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linux/Documentation/kernel-hacking/locking.rst
Alexander Sverdlin 379af13b31 docs: locking: Discourage from calling disable_irq() in atomic 
Correct the example in the documentation so that disable_irq() is not being
called in atomic context.

disable_irq() calls sleeping synchronize_irq(), it's not allowed to call
them in atomic context.

Signed-off-by: Alexander Sverdlin <alexander.sverdlin@siemens.com>
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
Reviewed-by: Manfred Spraul <manfred@colorfullife.com>
Cc: linux-doc@vger.kernel.org
Link: https://lore.kernel.org/lkml/87k02wbs2n.ffs@tglx/
Link: https://lore.kernel.org/r/20221212163715.830315-1-alexander.sverdlin@siemens.com
2023-01-11 19:45:26 +01:00

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.. _kernel_hacking_lock:
===========================
Unreliable Guide To Locking
===========================
:Author: Rusty Russell
Introduction
============
Welcome, to Rusty's Remarkably Unreliable Guide to Kernel Locking
issues. This document describes the locking systems in the Linux Kernel
in 2.6.
With the wide availability of HyperThreading, and preemption in the
Linux Kernel, everyone hacking on the kernel needs to know the
fundamentals of concurrency and locking for SMP.
The Problem With Concurrency
============================
(Skip this if you know what a Race Condition is).
In a normal program, you can increment a counter like so:
::
very_important_count++;
This is what they would expect to happen:
.. table:: Expected Results
+------------------------------------+------------------------------------+
| Instance 1 | Instance 2 |
+====================================+====================================+
| read very_important_count (5) | |
+------------------------------------+------------------------------------+
| add 1 (6) | |
+------------------------------------+------------------------------------+
| write very_important_count (6) | |
+------------------------------------+------------------------------------+
| | read very_important_count (6) |
+------------------------------------+------------------------------------+
| | add 1 (7) |
+------------------------------------+------------------------------------+
| | write very_important_count (7) |
+------------------------------------+------------------------------------+
This is what might happen:
.. table:: Possible Results
+------------------------------------+------------------------------------+
| Instance 1 | Instance 2 |
+====================================+====================================+
| read very_important_count (5) | |
+------------------------------------+------------------------------------+
| | read very_important_count (5) |
+------------------------------------+------------------------------------+
| add 1 (6) | |
+------------------------------------+------------------------------------+
| | add 1 (6) |
+------------------------------------+------------------------------------+
| write very_important_count (6) | |
+------------------------------------+------------------------------------+
| | write very_important_count (6) |
+------------------------------------+------------------------------------+
Race Conditions and Critical Regions
------------------------------------
This overlap, where the result depends on the relative timing of
multiple tasks, is called a race condition. The piece of code containing
the concurrency issue is called a critical region. And especially since
Linux starting running on SMP machines, they became one of the major
issues in kernel design and implementation.
Preemption can have the same effect, even if there is only one CPU: by
preempting one task during the critical region, we have exactly the same
race condition. In this case the thread which preempts might run the
critical region itself.
The solution is to recognize when these simultaneous accesses occur, and
use locks to make sure that only one instance can enter the critical
region at any time. There are many friendly primitives in the Linux
kernel to help you do this. And then there are the unfriendly
primitives, but I'll pretend they don't exist.
Locking in the Linux Kernel
===========================
If I could give you one piece of advice on locking: **keep it simple**.
Be reluctant to introduce new locks.
Two Main Types of Kernel Locks: Spinlocks and Mutexes
-----------------------------------------------------
There are two main types of kernel locks. The fundamental type is the
spinlock (``include/asm/spinlock.h``), which is a very simple
single-holder lock: if you can't get the spinlock, you keep trying
(spinning) until you can. Spinlocks are very small and fast, and can be
used anywhere.
The second type is a mutex (``include/linux/mutex.h``): it is like a
spinlock, but you may block holding a mutex. If you can't lock a mutex,
your task will suspend itself, and be woken up when the mutex is
released. This means the CPU can do something else while you are
waiting. There are many cases when you simply can't sleep (see
`What Functions Are Safe To Call From Interrupts?`_),
and so have to use a spinlock instead.
Neither type of lock is recursive: see
`Deadlock: Simple and Advanced`_.
Locks and Uniprocessor Kernels
------------------------------
For kernels compiled without ``CONFIG_SMP``, and without
``CONFIG_PREEMPT`` spinlocks do not exist at all. This is an excellent
design decision: when no-one else can run at the same time, there is no
reason to have a lock.
If the kernel is compiled without ``CONFIG_SMP``, but ``CONFIG_PREEMPT``
is set, then spinlocks simply disable preemption, which is sufficient to
prevent any races. For most purposes, we can think of preemption as
equivalent to SMP, and not worry about it separately.
You should always test your locking code with ``CONFIG_SMP`` and
``CONFIG_PREEMPT`` enabled, even if you don't have an SMP test box,
because it will still catch some kinds of locking bugs.
Mutexes still exist, because they are required for synchronization
between user contexts, as we will see below.
Locking Only In User Context
----------------------------
If you have a data structure which is only ever accessed from user
context, then you can use a simple mutex (``include/linux/mutex.h``) to
protect it. This is the most trivial case: you initialize the mutex.
Then you can call mutex_lock_interruptible() to grab the
mutex, and mutex_unlock() to release it. There is also a
mutex_lock(), which should be avoided, because it will
not return if a signal is received.
Example: ``net/netfilter/nf_sockopt.c`` allows registration of new
setsockopt() and getsockopt() calls, with
nf_register_sockopt(). Registration and de-registration
are only done on module load and unload (and boot time, where there is
no concurrency), and the list of registrations is only consulted for an
unknown setsockopt() or getsockopt() system
call. The ``nf_sockopt_mutex`` is perfect to protect this, especially
since the setsockopt and getsockopt calls may well sleep.
