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Message-Id: <1460474301-22998-7-git-send-email-paulmck@linux.vnet.ibm.com>
Date: Tue, 12 Apr 2016 08:18:18 -0700
From: "Paul E. McKenney" <paulmck@...ux.vnet.ibm.com>
To: linux-kernel@...r.kernel.org
Cc: mingo@...nel.org, jiangshanlai@...il.com, dipankar@...ibm.com,
akpm@...ux-foundation.org, mathieu.desnoyers@...icios.com,
josh@...htriplett.org, tglx@...utronix.de, peterz@...radead.org,
rostedt@...dmis.org, dhowells@...hat.com, edumazet@...gle.com,
dvhart@...ux.intel.com, fweisbec@...il.com, oleg@...hat.com,
bobby.prani@...il.com,
"Paul E. McKenney" <paulmck@...ux.vnet.ibm.com>
Subject: [PATCH tip/core/rcu 07/10] documentation: Get rid of duplicate .htmlx file
This commit uses colors to obscure the quick-quiz answers, thus getting
rid of the .htmlx file. Use your mouse to select the answer in order
to see the text. Alternatively, use your favorite scripting language
to remove all occurences of "<font color="ffffff">" from the file.
Reported-by: Linus Torvalds <torvalds@...ux-foundation.org>
Signed-off-by: Paul E. McKenney <paulmck@...ux.vnet.ibm.com>
---
.../RCU/Design/Requirements/Requirements.html | 889 +++---
.../RCU/Design/Requirements/Requirements.htmlx | 2872 --------------------
Documentation/RCU/Design/htmlqqz.sh | 108 -
3 files changed, 428 insertions(+), 3441 deletions(-)
delete mode 100644 Documentation/RCU/Design/Requirements/Requirements.htmlx
delete mode 100755 Documentation/RCU/Design/htmlqqz.sh
diff --git a/Documentation/RCU/Design/Requirements/Requirements.html b/Documentation/RCU/Design/Requirements/Requirements.html
index c67a96a2a389..acdad96f78e9 100644
--- a/Documentation/RCU/Design/Requirements/Requirements.html
+++ b/Documentation/RCU/Design/Requirements/Requirements.html
@@ -1,5 +1,3 @@
-<!-- DO NOT HAND EDIT. -->
-<!-- Instead, edit Requirements.htmlx and run 'sh htmlqqz.sh Requirements' -->
<!DOCTYPE HTML PUBLIC "-//W3C//DTD HTML 4.01 Transitional//EN"
"http://www.w3.org/TR/html4/loose.dtd">
<html>
@@ -65,8 +63,8 @@ All that aside, here are the categories of currently known RCU requirements:
<p>
This is followed by a <a href="#Summary">summary</a>,
-which is in turn followed by the inevitable
-<a href="#Answers to Quick Quizzes">answers to the quick quizzes</a>.
+however, the answers to each quick quiz immediately follows the quiz.
+Select the big white space with your mouse to see the answer.
<h2><a name="Fundamental Requirements">Fundamental Requirements</a></h2>
@@ -153,13 +151,27 @@ Therefore, the outcome:
</blockquote>
cannot happen.
-<p><a name="Quick Quiz 1"><b>Quick Quiz 1</b>:</a>
-Wait a minute!
-You said that updaters can make useful forward progress concurrently
-with readers, but pre-existing readers will block
-<tt>synchronize_rcu()</tt>!!!
-Just who are you trying to fool???
-<br><a href="#qq1answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ Wait a minute!
+ You said that updaters can make useful forward progress concurrently
+ with readers, but pre-existing readers will block
+ <tt>synchronize_rcu()</tt>!!!
+ Just who are you trying to fool???
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ First, if updaters do not wish to be blocked by readers, they can use
+ <tt>call_rcu()</tt> or <tt>kfree_rcu()</tt>, which will
+ be discussed later.
+ Second, even when using <tt>synchronize_rcu()</tt>, the other
+ update-side code does run concurrently with readers, whether
+ pre-existing or not.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
This scenario resembles one of the first uses of RCU in
@@ -210,9 +222,20 @@ to guarantee that <tt>do_something()</tt> never runs concurrently
with <tt>recovery()</tt>, but with little or no synchronization
overhead in <tt>do_something_dlm()</tt>.
-<p><a name="Quick Quiz 2"><b>Quick Quiz 2</b>:</a>
-Why is the <tt>synchronize_rcu()</tt> on line 28 needed?
-<br><a href="#qq2answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ Why is the <tt>synchronize_rcu()</tt> on line 28 needed?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ Without that extra grace period, memory reordering could result in
+ <tt>do_something_dlm()</tt> executing <tt>do_something()</tt>
+ concurrently with the last bits of <tt>recovery()</tt>.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
In order to avoid fatal problems such as deadlocks,
@@ -332,12 +355,27 @@ It also prevents any number of “interesting” compiler
optimizations, for example, the use of <tt>gp</tt> as a scratch
location immediately preceding the assignment.
-<p><a name="Quick Quiz 3"><b>Quick Quiz 3</b>:</a>
-But <tt>rcu_assign_pointer()</tt> does nothing to prevent the
-two assignments to <tt>p->a</tt> and <tt>p->b</tt>
-from being reordered.
-Can't that also cause problems?
-<br><a href="#qq3answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ But <tt>rcu_assign_pointer()</tt> does nothing to prevent the
+ two assignments to <tt>p->a</tt> and <tt>p->b</tt>
+ from being reordered.
+ Can't that also cause problems?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ No, it cannot.
+ The readers cannot see either of these two fields until
+ the assignment to <tt>gp</tt>, by which time both fields are
+ fully initialized.
+ So reordering the assignments
+ to <tt>p->a</tt> and <tt>p->b</tt> cannot possibly
+ cause any problems.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
It is tempting to assume that the reader need not do anything special
@@ -494,11 +532,42 @@ The <tt>rcu_access_pointer()</tt> on line 6 is similar to
code protected by the corresponding update-side lock.
</ol>
-<p><a name="Quick Quiz 4"><b>Quick Quiz 4</b>:</a>
-Without the <tt>rcu_dereference()</tt> or the
-<tt>rcu_access_pointer()</tt>, what destructive optimizations
-might the compiler make use of?
-<br><a href="#qq4answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ Without the <tt>rcu_dereference()</tt> or the
+ <tt>rcu_access_pointer()</tt>, what destructive optimizations
+ might the compiler make use of?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ Let's start with what happens to <tt>do_something_gp()</tt>
+ if it fails to use <tt>rcu_dereference()</tt>.
+ It could reuse a value formerly fetched from this same pointer.
+ It could also fetch the pointer from <tt>gp</tt> in a byte-at-a-time
+ manner, resulting in <i>load tearing</i>, in turn resulting a bytewise
+ mash-up of two distince pointer values.
+ It might even use value-speculation optimizations, where it makes
+ a wrong guess, but by the time it gets around to checking the
+ value, an update has changed the pointer to match the wrong guess.
+ Too bad about any dereferences that returned pre-initialization garbage
+ in the meantime!
+ </font>
+
+ <p><font color="ffffff">
+ For <tt>remove_gp_synchronous()</tt>, as long as all modifications
+ to <tt>gp</tt> are carried out while holding <tt>gp_lock</tt>,
+ the above optimizations are harmless.
+ However,
+ with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt>,
+ <tt>sparse</tt> will complain if you
+ define <tt>gp</tt> with <tt>__rcu</tt> and then
+ access it without using
+ either <tt>rcu_access_pointer()</tt> or <tt>rcu_dereference()</tt>.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
In short, RCU's publish-subscribe guarantee is provided by the combination
@@ -571,28 +640,156 @@ systems with more than one CPU:
<tt>synchronize_rcu()</tt> migrates in the meantime.
</ol>
-<p><a name="Quick Quiz 5"><b>Quick Quiz 5</b>:</a>
-Given that multiple CPUs can start RCU read-side critical sections
-at any time without any ordering whatsoever, how can RCU possibly tell whether
-or not a given RCU read-side critical section starts before a
-given instance of <tt>synchronize_rcu()</tt>?
-<br><a href="#qq5answer">Answer</a>
-
-<p><a name="Quick Quiz 6"><b>Quick Quiz 6</b>:</a>
-The first and second guarantees require unbelievably strict ordering!
-Are all these memory barriers <i> really</i> required?
-<br><a href="#qq6answer">Answer</a>
-
-<p><a name="Quick Quiz 7"><b>Quick Quiz 7</b>:</a>
-You claim that <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
-generate absolutely no code in some kernel builds.
-This means that the compiler might arbitrarily rearrange consecutive
-RCU read-side critical sections.
-Given such rearrangement, if a given RCU read-side critical section
-is done, how can you be sure that all prior RCU read-side critical
-sections are done?
-Won't the compiler rearrangements make that impossible to determine?
-<br><a href="#qq7answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ Given that multiple CPUs can start RCU read-side critical sections
+ at any time without any ordering whatsoever, how can RCU possibly
+ tell whether or not a given RCU read-side critical section starts
+ before a given instance of <tt>synchronize_rcu()</tt>?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ If RCU cannot tell whether or not a given
+ RCU read-side critical section starts before a
+ given instance of <tt>synchronize_rcu()</tt>,
+ then it must assume that the RCU read-side critical section
+ started first.
+ In other words, a given instance of <tt>synchronize_rcu()</tt>
+ can avoid waiting on a given RCU read-side critical section only
+ if it can prove that <tt>synchronize_rcu()</tt> started first.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
+
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ The first and second guarantees require unbelievably strict ordering!
+ Are all these memory barriers <i> really</i> required?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ Yes, they really are required.
+ To see why the first guarantee is required, consider the following
+ sequence of events:
+ </font>
+
+ <ol>
+ <li> <font color="ffffff">
+ CPU 1: <tt>rcu_read_lock()</tt>
+ </font>
+ <li> <font color="ffffff">
+ CPU 1: <tt>q = rcu_dereference(gp);
+ /* Very likely to return p. */</tt>
+ </font>
+ <li> <font color="ffffff">
+ CPU 0: <tt>list_del_rcu(p);</tt>
+ </font>
+ <li> <font color="ffffff">
+ CPU 0: <tt>synchronize_rcu()</tt> starts.
+ </font>
+ <li> <font color="ffffff">
+ CPU 1: <tt>do_something_with(q->a);
+ /* No smp_mb(), so might happen after kfree(). */</tt>
+ </font>
+ <li> <font color="ffffff">
+ CPU 1: <tt>rcu_read_unlock()</tt>
+ </font>
+ <li> <font color="ffffff">
+ CPU 0: <tt>synchronize_rcu()</tt> returns.
