task can exploit this new
+ scheduling discipline, even if it must be said that it is particularly
+ suited for periodic or sporadic real-time tasks that need guarantees on their
+ timing behavior, e.g., multimedia, streaming, control applications, etc.
+
+3.1 Definitions
+------------------------
+
+ A typical real-time task is composed of a repetition of computation phases
+ (task instances, or jobs) which are activated on a periodic or sporadic
+ fashion.
+ Each job J_j (where J_j is the j^th job of the task) is characterized by an
+ arrival time r_j (the time when the job starts), an amount of computation
+ time c_j needed to finish the job, and a job absolute deadline d_j, which
+ is the time within which the job should be finished. The maximum execution
+ time max{c_j} is called "Worst Case Execution Time" (WCET) for the task.
+ A real-time task can be periodic with period P if r_{j+1} = r_j + P, or
+ sporadic with minimum inter-arrival time P is r_{j+1} >= r_j + P. Finally,
+ d_j = r_j + D, where D is the task's relative deadline.
+ Summing up, a real-time task can be described as
+	Task = (WCET, D, P)
+
+ The utilization of a real-time task is defined as the ratio between its
+ WCET and its period (or minimum inter-arrival time), and represents
+ the fraction of CPU time needed to execute the task.
+
+ If the total utilization U=sum(WCET_i/P_i) is larger than M (with M equal
+ to the number of CPUs), then the scheduler is unable to respect all the
+ deadlines.
+ Note that total utilization is defined as the sum of the utilizations
+ WCET_i/P_i over all the real-time tasks in the system. When considering
+ multiple real-time tasks, the parameters of the i-th task are indicated
+ with the "_i" suffix.
+ Moreover, if the total utilization is larger than M, then we risk starving
+ non- real-time tasks by real-time tasks.
+ If, instead, the total utilization is smaller than M, then non real-time
+ tasks will not be starved and the system might be able to respect all the
+ deadlines.
+ As a matter of fact, in this case it is possible to provide an upper bound
+ for tardiness (defined as the maximum between 0 and the difference
+ between the finishing time of a job and its absolute deadline).
+ More precisely, it can be proven that using a global EDF scheduler the
+ maximum tardiness of each task is smaller or equal than
+	((M − 1) · WCET_max − WCET_min)/(M − (M − 2) · U_max) + WCET_max
+ where WCET_max = max{WCET_i} is the maximum WCET, WCET_min=min{WCET_i}
+ is the minimum WCET, and U_max = max{WCET_i/P_i} is the maximum
+ utilization[12].
+
+3.2 Schedulability Analysis for Uniprocessor Systems
+------------------------
+
+ If M=1 (uniprocessor system), or in case of partitioned scheduling (each
+ real-time task is statically assigned to one and only one CPU), it is
+ possible to formally check if all the deadlines are respected.
+ If D_i = P_i for all tasks, then EDF is able to respect all the deadlines
+ of all the tasks executing on a CPU if and only if the total utilization
+ of the tasks running on such a CPU is smaller or equal than 1.
+ If D_i != P_i for some task, then it is possible to define the density of
+ a task as WCET_i/min{D_i,P_i}, and EDF is able to respect all the deadlines
+ of all the tasks running on a CPU if the sum of the densities of the tasks
+ running on such a CPU is smaller or equal than 1:
+	sum(WCET_i / min{D_i, P_i}) <= 1
+ It is important to notice that this condition is only sufficient, and not
+ necessary: there are task sets that are schedulable, but do not respect the
+ condition. For example, consider the task set {Task_1,Task_2} composed by
+ Task_1=(50ms,50ms,100ms) and Task_2=(10ms,100ms,100ms).
+ EDF is clearly able to schedule the two tasks without missing any deadline
+ (Task_1 is scheduled as soon as it is released, and finishes just in time
+ to respect its deadline; Task_2 is scheduled immediately after Task_1, hence
+ its response time cannot be larger than 50ms + 10ms = 60ms) even if
+	50 / min{50,100} + 10 / min{100, 100} = 50 / 50 + 10 / 100 = 1.1
+ Of course it is possible to test the exact schedulability of tasks with
+ D_i != P_i (checking a condition that is both sufficient and necessary),
+ but this cannot be done by comparing the total utilization or density with
+ a constant. Instead, the so called "processor demand" approach can be used,
+ computing the total amount of CPU time h(t) needed by all the tasks to
+ respect all of their deadlines in a time interval of size t, and comparing
+ such a time with the interval size t. If h(t) is smaller than t (that is,
+ the amount of time needed by the tasks in a time interval of size t is
+ smaller than the size of the interval) for all the possible values of t, then
+ EDF is able to schedule the tasks respecting all of their deadlines. Since
+ performing this check for all possible values of t is impossible, it has been
+ proven[4,5,6] that it is sufficient to perform the test for values of t
+ between 0 and a maximum value L. The cited papers contain all of the
+ mathematical details and explain how to compute h(t) and L.
+ In any case, this kind of analysis is too complex as well as too
+ time-consuming to be performed on-line. Hence, as explained in Section
+ 4 Linux uses an admission test based on the tasks' utilizations.
+
+3.3 Schedulability Analysis for Multiprocessor Systems
+------------------------
+
+ On multiprocessor systems with global EDF scheduling (non partitioned
+ systems), a sufficient test for schedulability can not be based on the
+ utilizations or densities: it can be shown that even if D_i = P_i task
+ sets with utilizations slightly larger than 1 can miss deadlines regardless
+ of the number of CPUs.
+
+ Consider a set {Task_1,...Task_{M+1}} of M+1 tasks on a system with M
+ CPUs, with the first task Task_1=(P,P,P) having period, relative deadline
+ and WCET equal to P. The remaining M tasks Task_i=(e,P-1,P-1) have an
+ arbitrarily small worst case execution time (indicated as "e" here) and a
+ period smaller than the one of the first task. Hence, if all the tasks
+ activate at the same time t, global EDF schedules these M tasks first
+ (because their absolute deadlines are equal to t + P - 1, hence they are
+ smaller than the absolute deadline of Task_1, which is t + P). As a
+ result, Task_1 can be scheduled only at time t + e, and will finish at
+ time t + e + P, after its absolute deadline. The total utilization of the
+ task set is U = M · e / (P - 1) + P / P = M · e / (P - 1) + 1, and for small
+ values of e this can become very close to 1. This is known as "Dhall's
+ effect"[7]. Note: the example in the original paper by Dhall has been
+ slightly simplified here (for example, Dhall more correctly computed
+ lim_{e->0}U).
+
+ More complex schedulability tests for global EDF have been developed in
+ real-time literature[8,9], but they are not based on a simple comparison
+ between total utilization (or density) and a fixed constant. If all tasks
+ have D_i = P_i, a sufficient schedulability condition can be expressed in
+ a simple way:
+	sum(WCET_i / P_i) <= M - (M - 1) · U_max
+ where U_max = max{WCET_i / P_i}[10]. Notice that for U_max = 1,
+ M - (M - 1) · U_max becomes M - M + 1 = 1 and this schedulability condition
+ just confirms the Dhall's effect. A more complete survey of the literature
+ about schedulability tests for multi-processor real-time scheduling can be
+ found in [11].
+
+ As seen, enforcing that the total utilization is smaller than M does not
+ guarantee that global EDF schedules the tasks without missing any deadline
+ (in other words, global EDF is not an optimal scheduling algorithm). However,
+ a total utilization smaller than M is enough to guarantee that non real-time
+ tasks are not starved and that the tardiness of real-time tasks has an upper
+ bound[12] (as previously noted). Different bounds on the maximum tardiness
+ experienced by real-time tasks have been developed in various papers[13,14],
+ but the theoretical result that is important for SCHED_DEADLINE is that if
+ the total utilization is smaller or equal than M then the response times of
+ the tasks are limited.