Locking Between User Context and Softirqs
-----------------------------------------
If a softirq shares data with user context, you have two problems.
Firstly, the current user context can be interrupted by a softirq, and
secondly, the critical region could be entered from another CPU. This is
where spin_lock_bh() (``include/linux/spinlock.h``) is
used. It disables softirqs on that CPU, then grabs the lock.
spin_unlock_bh() does the reverse. (The '_bh' suffix is
a historical reference to "Bottom Halves", the old name for software
interrupts. It should really be called spin_lock_softirq()' in a
perfect world).
Note that you can also use spin_lock_irq() or
spin_lock_irqsave() here, which stop hardware interrupts
as well: see `Hard IRQ Context`_.
This works perfectly for UP as well: the spin lock vanishes, and this
macro simply becomes local_bh_disable()
(``include/linux/interrupt.h``), which protects you from the softirq
being run.
Locking Between User Context and Tasklets
-----------------------------------------
This is exactly the same as above, because tasklets are actually run
from a softirq.
Locking Between User Context and Timers
---------------------------------------
This, too, is exactly the same as above, because timers are actually run
from a softirq. From a locking point of view, tasklets and timers are
identical.
Locking Between Tasklets/Timers
-------------------------------
Sometimes a tasklet or timer might want to share data with another
tasklet or timer.
The Same Tasklet/Timer
~~~~~~~~~~~~~~~~~~~~~~
Since a tasklet is never run on two CPUs at once, you don't need to
worry about your tasklet being reentrant (running twice at once), even
on SMP.
Different Tasklets/Timers
~~~~~~~~~~~~~~~~~~~~~~~~~
If another tasklet/timer wants to share data with your tasklet or timer
, you will both need to use spin_lock() and
spin_unlock() calls. spin_lock_bh() is
unnecessary here, as you are already in a tasklet, and none will be run
on the same CPU.
Locking Between Softirqs
------------------------
Often a softirq might want to share data with itself or a tasklet/timer.
The Same Softirq
~~~~~~~~~~~~~~~~
The same softirq can run on the other CPUs: you can use a per-CPU array
(see `Per-CPU Data`_) for better performance. If you're
going so far as to use a softirq, you probably care about scalable
performance enough to justify the extra complexity.
You'll need to use spin_lock() and
spin_unlock() for shared data.
Different Softirqs
~~~~~~~~~~~~~~~~~~
You'll need to use spin_lock() and
spin_unlock() for shared data, whether it be a timer,
tasklet, different softirq or the same or another softirq: any of them
could be running on a different CPU.
Hard IRQ Context
================
Hardware interrupts usually communicate with a tasklet or softirq.
Frequently this involves putting work in a queue, which the softirq will
take out.
Locking Between Hard IRQ and Softirqs/Tasklets
----------------------------------------------
If a hardware irq handler shares data with a softirq, you have two
concerns. Firstly, the softirq processing can be interrupted by a
hardware interrupt, and secondly, the critical region could be entered
by a hardware interrupt on another CPU. This is where
spin_lock_irq() is used. It is defined to disable
interrupts on that cpu, then grab the lock.
spin_unlock_irq() does the reverse.
The irq handler does not need to use spin_lock_irq(), because
the softirq cannot run while the irq handler is running: it can use
spin_lock(), which is slightly faster. The only exception
would be if a different hardware irq handler uses the same lock:
spin_lock_irq() will stop that from interrupting us.
This works perfectly for UP as well: the spin lock vanishes, and this
macro simply becomes local_irq_disable()
(``include/asm/smp.h``), which protects you from the softirq/tasklet/BH
being run.
spin_lock_irqsave() (``include/linux/spinlock.h``) is a
variant which saves whether interrupts were on or off in a flags word,
which is passed to spin_unlock_irqrestore(). This means
that the same code can be used inside an hard irq handler (where
interrupts are already off) and in softirqs (where the irq disabling is
required).
Note that softirqs (and hence tasklets and timers) are run on return
from hardware interrupts, so spin_lock_irq() also stops
these. In that sense, spin_lock_irqsave() is the most
general and powerful locking function.
Locking Between Two Hard IRQ Handlers
-------------------------------------
It is rare to have to share data between two IRQ handlers, but if you
do, spin_lock_irqsave() should be used: it is
architecture-specific whether all interrupts are disabled inside irq
handlers themselves.
Cheat Sheet For Locking
=======================
Pete Zaitcev gives the following summary:
- If you are in a process context (any syscall) and want to lock other
process out, use a mutex. You can take a mutex and sleep
(``copy_from_user()`` or ``kmalloc(x,GFP_KERNEL)``).
- Otherwise (== data can be touched in an interrupt), use
spin_lock_irqsave() and
spin_unlock_irqrestore().
- Avoid holding spinlock for more than 5 lines of code and across any
function call (except accessors like readb()).
Table of Minimum Requirements
-----------------------------
The following table lists the **minimum** locking requirements between
various contexts. In some cases, the same context can only be running on
one CPU at a time, so no locking is required for that context (eg. a
particular thread can only run on one CPU at a time, but if it needs
shares data with another thread, locking is required).
Remember the advice above: you can always use
spin_lock_irqsave(), which is a superset of all other
spinlock primitives.