+ </font>
+ <li> <font color="ffffff">
+ CPU 0: <tt>kfree(p);</tt>
+ </font>
+ </ol>
+
+ <p><font color="ffffff">
+ Therefore, there absolutely must be a full memory barrier between the
+ end of the RCU read-side critical section and the end of the
+ grace period.
+ </font>
+
+ <p><font color="ffffff">
+ The sequence of events demonstrating the necessity of the second rule
+ is roughly similar:
+ </font>
+
+ <ol>
+ <li> <font color="ffffff">CPU 0: <tt>list_del_rcu(p);</tt>
+ </font>
+ <li> <font color="ffffff">CPU 0: <tt>synchronize_rcu()</tt> starts.
+ </font>
+ <li> <font color="ffffff">CPU 1: <tt>rcu_read_lock()</tt>
+ </font>
+ <li> <font color="ffffff">CPU 1: <tt>q = rcu_dereference(gp);
+ /* Might return p if no memory barrier. */</tt>
+ </font>
+ <li> <font color="ffffff">CPU 0: <tt>synchronize_rcu()</tt> returns.
+ </font>
+ <li> <font color="ffffff">CPU 0: <tt>kfree(p);</tt>
+ </font>
+ <li> <font color="ffffff">
+ CPU 1: <tt>do_something_with(q->a); /* Boom!!! */</tt>
+ </font>
+ <li> <font color="ffffff">CPU 1: <tt>rcu_read_unlock()</tt>
+ </font>
+ </ol>
+
+ <p><font color="ffffff">
+ And similarly, without a memory barrier between the beginning of the
+ grace period and the beginning of the RCU read-side critical section,
+ CPU 1 might end up accessing the freelist.
+ </font>
+
+ <p><font color="ffffff">
+ The “as if” rule of course applies, so that any
+ implementation that acts as if the appropriate memory barriers
+ were in place is a correct implementation.
+ That said, it is much easier to fool yourself into believing
+ that you have adhered to the as-if rule than it is to actually
+ adhere to it!
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
+
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ You claim that <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
+ generate absolutely no code in some kernel builds.
+ This means that the compiler might arbitrarily rearrange consecutive
+ RCU read-side critical sections.
+ Given such rearrangement, if a given RCU read-side critical section
+ is done, how can you be sure that all prior RCU read-side critical
+ sections are done?
+ Won't the compiler rearrangements make that impossible to determine?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ In cases where <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
+ generate absolutely no code, RCU infers quiescent states only at
+ special locations, for example, within the scheduler.
+ Because calls to <tt>schedule()</tt> had better prevent calling-code
+ accesses to shared variables from being rearranged across the call to
+ <tt>schedule()</tt>, if RCU detects the end of a given RCU read-side
+ critical section, it will necessarily detect the end of all prior
+ RCU read-side critical sections, no matter how aggressively the
+ compiler scrambles the code.
+ </font>
+
+ <p><font color="ffffff">
+ Again, this all assumes that the compiler cannot scramble code across
+ calls to the scheduler, out of interrupt handlers, into the idle loop,
+ into user-mode code, and so on.
+ But if your kernel build allows that sort of scrambling, you have broken
+ far more than just RCU!
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
Note that these memory-barrier requirements do not replace the fundamental
@@ -637,9 +834,19 @@ inconvenience can be avoided through use of the
<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt> API members
described later in this document.
-<p><a name="Quick Quiz 8"><b>Quick Quiz 8</b>:</a>
-But how does the upgrade-to-write operation exclude other readers?
-<br><a href="#qq8answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ But how does the upgrade-to-write operation exclude other readers?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ It doesn't, just like normal RCU updates, which also do not exclude
+ RCU readers.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
This guarantee allows lookup code to be shared between read-side
@@ -725,9 +932,20 @@ to do significant reordering.
This is by design: Any significant ordering constraints would slow down
these fast-path APIs.
-<p><a name="Quick Quiz 9"><b>Quick Quiz 9</b>:</a>
-Can't the compiler also reorder this code?
-<br><a href="#qq9answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ Can't the compiler also reorder this code?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ No, the volatile casts in <tt>READ_ONCE()</tt> and
+ <tt>WRITE_ONCE()</tt> prevent the compiler from reordering in
+ this particular case.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<h3><a name="Readers Do Not Exclude Updaters">Readers Do Not Exclude Updaters</a></h3>
@@ -780,10 +998,25 @@ new readers can start immediately after <tt>synchronize_rcu()</tt>
starts, and <tt>synchronize_rcu()</tt> is under no
obligation to wait for these new readers.
-<p><a name="Quick Quiz 10"><b>Quick Quiz 10</b>:</a>
-Suppose that synchronize_rcu() did wait until all readers had completed.
-Would the updater be able to rely on this?
-<br><a href="#qq10answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ Suppose that synchronize_rcu() did wait until all readers had completed.
+ Would the updater be able to rely on this?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ No.
+ Even if <tt>synchronize_rcu()</tt> were to wait until
+ all readers had completed, a new reader might start immediately after
+ <tt>synchronize_rcu()</tt> completed.
+ Therefore, the code following
+ <tt>synchronize_rcu()</tt> cannot rely on there being no readers
+ in any case.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<h3><a name="Grace Periods Don't Partition Read-Side Critical Sections">
Grace Periods Don't Partition Read-Side Critical Sections</a></h3>
@@ -980,11 +1213,24 @@ grace period.
As a result, an RCU read-side critical section cannot partition a pair
of RCU grace periods.
-<p><a name="Quick Quiz 11"><b>Quick Quiz 11</b>:</a>
-How long a sequence of grace periods, each separated by an RCU read-side
-critical section, would be required to partition the RCU read-side
-critical sections at the beginning and end of the chain?
-<br><a href="#qq11answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ How long a sequence of grace periods, each separated by an RCU
+ read-side critical section, would be required to partition the RCU
+ read-side critical sections at the beginning and end of the chain?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ In theory, an infinite number.
+ In practice, an unknown number that is sensitive to both implementation
+ details and timing considerations.
+ Therefore, even in practice, RCU users must abide by the
+ theoretical rather than the practical answer.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<h3><a name="Disabling Preemption Does Not Block Grace Periods">
Disabling Preemption Does Not Block Grace Periods</a></h3>
@@ -1153,9 +1399,43 @@ synchronization primitives be legal within RCU read-side critical sections,
including spinlocks, sequence locks, atomic operations, reference
counters, and memory barriers.
-<p><a name="Quick Quiz 12"><b>Quick Quiz 12</b>:</a>
-What about sleeping locks?
-<br><a href="#qq12answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ What about sleeping locks?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ These are forbidden within Linux-kernel RCU read-side critical
+ sections because it is not legal to place a quiescent state
+ (in this case, voluntary context switch) within an RCU read-side
+ critical section.
+ However, sleeping locks may be used within userspace RCU read-side
+ critical sections, and also within Linux-kernel sleepable RCU
+ <a href="#Sleepable RCU"><font color="ffffff">(SRCU)</font></a>
+ read-side critical sections.
+ In addition, the -rt patchset turns spinlocks into a
+ sleeping locks so that the corresponding critical sections
+ can be preempted, which also means that these sleeplockified
+ spinlocks (but not other sleeping locks!) may be acquire within
+ -rt-Linux-kernel RCU read-side critical sections.
+ </font>
+
+ <p><font color="ffffff">
+ Note that it <i>is</i> legal for a normal RCU read-side
+ critical section to conditionally acquire a sleeping locks
+ (as in <tt>mutex_trylock()</tt>), but only as long as it does
+ not loop indefinitely attempting to conditionally acquire that
+ sleeping locks.
+ The key point is that things like <tt>mutex_trylock()</tt>
+ either return with the mutex held, or return an error indication if
+ the mutex was not immediately available.
+ Either way, <tt>mutex_trylock()</tt> returns immediately without
+ sleeping.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
It often comes as a surprise that many algorithms do not require a
@@ -1378,12 +1658,27 @@ write an RCU callback function that takes too long.
Long-running operations should be relegated to separate threads or
(in the Linux kernel) workqueues.
-<p><a name="Quick Quiz 13"><b>Quick Quiz 13</b>:</a>
-Why does line 19 use <tt>rcu_access_pointer()</tt>?
-After all, <tt>call_rcu()</tt> on line 25 stores into the
-structure, which would interact badly with concurrent insertions.
-Doesn't this mean that <tt>rcu_dereference()</tt> is required?
-<br><a href="#qq13answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ Why does line 19 use <tt>rcu_access_pointer()</tt>?
+ After all, <tt>call_rcu()</tt> on line 25 stores into the
+ structure, which would interact badly with concurrent insertions.
+ Doesn't this mean that <tt>rcu_dereference()</tt> is required?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ Presumably the <tt>->gp_lock</tt> acquired on line 18 excludes
+ any changes, including any insertions that <tt>rcu_dereference()</tt>
+ would protect against.
+ Therefore, any insertions will be delayed until after
+ <tt>->gp_lock</tt>
+ is released on line 25, which in turn means that
+ <tt>rcu_access_pointer()</tt> suffices.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
However, all that <tt>remove_gp_cb()</tt> is doing is
@@ -1430,14 +1725,31 @@ This was due to the fact that RCU was not heavily used within DYNIX/ptx,
so the very few places that needed something like
<tt>synchronize_rcu()</tt> simply open-coded it.
-<p><a name="Quick Quiz 14"><b>Quick Quiz 14</b>:</a>
-Earlier it was claimed that <tt>call_rcu()</tt> and
-<tt>kfree_rcu()</tt> allowed updaters to avoid being blocked
-by readers.
-But how can that be correct, given that the invocation of the callback
-and the freeing of the memory (respectively) must still wait for
-a grace period to elapse?
-<br><a href="#qq14answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ Earlier it was claimed that <tt>call_rcu()</tt> and
+ <tt>kfree_rcu()</tt> allowed updaters to avoid being blocked
+ by readers.
+ But how can that be correct, given that the invocation of the callback
+ and the freeing of the memory (respectively) must still wait for
+ a grace period to elapse?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ We could define things this way, but keep in mind that this sort of
+ definition would say that updates in garbage-collected languages
+ cannot complete until the next time the garbage collector runs,
+ which does not seem at all reasonable.
+ The key point is that in most cases, an updater using either
+ <tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> can proceed to the
+ next update as soon as it has invoked <tt>call_rcu()</tt> or
+ <tt>kfree_rcu()</tt>, without having to wait for a subsequent
+ grace period.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
But what if the updater must wait for the completion of code to be
@@ -1862,11 +2174,26 @@ kthreads to be spawned.
Therefore, invoking <tt>synchronize_rcu()</tt> during scheduler
initialization can result in deadlock.