+
+3.4 Relationship with SCHED_DEADLINE Parameters
+------------------------
+
+ Finally, it is important to understand the relationship between the
+ SCHED_DEADLINE scheduling parameters described in Section 2 (runtime,
+ deadline and period) and the real-time task parameters (WCET, D, P)
+ described in this section. Note that the tasks' temporal constraints are
+ represented by its absolute deadlines d_j = r_j + D described above, while
+ SCHED_DEADLINE schedules the tasks according to scheduling deadlines (see
+ Section 2).
+ If an admission test is used to guarantee that the scheduling deadlines
+ are respected, then SCHED_DEADLINE can be used to schedule real-time tasks
+ guaranteeing that all the jobs' deadlines of a task are respected.
+ In order to do this, a task must be scheduled by setting:
+
+  - runtime >= WCET
+  - deadline = D
+  - period <= P
+
+ IOW, if runtime >= WCET and if period is <= P, then the scheduling deadlines
+ and the absolute deadlines (d_j) coincide, so a proper admission control
+ allows to respect the jobs' absolute deadlines for this task (this is what is
+ called "hard schedulability property" and is an extension of Lemma 1 of [2]).
+ Notice that if runtime > deadline the admission control will surely reject
+ this task, as it is not possible to respect its temporal constraints.
+
+ References:
+  1 - C. L. Liu and J. W. Layland. Scheduling algorithms for multiprogram-
+      ming in a hard-real-time environment. Journal of the Association for
+      Computing Machinery, 20(1), 1973.
+  2 - L. Abeni , G. Buttazzo. Integrating Multimedia Applications in Hard
+      Real-Time Systems. Proceedings of the 19th IEEE Real-time Systems
+      Symposium, 1998. http://retis.sssup.it/~giorgio/paps/1998/rtss98-cbs.pdf
+  3 - L. Abeni. Server Mechanisms for Multimedia Applications. ReTiS Lab
+      Technical Report. http://disi.unitn.it/~abeni/tr-98-01.pdf
+  4 - J. Y. Leung and M.L. Merril. A Note on Preemptive Scheduling of
+      Periodic, Real-Time Tasks. Information Processing Letters, vol. 11,
+      no. 3, pp. 115-118, 1980.
+  5 - S. K. Baruah, A. K. Mok and L. E. Rosier. Preemptively Scheduling
+      Hard-Real-Time Sporadic Tasks on One Processor. Proceedings of the
+      11th IEEE Real-time Systems Symposium, 1990.
+  6 - S. K. Baruah, L. E. Rosier and R. R. Howell. Algorithms and Complexity
+      Concerning the Preemptive Scheduling of Periodic Real-Time tasks on
+      One Processor. Real-Time Systems Journal, vol. 4, no. 2, pp 301-324,
+      1990.
+  7 - S. J. Dhall and C. L. Liu. On a real-time scheduling problem. Operations
+      research, vol. 26, no. 1, pp 127-140, 1978.
+  8 - T. Baker. Multiprocessor EDF and Deadline Monotonic Schedulability
+      Analysis. Proceedings of the 24th IEEE Real-Time Systems Symposium, 2003.
+  9 - T. Baker. An Analysis of EDF Schedulability on a Multiprocessor.
+      IEEE Transactions on Parallel and Distributed Systems, vol. 16, no. 8,
+      pp 760-768, 2005.
+  10 - J. Goossens, S. Funk and S. Baruah, Priority-Driven Scheduling of
+       Periodic Task Systems on Multiprocessors. Real-Time Systems Journal,
+       vol. 25, no. 2–3, pp. 187–205, 2003.
+  11 - R. Davis and A. Burns. A Survey of Hard Real-Time Scheduling for
+       Multiprocessor Systems. ACM Computing Surveys, vol. 43, no. 4, 2011.
+       http://www-users.cs.york.ac.uk/~robdavis/papers/MPSurveyv5.0.pdf
+  12 - U. C. Devi and J. H. Anderson. Tardiness Bounds under Global EDF
+       Scheduling on a Multiprocessor. Real-Time Systems Journal, vol. 32,
+       no. 2, pp 133-189, 2008.
+  13 - P. Valente and G. Lipari. An Upper Bound to the Lateness of Soft
+       Real-Time Tasks Scheduled by EDF on Multiprocessors. Proceedings of
+       the 26th IEEE Real-Time Systems Symposium, 2005.
+  14 - J. Erickson, U. Devi and S. Baruah. Improved tardiness bounds for
+       Global EDF. Proceedings of the 22nd Euromicro Conference on
+       Real-Time Systems, 2010.
+  15 - G. Lipari, S. Baruah, Greedy reclamation of unused bandwidth in
+       constant-bandwidth servers, 12th IEEE Euromicro Conference on Real-Time
+       Systems, 2000.
+  16 - L. Abeni, J. Lelli, C. Scordino, L. Palopoli, Greedy CPU reclaiming for
+       SCHED DEADLINE. In Proceedings of the Real-Time Linux Workshop (RTLWS),
+       Dusseldorf, Germany, 2014.
+  17 - L. Abeni, G. Lipari, A. Parri, Y. Sun, Multicore CPU reclaiming: parallel
+       or sequential?. In Proceedings of the 31st Annual ACM Symposium on Applied
+       Computing, 2016.
+  18 - J. Lelli, C. Scordino, L. Abeni, D. Faggioli, Deadline scheduling in the
+       Linux kernel, Software: Practice and Experience, 46(6): 821-839, June
+       2016.
+  19 - C. Scordino, L. Abeni, J. Lelli, Energy-Aware Real-Time Scheduling in
+       the Linux Kernel, 33rd ACM/SIGAPP Symposium On Applied Computing (SAC
+       2018), Pau, France, April 2018.
+
+
+4. Bandwidth management
+=======================
+
+ As previously mentioned, in order for -deadline scheduling to be
+ effective and useful (that is, to be able to provide "runtime" time units
+ within "deadline"), it is important to have some method to keep the allocation
+ of the available fractions of CPU time to the various tasks under control.
+ This is usually called "admission control" and if it is not performed, then
+ no guarantee can be given on the actual scheduling of the -deadline tasks.
+
+ As already stated in Section 3, a necessary condition to be respected to
+ correctly schedule a set of real-time tasks is that the total utilization
+ is smaller than M. When talking about -deadline tasks, this requires that
+ the sum of the ratio between runtime and period for all tasks is smaller
+ than M. Notice that the ratio runtime/period is equivalent to the utilization
+ of a "traditional" real-time task, and is also often referred to as
+ "bandwidth".
+ The interface used to control the CPU bandwidth that can be allocated
+ to -deadline tasks is similar to the one already used for -rt
+ tasks with real-time group scheduling (a.k.a. RT-throttling - see
+ Documentation/scheduler/sched-rt-group.txt), and is based on readable/
+ writable control files located in procfs (for system wide settings).
+ Notice that per-group settings (controlled through cgroupfs) are still not
+ defined for -deadline tasks, because more discussion is needed in order to
+ figure out how we want to manage SCHED_DEADLINE bandwidth at the task group
+ level.
+
+ A main difference between deadline bandwidth management and RT-throttling
+ is that -deadline tasks have bandwidth on their own (while -rt ones don't!),
+ and thus we don't need a higher level throttling mechanism to enforce the
+ desired bandwidth. In other words, this means that interface parameters are
+ only used at admission control time (i.e., when the user calls
+ sched_setattr()). Scheduling is then performed considering actual tasks'
+ parameters, so that CPU bandwidth is allocated to SCHED_DEADLINE tasks
+ respecting their needs in terms of granularity. Therefore, using this simple
+ interface we can put a cap on total utilization of -deadline tasks (i.e.,
+ \Sum (runtime_i / period_i) < global_dl_utilization_cap).
+
+4.1 System wide settings
+------------------------
+
+ The system wide settings are configured under the /proc virtual file system.