============== ============= ============= ========= ========= ========= ========= ======= ======= ============== ==============
. IRQ Handler A IRQ Handler B Softirq A Softirq B Tasklet A Tasklet B Timer A Timer B User Context A User Context B
============== ============= ============= ========= ========= ========= ========= ======= ======= ============== ==============
IRQ Handler A None
IRQ Handler B SLIS None
Softirq A SLI SLI SL
Softirq B SLI SLI SL SL
Tasklet A SLI SLI SL SL None
Tasklet B SLI SLI SL SL SL None
Timer A SLI SLI SL SL SL SL None
Timer B SLI SLI SL SL SL SL SL None
User Context A SLI SLI SLBH SLBH SLBH SLBH SLBH SLBH None
User Context B SLI SLI SLBH SLBH SLBH SLBH SLBH SLBH MLI None
============== ============= ============= ========= ========= ========= ========= ======= ======= ============== ==============
Table: Table of Locking Requirements
+--------+----------------------------+
| SLIS | spin_lock_irqsave |
+--------+----------------------------+
| SLI | spin_lock_irq |
+--------+----------------------------+
| SL | spin_lock |
+--------+----------------------------+
| SLBH | spin_lock_bh |
+--------+----------------------------+
| MLI | mutex_lock_interruptible |
+--------+----------------------------+
Table: Legend for Locking Requirements Table
The trylock Functions
=====================
There are functions that try to acquire a lock only once and immediately
return a value telling about success or failure to acquire the lock.
They can be used if you need no access to the data protected with the
lock when some other thread is holding the lock. You should acquire the
lock later if you then need access to the data protected with the lock.
spin_trylock() does not spin but returns non-zero if it
acquires the spinlock on the first try or 0 if not. This function can be
used in all contexts like spin_lock(): you must have
disabled the contexts that might interrupt you and acquire the spin
lock.
mutex_trylock() does not suspend your task but returns
non-zero if it could lock the mutex on the first try or 0 if not. This
function cannot be safely used in hardware or software interrupt
contexts despite not sleeping.
Common Examples
===============
Let's step through a simple example: a cache of number to name mappings.
The cache keeps a count of how often each of the objects is used, and
when it gets full, throws out the least used one.
All In User Context
-------------------
For our first example, we assume that all operations are in user context
(ie. from system calls), so we can sleep. This means we can use a mutex
to protect the cache and all the objects within it. Here's the code::
#include <linux/list.h>
#include <linux/slab.h>
#include <linux/string.h>
#include <linux/mutex.h>
#include <asm/errno.h>
struct object
{
struct list_head list;
int id;
char name[32];
int popularity;
};
/* Protects the cache, cache_num, and the objects within it */
static DEFINE_MUTEX(cache_lock);
static LIST_HEAD(cache);
static unsigned int cache_num = 0;
#define MAX_CACHE_SIZE 10
/* Must be holding cache_lock */
static struct object *__cache_find(int id)
{
struct object *i;
list_for_each_entry(i, &cache, list)
if (i->id == id) {
i->popularity++;
return i;
}
return NULL;
}
/* Must be holding cache_lock */
static void __cache_delete(struct object *obj)
{
BUG_ON(!obj);
list_del(&obj->list);
kfree(obj);
cache_num--;
}
/* Must be holding cache_lock */
static void __cache_add(struct object *obj)
{
list_add(&obj->list, &cache);
if (++cache_num > MAX_CACHE_SIZE) {
struct object *i, *outcast = NULL;
list_for_each_entry(i, &cache, list) {
if (!outcast || i->popularity < outcast->popularity)
outcast = i;
}
__cache_delete(outcast);
}
}
int cache_add(int id, const char *name)
{
struct object *obj;
if ((obj = kmalloc(sizeof(*obj), GFP_KERNEL)) == NULL)
return -ENOMEM;
strscpy(obj->name, name, sizeof(obj->name));
obj->id = id;
obj->popularity = 0;
mutex_lock(&cache_lock);
__cache_add(obj);
mutex_unlock(&cache_lock);
return 0;
}
void cache_delete(int id)
{
mutex_lock(&cache_lock);
__cache_delete(__cache_find(id));
mutex_unlock(&cache_lock);
}
int cache_find(int id, char *name)
{
struct object *obj;
int ret = -ENOENT;
mutex_lock(&cache_lock);
obj = __cache_find(id);
if (obj) {
ret = 0;
strcpy(name, obj->name);
}
mutex_unlock(&cache_lock);
return ret;
}
Note that we always make sure we have the cache_lock when we add,
delete, or look up the cache: both the cache infrastructure itself and
the contents of the objects are protected by the lock. In this case it's
easy, since we copy the data for the user, and never let them access the
objects directly.
There is a slight (and common) optimization here: in
cache_add() we set up the fields of the object before
grabbing the lock. This is safe, as no-one else can access it until we
put it in cache.
Accessing From Interrupt Context
--------------------------------
Now consider the case where cache_find() can be called
from interrupt context: either a hardware interrupt or a softirq. An
example would be a timer which deletes object from the cache.
The change is shown below, in standard patch format: the ``-`` are lines
which are taken away, and the ``+`` are lines which are added.