-<p><a name="Quick Quiz 15"><b>Quick Quiz 15</b>:</a>
-So what happens with <tt>synchronize_rcu()</tt> during
-scheduler initialization for <tt>CONFIG_PREEMPT=n</tt>
-kernels?
-<br><a href="#qq15answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ So what happens with <tt>synchronize_rcu()</tt> during
+ scheduler initialization for <tt>CONFIG_PREEMPT=n</tt>
+ kernels?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ In <tt>CONFIG_PREEMPT=n</tt> kernel, <tt>synchronize_rcu()</tt>
+ maps directly to <tt>synchronize_sched()</tt>.
+ Therefore, <tt>synchronize_rcu()</tt> works normally throughout
+ boot in <tt>CONFIG_PREEMPT=n</tt> kernels.
+ However, your code must also work in <tt>CONFIG_PREEMPT=y</tt> kernels,
+ so it is still necessary to avoid invoking <tt>synchronize_rcu()</tt>
+ during scheduler initialization.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
I learned of these boot-time requirements as a result of a series of
@@ -2571,10 +2898,23 @@ If you needed to wait on multiple different flavors of SRCU
(but why???), you would need to create a wrapper function resembling
<tt>call_my_srcu()</tt> for each SRCU flavor.
-<p><a name="Quick Quiz 16"><b>Quick Quiz 16</b>:</a>
-But what if I need to wait for multiple RCU flavors, but I also need
-the grace periods to be expedited?
-<br><a href="#qq16answer">Answer</a>
+<table>
+<tr><th> </th></tr>
+<tr><th align="left">Quick Quiz:</th></tr>
+<tr><td>
+ But what if I need to wait for multiple RCU flavors, but I also need
+ the grace periods to be expedited?
+</td></tr>
+<tr><th align="left">Answer:</th></tr>
+<tr><td bgcolor="#ffffff"><font color="ffffff">
+ If you are using expedited grace periods, there should be less penalty
+ for waiting on them in succession.
+ But if that is nevertheless a problem, you can use workqueues
+ or multiple kthreads to wait on the various expedited grace
+ periods concurrently.
+</font></td></tr>
+<tr><td> </td></tr>
+</table>
<p>
Again, it is usually better to adjust the RCU read-side critical sections
@@ -2678,377 +3018,4 @@ and is provided
under the terms of the Creative Commons Attribution-Share Alike 3.0
United States license.
-<h3><a name="Answers to Quick Quizzes">
-Answers to Quick Quizzes</a></h3>
-
-<a name="qq1answer"></a>
-<p><b>Quick Quiz 1</b>:
-Wait a minute!
-You said that updaters can make useful forward progress concurrently
-with readers, but pre-existing readers will block
-<tt>synchronize_rcu()</tt>!!!
-Just who are you trying to fool???
-
-
-</p><p><b>Answer</b>:
-First, if updaters do not wish to be blocked by readers, they can use
-<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt>, which will
-be discussed later.
-Second, even when using <tt>synchronize_rcu()</tt>, the other
-update-side code does run concurrently with readers, whether pre-existing
-or not.
-
-
-</p><p><a href="#Quick%20Quiz%201"><b>Back to Quick Quiz 1</b>.</a>
-
-<a name="qq2answer"></a>
-<p><b>Quick Quiz 2</b>:
-Why is the <tt>synchronize_rcu()</tt> on line 28 needed?
-
-
-</p><p><b>Answer</b>:
-Without that extra grace period, memory reordering could result in
-<tt>do_something_dlm()</tt> executing <tt>do_something()</tt>
-concurrently with the last bits of <tt>recovery()</tt>.
-
-
-</p><p><a href="#Quick%20Quiz%202"><b>Back to Quick Quiz 2</b>.</a>
-
-<a name="qq3answer"></a>
-<p><b>Quick Quiz 3</b>:
-But <tt>rcu_assign_pointer()</tt> does nothing to prevent the
-two assignments to <tt>p->a</tt> and <tt>p->b</tt>
-from being reordered.
-Can't that also cause problems?
-
-
-</p><p><b>Answer</b>:
-No, it cannot.
-The readers cannot see either of these two fields until
-the assignment to <tt>gp</tt>, by which time both fields are
-fully initialized.
-So reordering the assignments
-to <tt>p->a</tt> and <tt>p->b</tt> cannot possibly
-cause any problems.
-
-
-</p><p><a href="#Quick%20Quiz%203"><b>Back to Quick Quiz 3</b>.</a>
-
-<a name="qq4answer"></a>
-<p><b>Quick Quiz 4</b>:
-Without the <tt>rcu_dereference()</tt> or the
-<tt>rcu_access_pointer()</tt>, what destructive optimizations
-might the compiler make use of?
-
-
-</p><p><b>Answer</b>:
-Let's start with what happens to <tt>do_something_gp()</tt>
-if it fails to use <tt>rcu_dereference()</tt>.
-It could reuse a value formerly fetched from this same pointer.
-It could also fetch the pointer from <tt>gp</tt> in a byte-at-a-time
-manner, resulting in <i>load tearing</i>, in turn resulting a bytewise
-mash-up of two distince pointer values.
-It might even use value-speculation optimizations, where it makes a wrong
-guess, but by the time it gets around to checking the value, an update
-has changed the pointer to match the wrong guess.
-Too bad about any dereferences that returned pre-initialization garbage
-in the meantime!
-
-<p>
-For <tt>remove_gp_synchronous()</tt>, as long as all modifications
-to <tt>gp</tt> are carried out while holding <tt>gp_lock</tt>,
-the above optimizations are harmless.
-However,
-with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt>,
-<tt>sparse</tt> will complain if you
-define <tt>gp</tt> with <tt>__rcu</tt> and then
-access it without using
-either <tt>rcu_access_pointer()</tt> or <tt>rcu_dereference()</tt>.
-
-
-</p><p><a href="#Quick%20Quiz%204"><b>Back to Quick Quiz 4</b>.</a>
-
-<a name="qq5answer"></a>
-<p><b>Quick Quiz 5</b>:
-Given that multiple CPUs can start RCU read-side critical sections
-at any time without any ordering whatsoever, how can RCU possibly tell whether
-or not a given RCU read-side critical section starts before a
-given instance of <tt>synchronize_rcu()</tt>?
-
-
-</p><p><b>Answer</b>:
-If RCU cannot tell whether or not a given
-RCU read-side critical section starts before a
-given instance of <tt>synchronize_rcu()</tt>,
-then it must assume that the RCU read-side critical section
-started first.
-In other words, a given instance of <tt>synchronize_rcu()</tt>
-can avoid waiting on a given RCU read-side critical section only
-if it can prove that <tt>synchronize_rcu()</tt> started first.
-
-
-</p><p><a href="#Quick%20Quiz%205"><b>Back to Quick Quiz 5</b>.</a>
-
-<a name="qq6answer"></a>
-<p><b>Quick Quiz 6</b>:
-The first and second guarantees require unbelievably strict ordering!
-Are all these memory barriers <i> really</i> required?
-
-
-</p><p><b>Answer</b>:
-Yes, they really are required.
-To see why the first guarantee is required, consider the following
-sequence of events:
-
-<ol>
-<li> CPU 1: <tt>rcu_read_lock()</tt>
-<li> CPU 1: <tt>q = rcu_dereference(gp);
- /* Very likely to return p. */</tt>
-<li> CPU 0: <tt>list_del_rcu(p);</tt>
-<li> CPU 0: <tt>synchronize_rcu()</tt> starts.
-<li> CPU 1: <tt>do_something_with(q->a);
- /* No smp_mb(), so might happen after kfree(). */</tt>
-<li> CPU 1: <tt>rcu_read_unlock()</tt>
-<li> CPU 0: <tt>synchronize_rcu()</tt> returns.
-<li> CPU 0: <tt>kfree(p);</tt>
-</ol>
-
-<p>
-Therefore, there absolutely must be a full memory barrier between the
-end of the RCU read-side critical section and the end of the
-grace period.
-
-<p>
-The sequence of events demonstrating the necessity of the second rule
-is roughly similar:
-
-<ol>
-<li> CPU 0: <tt>list_del_rcu(p);</tt>
-<li> CPU 0: <tt>synchronize_rcu()</tt> starts.
-<li> CPU 1: <tt>rcu_read_lock()</tt>
-<li> CPU 1: <tt>q = rcu_dereference(gp);
- /* Might return p if no memory barrier. */</tt>
-<li> CPU 0: <tt>synchronize_rcu()</tt> returns.
-<li> CPU 0: <tt>kfree(p);</tt>
-<li> CPU 1: <tt>do_something_with(q->a); /* Boom!!! */</tt>
-<li> CPU 1: <tt>rcu_read_unlock()</tt>
-</ol>
-
-<p>
-And similarly, without a memory barrier between the beginning of the
-grace period and the beginning of the RCU read-side critical section,
-CPU 1 might end up accessing the freelist.
-
-<p>
-The “as if” rule of course applies, so that any implementation
-that acts as if the appropriate memory barriers were in place is a
-correct implementation.
-That said, it is much easier to fool yourself into believing that you have
-adhered to the as-if rule than it is to actually adhere to it!
-
-
-</p><p><a href="#Quick%20Quiz%206"><b>Back to Quick Quiz 6</b>.</a>
-
-<a name="qq7answer"></a>
-<p><b>Quick Quiz 7</b>:
-You claim that <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
-generate absolutely no code in some kernel builds.
-This means that the compiler might arbitrarily rearrange consecutive
-RCU read-side critical sections.
-Given such rearrangement, if a given RCU read-side critical section
-is done, how can you be sure that all prior RCU read-side critical
-sections are done?
-Won't the compiler rearrangements make that impossible to determine?
-
-
-</p><p><b>Answer</b>:
-In cases where <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
-generate absolutely no code, RCU infers quiescent states only at
-special locations, for example, within the scheduler.
-Because calls to <tt>schedule()</tt> had better prevent calling-code
-accesses to shared variables from being rearranged across the call to
-<tt>schedule()</tt>, if RCU detects the end of a given RCU read-side
-critical section, it will necessarily detect the end of all prior
-RCU read-side critical sections, no matter how aggressively the
-compiler scrambles the code.
-
-<p>
-Again, this all assumes that the compiler cannot scramble code across
-calls to the scheduler, out of interrupt handlers, into the idle loop,
-into user-mode code, and so on.
-But if your kernel build allows that sort of scrambling, you have broken
-far more than just RCU!
-
-
-</p><p><a href="#Quick%20Quiz%207"><b>Back to Quick Quiz 7</b>.</a>
-
-<a name="qq8answer"></a>
-<p><b>Quick Quiz 8</b>:
-But how does the upgrade-to-write operation exclude other readers?
-
-
-</p><p><b>Answer</b>:
-It doesn't, just like normal RCU updates, which also do not exclude
-RCU readers.