+
+ For now the -rt knobs are used for -deadline admission control and the
+ -deadline runtime is accounted against the -rt runtime. We realize that this
+ isn't entirely desirable; however, it is better to have a small interface for
+ now, and be able to change it easily later. The ideal situation (see 5.) is to
+ run -rt tasks from a -deadline server; in which case the -rt bandwidth is a
+ direct subset of dl_bw.
+
+ This means that, for a root_domain comprising M CPUs, -deadline tasks
+ can be created while the sum of their bandwidths stays below:
+
+   M * (sched_rt_runtime_us / sched_rt_period_us)
+
+ It is also possible to disable this bandwidth management logic, and
+ be thus free of oversubscribing the system up to any arbitrary level.
+ This is done by writing -1 in /proc/sys/kernel/sched_rt_runtime_us.
+
+
+4.2 Task interface
+------------------
+
+ Specifying a periodic/sporadic task that executes for a given amount of
+ runtime at each instance, and that is scheduled according to the urgency of
+ its own timing constraints needs, in general, a way of declaring:
+  - a (maximum/typical) instance execution time,
+  - a minimum interval between consecutive instances,
+  - a time constraint by which each instance must be completed.
+
+ Therefore:
+  * a new struct sched_attr, containing all the necessary fields is
+    provided;
+  * the new scheduling related syscalls that manipulate it, i.e.,
+    sched_setattr() and sched_getattr() are implemented.
+
+ For debugging purposes, the leftover runtime and absolute deadline of a
+ SCHED_DEADLINE task can be retrieved through /proc/<pid>/sched (entries
+ dl.runtime and dl.deadline, both values in ns). A programmatic way to
+ retrieve these values from production code is under discussion.
+
+
+4.3 Default behavior
+---------------------
+
+ The default value for SCHED_DEADLINE bandwidth is to have rt_runtime equal to
+ 950000. With rt_period equal to 1000000, by default, it means that -deadline
+ tasks can use at most 95%, multiplied by the number of CPUs that compose the
+ root_domain, for each root_domain.
+ This means that non -deadline tasks will receive at least 5% of the CPU time,
+ and that -deadline tasks will receive their runtime with a guaranteed
+ worst-case delay respect to the "deadline" parameter. If "deadline" = "period"
+ and the cpuset mechanism is used to implement partitioned scheduling (see
+ Section 5), then this simple setting of the bandwidth management is able to
+ deterministically guarantee that -deadline tasks will receive their runtime
+ in a period.
+
+ Finally, notice that in order not to jeopardize the admission control a
+ -deadline task cannot fork.
+
+
+4.4 Behavior of sched_yield()
+-----------------------------
+
+ When a SCHED_DEADLINE task calls sched_yield(), it gives up its
+ remaining runtime and is immediately throttled, until the next
+ period, when its runtime will be replenished (a special flag
+ dl_yielded is set and used to handle correctly throttling and runtime
+ replenishment after a call to sched_yield()).
+
+ This behavior of sched_yield() allows the task to wake-up exactly at
+ the beginning of the next period. Also, this may be useful in the
+ future with bandwidth reclaiming mechanisms, where sched_yield() will
+ make the leftoever runtime available for reclamation by other
+ SCHED_DEADLINE tasks.
+
+
+5. Tasks CPU affinity
+=====================
+
+ -deadline tasks cannot have an affinity mask smaller that the entire
+ root_domain they are created on. However, affinities can be specified
+ through the cpuset facility (Documentation/cgroup-v1/cpusets.txt).
+
+5.1 SCHED_DEADLINE and cpusets HOWTO
+------------------------------------
+
+ An example of a simple configuration (pin a -deadline task to CPU0)
+ follows (rt-app is used to create a -deadline task).
+
+ mkdir /dev/cpuset
+ mount -t cgroup -o cpuset cpuset /dev/cpuset
+ cd /dev/cpuset
+ mkdir cpu0
+ echo 0 > cpu0/cpuset.cpus
+ echo 0 > cpu0/cpuset.mems
+ echo 1 > cpuset.cpu_exclusive
+ echo 0 > cpuset.sched_load_balance
+ echo 1 > cpu0/cpuset.cpu_exclusive
+ echo 1 > cpu0/cpuset.mem_exclusive
+ echo $$ > cpu0/tasks
+ rt-app -t 100000:10000:d:0 -D5 (it is now actually superfluous to specify
+ task affinity)
+
+6. Future plans
+===============
+
+ Still missing:
+
+  - programmatic way to retrieve current runtime and absolute deadline
+  - refinements to deadline inheritance, especially regarding the possibility
+    of retaining bandwidth isolation among non-interacting tasks. This is
+    being studied from both theoretical and practical points of view, and
+    hopefully we should be able to produce some demonstrative code soon;
+  - (c)group based bandwidth management, and maybe scheduling;
+  - access control for non-root users (and related security concerns to
+    address), which is the best way to allow unprivileged use of the mechanisms
+    and how to prevent non-root users "cheat" the system?
+
+ As already discussed, we are planning also to merge this work with the EDF
+ throttling patches [https://lkml.org/lkml/2010/2/23/239] but we still are in
+ the preliminary phases of the merge and we really seek feedback that would
+ help us decide on the direction it should take.
+
+Appendix A. Test suite
+======================
+
+ The SCHED_DEADLINE policy can be easily tested using two applications that
+ are part of a wider Linux Scheduler validation suite. The suite is
+ available as a GitHub repository: https://github.com/scheduler-tools.
+
+ The first testing application is called rt-app and can be used to
+ start multiple threads with specific parameters. rt-app supports
+ SCHED_{OTHER,FIFO,RR,DEADLINE} scheduling policies and their related
+ parameters (e.g., niceness, priority, runtime/deadline/period). rt-app
+ is a valuable tool, as it can be used to synthetically recreate certain
+ workloads (maybe mimicking real use-cases) and evaluate how the scheduler
+ behaves under such workloads. In this way, results are easily reproducible.
+ rt-app is available at: https://github.com/scheduler-tools/rt-app.
+
+ Thread parameters can be specified from the command line, with something like
+ this:
+
+  # rt-app -t 100000:10000:d -t 150000:20000:f:10 -D5
+
+ The above creates 2 threads. The first one, scheduled by SCHED_DEADLINE,
+ executes for 10ms every 100ms. The second one, scheduled at SCHED_FIFO
+ priority 10, executes for 20ms every 150ms. The test will run for a total
+ of 5 seconds.
+
+ More interestingly, configurations can be described with a json file that
+ can be passed as input to rt-app with something like this:
+
+  # rt-app my_config.json
+
+ The parameters that can be specified with the second method are a superset
+ of the command line options. Please refer to rt-app documentation for more
+ details (<rt-app-sources>/doc/*.json).
+
+ The second testing application is a modification of schedtool, called
+ schedtool-dl, which can be used to setup SCHED_DEADLINE parameters for a
+ certain pid/application. schedtool-dl is available at:
+ https://github.com/scheduler-tools/schedtool-dl.git.
+
+ The usage is straightforward:
+
+  # schedtool -E -t 10000000:100000000 -e ./my_cpuhog_app
+
+ With this, my_cpuhog_app is put to run inside a SCHED_DEADLINE reservation
+ of 10ms every 100ms (note that parameters are expressed in microseconds).
+ You can also use schedtool to create a reservation for an already running
+ application, given that you know its pid:
+
+  # schedtool -E -t 10000000:100000000 my_app_pid
+
+Appendix B. Minimal main()
+==========================
+
+ We provide in what follows a simple (ugly) self-contained code snippet
+ showing how SCHED_DEADLINE reservations can be created by a real-time
+ application developer.