::
--- cache.c.usercontext 2003-12-09 13:58:54.000000000 +1100
+++ cache.c.interrupt 2003-12-09 14:07:49.000000000 +1100
@@ -12,7 +12,7 @@
int popularity;
};
-static DEFINE_MUTEX(cache_lock);
+static DEFINE_SPINLOCK(cache_lock);
static LIST_HEAD(cache);
static unsigned int cache_num = 0;
#define MAX_CACHE_SIZE 10
@@ -55,6 +55,7 @@
int cache_add(int id, const char *name)
{
struct object *obj;
+ unsigned long flags;
if ((obj = kmalloc(sizeof(*obj), GFP_KERNEL)) == NULL)
return -ENOMEM;
@@ -63,30 +64,33 @@
obj->id = id;
obj->popularity = 0;
- mutex_lock(&cache_lock);
+ spin_lock_irqsave(&cache_lock, flags);
__cache_add(obj);
- mutex_unlock(&cache_lock);
+ spin_unlock_irqrestore(&cache_lock, flags);
return 0;
}
void cache_delete(int id)
{
- mutex_lock(&cache_lock);
+ unsigned long flags;
+
+ spin_lock_irqsave(&cache_lock, flags);
__cache_delete(__cache_find(id));
- mutex_unlock(&cache_lock);
+ spin_unlock_irqrestore(&cache_lock, flags);
}
int cache_find(int id, char *name)
{
struct object *obj;
int ret = -ENOENT;
+ unsigned long flags;
- mutex_lock(&cache_lock);
+ spin_lock_irqsave(&cache_lock, flags);
obj = __cache_find(id);
if (obj) {
ret = 0;
strcpy(name, obj->name);
}
- mutex_unlock(&cache_lock);
+ spin_unlock_irqrestore(&cache_lock, flags);
return ret;
}
Note that the spin_lock_irqsave() will turn off
interrupts if they are on, otherwise does nothing (if we are already in
an interrupt handler), hence these functions are safe to call from any
context.
Unfortunately, cache_add() calls kmalloc()
with the ``GFP_KERNEL`` flag, which is only legal in user context. I
have assumed that cache_add() is still only called in
user context, otherwise this should become a parameter to
cache_add().
Exposing Objects Outside This File
----------------------------------
If our objects contained more information, it might not be sufficient to
copy the information in and out: other parts of the code might want to
keep pointers to these objects, for example, rather than looking up the
id every time. This produces two problems.
The first problem is that we use the ``cache_lock`` to protect objects:
we'd need to make this non-static so the rest of the code can use it.
This makes locking trickier, as it is no longer all in one place.
The second problem is the lifetime problem: if another structure keeps a
pointer to an object, it presumably expects that pointer to remain
valid. Unfortunately, this is only guaranteed while you hold the lock,
otherwise someone might call cache_delete() and even
worse, add another object, re-using the same address.
As there is only one lock, you can't hold it forever: no-one else would
get any work done.
The solution to this problem is to use a reference count: everyone who
has a pointer to the object increases it when they first get the object,
and drops the reference count when they're finished with it. Whoever
drops it to zero knows it is unused, and can actually delete it.
Here is the code::
--- cache.c.interrupt 2003-12-09 14:25:43.000000000 +1100
+++ cache.c.refcnt 2003-12-09 14:33:05.000000000 +1100
@@ -7,6 +7,7 @@
struct object
{
struct list_head list;
+ unsigned int refcnt;
int id;
char name[32];
int popularity;
@@ -17,6 +18,35 @@
static unsigned int cache_num = 0;
#define MAX_CACHE_SIZE 10
+static void __object_put(struct object *obj)
+{
+ if (--obj->refcnt == 0)
+ kfree(obj);
+}
+
+static void __object_get(struct object *obj)
+{
+ obj->refcnt++;
+}
+
+void object_put(struct object *obj)
+{
+ unsigned long flags;
+
+ spin_lock_irqsave(&cache_lock, flags);
+ __object_put(obj);
+ spin_unlock_irqrestore(&cache_lock, flags);
+}
+
+void object_get(struct object *obj)
+{
+ unsigned long flags;
+
+ spin_lock_irqsave(&cache_lock, flags);
+ __object_get(obj);
+ spin_unlock_irqrestore(&cache_lock, flags);
+}
+
/* Must be holding cache_lock */
static struct object *__cache_find(int id)
{
@@ -35,6 +65,7 @@
{
BUG_ON(!obj);
list_del(&obj->list);
+ __object_put(obj);
cache_num--;
}
@@ -63,6 +94,7 @@
strscpy(obj->name, name, sizeof(obj->name));
obj->id = id;
obj->popularity = 0;
+ obj->refcnt = 1; /* The cache holds a reference */
spin_lock_irqsave(&cache_lock, flags);
__cache_add(obj);
@@ -79,18 +111,15 @@
spin_unlock_irqrestore(&cache_lock, flags);
}
-int cache_find(int id, char *name)
+struct object *cache_find(int id)
{
struct object *obj;
- int ret = -ENOENT;
unsigned long flags;
spin_lock_irqsave(&cache_lock, flags);
obj = __cache_find(id);
- if (obj) {
- ret = 0;
- strcpy(name, obj->name);
- }
+ if (obj)
+ __object_get(obj);
spin_unlock_irqrestore(&cache_lock, flags);
- return ret;
+ return obj;
}
We encapsulate the reference counting in the standard 'get' and 'put'
functions. Now we can return the object itself from
cache_find() which has the advantage that the user can
now sleep holding the object (eg. to copy_to_user() to
name to userspace).
The other point to note is that I said a reference should be held for
every pointer to the object: thus the reference count is 1 when first
inserted into the cache. In some versions the framework does not hold a
reference count, but they are more complicated.
Using Atomic Operations For The Reference Count
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
In practice, :c:type:`atomic_t` would usually be used for refcnt. There are a
number of atomic operations defined in ``include/asm/atomic.h``: these
are guaranteed to be seen atomically from all CPUs in the system, so no
lock is required. In this case, it is simpler than using spinlocks,
although for anything non-trivial using spinlocks is clearer. The
atomic_inc() and atomic_dec_and_test()
are used instead of the standard increment and decrement operators, and
the lock is no longer used to protect the reference count itself.