-
-
-</p><p><a href="#Quick%20Quiz%208"><b>Back to Quick Quiz 8</b>.</a>
-
-<a name="qq9answer"></a>
-<p><b>Quick Quiz 9</b>:
-Can't the compiler also reorder this code?
-
-
-</p><p><b>Answer</b>:
-No, the volatile casts in <tt>READ_ONCE()</tt> and
-<tt>WRITE_ONCE()</tt> prevent the compiler from reordering in
-this particular case.
-
-
-</p><p><a href="#Quick%20Quiz%209"><b>Back to Quick Quiz 9</b>.</a>
-
-<a name="qq10answer"></a>
-<p><b>Quick Quiz 10</b>:
-Suppose that synchronize_rcu() did wait until all readers had completed.
-Would the updater be able to rely on this?
-
-
-</p><p><b>Answer</b>:
-No.
-Even if <tt>synchronize_rcu()</tt> were to wait until
-all readers had completed, a new reader might start immediately after
-<tt>synchronize_rcu()</tt> completed.
-Therefore, the code following
-<tt>synchronize_rcu()</tt> cannot rely on there being no readers
-in any case.
-
-
-</p><p><a href="#Quick%20Quiz%2010"><b>Back to Quick Quiz 10</b>.</a>
-
-<a name="qq11answer"></a>
-<p><b>Quick Quiz 11</b>:
-How long a sequence of grace periods, each separated by an RCU read-side
-critical section, would be required to partition the RCU read-side
-critical sections at the beginning and end of the chain?
-
-
-</p><p><b>Answer</b>:
-In theory, an infinite number.
-In practice, an unknown number that is sensitive to both implementation
-details and timing considerations.
-Therefore, even in practice, RCU users must abide by the theoretical rather
-than the practical answer.
-
-
-</p><p><a href="#Quick%20Quiz%2011"><b>Back to Quick Quiz 11</b>.</a>
-
-<a name="qq12answer"></a>
-<p><b>Quick Quiz 12</b>:
-What about sleeping locks?
-
-
-</p><p><b>Answer</b>:
-These are forbidden within Linux-kernel RCU read-side critical sections
-because it is not legal to place a quiescent state (in this case,
-voluntary context switch) within an RCU read-side critical section.
-However, sleeping locks may be used within userspace RCU read-side critical
-sections, and also within Linux-kernel sleepable RCU
-<a href="#Sleepable RCU">(SRCU)</a>
-read-side critical sections.
-In addition, the -rt patchset turns spinlocks into a sleeping locks so
-that the corresponding critical sections can be preempted, which
-also means that these sleeplockified spinlocks (but not other sleeping locks!)
-may be acquire within -rt-Linux-kernel RCU read-side critical sections.
-
-<p>
-Note that it <i>is</i> legal for a normal RCU read-side critical section
-to conditionally acquire a sleeping locks (as in <tt>mutex_trylock()</tt>),
-but only as long as it does not loop indefinitely attempting to
-conditionally acquire that sleeping locks.
-The key point is that things like <tt>mutex_trylock()</tt>
-either return with the mutex held, or return an error indication if
-the mutex was not immediately available.
-Either way, <tt>mutex_trylock()</tt> returns immediately without sleeping.
-
-
-</p><p><a href="#Quick%20Quiz%2012"><b>Back to Quick Quiz 12</b>.</a>
-
-<a name="qq13answer"></a>
-<p><b>Quick Quiz 13</b>:
-Why does line 19 use <tt>rcu_access_pointer()</tt>?
-After all, <tt>call_rcu()</tt> on line 25 stores into the
-structure, which would interact badly with concurrent insertions.
-Doesn't this mean that <tt>rcu_dereference()</tt> is required?
-
-
-</p><p><b>Answer</b>:
-Presumably the <tt>->gp_lock</tt> acquired on line 18 excludes
-any changes, including any insertions that <tt>rcu_dereference()</tt>
-would protect against.
-Therefore, any insertions will be delayed until after <tt>->gp_lock</tt>
-is released on line 25, which in turn means that
-<tt>rcu_access_pointer()</tt> suffices.
-
-
-</p><p><a href="#Quick%20Quiz%2013"><b>Back to Quick Quiz 13</b>.</a>
-
-<a name="qq14answer"></a>
-<p><b>Quick Quiz 14</b>:
-Earlier it was claimed that <tt>call_rcu()</tt> and
-<tt>kfree_rcu()</tt> allowed updaters to avoid being blocked
-by readers.
-But how can that be correct, given that the invocation of the callback
-and the freeing of the memory (respectively) must still wait for
-a grace period to elapse?
-
-
-</p><p><b>Answer</b>:
-We could define things this way, but keep in mind that this sort of
-definition would say that updates in garbage-collected languages
-cannot complete until the next time the garbage collector runs,
-which does not seem at all reasonable.
-The key point is that in most cases, an updater using either
-<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> can proceed to the
-next update as soon as it has invoked <tt>call_rcu()</tt> or
-<tt>kfree_rcu()</tt>, without having to wait for a subsequent
-grace period.
-
-
-</p><p><a href="#Quick%20Quiz%2014"><b>Back to Quick Quiz 14</b>.</a>
-
-<a name="qq15answer"></a>
-<p><b>Quick Quiz 15</b>:
-So what happens with <tt>synchronize_rcu()</tt> during
-scheduler initialization for <tt>CONFIG_PREEMPT=n</tt>
-kernels?
-
-
-</p><p><b>Answer</b>:
-In <tt>CONFIG_PREEMPT=n</tt> kernel, <tt>synchronize_rcu()</tt>
-maps directly to <tt>synchronize_sched()</tt>.
-Therefore, <tt>synchronize_rcu()</tt> works normally throughout
-boot in <tt>CONFIG_PREEMPT=n</tt> kernels.
-However, your code must also work in <tt>CONFIG_PREEMPT=y</tt> kernels,
-so it is still necessary to avoid invoking <tt>synchronize_rcu()</tt>
-during scheduler initialization.
-
-
-</p><p><a href="#Quick%20Quiz%2015"><b>Back to Quick Quiz 15</b>.</a>
-
-<a name="qq16answer"></a>
-<p><b>Quick Quiz 16</b>:
-But what if I need to wait for multiple RCU flavors, but I also need
-the grace periods to be expedited?
-
-
-</p><p><b>Answer</b>:
-If you are using expedited grace periods, there should be less penalty
-for waiting on them in succession.
-But if that is nevertheless a problem, you can use workqueues or multiple
-kthreads to wait on the various expedited grace periods concurrently.
-
-
-</p><p><a href="#Quick%20Quiz%2016"><b>Back to Quick Quiz 16</b>.</a>
-
-
</body></html>
diff --git a/Documentation/RCU/Design/Requirements/Requirements.htmlx b/Documentation/RCU/Design/Requirements/Requirements.htmlx
deleted file mode 100644
index d6a84f3e0451..000000000000
--- a/Documentation/RCU/Design/Requirements/Requirements.htmlx
+++ /dev/null
@@ -1,2872 +0,0 @@
-<!DOCTYPE HTML PUBLIC "-//W3C//DTD HTML 4.01 Transitional//EN"
- "http://www.w3.org/TR/html4/loose.dtd">
- <html>
- <head><title>A Tour Through RCU's Requirements [LWN.net]</title>
- <meta HTTP-EQUIV="Content-Type" CONTENT="text/html; charset=utf-8">
-
-<h1>A Tour Through RCU's Requirements</h1>
-
-<p>Copyright IBM Corporation, 2015</p>
-<p>Author: Paul E. McKenney</p>
-<p><i>The initial version of this document appeared in the
-<a href="https://lwn.net/">LWN</a> articles
-<a href="https://lwn.net/Articles/652156/">here</a>,
-<a href="https://lwn.net/Articles/652677/">here</a>, and
-<a href="https://lwn.net/Articles/653326/">here</a>.</i></p>
-
-<h2>Introduction</h2>
-
-<p>
-Read-copy update (RCU) is a synchronization mechanism that is often
-used as a replacement for reader-writer locking.
-RCU is unusual in that updaters do not block readers,
-which means that RCU's read-side primitives can be exceedingly fast
-and scalable.
-In addition, updaters can make useful forward progress concurrently
-with readers.
-However, all this concurrency between RCU readers and updaters does raise
-the question of exactly what RCU readers are doing, which in turn
-raises the question of exactly what RCU's requirements are.
-
-<p>
-This document therefore summarizes RCU's requirements, and can be thought
-of as an informal, high-level specification for RCU.
-It is important to understand that RCU's specification is primarily
-empirical in nature;
-in fact, I learned about many of these requirements the hard way.
-This situation might cause some consternation, however, not only
-has this learning process been a lot of fun, but it has also been
-a great privilege to work with so many people willing to apply
-technologies in interesting new ways.
-
-<p>
-All that aside, here are the categories of currently known RCU requirements:
-</p>
-
-<ol>
-<li> <a href="#Fundamental Requirements">
- Fundamental Requirements</a>
-<li> <a href="#Fundamental Non-Requirements">Fundamental Non-Requirements</a>
-<li> <a href="#Parallelism Facts of Life">
- Parallelism Facts of Life</a>
-<li> <a href="#Quality-of-Implementation Requirements">
- Quality-of-Implementation Requirements</a>
-<li> <a href="#Linux Kernel Complications">
- Linux Kernel Complications</a>
-<li> <a href="#Software-Engineering Requirements">
- Software-Engineering Requirements</a>
-<li> <a href="#Other RCU Flavors">
- Other RCU Flavors</a>
-<li> <a href="#Possible Future Changes">
- Possible Future Changes</a>
-</ol>
-
-<p>
-This is followed by a <a href="#Summary">summary</a>,
-which is in turn followed by the inevitable
-<a href="#Answers to Quick Quizzes">answers to the quick quizzes</a>.
-
-<h2><a name="Fundamental Requirements">Fundamental Requirements</a></h2>
-
-<p>
-RCU's fundamental requirements are the closest thing RCU has to hard
-mathematical requirements.
-These are:
-
-<ol>
-<li> <a href="#Grace-Period Guarantee">
- Grace-Period Guarantee</a>
-<li> <a href="#Publish-Subscribe Guarantee">
- Publish-Subscribe Guarantee</a>
-<li> <a href="#Memory-Barrier Guarantees">
- Memory-Barrier Guarantees</a>
-<li> <a href="#RCU Primitives Guaranteed to Execute Unconditionally">
- RCU Primitives Guaranteed to Execute Unconditionally</a>
-<li> <a href="#Guaranteed Read-to-Write Upgrade">
- Guaranteed Read-to-Write Upgrade</a>
-</ol>
-
-<h3><a name="Grace-Period Guarantee">Grace-Period Guarantee</a></h3>
-
-<p>
-RCU's grace-period guarantee is unusual in being premeditated:
-Jack Slingwine and I had this guarantee firmly in mind when we started
-work on RCU (then called “rclock”) in the early 1990s.