+
+ #define _GNU_SOURCE
+ #include <unistd.h>
+ #include <stdio.h>
+ #include <stdlib.h>
+ #include <string.h>
+ #include <time.h>
+ #include <linux/unistd.h>
+ #include <linux/kernel.h>
+ #include <linux/types.h>
+ #include <sys/syscall.h>
+ #include <pthread.h>
+
+ #define gettid() syscall(__NR_gettid)
+
+ #define SCHED_DEADLINE	6
+
+ /* XXX use the proper syscall numbers */
+ #ifdef __x86_64__
+ #define __NR_sched_setattr		314
+ #define __NR_sched_getattr		315
+ #endif
+
+ #ifdef __i386__
+ #define __NR_sched_setattr		351
+ #define __NR_sched_getattr		352
+ #endif
+
+ #ifdef __arm__
+ #define __NR_sched_setattr		380
+ #define __NR_sched_getattr		381
+ #endif
+
+ static volatile int done;
+
+ struct sched_attr {
+	__u32 size;
+
+	__u32 sched_policy;
+	__u64 sched_flags;
+
+	/* SCHED_NORMAL, SCHED_BATCH */
+	__s32 sched_nice;
+
+	/* SCHED_FIFO, SCHED_RR */
+	__u32 sched_priority;
+
+	/* SCHED_DEADLINE (nsec) */
+	__u64 sched_runtime;
+	__u64 sched_deadline;
+	__u64 sched_period;
+ };
+
+ int sched_setattr(pid_t pid,
+		  const struct sched_attr *attr,
+		  unsigned int flags)
+ {
+	return syscall(__NR_sched_setattr, pid, attr, flags);
+ }
+
+ int sched_getattr(pid_t pid,
+		  struct sched_attr *attr,
+		  unsigned int size,
+		  unsigned int flags)
+ {
+	return syscall(__NR_sched_getattr, pid, attr, size, flags);
+ }
+
+ void *run_deadline(void *data)
+ {
+	struct sched_attr attr;
+	int x = 0;
+	int ret;
+	unsigned int flags = 0;
+
+	printf("deadline thread started [%ld]\n", gettid());
+
+	attr.size = sizeof(attr);
+	attr.sched_flags = 0;
+	attr.sched_nice = 0;
+	attr.sched_priority = 0;
+
+	/* This creates a 10ms/30ms reservation */
+	attr.sched_policy = SCHED_DEADLINE;
+	attr.sched_runtime = 10 * 1000 * 1000;
+	attr.sched_period = attr.sched_deadline = 30 * 1000 * 1000;
+
+	ret = sched_setattr(0, &attr, flags);
+	if (ret < 0) {
+		done = 0;
+		perror("sched_setattr");
+		exit(-1);
+	}
+
+	while (!done) {
+		x++;
+	}
+
+	printf("deadline thread dies [%ld]\n", gettid());
+	return NULL;
+ }
+
+ int main (int argc, char **argv)
+ {
+	pthread_t thread;
+
+	printf("main thread [%ld]\n", gettid());
+
+	pthread_create(&thread, NULL, run_deadline, NULL);
+
+	sleep(10);
+
+	done = 1;
+	pthread_join(thread, NULL);
+
+	printf("main dies [%ld]\n", gettid());
+	return 0;
+ }
diff --git a/Documentation/scheduler/sched-design-CFS.txt b/Documentation/scheduler/sched-design-CFS.txt
new file mode 100644
index 000000000..edd861c94
--- /dev/null
+++ b/Documentation/scheduler/sched-design-CFS.txt
@@ -0,0 +1,242 @@
+                      =============
+                      CFS Scheduler
+                      =============
+
+
+1.  OVERVIEW
+
+CFS stands for "Completely Fair Scheduler," and is the new "desktop" process
+scheduler implemented by Ingo Molnar and merged in Linux 2.6.23.  It is the
+replacement for the previous vanilla scheduler's SCHED_OTHER interactivity
+code.
+
+80% of CFS's design can be summed up in a single sentence: CFS basically models
+an "ideal, precise multi-tasking CPU" on real hardware.
+
+"Ideal multi-tasking CPU" is a (non-existent  :-)) CPU that has 100% physical
+power and which can run each task at precise equal speed, in parallel, each at
+1/nr_running speed.  For example: if there are 2 tasks running, then it runs
+each at 50% physical power --- i.e., actually in parallel.
+
+On real hardware, we can run only a single task at once, so we have to
+introduce the concept of "virtual runtime."  The virtual runtime of a task
+specifies when its next timeslice would start execution on the ideal
+multi-tasking CPU described above.  In practice, the virtual runtime of a task
+is its actual runtime normalized to the total number of running tasks.
+
+
+
+2.  FEW IMPLEMENTATION DETAILS
+
+In CFS the virtual runtime is expressed and tracked via the per-task
+p->se.vruntime (nanosec-unit) value.  This way, it's possible to accurately
+timestamp and measure the "expected CPU time" a task should have gotten.
+
+[ small detail: on "ideal" hardware, at any time all tasks would have the same
+  p->se.vruntime value --- i.e., tasks would execute simultaneously and no task
+  would ever get "out of balance" from the "ideal" share of CPU time.  ]
+
+CFS's task picking logic is based on this p->se.vruntime value and it is thus
+very simple: it always tries to run the task with the smallest p->se.vruntime
+value (i.e., the task which executed least so far).  CFS always tries to split
+up CPU time between runnable tasks as close to "ideal multitasking hardware" as
+possible.
+
+Most of the rest of CFS's design just falls out of this really simple concept,
+with a few add-on embellishments like nice levels, multiprocessing and various
+algorithm variants to recognize sleepers.
+
+
+
+3.  THE RBTREE
+
+CFS's design is quite radical: it does not use the old data structures for the
+runqueues, but it uses a time-ordered rbtree to build a "timeline" of future
+task execution, and thus has no "array switch" artifacts (by which both the
+previous vanilla scheduler and RSDL/SD are affected).
+
+CFS also maintains the rq->cfs.min_vruntime value, which is a monotonic
+increasing value tracking the smallest vruntime among all tasks in the
+runqueue.  The total amount of work done by the system is tracked using
+min_vruntime; that value is used to place newly activated entities on the left
+side of the tree as much as possible.
+
+The total number of running tasks in the runqueue is accounted through the
+rq->cfs.load value, which is the sum of the weights of the tasks queued on the
+runqueue.
+
+CFS maintains a time-ordered rbtree, where all runnable tasks are sorted by the
+p->se.vruntime key. CFS picks the "leftmost" task from this tree and sticks to it.
+As the system progresses forwards, the executed tasks are put into the tree
+more and more to the right --- slowly but surely giving a chance for every task
+to become the "leftmost task" and thus get on the CPU within a deterministic
+amount of time.
+
+Summing up, CFS works like this: it runs a task a bit, and when the task
+schedules (or a scheduler tick happens) the task's CPU usage is "accounted
+for": the (small) time it just spent using the physical CPU is added to
+p->se.vruntime.  Once p->se.vruntime gets high enough so that another task
+becomes the "leftmost task" of the time-ordered rbtree it maintains (plus a
+small amount of "granularity" distance relative to the leftmost task so that we
+do not over-schedule tasks and trash the cache), then the new leftmost task is
+picked and the current task is preempted.
+
+
+
+4.  SOME FEATURES OF CFS
+
+CFS uses nanosecond granularity accounting and does not rely on any jiffies or
+other HZ detail.  Thus the CFS scheduler has no notion of "timeslices" in the
+way the previous scheduler had, and has no heuristics whatsoever.  There is
+only one central tunable (you have to switch on CONFIG_SCHED_DEBUG):
+
+   /proc/sys/kernel/sched_min_granularity_ns
+
+which can be used to tune the scheduler from "desktop" (i.e., low latencies) to
+"server" (i.e., good batching) workloads.  It defaults to a setting suitable
+for desktop workloads.  SCHED_BATCH is handled by the CFS scheduler module too.
+
+Due to its design, the CFS scheduler is not prone to any of the "attacks" that
+exist today against the heuristics of the stock scheduler: fiftyp.c, thud.c,
+chew.c, ring-test.c, massive_intr.c all work fine and do not impact
+interactivity and produce the expected behavior.