::
--- cache.c.refcnt 2003-12-09 15:00:35.000000000 +1100
+++ cache.c.refcnt-atomic 2003-12-11 15:49:42.000000000 +1100
@@ -7,7 +7,7 @@
struct object
{
struct list_head list;
- unsigned int refcnt;
+ atomic_t refcnt;
int id;
char name[32];
int popularity;
@@ -18,33 +18,15 @@
static unsigned int cache_num = 0;
#define MAX_CACHE_SIZE 10
-static void __object_put(struct object *obj)
-{
- if (--obj->refcnt == 0)
- kfree(obj);
-}
-
-static void __object_get(struct object *obj)
-{
- obj->refcnt++;
-}
-
void object_put(struct object *obj)
{
- unsigned long flags;
-
- spin_lock_irqsave(&cache_lock, flags);
- __object_put(obj);
- spin_unlock_irqrestore(&cache_lock, flags);
+ if (atomic_dec_and_test(&obj->refcnt))
+ kfree(obj);
}
void object_get(struct object *obj)
{
- unsigned long flags;
-
- spin_lock_irqsave(&cache_lock, flags);
- __object_get(obj);
- spin_unlock_irqrestore(&cache_lock, flags);
+ atomic_inc(&obj->refcnt);
}
/* Must be holding cache_lock */
@@ -65,7 +47,7 @@
{
BUG_ON(!obj);
list_del(&obj->list);
- __object_put(obj);
+ object_put(obj);
cache_num--;
}
@@ -94,7 +76,7 @@
strscpy(obj->name, name, sizeof(obj->name));
obj->id = id;
obj->popularity = 0;
- obj->refcnt = 1; /* The cache holds a reference */
+ atomic_set(&obj->refcnt, 1); /* The cache holds a reference */
spin_lock_irqsave(&cache_lock, flags);
__cache_add(obj);
@@ -119,7 +101,7 @@
spin_lock_irqsave(&cache_lock, flags);
obj = __cache_find(id);
if (obj)
- __object_get(obj);
+ object_get(obj);
spin_unlock_irqrestore(&cache_lock, flags);
return obj;
}
Protecting The Objects Themselves
---------------------------------
In these examples, we assumed that the objects (except the reference
counts) never changed once they are created. If we wanted to allow the
name to change, there are three possibilities:
- You can make ``cache_lock`` non-static, and tell people to grab that
lock before changing the name in any object.
- You can provide a cache_obj_rename() which grabs this
lock and changes the name for the caller, and tell everyone to use
that function.
- You can make the ``cache_lock`` protect only the cache itself, and
use another lock to protect the name.
Theoretically, you can make the locks as fine-grained as one lock for
every field, for every object. In practice, the most common variants
are:
- One lock which protects the infrastructure (the ``cache`` list in
this example) and all the objects. This is what we have done so far.
- One lock which protects the infrastructure (including the list
pointers inside the objects), and one lock inside the object which
protects the rest of that object.
- Multiple locks to protect the infrastructure (eg. one lock per hash
chain), possibly with a separate per-object lock.
Here is the "lock-per-object" implementation:
::
--- cache.c.refcnt-atomic 2003-12-11 15:50:54.000000000 +1100
+++ cache.c.perobjectlock 2003-12-11 17:15:03.000000000 +1100
@@ -6,11 +6,17 @@
struct object
{
+ /* These two protected by cache_lock. */
struct list_head list;
+ int popularity;
+
atomic_t refcnt;
+
+ /* Doesn't change once created. */
int id;
+
+ spinlock_t lock; /* Protects the name */
char name[32];
- int popularity;
};
static DEFINE_SPINLOCK(cache_lock);
@@ -77,6 +84,7 @@
obj->id = id;
obj->popularity = 0;
atomic_set(&obj->refcnt, 1); /* The cache holds a reference */
+ spin_lock_init(&obj->lock);
spin_lock_irqsave(&cache_lock, flags);
__cache_add(obj);
Note that I decide that the popularity count should be protected by the
``cache_lock`` rather than the per-object lock: this is because it (like
the :c:type:`struct list_head <list_head>` inside the object)
is logically part of the infrastructure. This way, I don't need to grab
the lock of every object in __cache_add() when seeking
the least popular.
I also decided that the id member is unchangeable, so I don't need to
grab each object lock in __cache_find() to examine the
id: the object lock is only used by a caller who wants to read or write
the name field.
Note also that I added a comment describing what data was protected by
which locks. This is extremely important, as it describes the runtime
behavior of the code, and can be hard to gain from just reading. And as
Alan Cox says, “Lock data, not code”.
Common Problems
===============
Deadlock: Simple and Advanced
-----------------------------
There is a coding bug where a piece of code tries to grab a spinlock
twice: it will spin forever, waiting for the lock to be released
(spinlocks, rwlocks and mutexes are not recursive in Linux). This is
trivial to diagnose: not a
stay-up-five-nights-talk-to-fluffy-code-bunnies kind of problem.
For a slightly more complex case, imagine you have a region shared by a
softirq and user context. If you use a spin_lock() call
to protect it, it is possible that the user context will be interrupted
by the softirq while it holds the lock, and the softirq will then spin
forever trying to get the same lock.
Both of these are called deadlock, and as shown above, it can occur even
with a single CPU (although not on UP compiles, since spinlocks vanish
on kernel compiles with ``CONFIG_SMP``\ =n. You'll still get data
corruption in the second example).
This complete lockup is easy to diagnose: on SMP boxes the watchdog
timer or compiling with ``DEBUG_SPINLOCK`` set
(``include/linux/spinlock.h``) will show this up immediately when it
happens.
A more complex problem is the so-called 'deadly embrace', involving two
or more locks. Say you have a hash table: each entry in the table is a
spinlock, and a chain of hashed objects. Inside a softirq handler, you
sometimes want to alter an object from one place in the hash to another:
you grab the spinlock of the old hash chain and the spinlock of the new
hash chain, and delete the object from the old one, and insert it in the
new one.