-That said, the past two decades of experience with RCU have produced
-a much more detailed understanding of this guarantee.
-
-<p>
-RCU's grace-period guarantee allows updaters to wait for the completion
-of all pre-existing RCU read-side critical sections.
-An RCU read-side critical section
-begins with the marker <tt>rcu_read_lock()</tt> and ends with
-the marker <tt>rcu_read_unlock()</tt>.
-These markers may be nested, and RCU treats a nested set as one
-big RCU read-side critical section.
-Production-quality implementations of <tt>rcu_read_lock()</tt> and
-<tt>rcu_read_unlock()</tt> are extremely lightweight, and in
-fact have exactly zero overhead in Linux kernels built for production
-use with <tt>CONFIG_PREEMPT=n</tt>.
-
-<p>
-This guarantee allows ordering to be enforced with extremely low
-overhead to readers, for example:
-
-<blockquote>
-<pre>
- 1 int x, y;
- 2
- 3 void thread0(void)
- 4 {
- 5 rcu_read_lock();
- 6 r1 = READ_ONCE(x);
- 7 r2 = READ_ONCE(y);
- 8 rcu_read_unlock();
- 9 }
-10
-11 void thread1(void)
-12 {
-13 WRITE_ONCE(x, 1);
-14 synchronize_rcu();
-15 WRITE_ONCE(y, 1);
-16 }
-</pre>
-</blockquote>
-
-<p>
-Because the <tt>synchronize_rcu()</tt> on line 14 waits for
-all pre-existing readers, any instance of <tt>thread0()</tt> that
-loads a value of zero from <tt>x</tt> must complete before
-<tt>thread1()</tt> stores to <tt>y</tt>, so that instance must
-also load a value of zero from <tt>y</tt>.
-Similarly, any instance of <tt>thread0()</tt> that loads a value of
-one from <tt>y</tt> must have started after the
-<tt>synchronize_rcu()</tt> started, and must therefore also load
-a value of one from <tt>x</tt>.
-Therefore, the outcome:
-<blockquote>
-<pre>
-(r1 == 0 && r2 == 1)
-</pre>
-</blockquote>
-cannot happen.
-
-<p>@@QQ@@
-Wait a minute!
-You said that updaters can make useful forward progress concurrently
-with readers, but pre-existing readers will block
-<tt>synchronize_rcu()</tt>!!!
-Just who are you trying to fool???
-<p>@@QQA@@
-First, if updaters do not wish to be blocked by readers, they can use
-<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt>, which will
-be discussed later.
-Second, even when using <tt>synchronize_rcu()</tt>, the other
-update-side code does run concurrently with readers, whether pre-existing
-or not.
-<p>@@QQE@@
-
-<p>
-This scenario resembles one of the first uses of RCU in
-<a href="https://en.wikipedia.org/wiki/DYNIX">DYNIX/ptx</a>,
-which managed a distributed lock manager's transition into
-a state suitable for handling recovery from node failure,
-more or less as follows:
-
-<blockquote>
-<pre>
- 1 #define STATE_NORMAL 0
- 2 #define STATE_WANT_RECOVERY 1
- 3 #define STATE_RECOVERING 2
- 4 #define STATE_WANT_NORMAL 3
- 5
- 6 int state = STATE_NORMAL;
- 7
- 8 void do_something_dlm(void)
- 9 {
-10 int state_snap;
-11
-12 rcu_read_lock();
-13 state_snap = READ_ONCE(state);
-14 if (state_snap == STATE_NORMAL)
-15 do_something();
-16 else
-17 do_something_carefully();
-18 rcu_read_unlock();
-19 }
-20
-21 void start_recovery(void)
-22 {
-23 WRITE_ONCE(state, STATE_WANT_RECOVERY);
-24 synchronize_rcu();
-25 WRITE_ONCE(state, STATE_RECOVERING);
-26 recovery();
-27 WRITE_ONCE(state, STATE_WANT_NORMAL);
-28 synchronize_rcu();
-29 WRITE_ONCE(state, STATE_NORMAL);
-30 }
-</pre>
-</blockquote>
-
-<p>
-The RCU read-side critical section in <tt>do_something_dlm()</tt>
-works with the <tt>synchronize_rcu()</tt> in <tt>start_recovery()</tt>
-to guarantee that <tt>do_something()</tt> never runs concurrently
-with <tt>recovery()</tt>, but with little or no synchronization
-overhead in <tt>do_something_dlm()</tt>.
-
-<p>@@QQ@@
-Why is the <tt>synchronize_rcu()</tt> on line 28 needed?
-<p>@@QQA@@
-Without that extra grace period, memory reordering could result in
-<tt>do_something_dlm()</tt> executing <tt>do_something()</tt>
-concurrently with the last bits of <tt>recovery()</tt>.
-<p>@@QQE@@
-
-<p>
-In order to avoid fatal problems such as deadlocks,
-an RCU read-side critical section must not contain calls to
-<tt>synchronize_rcu()</tt>.
-Similarly, an RCU read-side critical section must not
-contain anything that waits, directly or indirectly, on completion of
-an invocation of <tt>synchronize_rcu()</tt>.
-
-<p>
-Although RCU's grace-period guarantee is useful in and of itself, with
-<a href="https://lwn.net/Articles/573497/">quite a few use cases</a>,
-it would be good to be able to use RCU to coordinate read-side
-access to linked data structures.
-For this, the grace-period guarantee is not sufficient, as can
-be seen in function <tt>add_gp_buggy()</tt> below.
-We will look at the reader's code later, but in the meantime, just think of
-the reader as locklessly picking up the <tt>gp</tt> pointer,
-and, if the value loaded is non-<tt>NULL</tt>, locklessly accessing the
-<tt>->a</tt> and <tt>->b</tt> fields.
-
-<blockquote>
-<pre>
- 1 bool add_gp_buggy(int a, int b)
- 2 {
- 3 p = kmalloc(sizeof(*p), GFP_KERNEL);
- 4 if (!p)
- 5 return -ENOMEM;
- 6 spin_lock(&gp_lock);
- 7 if (rcu_access_pointer(gp)) {
- 8 spin_unlock(&gp_lock);
- 9 return false;
-10 }
-11 p->a = a;
-12 p->b = a;
-13 gp = p; /* ORDERING BUG */
-14 spin_unlock(&gp_lock);
-15 return true;
-16 }
-</pre>
-</blockquote>
-
-<p>
-The problem is that both the compiler and weakly ordered CPUs are within
-their rights to reorder this code as follows:
-
-<blockquote>
-<pre>
- 1 bool add_gp_buggy_optimized(int a, int b)
- 2 {
- 3 p = kmalloc(sizeof(*p), GFP_KERNEL);
- 4 if (!p)
- 5 return -ENOMEM;
- 6 spin_lock(&gp_lock);
- 7 if (rcu_access_pointer(gp)) {
- 8 spin_unlock(&gp_lock);
- 9 return false;
-10 }
-<b>11 gp = p; /* ORDERING BUG */
-12 p->a = a;
-13 p->b = a;</b>
-14 spin_unlock(&gp_lock);
-15 return true;
-16 }
-</pre>
-</blockquote>
-
-<p>
-If an RCU reader fetches <tt>gp</tt> just after
-<tt>add_gp_buggy_optimized</tt> executes line 11,
-it will see garbage in the <tt>->a</tt> and <tt>->b</tt>
-fields.
-And this is but one of many ways in which compiler and hardware optimizations
-could cause trouble.
-Therefore, we clearly need some way to prevent the compiler and the CPU from
-reordering in this manner, which brings us to the publish-subscribe
-guarantee discussed in the next section.
-
-<h3><a name="Publish-Subscribe Guarantee">Publish/Subscribe Guarantee</a></h3>
-
-<p>
-RCU's publish-subscribe guarantee allows data to be inserted
-into a linked data structure without disrupting RCU readers.
-The updater uses <tt>rcu_assign_pointer()</tt> to insert the
-new data, and readers use <tt>rcu_dereference()</tt> to
-access data, whether new or old.
-The following shows an example of insertion:
-
-<blockquote>
-<pre>
- 1 bool add_gp(int a, int b)
- 2 {
- 3 p = kmalloc(sizeof(*p), GFP_KERNEL);
- 4 if (!p)
- 5 return -ENOMEM;
- 6 spin_lock(&gp_lock);
- 7 if (rcu_access_pointer(gp)) {
- 8 spin_unlock(&gp_lock);
- 9 return false;
-10 }
-11 p->a = a;
-12 p->b = a;
-13 rcu_assign_pointer(gp, p);
-14 spin_unlock(&gp_lock);
-15 return true;
-16 }
-</pre>
-</blockquote>
-
-<p>
-The <tt>rcu_assign_pointer()</tt> on line 13 is conceptually
-equivalent to a simple assignment statement, but also guarantees
-that its assignment will
-happen after the two assignments in lines 11 and 12,
-similar to the C11 <tt>memory_order_release</tt> store operation.
-It also prevents any number of “interesting” compiler
-optimizations, for example, the use of <tt>gp</tt> as a scratch
-location immediately preceding the assignment.
-
-<p>@@QQ@@
-But <tt>rcu_assign_pointer()</tt> does nothing to prevent the
-two assignments to <tt>p->a</tt> and <tt>p->b</tt>
-from being reordered.
-Can't that also cause problems?
-<p>@@QQA@@
-No, it cannot.
-The readers cannot see either of these two fields until
-the assignment to <tt>gp</tt>, by which time both fields are
-fully initialized.
-So reordering the assignments
-to <tt>p->a</tt> and <tt>p->b</tt> cannot possibly
-cause any problems.
-<p>@@QQE@@
-
-<p>
-It is tempting to assume that the reader need not do anything special
-to control its accesses to the RCU-protected data,
-as shown in <tt>do_something_gp_buggy()</tt> below:
-
-<blockquote>
-<pre>
- 1 bool do_something_gp_buggy(void)
- 2 {
- 3 rcu_read_lock();
- 4 p = gp; /* OPTIMIZATIONS GALORE!!! */
- 5 if (p) {
- 6 do_something(p->a, p->b);
- 7 rcu_read_unlock();
- 8 return true;
- 9 }
-10 rcu_read_unlock();
-11 return false;
-12 }
-</pre>
-</blockquote>
-
-<p>
-However, this temptation must be resisted because there are a
-surprisingly large number of ways that the compiler
-(to say nothing of
-<a href="https://h71000.www7.hp.com/wizard/wiz_2637.html">DEC Alpha CPUs</a>)
-can trip this code up.