+
+The CFS scheduler has a much stronger handling of nice levels and SCHED_BATCH
+than the previous vanilla scheduler: both types of workloads are isolated much
+more aggressively.
+
+SMP load-balancing has been reworked/sanitized: the runqueue-walking
+assumptions are gone from the load-balancing code now, and iterators of the
+scheduling modules are used.  The balancing code got quite a bit simpler as a
+result.
+
+
+
+5. Scheduling policies
+
+CFS implements three scheduling policies:
+
+  - SCHED_NORMAL (traditionally called SCHED_OTHER): The scheduling
+    policy that is used for regular tasks.
+
+  - SCHED_BATCH: Does not preempt nearly as often as regular tasks
+    would, thereby allowing tasks to run longer and make better use of
+    caches but at the cost of interactivity. This is well suited for
+    batch jobs.
+
+  - SCHED_IDLE: This is even weaker than nice 19, but its not a true
+    idle timer scheduler in order to avoid to get into priority
+    inversion problems which would deadlock the machine.
+
+SCHED_FIFO/_RR are implemented in sched/rt.c and are as specified by
+POSIX.
+
+The command chrt from util-linux-ng 2.13.1.1 can set all of these except
+SCHED_IDLE.
+
+
+
+6.  SCHEDULING CLASSES
+
+The new CFS scheduler has been designed in such a way to introduce "Scheduling
+Classes," an extensible hierarchy of scheduler modules.  These modules
+encapsulate scheduling policy details and are handled by the scheduler core
+without the core code assuming too much about them.
+
+sched/fair.c implements the CFS scheduler described above.
+
+sched/rt.c implements SCHED_FIFO and SCHED_RR semantics, in a simpler way than
+the previous vanilla scheduler did.  It uses 100 runqueues (for all 100 RT
+priority levels, instead of 140 in the previous scheduler) and it needs no
+expired array.
+
+Scheduling classes are implemented through the sched_class structure, which
+contains hooks to functions that must be called whenever an interesting event
+occurs.
+
+This is the (partial) list of the hooks:
+
+ - enqueue_task(...)
+
+   Called when a task enters a runnable state.
+   It puts the scheduling entity (task) into the red-black tree and
+   increments the nr_running variable.
+
+ - dequeue_task(...)
+
+   When a task is no longer runnable, this function is called to keep the
+   corresponding scheduling entity out of the red-black tree.  It decrements
+   the nr_running variable.
+
+ - yield_task(...)
+
+   This function is basically just a dequeue followed by an enqueue, unless the
+   compat_yield sysctl is turned on; in that case, it places the scheduling
+   entity at the right-most end of the red-black tree.
+
+ - check_preempt_curr(...)
+
+   This function checks if a task that entered the runnable state should
+   preempt the currently running task.
+
+ - pick_next_task(...)
+
+   This function chooses the most appropriate task eligible to run next.
+
+ - set_curr_task(...)
+
+   This function is called when a task changes its scheduling class or changes
+   its task group.
+
+ - task_tick(...)
+
+   This function is mostly called from time tick functions; it might lead to
+   process switch.  This drives the running preemption.
+
+
+
+
+7.  GROUP SCHEDULER EXTENSIONS TO CFS
+
+Normally, the scheduler operates on individual tasks and strives to provide
+fair CPU time to each task.  Sometimes, it may be desirable to group tasks and
+provide fair CPU time to each such task group.  For example, it may be
+desirable to first provide fair CPU time to each user on the system and then to
+each task belonging to a user.
+
+CONFIG_CGROUP_SCHED strives to achieve exactly that.  It lets tasks to be
+grouped and divides CPU time fairly among such groups.
+
+CONFIG_RT_GROUP_SCHED permits to group real-time (i.e., SCHED_FIFO and
+SCHED_RR) tasks.
+
+CONFIG_FAIR_GROUP_SCHED permits to group CFS (i.e., SCHED_NORMAL and
+SCHED_BATCH) tasks.
+
+   These options need CONFIG_CGROUPS to be defined, and let the administrator
+   create arbitrary groups of tasks, using the "cgroup" pseudo filesystem.  See
+   Documentation/cgroup-v1/cgroups.txt for more information about this filesystem.
+
+When CONFIG_FAIR_GROUP_SCHED is defined, a "cpu.shares" file is created for each
+group created using the pseudo filesystem.  See example steps below to create
+task groups and modify their CPU share using the "cgroups" pseudo filesystem.
+
+	# mount -t tmpfs cgroup_root /sys/fs/cgroup
+	# mkdir /sys/fs/cgroup/cpu
+	# mount -t cgroup -ocpu none /sys/fs/cgroup/cpu
+	# cd /sys/fs/cgroup/cpu
+
+	# mkdir multimedia	# create "multimedia" group of tasks
+	# mkdir browser		# create "browser" group of tasks
+
+	# #Configure the multimedia group to receive twice the CPU bandwidth
+	# #that of browser group
+
+	# echo 2048 > multimedia/cpu.shares
+	# echo 1024 > browser/cpu.shares
+
+	# firefox &	# Launch firefox and move it to "browser" group
+	# echo <firefox_pid> > browser/tasks
+
+	# #Launch gmplayer (or your favourite movie player)
+	# echo <movie_player_pid> > multimedia/tasks
diff --git a/Documentation/scheduler/sched-domains.txt b/Documentation/scheduler/sched-domains.txt
new file mode 100644
index 000000000..4af80b1c0
--- /dev/null
+++ b/Documentation/scheduler/sched-domains.txt
@@ -0,0 +1,77 @@
+Each CPU has a "base" scheduling domain (struct sched_domain). The domain
+hierarchy is built from these base domains via the ->parent pointer. ->parent
+MUST be NULL terminated, and domain structures should be per-CPU as they are
+locklessly updated.
+
+Each scheduling domain spans a number of CPUs (stored in the ->span field).
+A domain's span MUST be a superset of it child's span (this restriction could
+be relaxed if the need arises), and a base domain for CPU i MUST span at least
+i. The top domain for each CPU will generally span all CPUs in the system
+although strictly it doesn't have to, but this could lead to a case where some
+CPUs will never be given tasks to run unless the CPUs allowed mask is
+explicitly set. A sched domain's span means "balance process load among these
+CPUs".
+
+Each scheduling domain must have one or more CPU groups (struct sched_group)
+which are organised as a circular one way linked list from the ->groups
+pointer. The union of cpumasks of these groups MUST be the same as the
+domain's span. The intersection of cpumasks from any two of these groups
+MUST be the empty set. The group pointed to by the ->groups pointer MUST
+contain the CPU to which the domain belongs. Groups may be shared among
+CPUs as they contain read only data after they have been set up.
+
+Balancing within a sched domain occurs between groups. That is, each group
+is treated as one entity. The load of a group is defined as the sum of the
+load of each of its member CPUs, and only when the load of a group becomes
+out of balance are tasks moved between groups.
+
+In kernel/sched/core.c, trigger_load_balance() is run periodically on each CPU
+through scheduler_tick(). It raises a softirq after the next regularly scheduled
+rebalancing event for the current runqueue has arrived. The actual load
+balancing workhorse, run_rebalance_domains()->rebalance_domains(), is then run
+in softirq context (SCHED_SOFTIRQ).
+
+The latter function takes two arguments: the current CPU and whether it was idle
+at the time the scheduler_tick() happened and iterates over all sched domains
+our CPU is on, starting from its base domain and going up the ->parent chain.
+While doing that, it checks to see if the current domain has exhausted its
+rebalance interval. If so, it runs load_balance() on that domain. It then checks
+the parent sched_domain (if it exists), and the parent of the parent and so
+forth.
+
+Initially, load_balance() finds the busiest group in the current sched domain.
+If it succeeds, it looks for the busiest runqueue of all the CPUs' runqueues in
+that group. If it manages to find such a runqueue, it locks both our initial
+CPU's runqueue and the newly found busiest one and starts moving tasks from it
+to our runqueue. The exact number of tasks amounts to an imbalance previously
+computed while iterating over this sched domain's groups.