There are two problems here. First, if your code ever tries to move the
object to the same chain, it will deadlock with itself as it tries to
lock it twice. Secondly, if the same softirq on another CPU is trying to
move another object in the reverse direction, the following could
happen:
+-----------------------+-----------------------+
| CPU 1 | CPU 2 |
+=======================+=======================+
| Grab lock A -> OK | Grab lock B -> OK |
+-----------------------+-----------------------+
| Grab lock B -> spin | Grab lock A -> spin |
+-----------------------+-----------------------+
Table: Consequences
The two CPUs will spin forever, waiting for the other to give up their
lock. It will look, smell, and feel like a crash.
Preventing Deadlock
-------------------
Textbooks will tell you that if you always lock in the same order, you
will never get this kind of deadlock. Practice will tell you that this
approach doesn't scale: when I create a new lock, I don't understand
enough of the kernel to figure out where in the 5000 lock hierarchy it
will fit.
The best locks are encapsulated: they never get exposed in headers, and
are never held around calls to non-trivial functions outside the same
file. You can read through this code and see that it will never
deadlock, because it never tries to grab another lock while it has that
one. People using your code don't even need to know you are using a
lock.
A classic problem here is when you provide callbacks or hooks: if you
call these with the lock held, you risk simple deadlock, or a deadly
embrace (who knows what the callback will do?).
Overzealous Prevention Of Deadlocks
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Deadlocks are problematic, but not as bad as data corruption. Code which
grabs a read lock, searches a list, fails to find what it wants, drops
the read lock, grabs a write lock and inserts the object has a race
condition.
Racing Timers: A Kernel Pastime
-------------------------------
Timers can produce their own special problems with races. Consider a
collection of objects (list, hash, etc) where each object has a timer
which is due to destroy it.
If you want to destroy the entire collection (say on module removal),
you might do the following::
/* THIS CODE BAD BAD BAD BAD: IF IT WAS ANY WORSE IT WOULD USE
HUNGARIAN NOTATION */
spin_lock_bh(&list_lock);
while (list) {
struct foo *next = list->next;
timer_delete(&list->timer);
kfree(list);
list = next;
}
spin_unlock_bh(&list_lock);
Sooner or later, this will crash on SMP, because a timer can have just
gone off before the spin_lock_bh(), and it will only get
the lock after we spin_unlock_bh(), and then try to free
the element (which has already been freed!).
This can be avoided by checking the result of
timer_delete(): if it returns 1, the timer has been deleted.
If 0, it means (in this case) that it is currently running, so we can
do::
retry:
spin_lock_bh(&list_lock);
while (list) {
struct foo *next = list->next;
if (!timer_delete(&list->timer)) {
/* Give timer a chance to delete this */
spin_unlock_bh(&list_lock);
goto retry;
}
kfree(list);
list = next;
}
spin_unlock_bh(&list_lock);
Another common problem is deleting timers which restart themselves (by
calling add_timer() at the end of their timer function).
Because this is a fairly common case which is prone to races, you should
use timer_delete_sync() (``include/linux/timer.h``) to handle this case.
Before freeing a timer, timer_shutdown() or timer_shutdown_sync() should be
called which will keep it from being rearmed. Any subsequent attempt to
rearm the timer will be silently ignored by the core code.
Locking Speed
=============
There are three main things to worry about when considering speed of
some code which does locking. First is concurrency: how many things are
going to be waiting while someone else is holding a lock. Second is the
time taken to actually acquire and release an uncontended lock. Third is
using fewer, or smarter locks. I'm assuming that the lock is used fairly
often: otherwise, you wouldn't be concerned about efficiency.
Concurrency depends on how long the lock is usually held: you should
hold the lock for as long as needed, but no longer. In the cache
example, we always create the object without the lock held, and then
grab the lock only when we are ready to insert it in the list.
Acquisition times depend on how much damage the lock operations do to
the pipeline (pipeline stalls) and how likely it is that this CPU was
the last one to grab the lock (ie. is the lock cache-hot for this CPU):
on a machine with more CPUs, this likelihood drops fast. Consider a
700MHz Intel Pentium III: an instruction takes about 0.7ns, an atomic
increment takes about 58ns, a lock which is cache-hot on this CPU takes
160ns, and a cacheline transfer from another CPU takes an additional 170
to 360ns. (These figures from Paul McKenney's `Linux Journal RCU
article <http://www.linuxjournal.com/article.php?sid=6993>`__).
These two aims conflict: holding a lock for a short time might be done
by splitting locks into parts (such as in our final per-object-lock
example), but this increases the number of lock acquisitions, and the
results are often slower than having a single lock. This is another
reason to advocate locking simplicity.
The third concern is addressed below: there are some methods to reduce
the amount of locking which needs to be done.
Read/Write Lock Variants
------------------------
Both spinlocks and mutexes have read/write variants: ``rwlock_t`` and
:c:type:`struct rw_semaphore <rw_semaphore>`. These divide
users into two classes: the readers and the writers. If you are only
reading the data, you can get a read lock, but to write to the data you
need the write lock. Many people can hold a read lock, but a writer must
be sole holder.
If your code divides neatly along reader/writer lines (as our cache code
does), and the lock is held by readers for significant lengths of time,
using these locks can help. They are slightly slower than the normal
locks though, so in practice ``rwlock_t`` is not usually worthwhile.
Avoiding Locks: Read Copy Update
--------------------------------
There is a special method of read/write locking called Read Copy Update.
Using RCU, the readers can avoid taking a lock altogether: as we expect
our cache to be read more often than updated (otherwise the cache is a
waste of time), it is a candidate for this optimization.