-For but one example, if the compiler were short of registers, it
-might choose to refetch from <tt>gp</tt> rather than keeping
-a separate copy in <tt>p</tt> as follows:
-
-<blockquote>
-<pre>
- 1 bool do_something_gp_buggy_optimized(void)
- 2 {
- 3 rcu_read_lock();
- 4 if (gp) { /* OPTIMIZATIONS GALORE!!! */
-<b> 5 do_something(gp->a, gp->b);</b>
- 6 rcu_read_unlock();
- 7 return true;
- 8 }
- 9 rcu_read_unlock();
-10 return false;
-11 }
-</pre>
-</blockquote>
-
-<p>
-If this function ran concurrently with a series of updates that
-replaced the current structure with a new one,
-the fetches of <tt>gp->a</tt>
-and <tt>gp->b</tt> might well come from two different structures,
-which could cause serious confusion.
-To prevent this (and much else besides), <tt>do_something_gp()</tt> uses
-<tt>rcu_dereference()</tt> to fetch from <tt>gp</tt>:
-
-<blockquote>
-<pre>
- 1 bool do_something_gp(void)
- 2 {
- 3 rcu_read_lock();
- 4 p = rcu_dereference(gp);
- 5 if (p) {
- 6 do_something(p->a, p->b);
- 7 rcu_read_unlock();
- 8 return true;
- 9 }
-10 rcu_read_unlock();
-11 return false;
-12 }
-</pre>
-</blockquote>
-
-<p>
-The <tt>rcu_dereference()</tt> uses volatile casts and (for DEC Alpha)
-memory barriers in the Linux kernel.
-Should a
-<a href="http://www.rdrop.com/users/paulmck/RCU/consume.2015.07.13a.pdf">high-quality implementation of C11 <tt>memory_order_consume</tt> [PDF]</a>
-ever appear, then <tt>rcu_dereference()</tt> could be implemented
-as a <tt>memory_order_consume</tt> load.
-Regardless of the exact implementation, a pointer fetched by
-<tt>rcu_dereference()</tt> may not be used outside of the
-outermost RCU read-side critical section containing that
-<tt>rcu_dereference()</tt>, unless protection of
-the corresponding data element has been passed from RCU to some
-other synchronization mechanism, most commonly locking or
-<a href="https://www.kernel.org/doc/Documentation/RCU/rcuref.txt">reference counting</a>.
-
-<p>
-In short, updaters use <tt>rcu_assign_pointer()</tt> and readers
-use <tt>rcu_dereference()</tt>, and these two RCU API elements
-work together to ensure that readers have a consistent view of
-newly added data elements.
-
-<p>
-Of course, it is also necessary to remove elements from RCU-protected
-data structures, for example, using the following process:
-
-<ol>
-<li> Remove the data element from the enclosing structure.
-<li> Wait for all pre-existing RCU read-side critical sections
- to complete (because only pre-existing readers can possibly have
- a reference to the newly removed data element).
-<li> At this point, only the updater has a reference to the
- newly removed data element, so it can safely reclaim
- the data element, for example, by passing it to <tt>kfree()</tt>.
-</ol>
-
-This process is implemented by <tt>remove_gp_synchronous()</tt>:
-
-<blockquote>
-<pre>
- 1 bool remove_gp_synchronous(void)
- 2 {
- 3 struct foo *p;
- 4
- 5 spin_lock(&gp_lock);
- 6 p = rcu_access_pointer(gp);
- 7 if (!p) {
- 8 spin_unlock(&gp_lock);
- 9 return false;
-10 }
-11 rcu_assign_pointer(gp, NULL);
-12 spin_unlock(&gp_lock);
-13 synchronize_rcu();
-14 kfree(p);
-15 return true;
-16 }
-</pre>
-</blockquote>
-
-<p>
-This function is straightforward, with line 13 waiting for a grace
-period before line 14 frees the old data element.
-This waiting ensures that readers will reach line 7 of
-<tt>do_something_gp()</tt> before the data element referenced by
-<tt>p</tt> is freed.
-The <tt>rcu_access_pointer()</tt> on line 6 is similar to
-<tt>rcu_dereference()</tt>, except that:
-
-<ol>
-<li> The value returned by <tt>rcu_access_pointer()</tt>
- cannot be dereferenced.
- If you want to access the value pointed to as well as
- the pointer itself, use <tt>rcu_dereference()</tt>
- instead of <tt>rcu_access_pointer()</tt>.
-<li> The call to <tt>rcu_access_pointer()</tt> need not be
- protected.
- In contrast, <tt>rcu_dereference()</tt> must either be
- within an RCU read-side critical section or in a code
- segment where the pointer cannot change, for example, in
- code protected by the corresponding update-side lock.
-</ol>
-
-<p>@@QQ@@
-Without the <tt>rcu_dereference()</tt> or the
-<tt>rcu_access_pointer()</tt>, what destructive optimizations
-might the compiler make use of?
-<p>@@QQA@@
-Let's start with what happens to <tt>do_something_gp()</tt>
-if it fails to use <tt>rcu_dereference()</tt>.
-It could reuse a value formerly fetched from this same pointer.
-It could also fetch the pointer from <tt>gp</tt> in a byte-at-a-time
-manner, resulting in <i>load tearing</i>, in turn resulting a bytewise
-mash-up of two distince pointer values.
-It might even use value-speculation optimizations, where it makes a wrong
-guess, but by the time it gets around to checking the value, an update
-has changed the pointer to match the wrong guess.
-Too bad about any dereferences that returned pre-initialization garbage
-in the meantime!
-
-<p>
-For <tt>remove_gp_synchronous()</tt>, as long as all modifications
-to <tt>gp</tt> are carried out while holding <tt>gp_lock</tt>,
-the above optimizations are harmless.
-However,
-with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt>,
-<tt>sparse</tt> will complain if you
-define <tt>gp</tt> with <tt>__rcu</tt> and then
-access it without using
-either <tt>rcu_access_pointer()</tt> or <tt>rcu_dereference()</tt>.
-<p>@@QQE@@
-
-<p>
-In short, RCU's publish-subscribe guarantee is provided by the combination
-of <tt>rcu_assign_pointer()</tt> and <tt>rcu_dereference()</tt>.
-This guarantee allows data elements to be safely added to RCU-protected
-linked data structures without disrupting RCU readers.
-This guarantee can be used in combination with the grace-period
-guarantee to also allow data elements to be removed from RCU-protected
-linked data structures, again without disrupting RCU readers.
-
-<p>
-This guarantee was only partially premeditated.
-DYNIX/ptx used an explicit memory barrier for publication, but had nothing
-resembling <tt>rcu_dereference()</tt> for subscription, nor did it
-have anything resembling the <tt>smp_read_barrier_depends()</tt>
-that was later subsumed into <tt>rcu_dereference()</tt>.
-The need for these operations made itself known quite suddenly at a
-late-1990s meeting with the DEC Alpha architects, back in the days when
-DEC was still a free-standing company.
-It took the Alpha architects a good hour to convince me that any sort
-of barrier would ever be needed, and it then took me a good <i>two</i> hours
-to convince them that their documentation did not make this point clear.
-More recent work with the C and C++ standards committees have provided
-much education on tricks and traps from the compiler.
-In short, compilers were much less tricky in the early 1990s, but in
-2015, don't even think about omitting <tt>rcu_dereference()</tt>!
-
-<h3><a name="Memory-Barrier Guarantees">Memory-Barrier Guarantees</a></h3>
-
-<p>
-The previous section's simple linked-data-structure scenario clearly
-demonstrates the need for RCU's stringent memory-ordering guarantees on
-systems with more than one CPU:
-
-<ol>
-<li> Each CPU that has an RCU read-side critical section that
- begins before <tt>synchronize_rcu()</tt> starts is
- guaranteed to execute a full memory barrier between the time
- that the RCU read-side critical section ends and the time that
- <tt>synchronize_rcu()</tt> returns.
- Without this guarantee, a pre-existing RCU read-side critical section
- might hold a reference to the newly removed <tt>struct foo</tt>
- after the <tt>kfree()</tt> on line 14 of
- <tt>remove_gp_synchronous()</tt>.
-<li> Each CPU that has an RCU read-side critical section that ends
- after <tt>synchronize_rcu()</tt> returns is guaranteed
- to execute a full memory barrier between the time that
- <tt>synchronize_rcu()</tt> begins and the time that the RCU
- read-side critical section begins.
- Without this guarantee, a later RCU read-side critical section
- running after the <tt>kfree()</tt> on line 14 of
- <tt>remove_gp_synchronous()</tt> might
- later run <tt>do_something_gp()</tt> and find the
- newly deleted <tt>struct foo</tt>.
-<li> If the task invoking <tt>synchronize_rcu()</tt> remains
- on a given CPU, then that CPU is guaranteed to execute a full
- memory barrier sometime during the execution of
- <tt>synchronize_rcu()</tt>.
- This guarantee ensures that the <tt>kfree()</tt> on
- line 14 of <tt>remove_gp_synchronous()</tt> really does
- execute after the removal on line 11.
-<li> If the task invoking <tt>synchronize_rcu()</tt> migrates
- among a group of CPUs during that invocation, then each of the
- CPUs in that group is guaranteed to execute a full memory barrier
- sometime during the execution of <tt>synchronize_rcu()</tt>.
- This guarantee also ensures that the <tt>kfree()</tt> on
- line 14 of <tt>remove_gp_synchronous()</tt> really does
- execute after the removal on
- line 11, but also in the case where the thread executing the
- <tt>synchronize_rcu()</tt> migrates in the meantime.
-</ol>
-
-<p>@@QQ@@
-Given that multiple CPUs can start RCU read-side critical sections
-at any time without any ordering whatsoever, how can RCU possibly tell whether
-or not a given RCU read-side critical section starts before a
-given instance of <tt>synchronize_rcu()</tt>?
-<p>@@QQA@@
-If RCU cannot tell whether or not a given
-RCU read-side critical section starts before a
-given instance of <tt>synchronize_rcu()</tt>,
-then it must assume that the RCU read-side critical section
-started first.
-In other words, a given instance of <tt>synchronize_rcu()</tt>
-can avoid waiting on a given RCU read-side critical section only
-if it can prove that <tt>synchronize_rcu()</tt> started first.
-<p>@@QQE@@
-
-<p>@@QQ@@
-The first and second guarantees require unbelievably strict ordering!
-Are all these memory barriers <i> really</i> required?
-<p>@@QQA@@
-Yes, they really are required.
-To see why the first guarantee is required, consider the following
-sequence of events:
-
-<ol>
-<li> CPU 1: <tt>rcu_read_lock()</tt>
-<li> CPU 1: <tt>q = rcu_dereference(gp);
- /* Very likely to return p. */</tt>
-<li> CPU 0: <tt>list_del_rcu(p);</tt>
-<li> CPU 0: <tt>synchronize_rcu()</tt> starts.