+
+*** Implementing sched domains ***
+The "base" domain will "span" the first level of the hierarchy. In the case
+of SMT, you'll span all siblings of the physical CPU, with each group being
+a single virtual CPU.
+
+In SMP, the parent of the base domain will span all physical CPUs in the
+node. Each group being a single physical CPU. Then with NUMA, the parent
+of the SMP domain will span the entire machine, with each group having the
+cpumask of a node. Or, you could do multi-level NUMA or Opteron, for example,
+might have just one domain covering its one NUMA level.
+
+The implementor should read comments in include/linux/sched.h:
+struct sched_domain fields, SD_FLAG_*, SD_*_INIT to get an idea of
+the specifics and what to tune.
+
+Architectures may retain the regular override the default SD_*_INIT flags
+while using the generic domain builder in kernel/sched/core.c if they wish to
+retain the traditional SMT->SMP->NUMA topology (or some subset of that). This
+can be done by #define'ing ARCH_HASH_SCHED_TUNE.
+
+Alternatively, the architecture may completely override the generic domain
+builder by #define'ing ARCH_HASH_SCHED_DOMAIN, and exporting your
+arch_init_sched_domains function. This function will attach domains to all
+CPUs using cpu_attach_domain.
+
+The sched-domains debugging infrastructure can be enabled by enabling
+CONFIG_SCHED_DEBUG. This enables an error checking parse of the sched domains
+which should catch most possible errors (described above). It also prints out
+the domain structure in a visual format.
diff --git a/Documentation/scheduler/sched-nice-design.txt b/Documentation/scheduler/sched-nice-design.txt
new file mode 100644
index 000000000..3ac1e46d5
--- /dev/null
+++ b/Documentation/scheduler/sched-nice-design.txt
@@ -0,0 +1,108 @@
+This document explains the thinking about the revamped and streamlined
+nice-levels implementation in the new Linux scheduler.
+
+Nice levels were always pretty weak under Linux and people continuously
+pestered us to make nice +19 tasks use up much less CPU time.
+
+Unfortunately that was not that easy to implement under the old
+scheduler, (otherwise we'd have done it long ago) because nice level
+support was historically coupled to timeslice length, and timeslice
+units were driven by the HZ tick, so the smallest timeslice was 1/HZ.
+
+In the O(1) scheduler (in 2003) we changed negative nice levels to be
+much stronger than they were before in 2.4 (and people were happy about
+that change), and we also intentionally calibrated the linear timeslice
+rule so that nice +19 level would be _exactly_ 1 jiffy. To better
+understand it, the timeslice graph went like this (cheesy ASCII art
+alert!):
+
+
+                   A
+             \     | [timeslice length]
+              \    |
+               \   |
+                \  |
+                 \ |
+                  \|___100msecs
+                   |^ . _
+                   |      ^ . _
+                   |            ^ . _
+ -*----------------------------------*-----> [nice level]
+ -20               |                +19
+                   |
+                   |
+
+So that if someone wanted to really renice tasks, +19 would give a much
+bigger hit than the normal linear rule would do. (The solution of
+changing the ABI to extend priorities was discarded early on.)
+
+This approach worked to some degree for some time, but later on with
+HZ=1000 it caused 1 jiffy to be 1 msec, which meant 0.1% CPU usage which
+we felt to be a bit excessive. Excessive _not_ because it's too small of
+a CPU utilization, but because it causes too frequent (once per
+millisec) rescheduling. (and would thus trash the cache, etc. Remember,
+this was long ago when hardware was weaker and caches were smaller, and
+people were running number crunching apps at nice +19.)
+
+So for HZ=1000 we changed nice +19 to 5msecs, because that felt like the
+right minimal granularity - and this translates to 5% CPU utilization.
+But the fundamental HZ-sensitive property for nice+19 still remained,
+and we never got a single complaint about nice +19 being too _weak_ in
+terms of CPU utilization, we only got complaints about it (still) being
+too _strong_ :-)
+
+To sum it up: we always wanted to make nice levels more consistent, but
+within the constraints of HZ and jiffies and their nasty design level
+coupling to timeslices and granularity it was not really viable.
+
+The second (less frequent but still periodically occurring) complaint
+about Linux's nice level support was its assymetry around the origo
+(which you can see demonstrated in the picture above), or more
+accurately: the fact that nice level behavior depended on the _absolute_
+nice level as well, while the nice API itself is fundamentally
+"relative":
+
+   int nice(int inc);
+
+   asmlinkage long sys_nice(int increment)
+
+(the first one is the glibc API, the second one is the syscall API.)
+Note that the 'inc' is relative to the current nice level. Tools like
+bash's "nice" command mirror this relative API.
+
+With the old scheduler, if you for example started a niced task with +1
+and another task with +2, the CPU split between the two tasks would
+depend on the nice level of the parent shell - if it was at nice -10 the
+CPU split was different than if it was at +5 or +10.
+
+A third complaint against Linux's nice level support was that negative
+nice levels were not 'punchy enough', so lots of people had to resort to
+run audio (and other multimedia) apps under RT priorities such as
+SCHED_FIFO. But this caused other problems: SCHED_FIFO is not starvation
+proof, and a buggy SCHED_FIFO app can also lock up the system for good.
+
+The new scheduler in v2.6.23 addresses all three types of complaints:
+
+To address the first complaint (of nice levels being not "punchy"
+enough), the scheduler was decoupled from 'time slice' and HZ concepts
+(and granularity was made a separate concept from nice levels) and thus
+it was possible to implement better and more consistent nice +19
+support: with the new scheduler nice +19 tasks get a HZ-independent
+1.5%, instead of the variable 3%-5%-9% range they got in the old
+scheduler.
+
+To address the second complaint (of nice levels not being consistent),
+the new scheduler makes nice(1) have the same CPU utilization effect on
+tasks, regardless of their absolute nice levels. So on the new
+scheduler, running a nice +10 and a nice 11 task has the same CPU
+utilization "split" between them as running a nice -5 and a nice -4
+task. (one will get 55% of the CPU, the other 45%.) That is why nice
+levels were changed to be "multiplicative" (or exponential) - that way
+it does not matter which nice level you start out from, the 'relative
+result' will always be the same.
+
+The third complaint (of negative nice levels not being "punchy" enough
+and forcing audio apps to run under the more dangerous SCHED_FIFO
+scheduling policy) is addressed by the new scheduler almost
+automatically: stronger negative nice levels are an automatic
+side-effect of the recalibrated dynamic range of nice levels.
diff --git a/Documentation/scheduler/sched-pelt.c b/Documentation/scheduler/sched-pelt.c
new file mode 100644
index 000000000..7238b3559
--- /dev/null
+++ b/Documentation/scheduler/sched-pelt.c
@@ -0,0 +1,109 @@
+/*
+ * The following program is used to generate the constants for
+ * computing sched averages.