How do we get rid of read locks? Getting rid of read locks means that
writers may be changing the list underneath the readers. That is
actually quite simple: we can read a linked list while an element is
being added if the writer adds the element very carefully. For example,
adding ``new`` to a single linked list called ``list``::
new->next = list->next;
wmb();
list->next = new;
The wmb() is a write memory barrier. It ensures that the
first operation (setting the new element's ``next`` pointer) is complete
and will be seen by all CPUs, before the second operation is (putting
the new element into the list). This is important, since modern
compilers and modern CPUs can both reorder instructions unless told
otherwise: we want a reader to either not see the new element at all, or
see the new element with the ``next`` pointer correctly pointing at the
rest of the list.
Fortunately, there is a function to do this for standard
:c:type:`struct list_head <list_head>` lists:
list_add_rcu() (``include/linux/list.h``).
Removing an element from the list is even simpler: we replace the
pointer to the old element with a pointer to its successor, and readers
will either see it, or skip over it.
::
list->next = old->next;
There is list_del_rcu() (``include/linux/list.h``) which
does this (the normal version poisons the old object, which we don't
want).
The reader must also be careful: some CPUs can look through the ``next``
pointer to start reading the contents of the next element early, but
don't realize that the pre-fetched contents is wrong when the ``next``
pointer changes underneath them. Once again, there is a
list_for_each_entry_rcu() (``include/linux/list.h``)
to help you. Of course, writers can just use
list_for_each_entry(), since there cannot be two
simultaneous writers.
Our final dilemma is this: when can we actually destroy the removed
element? Remember, a reader might be stepping through this element in
the list right now: if we free this element and the ``next`` pointer
changes, the reader will jump off into garbage and crash. We need to
wait until we know that all the readers who were traversing the list
when we deleted the element are finished. We use
call_rcu() to register a callback which will actually
destroy the object once all pre-existing readers are finished.
Alternatively, synchronize_rcu() may be used to block
until all pre-existing are finished.
But how does Read Copy Update know when the readers are finished? The
method is this: firstly, the readers always traverse the list inside
rcu_read_lock()/rcu_read_unlock() pairs:
these simply disable preemption so the reader won't go to sleep while
reading the list.
RCU then waits until every other CPU has slept at least once: since
readers cannot sleep, we know that any readers which were traversing the
list during the deletion are finished, and the callback is triggered.
The real Read Copy Update code is a little more optimized than this, but
this is the fundamental idea.
::
--- cache.c.perobjectlock 2003-12-11 17:15:03.000000000 +1100
+++ cache.c.rcupdate 2003-12-11 17:55:14.000000000 +1100
@@ -1,15 +1,18 @@
#include <linux/list.h>
#include <linux/slab.h>
#include <linux/string.h>
+#include <linux/rcupdate.h>
#include <linux/mutex.h>
#include <asm/errno.h>
struct object
{
- /* These two protected by cache_lock. */
+ /* This is protected by RCU */
struct list_head list;
int popularity;
+ struct rcu_head rcu;
+
atomic_t refcnt;
/* Doesn't change once created. */
@@ -40,7 +43,7 @@
{
struct object *i;
- list_for_each_entry(i, &cache, list) {
+ list_for_each_entry_rcu(i, &cache, list) {
if (i->id == id) {
i->popularity++;
return i;
@@ -49,19 +52,25 @@
return NULL;
}
+/* Final discard done once we know no readers are looking. */
+static void cache_delete_rcu(void *arg)
+{
+ object_put(arg);
+}
+
/* Must be holding cache_lock */
static void __cache_delete(struct object *obj)
{
BUG_ON(!obj);
- list_del(&obj->list);
- object_put(obj);
+ list_del_rcu(&obj->list);
cache_num--;
+ call_rcu(&obj->rcu, cache_delete_rcu);
}
/* Must be holding cache_lock */
static void __cache_add(struct object *obj)
{
- list_add(&obj->list, &cache);
+ list_add_rcu(&obj->list, &cache);
if (++cache_num > MAX_CACHE_SIZE) {
struct object *i, *outcast = NULL;
list_for_each_entry(i, &cache, list) {
@@ -104,12 +114,11 @@
struct object *cache_find(int id)
{
struct object *obj;
- unsigned long flags;
- spin_lock_irqsave(&cache_lock, flags);
+ rcu_read_lock();
obj = __cache_find(id);
if (obj)
object_get(obj);
- spin_unlock_irqrestore(&cache_lock, flags);
+ rcu_read_unlock();
return obj;
}
Note that the reader will alter the popularity member in
__cache_find(), and now it doesn't hold a lock. One
solution would be to make it an ``atomic_t``, but for this usage, we
don't really care about races: an approximate result is good enough, so
I didn't change it.
The result is that cache_find() requires no
synchronization with any other functions, so is almost as fast on SMP as
it would be on UP.
There is a further optimization possible here: remember our original
cache code, where there were no reference counts and the caller simply
held the lock whenever using the object? This is still possible: if you
hold the lock, no one can delete the object, so you don't need to get
and put the reference count.
Now, because the 'read lock' in RCU is simply disabling preemption, a
caller which always has preemption disabled between calling
cache_find() and object_put() does not
need to actually get and put the reference count: we could expose
__cache_find() by making it non-static, and such
callers could simply call that.
The benefit here is that the reference count is not written to: the
object is not altered in any way, which is much faster on SMP machines
due to caching.
Per-CPU Data
------------
Another technique for avoiding locking which is used fairly widely is to
duplicate information for each CPU. For example, if you wanted to keep a
count of a common condition, you could use a spin lock and a single
counter. Nice and simple.
If that was too slow (it's usually not, but if you've got a really big
machine to test on and can show that it is), you could instead use a
counter for each CPU, then none of them need an exclusive lock. See
DEFINE_PER_CPU(), get_cpu_var() and
put_cpu_var() (``include/linux/percpu.h``).