-<li> CPU 1: <tt>do_something_with(q->a);
- /* No smp_mb(), so might happen after kfree(). */</tt>
-<li> CPU 1: <tt>rcu_read_unlock()</tt>
-<li> CPU 0: <tt>synchronize_rcu()</tt> returns.
-<li> CPU 0: <tt>kfree(p);</tt>
-</ol>
-
-<p>
-Therefore, there absolutely must be a full memory barrier between the
-end of the RCU read-side critical section and the end of the
-grace period.
-
-<p>
-The sequence of events demonstrating the necessity of the second rule
-is roughly similar:
-
-<ol>
-<li> CPU 0: <tt>list_del_rcu(p);</tt>
-<li> CPU 0: <tt>synchronize_rcu()</tt> starts.
-<li> CPU 1: <tt>rcu_read_lock()</tt>
-<li> CPU 1: <tt>q = rcu_dereference(gp);
- /* Might return p if no memory barrier. */</tt>
-<li> CPU 0: <tt>synchronize_rcu()</tt> returns.
-<li> CPU 0: <tt>kfree(p);</tt>
-<li> CPU 1: <tt>do_something_with(q->a); /* Boom!!! */</tt>
-<li> CPU 1: <tt>rcu_read_unlock()</tt>
-</ol>
-
-<p>
-And similarly, without a memory barrier between the beginning of the
-grace period and the beginning of the RCU read-side critical section,
-CPU 1 might end up accessing the freelist.
-
-<p>
-The “as if” rule of course applies, so that any implementation
-that acts as if the appropriate memory barriers were in place is a
-correct implementation.
-That said, it is much easier to fool yourself into believing that you have
-adhered to the as-if rule than it is to actually adhere to it!
-<p>@@QQE@@
-
-<p>@@QQ@@
-You claim that <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
-generate absolutely no code in some kernel builds.
-This means that the compiler might arbitrarily rearrange consecutive
-RCU read-side critical sections.
-Given such rearrangement, if a given RCU read-side critical section
-is done, how can you be sure that all prior RCU read-side critical
-sections are done?
-Won't the compiler rearrangements make that impossible to determine?
-<p>@@QQA@@
-In cases where <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
-generate absolutely no code, RCU infers quiescent states only at
-special locations, for example, within the scheduler.
-Because calls to <tt>schedule()</tt> had better prevent calling-code
-accesses to shared variables from being rearranged across the call to
-<tt>schedule()</tt>, if RCU detects the end of a given RCU read-side
-critical section, it will necessarily detect the end of all prior
-RCU read-side critical sections, no matter how aggressively the
-compiler scrambles the code.
-
-<p>
-Again, this all assumes that the compiler cannot scramble code across
-calls to the scheduler, out of interrupt handlers, into the idle loop,
-into user-mode code, and so on.
-But if your kernel build allows that sort of scrambling, you have broken
-far more than just RCU!
-<p>@@QQE@@
-
-<p>
-Note that these memory-barrier requirements do not replace the fundamental
-RCU requirement that a grace period wait for all pre-existing readers.
-On the contrary, the memory barriers called out in this section must operate in
-such a way as to <i>enforce</i> this fundamental requirement.
-Of course, different implementations enforce this requirement in different
-ways, but enforce it they must.
-
-<h3><a name="RCU Primitives Guaranteed to Execute Unconditionally">RCU Primitives Guaranteed to Execute Unconditionally</a></h3>
-
-<p>
-The common-case RCU primitives are unconditional.
-They are invoked, they do their job, and they return, with no possibility
-of error, and no need to retry.
-This is a key RCU design philosophy.
-
-<p>
-However, this philosophy is pragmatic rather than pigheaded.
-If someone comes up with a good justification for a particular conditional
-RCU primitive, it might well be implemented and added.
-After all, this guarantee was reverse-engineered, not premeditated.
-The unconditional nature of the RCU primitives was initially an
-accident of implementation, and later experience with synchronization
-primitives with conditional primitives caused me to elevate this
-accident to a guarantee.
-Therefore, the justification for adding a conditional primitive to
-RCU would need to be based on detailed and compelling use cases.
-
-<h3><a name="Guaranteed Read-to-Write Upgrade">Guaranteed Read-to-Write Upgrade</a></h3>
-
-<p>
-As far as RCU is concerned, it is always possible to carry out an
-update within an RCU read-side critical section.
-For example, that RCU read-side critical section might search for
-a given data element, and then might acquire the update-side
-spinlock in order to update that element, all while remaining
-in that RCU read-side critical section.
-Of course, it is necessary to exit the RCU read-side critical section
-before invoking <tt>synchronize_rcu()</tt>, however, this
-inconvenience can be avoided through use of the
-<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt> API members
-described later in this document.
-
-<p>@@QQ@@
-But how does the upgrade-to-write operation exclude other readers?
-<p>@@QQA@@
-It doesn't, just like normal RCU updates, which also do not exclude
-RCU readers.
-<p>@@QQE@@
-
-<p>
-This guarantee allows lookup code to be shared between read-side
-and update-side code, and was premeditated, appearing in the earliest
-DYNIX/ptx RCU documentation.
-
-<h2><a name="Fundamental Non-Requirements">Fundamental Non-Requirements</a></h2>
-
-<p>
-RCU provides extremely lightweight readers, and its read-side guarantees,
-though quite useful, are correspondingly lightweight.
-It is therefore all too easy to assume that RCU is guaranteeing more
-than it really is.
-Of course, the list of things that RCU does not guarantee is infinitely
-long, however, the following sections list a few non-guarantees that
-have caused confusion.
-Except where otherwise noted, these non-guarantees were premeditated.
-
-<ol>
-<li> <a href="#Readers Impose Minimal Ordering">
- Readers Impose Minimal Ordering</a>
-<li> <a href="#Readers Do Not Exclude Updaters">
- Readers Do Not Exclude Updaters</a>
-<li> <a href="#Updaters Only Wait For Old Readers">
- Updaters Only Wait For Old Readers</a>
-<li> <a href="#Grace Periods Don't Partition Read-Side Critical Sections">
- Grace Periods Don't Partition Read-Side Critical Sections</a>
-<li> <a href="#Read-Side Critical Sections Don't Partition Grace Periods">
- Read-Side Critical Sections Don't Partition Grace Periods</a>
-<li> <a href="#Disabling Preemption Does Not Block Grace Periods">
- Disabling Preemption Does Not Block Grace Periods</a>
-</ol>
-
-<h3><a name="Readers Impose Minimal Ordering">Readers Impose Minimal Ordering</a></h3>
-
-<p>
-Reader-side markers such as <tt>rcu_read_lock()</tt> and
-<tt>rcu_read_unlock()</tt> provide absolutely no ordering guarantees
-except through their interaction with the grace-period APIs such as
-<tt>synchronize_rcu()</tt>.
-To see this, consider the following pair of threads:
-
-<blockquote>
-<pre>
- 1 void thread0(void)
- 2 {
- 3 rcu_read_lock();
- 4 WRITE_ONCE(x, 1);
- 5 rcu_read_unlock();
- 6 rcu_read_lock();
- 7 WRITE_ONCE(y, 1);
- 8 rcu_read_unlock();
- 9 }
-10
-11 void thread1(void)
-12 {
-13 rcu_read_lock();
-14 r1 = READ_ONCE(y);
-15 rcu_read_unlock();
-16 rcu_read_lock();
-17 r2 = READ_ONCE(x);
-18 rcu_read_unlock();
-19 }
-</pre>
-</blockquote>
-
-<p>
-After <tt>thread0()</tt> and <tt>thread1()</tt> execute
-concurrently, it is quite possible to have
-
-<blockquote>
-<pre>
-(r1 == 1 && r2 == 0)
-</pre>
-</blockquote>
-
-(that is, <tt>y</tt> appears to have been assigned before <tt>x</tt>),
-which would not be possible if <tt>rcu_read_lock()</tt> and
-<tt>rcu_read_unlock()</tt> had much in the way of ordering
-properties.
-But they do not, so the CPU is within its rights
-to do significant reordering.
-This is by design: Any significant ordering constraints would slow down
-these fast-path APIs.
-
-<p>@@QQ@@
-Can't the compiler also reorder this code?
-<p>@@QQA@@
-No, the volatile casts in <tt>READ_ONCE()</tt> and
-<tt>WRITE_ONCE()</tt> prevent the compiler from reordering in
-this particular case.
-<p>@@QQE@@
-
-<h3><a name="Readers Do Not Exclude Updaters">Readers Do Not Exclude Updaters</a></h3>
-
-<p>
-Neither <tt>rcu_read_lock()</tt> nor <tt>rcu_read_unlock()</tt>
-exclude updates.
-All they do is to prevent grace periods from ending.
-The following example illustrates this:
-
-<blockquote>
-<pre>
- 1 void thread0(void)
- 2 {
- 3 rcu_read_lock();
- 4 r1 = READ_ONCE(y);
- 5 if (r1) {
- 6 do_something_with_nonzero_x();
- 7 r2 = READ_ONCE(x);
- 8 WARN_ON(!r2); /* BUG!!! */
- 9 }
-10 rcu_read_unlock();
-11 }
-12
-13 void thread1(void)
-14 {
-15 spin_lock(&my_lock);
-16 WRITE_ONCE(x, 1);
-17 WRITE_ONCE(y, 1);
-18 spin_unlock(&my_lock);
-19 }
-</pre>
-</blockquote>
-
-<p>
-If the <tt>thread0()</tt> function's <tt>rcu_read_lock()</tt>
-excluded the <tt>thread1()</tt> function's update,
-the <tt>WARN_ON()</tt> could never fire.
-But the fact is that <tt>rcu_read_lock()</tt> does not exclude
-much of anything aside from subsequent grace periods, of which
-<tt>thread1()</tt> has none, so the
-<tt>WARN_ON()</tt> can and does fire.
-
-<h3><a name="Updaters Only Wait For Old Readers">Updaters Only Wait For Old Readers</a></h3>
-
-<p>
-It might be tempting to assume that after <tt>synchronize_rcu()</tt>
-completes, there are no readers executing.
-This temptation must be avoided because
-new readers can start immediately after <tt>synchronize_rcu()</tt>
-starts, and <tt>synchronize_rcu()</tt> is under no
-obligation to wait for these new readers.
-
-<p>@@QQ@@
-Suppose that synchronize_rcu() did wait until all readers had completed.
-Would the updater be able to rely on this?
-<p>@@QQA@@
-No.
-Even if <tt>synchronize_rcu()</tt> were to wait until
-all readers had completed, a new reader might start immediately after
-<tt>synchronize_rcu()</tt> completed.