+ *
+ * ==============================================================
+ *		C program (compile with -lm)
+ * ==============================================================
+ */
+
+#include <math.h>
+#include <stdio.h>
+
+#define HALFLIFE 32
+#define SHIFT 32
+
+double y;
+
+void calc_runnable_avg_yN_inv(void)
+{
+	int i;
+	unsigned int x;
+
+	/* To silence -Wunused-but-set-variable warnings. */
+	printf("static const u32 runnable_avg_yN_inv[] __maybe_unused = {");
+	for (i = 0; i < HALFLIFE; i++) {
+		x = ((1UL<<32)-1)*pow(y, i);
+
+		if (i % 6 == 0) printf("\n\t");
+		printf("0x%8x, ", x);
+	}
+	printf("\n};\n\n");
+}
+
+int sum = 1024;
+
+void calc_runnable_avg_yN_sum(void)
+{
+	int i;
+
+	printf("static const u32 runnable_avg_yN_sum[] = {\n\t    0,");
+	for (i = 1; i <= HALFLIFE; i++) {
+		if (i == 1)
+			sum *= y;
+		else
+			sum = sum*y + 1024*y;
+
+		if (i % 11 == 0)
+			printf("\n\t");
+
+		printf("%5d,", sum);
+	}
+	printf("\n};\n\n");
+}
+
+int n = -1;
+/* first period */
+long max = 1024;
+
+void calc_converged_max(void)
+{
+	long last = 0, y_inv = ((1UL<<32)-1)*y;
+
+	for (; ; n++) {
+		if (n > -1)
+			max = ((max*y_inv)>>SHIFT) + 1024;
+			/*
+			 * This is the same as:
+			 * max = max*y + 1024;
+			 */
+
+		if (last == max)
+			break;
+
+		last = max;
+	}
+	n--;
+	printf("#define LOAD_AVG_PERIOD %d\n", HALFLIFE);
+	printf("#define LOAD_AVG_MAX %ld\n", max);
+//	printf("#define LOAD_AVG_MAX_N %d\n\n", n);
+}
+
+void calc_accumulated_sum_32(void)
+{
+	int i, x = sum;
+
+	printf("static const u32 __accumulated_sum_N32[] = {\n\t     0,");
+	for (i = 1; i <= n/HALFLIFE+1; i++) {
+		if (i > 1)
+			x = x/2 + sum;
+
+		if (i % 6 == 0)
+			printf("\n\t");
+
+		printf("%6d,", x);
+	}
+	printf("\n};\n\n");
+}
+
+void main(void)
+{
+	printf("/* Generated by Documentation/scheduler/sched-pelt; do not modify. */\n\n");
+
+	y = pow(0.5, 1/(double)HALFLIFE);
+
+	calc_runnable_avg_yN_inv();
+//	calc_runnable_avg_yN_sum();
+	calc_converged_max();
+//	calc_accumulated_sum_32();
+}
diff --git a/Documentation/scheduler/sched-rt-group.txt b/Documentation/scheduler/sched-rt-group.txt
new file mode 100644
index 000000000..d8fce3e78
--- /dev/null
+++ b/Documentation/scheduler/sched-rt-group.txt
@@ -0,0 +1,183 @@
+				Real-Time group scheduling
+				--------------------------
+
+CONTENTS
+========
+
+0. WARNING
+1. Overview
+  1.1 The problem
+  1.2 The solution
+2. The interface
+  2.1 System-wide settings
+  2.2 Default behaviour
+  2.3 Basis for grouping tasks
+3. Future plans
+
+
+0. WARNING
+==========
+
+ Fiddling with these settings can result in an unstable system, the knobs are
+ root only and assumes root knows what he is doing.
+
+Most notable:
+
+ * very small values in sched_rt_period_us can result in an unstable
+   system when the period is smaller than either the available hrtimer
+   resolution, or the time it takes to handle the budget refresh itself.
+
+ * very small values in sched_rt_runtime_us can result in an unstable
+   system when the runtime is so small the system has difficulty making
+   forward progress (NOTE: the migration thread and kstopmachine both
+   are real-time processes).
+
+1. Overview
+===========
+
+
+1.1 The problem
+---------------
+
+Realtime scheduling is all about determinism, a group has to be able to rely on
+the amount of bandwidth (eg. CPU time) being constant. In order to schedule
+multiple groups of realtime tasks, each group must be assigned a fixed portion
+of the CPU time available.  Without a minimum guarantee a realtime group can
+obviously fall short. A fuzzy upper limit is of no use since it cannot be
+relied upon. Which leaves us with just the single fixed portion.
+
+1.2 The solution
+----------------
+
+CPU time is divided by means of specifying how much time can be spent running
+in a given period. We allocate this "run time" for each realtime group which
+the other realtime groups will not be permitted to use.
+
+Any time not allocated to a realtime group will be used to run normal priority
+tasks (SCHED_OTHER). Any allocated run time not used will also be picked up by
+SCHED_OTHER.
+
+Let's consider an example: a frame fixed realtime renderer must deliver 25
+frames a second, which yields a period of 0.04s per frame. Now say it will also
+have to play some music and respond to input, leaving it with around 80% CPU
+time dedicated for the graphics. We can then give this group a run time of 0.8
+* 0.04s = 0.032s.
+
+This way the graphics group will have a 0.04s period with a 0.032s run time
+limit. Now if the audio thread needs to refill the DMA buffer every 0.005s, but
+needs only about 3% CPU time to do so, it can do with a 0.03 * 0.005s =
+0.00015s. So this group can be scheduled with a period of 0.005s and a run time
+of 0.00015s.
+
+The remaining CPU time will be used for user input and other tasks. Because
+realtime tasks have explicitly allocated the CPU time they need to perform
+their tasks, buffer underruns in the graphics or audio can be eliminated.
+
+NOTE: the above example is not fully implemented yet. We still
+lack an EDF scheduler to make non-uniform periods usable.
+
+
+2. The Interface
+================
+
+
+2.1 System wide settings
+------------------------
+
+The system wide settings are configured under the /proc virtual file system:
+
+/proc/sys/kernel/sched_rt_period_us:
+  The scheduling period that is equivalent to 100% CPU bandwidth
+
+/proc/sys/kernel/sched_rt_runtime_us:
+  A global limit on how much time realtime scheduling may use.  Even without
+  CONFIG_RT_GROUP_SCHED enabled, this will limit time reserved to realtime
+  processes. With CONFIG_RT_GROUP_SCHED it signifies the total bandwidth
+  available to all realtime groups.
+
+  * Time is specified in us because the interface is s32. This gives an
+    operating range from 1us to about 35 minutes.
+  * sched_rt_period_us takes values from 1 to INT_MAX.
+  * sched_rt_runtime_us takes values from -1 to (INT_MAX - 1).
+  * A run time of -1 specifies runtime == period, ie. no limit.
+
+
+2.2 Default behaviour
+---------------------
+
+The default values for sched_rt_period_us (1000000 or 1s) and
+sched_rt_runtime_us (950000 or 0.95s).  This gives 0.05s to be used by
+SCHED_OTHER (non-RT tasks). These defaults were chosen so that a run-away
+realtime tasks will not lock up the machine but leave a little time to recover
+it.  By setting runtime to -1 you'd get the old behaviour back.
+
+By default all bandwidth is assigned to the root group and new groups get the
+period from /proc/sys/kernel/sched_rt_period_us and a run time of 0. If you
+want to assign bandwidth to another group, reduce the root group's bandwidth
+and assign some or all of the difference to another group.
+
+Realtime group scheduling means you have to assign a portion of total CPU
+bandwidth to the group before it will accept realtime tasks. Therefore you will
+not be able to run realtime tasks as any user other than root until you have
+done that, even if the user has the rights to run processes with realtime
+priority!
+
+
+2.3 Basis for grouping tasks
+----------------------------
+
+Enabling CONFIG_RT_GROUP_SCHED lets you explicitly allocate real
+CPU bandwidth to task groups.
+
+This uses the cgroup virtual file system and "<cgroup>/cpu.rt_runtime_us"
+to control the CPU time reserved for each control group.
+
+For more information on working with control groups, you should read
+Documentation/cgroup-v1/cgroups.txt as well.
+
+Group settings are checked against the following limits in order to keep the
+configuration schedulable:
+
+   \Sum_{i} runtime_{i} / global_period <= global_runtime / global_period
+
+For now, this can be simplified to just the following (but see Future plans):
+
+   \Sum_{i} runtime_{i} <= global_runtime
+
+
+3. Future plans
+===============
+
+There is work in progress to make the scheduling period for each group
+("<cgroup>/cpu.rt_period_us") configurable as well.
+
+The constraint on the period is that a subgroup must have a smaller or
+equal period to its parent. But realistically its not very useful _yet_
+as its prone to starvation without deadline scheduling.