Of particular use for simple per-cpu counters is the ``local_t`` type,
and the cpu_local_inc() and related functions, which are
more efficient than simple code on some architectures
(``include/asm/local.h``).
Note that there is no simple, reliable way of getting an exact value of
such a counter, without introducing more locks. This is not a problem
for some uses.
Data Which Mostly Used By An IRQ Handler
----------------------------------------
If data is always accessed from within the same IRQ handler, you don't
need a lock at all: the kernel already guarantees that the irq handler
will not run simultaneously on multiple CPUs.
Manfred Spraul points out that you can still do this, even if the data
is very occasionally accessed in user context or softirqs/tasklets. The
irq handler doesn't use a lock, and all other accesses are done as so::
mutex_lock(&lock);
disable_irq(irq);
...
enable_irq(irq);
mutex_unlock(&lock);
The disable_irq() prevents the irq handler from running
(and waits for it to finish if it's currently running on other CPUs).
The spinlock prevents any other accesses happening at the same time.
Naturally, this is slower than just a spin_lock_irq()
call, so it only makes sense if this type of access happens extremely
rarely.
What Functions Are Safe To Call From Interrupts?
================================================
Many functions in the kernel sleep (ie. call schedule()) directly or
indirectly: you can never call them while holding a spinlock, or with
preemption disabled. This also means you need to be in user context:
calling them from an interrupt is illegal.
Some Functions Which Sleep
--------------------------
The most common ones are listed below, but you usually have to read the
code to find out if other calls are safe. If everyone else who calls it
can sleep, you probably need to be able to sleep, too. In particular,
registration and deregistration functions usually expect to be called
from user context, and can sleep.
- Accesses to userspace:
- copy_from_user()
- copy_to_user()
- get_user()
- put_user()
- kmalloc(GP_KERNEL) <kmalloc>`
- mutex_lock_interruptible() and
mutex_lock()
There is a mutex_trylock() which does not sleep.
Still, it must not be used inside interrupt context since its
implementation is not safe for that. mutex_unlock()
will also never sleep. It cannot be used in interrupt context either
since a mutex must be released by the same task that acquired it.
Some Functions Which Don't Sleep
--------------------------------
Some functions are safe to call from any context, or holding almost any
lock.
- printk()
- kfree()
- add_timer() and timer_delete()
Mutex API reference
===================
.. kernel-doc:: include/linux/mutex.h
:internal:
.. kernel-doc:: kernel/locking/mutex.c
:export:
Futex API reference
===================
.. kernel-doc:: kernel/futex/core.c
:internal:
.. kernel-doc:: kernel/futex/futex.h
:internal:
.. kernel-doc:: kernel/futex/pi.c
:internal:
.. kernel-doc:: kernel/futex/requeue.c
:internal:
.. kernel-doc:: kernel/futex/waitwake.c
:internal:
Further reading
===============
- ``Documentation/locking/spinlocks.rst``: Linus Torvalds' spinlocking
tutorial in the kernel sources.
- Unix Systems for Modern Architectures: Symmetric Multiprocessing and
Caching for Kernel Programmers:
Curt Schimmel's very good introduction to kernel level locking (not
written for Linux, but nearly everything applies). The book is
expensive, but really worth every penny to understand SMP locking.
[ISBN: 0201633388]
Thanks
======
Thanks to Telsa Gwynne for DocBooking, neatening and adding style.
Thanks to Martin Pool, Philipp Rumpf, Stephen Rothwell, Paul Mackerras,
Ruedi Aschwanden, Alan Cox, Manfred Spraul, Tim Waugh, Pete Zaitcev,
James Morris, Robert Love, Paul McKenney, John Ashby for proofreading,
correcting, flaming, commenting.
Thanks to the cabal for having no influence on this document.
Glossary
========
preemption
Prior to 2.5, or when ``CONFIG_PREEMPT`` is unset, processes in user
context inside the kernel would not preempt each other (ie. you had that
CPU until you gave it up, except for interrupts). With the addition of
``CONFIG_PREEMPT`` in 2.5.4, this changed: when in user context, higher
priority tasks can "cut in": spinlocks were changed to disable
preemption, even on UP.
bh
Bottom Half: for historical reasons, functions with '_bh' in them often
now refer to any software interrupt, e.g. spin_lock_bh()
blocks any software interrupt on the current CPU. Bottom halves are
deprecated, and will eventually be replaced by tasklets. Only one bottom
half will be running at any time.
Hardware Interrupt / Hardware IRQ
Hardware interrupt request. in_hardirq() returns true in a
hardware interrupt handler.
Interrupt Context
Not user context: processing a hardware irq or software irq. Indicated
by the in_interrupt() macro returning true.
SMP
Symmetric Multi-Processor: kernels compiled for multiple-CPU machines.
(``CONFIG_SMP=y``).
Software Interrupt / softirq
Software interrupt handler. in_hardirq() returns false;
in_softirq() returns true. Tasklets and softirqs both
fall into the category of 'software interrupts'.
Strictly speaking a softirq is one of up to 32 enumerated software
interrupts which can run on multiple CPUs at once. Sometimes used to
refer to tasklets as well (ie. all software interrupts).
tasklet
A dynamically-registrable software interrupt, which is guaranteed to
only run on one CPU at a time.
timer
A dynamically-registrable software interrupt, which is run at (or close
to) a given time. When running, it is just like a tasklet (in fact, they
are called from the ``TIMER_SOFTIRQ``).
UP
Uni-Processor: Non-SMP. (``CONFIG_SMP=n``).
User Context
The kernel executing on behalf of a particular process (ie. a system
call or trap) or kernel thread. You can tell which process with the
``current`` macro.) Not to be confused with userspace. Can be
interrupted by software or hardware interrupts.
Userspace
A process executing its own code outside the kernel.
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