+
+Consider two sibling groups A and B; both have 50% bandwidth, but A's
+period is twice the length of B's.
+
+* group A: period=100000us, runtime=50000us
+	- this runs for 0.05s once every 0.1s
+
+* group B: period= 50000us, runtime=25000us
+	- this runs for 0.025s twice every 0.1s (or once every 0.05 sec).
+
+This means that currently a while (1) loop in A will run for the full period of
+B and can starve B's tasks (assuming they are of lower priority) for a whole
+period.
+
+The next project will be SCHED_EDF (Earliest Deadline First scheduling) to bring
+full deadline scheduling to the linux kernel. Deadline scheduling the above
+groups and treating end of the period as a deadline will ensure that they both
+get their allocated time.
+
+Implementing SCHED_EDF might take a while to complete. Priority Inheritance is
+the biggest challenge as the current linux PI infrastructure is geared towards
+the limited static priority levels 0-99. With deadline scheduling you need to
+do deadline inheritance (since priority is inversely proportional to the
+deadline delta (deadline - now)).
+
+This means the whole PI machinery will have to be reworked - and that is one of
+the most complex pieces of code we have.
diff --git a/Documentation/scheduler/sched-stats.txt b/Documentation/scheduler/sched-stats.txt
new file mode 100644
index 000000000..8259b34a6
--- /dev/null
+++ b/Documentation/scheduler/sched-stats.txt
@@ -0,0 +1,154 @@
+Version 15 of schedstats dropped counters for some sched_yield:
+yld_exp_empty, yld_act_empty and yld_both_empty. Otherwise, it is
+identical to version 14.
+
+Version 14 of schedstats includes support for sched_domains, which hit the
+mainline kernel in 2.6.20 although it is identical to the stats from version
+12 which was in the kernel from 2.6.13-2.6.19 (version 13 never saw a kernel
+release).  Some counters make more sense to be per-runqueue; other to be
+per-domain.  Note that domains (and their associated information) will only
+be pertinent and available on machines utilizing CONFIG_SMP.
+
+In version 14 of schedstat, there is at least one level of domain
+statistics for each cpu listed, and there may well be more than one
+domain.  Domains have no particular names in this implementation, but
+the highest numbered one typically arbitrates balancing across all the
+cpus on the machine, while domain0 is the most tightly focused domain,
+sometimes balancing only between pairs of cpus.  At this time, there
+are no architectures which need more than three domain levels. The first
+field in the domain stats is a bit map indicating which cpus are affected
+by that domain.
+
+These fields are counters, and only increment.  Programs which make use
+of these will need to start with a baseline observation and then calculate
+the change in the counters at each subsequent observation.  A perl script
+which does this for many of the fields is available at
+
+    http://eaglet.rain.com/rick/linux/schedstat/
+
+Note that any such script will necessarily be version-specific, as the main
+reason to change versions is changes in the output format.  For those wishing
+to write their own scripts, the fields are described here.
+
+CPU statistics
+--------------
+cpu<N> 1 2 3 4 5 6 7 8 9
+
+First field is a sched_yield() statistic:
+     1) # of times sched_yield() was called
+
+Next three are schedule() statistics:
+     2) This field is a legacy array expiration count field used in the O(1)
+	scheduler. We kept it for ABI compatibility, but it is always set to zero.
+     3) # of times schedule() was called
+     4) # of times schedule() left the processor idle
+
+Next two are try_to_wake_up() statistics:
+     5) # of times try_to_wake_up() was called
+     6) # of times try_to_wake_up() was called to wake up the local cpu
+
+Next three are statistics describing scheduling latency:
+     7) sum of all time spent running by tasks on this processor (in jiffies)
+     8) sum of all time spent waiting to run by tasks on this processor (in
+        jiffies)
+     9) # of timeslices run on this cpu
+
+
+Domain statistics
+-----------------
+One of these is produced per domain for each cpu described. (Note that if
+CONFIG_SMP is not defined, *no* domains are utilized and these lines
+will not appear in the output.)
+
+domain<N> <cpumask> 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25 26 27 28 29 30 31 32 33 34 35 36
+
+The first field is a bit mask indicating what cpus this domain operates over.
+
+The next 24 are a variety of load_balance() statistics in grouped into types
+of idleness (idle, busy, and newly idle):
+
+     1) # of times in this domain load_balance() was called when the
+        cpu was idle
+     2) # of times in this domain load_balance() checked but found
+        the load did not require balancing when the cpu was idle
+     3) # of times in this domain load_balance() tried to move one or
+        more tasks and failed, when the cpu was idle
+     4) sum of imbalances discovered (if any) with each call to
+        load_balance() in this domain when the cpu was idle
+     5) # of times in this domain pull_task() was called when the cpu
+        was idle
+     6) # of times in this domain pull_task() was called even though
+        the target task was cache-hot when idle
+     7) # of times in this domain load_balance() was called but did
+        not find a busier queue while the cpu was idle
+     8) # of times in this domain a busier queue was found while the
+        cpu was idle but no busier group was found
+
+     9) # of times in this domain load_balance() was called when the
+        cpu was busy
+    10) # of times in this domain load_balance() checked but found the
+        load did not require balancing when busy
+    11) # of times in this domain load_balance() tried to move one or
+        more tasks and failed, when the cpu was busy
+    12) sum of imbalances discovered (if any) with each call to
+        load_balance() in this domain when the cpu was busy
+    13) # of times in this domain pull_task() was called when busy
+    14) # of times in this domain pull_task() was called even though the
+        target task was cache-hot when busy
+    15) # of times in this domain load_balance() was called but did not
+        find a busier queue while the cpu was busy
+    16) # of times in this domain a busier queue was found while the cpu
+        was busy but no busier group was found
+
+    17) # of times in this domain load_balance() was called when the
+        cpu was just becoming idle
+    18) # of times in this domain load_balance() checked but found the
+        load did not require balancing when the cpu was just becoming idle
+    19) # of times in this domain load_balance() tried to move one or more
+        tasks and failed, when the cpu was just becoming idle
+    20) sum of imbalances discovered (if any) with each call to
+        load_balance() in this domain when the cpu was just becoming idle
+    21) # of times in this domain pull_task() was called when newly idle
+    22) # of times in this domain pull_task() was called even though the
+        target task was cache-hot when just becoming idle
+    23) # of times in this domain load_balance() was called but did not
+        find a busier queue while the cpu was just becoming idle
+    24) # of times in this domain a busier queue was found while the cpu
+        was just becoming idle but no busier group was found
+
+   Next three are active_load_balance() statistics:
+    25) # of times active_load_balance() was called
+    26) # of times active_load_balance() tried to move a task and failed
+    27) # of times active_load_balance() successfully moved a task
+
+   Next three are sched_balance_exec() statistics:
+    28) sbe_cnt is not used
+    29) sbe_balanced is not used
+    30) sbe_pushed is not used
+
+   Next three are sched_balance_fork() statistics:
+    31) sbf_cnt is not used
+    32) sbf_balanced is not used
+    33) sbf_pushed is not used
+
+   Next three are try_to_wake_up() statistics:
+    34) # of times in this domain try_to_wake_up() awoke a task that
+        last ran on a different cpu in this domain
+    35) # of times in this domain try_to_wake_up() moved a task to the
+        waking cpu because it was cache-cold on its own cpu anyway
+    36) # of times in this domain try_to_wake_up() started passive balancing
+
+/proc/<pid>/schedstat
+----------------
+schedstats also adds a new /proc/<pid>/schedstat file to include some of
+the same information on a per-process level.  There are three fields in
+this file correlating for that process to:
+     1) time spent on the cpu
+     2) time spent waiting on a runqueue
+     3) # of timeslices run on this cpu
+
+A program could be easily written to make use of these extra fields to
+report on how well a particular process or set of processes is faring
+under the scheduler's policies.  A simple version of such a program is
+available at
+    http://eaglet.rain.com/rick/linux/schedstat/v12/latency.c