-Therefore, the code following
-<tt>synchronize_rcu()</tt> cannot rely on there being no readers
-in any case.
-<p>@@QQE@@
-
-<h3><a name="Grace Periods Don't Partition Read-Side Critical Sections">
-Grace Periods Don't Partition Read-Side Critical Sections</a></h3>
-
-<p>
-It is tempting to assume that if any part of one RCU read-side critical
-section precedes a given grace period, and if any part of another RCU
-read-side critical section follows that same grace period, then all of
-the first RCU read-side critical section must precede all of the second.
-However, this just isn't the case: A single grace period does not
-partition the set of RCU read-side critical sections.
-An example of this situation can be illustrated as follows, where
-<tt>x</tt>, <tt>y</tt>, and <tt>z</tt> are initially all zero:
-
-<blockquote>
-<pre>
- 1 void thread0(void)
- 2 {
- 3 rcu_read_lock();
- 4 WRITE_ONCE(a, 1);
- 5 WRITE_ONCE(b, 1);
- 6 rcu_read_unlock();
- 7 }
- 8
- 9 void thread1(void)
-10 {
-11 r1 = READ_ONCE(a);
-12 synchronize_rcu();
-13 WRITE_ONCE(c, 1);
-14 }
-15
-16 void thread2(void)
-17 {
-18 rcu_read_lock();
-19 r2 = READ_ONCE(b);
-20 r3 = READ_ONCE(c);
-21 rcu_read_unlock();
-22 }
-</pre>
-</blockquote>
-
-<p>
-It turns out that the outcome:
-
-<blockquote>
-<pre>
-(r1 == 1 && r2 == 0 && r3 == 1)
-</pre>
-</blockquote>
-
-is entirely possible.
-The following figure show how this can happen, with each circled
-<tt>QS</tt> indicating the point at which RCU recorded a
-<i>quiescent state</i> for each thread, that is, a state in which
-RCU knows that the thread cannot be in the midst of an RCU read-side
-critical section that started before the current grace period:
-
-<p><img src="GPpartitionReaders1.svg" alt="GPpartitionReaders1.svg" width="60%"></p>
-
-<p>
-If it is necessary to partition RCU read-side critical sections in this
-manner, it is necessary to use two grace periods, where the first
-grace period is known to end before the second grace period starts:
-
-<blockquote>
-<pre>
- 1 void thread0(void)
- 2 {
- 3 rcu_read_lock();
- 4 WRITE_ONCE(a, 1);
- 5 WRITE_ONCE(b, 1);
- 6 rcu_read_unlock();
- 7 }
- 8
- 9 void thread1(void)
-10 {
-11 r1 = READ_ONCE(a);
-12 synchronize_rcu();
-13 WRITE_ONCE(c, 1);
-14 }
-15
-16 void thread2(void)
-17 {
-18 r2 = READ_ONCE(c);
-19 synchronize_rcu();
-20 WRITE_ONCE(d, 1);
-21 }
-22
-23 void thread3(void)
-24 {
-25 rcu_read_lock();
-26 r3 = READ_ONCE(b);
-27 r4 = READ_ONCE(d);
-28 rcu_read_unlock();
-29 }
-</pre>
-</blockquote>
-
-<p>
-Here, if <tt>(r1 == 1)</tt>, then
-<tt>thread0()</tt>'s write to <tt>b</tt> must happen
-before the end of <tt>thread1()</tt>'s grace period.
-If in addition <tt>(r4 == 1)</tt>, then
-<tt>thread3()</tt>'s read from <tt>b</tt> must happen
-after the beginning of <tt>thread2()</tt>'s grace period.
-If it is also the case that <tt>(r2 == 1)</tt>, then the
-end of <tt>thread1()</tt>'s grace period must precede the
-beginning of <tt>thread2()</tt>'s grace period.
-This mean that the two RCU read-side critical sections cannot overlap,
-guaranteeing that <tt>(r3 == 1)</tt>.
-As a result, the outcome:
-
-<blockquote>
-<pre>
-(r1 == 1 && r2 == 1 && r3 == 0 && r4 == 1)
-</pre>
-</blockquote>
-
-cannot happen.
-
-<p>
-This non-requirement was also non-premeditated, but became apparent
-when studying RCU's interaction with memory ordering.
-
-<h3><a name="Read-Side Critical Sections Don't Partition Grace Periods">
-Read-Side Critical Sections Don't Partition Grace Periods</a></h3>
-
-<p>
-It is also tempting to assume that if an RCU read-side critical section
-happens between a pair of grace periods, then those grace periods cannot
-overlap.
-However, this temptation leads nowhere good, as can be illustrated by
-the following, with all variables initially zero:
-
-<blockquote>
-<pre>
- 1 void thread0(void)
- 2 {
- 3 rcu_read_lock();
- 4 WRITE_ONCE(a, 1);
- 5 WRITE_ONCE(b, 1);
- 6 rcu_read_unlock();
- 7 }
- 8
- 9 void thread1(void)
-10 {
-11 r1 = READ_ONCE(a);
-12 synchronize_rcu();
-13 WRITE_ONCE(c, 1);
-14 }
-15
-16 void thread2(void)
-17 {
-18 rcu_read_lock();
-19 WRITE_ONCE(d, 1);
-20 r2 = READ_ONCE(c);
-21 rcu_read_unlock();
-22 }
-23
-24 void thread3(void)
-25 {
-26 r3 = READ_ONCE(d);
-27 synchronize_rcu();
-28 WRITE_ONCE(e, 1);
-29 }
-30
-31 void thread4(void)
-32 {
-33 rcu_read_lock();
-34 r4 = READ_ONCE(b);
-35 r5 = READ_ONCE(e);
-36 rcu_read_unlock();
-37 }
-</pre>
-</blockquote>
-
-<p>
-In this case, the outcome:
-
-<blockquote>
-<pre>
-(r1 == 1 && r2 == 1 && r3 == 1 && r4 == 0 && r5 == 1)
-</pre>
-</blockquote>
-
-is entirely possible, as illustrated below:
-
-<p><img src="ReadersPartitionGP1.svg" alt="ReadersPartitionGP1.svg" width="100%"></p>
-
-<p>
-Again, an RCU read-side critical section can overlap almost all of a
-given grace period, just so long as it does not overlap the entire
-grace period.
-As a result, an RCU read-side critical section cannot partition a pair
-of RCU grace periods.
-
-<p>@@QQ@@
-How long a sequence of grace periods, each separated by an RCU read-side
-critical section, would be required to partition the RCU read-side
-critical sections at the beginning and end of the chain?
-<p>@@QQA@@
-In theory, an infinite number.
-In practice, an unknown number that is sensitive to both implementation
-details and timing considerations.
-Therefore, even in practice, RCU users must abide by the theoretical rather
-than the practical answer.
-<p>@@QQE@@
-
-<h3><a name="Disabling Preemption Does Not Block Grace Periods">
-Disabling Preemption Does Not Block Grace Periods</a></h3>
-
-<p>
-There was a time when disabling preemption on any given CPU would block
-subsequent grace periods.
-However, this was an accident of implementation and is not a requirement.
-And in the current Linux-kernel implementation, disabling preemption
-on a given CPU in fact does not block grace periods, as Oleg Nesterov
-<a href="https://lkml.kernel.org/g/20150614193825.GA19582@redhat.com">demonstrated</a>.
-
-<p>
-If you need a preempt-disable region to block grace periods, you need to add
-<tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>, for example
-as follows:
-
-<blockquote>
-<pre>
- 1 preempt_disable();
- 2 rcu_read_lock();
- 3 do_something();
- 4 rcu_read_unlock();
- 5 preempt_enable();
- 6
- 7 /* Spinlocks implicitly disable preemption. */
- 8 spin_lock(&mylock);
- 9 rcu_read_lock();
-10 do_something();
-11 rcu_read_unlock();
-12 spin_unlock(&mylock);
-</pre>
-</blockquote>
-
-<p>
-In theory, you could enter the RCU read-side critical section first,
-but it is more efficient to keep the entire RCU read-side critical
-section contained in the preempt-disable region as shown above.
-Of course, RCU read-side critical sections that extend outside of
-preempt-disable regions will work correctly, but such critical sections
-can be preempted, which forces <tt>rcu_read_unlock()</tt> to do
-more work.
-And no, this is <i>not</i> an invitation to enclose all of your RCU
-read-side critical sections within preempt-disable regions, because
-doing so would degrade real-time response.
-
-<p>
-This non-requirement appeared with preemptible RCU.
-If you need a grace period that waits on non-preemptible code regions, use
-<a href="#Sched Flavor">RCU-sched</a>.
-
-<h2><a name="Parallelism Facts of Life">Parallelism Facts of Life</a></h2>
-
-<p>
-These parallelism facts of life are by no means specific to RCU, but
-the RCU implementation must abide by them.
-They therefore bear repeating:
-
-<ol>
-<li> Any CPU or task may be delayed at any time,
- and any attempts to avoid these delays by disabling
- preemption, interrupts, or whatever are completely futile.
- This is most obvious in preemptible user-level
- environments and in virtualized environments (where
- a given guest OS's VCPUs can be preempted at any time by
- the underlying hypervisor), but can also happen in bare-metal
- environments due to ECC errors, NMIs, and other hardware
- events.
- Although a delay of more than about 20 seconds can result
- in splats, the RCU implementation is obligated to use
- algorithms that can tolerate extremely long delays, but where
- “extremely long” is not long enough to allow
- wrap-around when incrementing a 64-bit counter.
-<li> Both the compiler and the CPU can reorder memory accesses.
- Where it matters, RCU must use compiler directives and
- memory-barrier instructions to preserve ordering.
-<li> Conflicting writes to memory locations in any given cache line
- will result in expensive cache misses.
- Greater numbers of concurrent writes and more-frequent
- concurrent writes will result in more dramatic slowdowns.
- RCU is therefore obligated to use algorithms that have
- sufficient locality to avoid significant performance and
- scalability problems.
-<li> As a rough rule of thumb, only one CPU's worth of processing
- may be carried out under the protection of any given exclusive
- lock.
- RCU must therefore use scalable locking designs.
-<li> Counters are finite, especially on 32-bit systems.
- RCU's use of counters must therefore tolerate counter wrap,
- or be designed such that counter wrap would take way more
- time than a single system is likely to run.
- An uptime of ten years is quite possible, a runtime
- of a century much less so.
- As an example of the latter, RCU's dyntick-idle nesting counter
- allows 54 bits for interrupt nesting level (this counter
- is 64 bits even on a 32-bit system).
- Overflowing this counter requires 2<sup>54</sup>
- half-interrupts on a given CPU without that CPU ever going idle.
- If a half-interrupt happened every microsecond, it would take
- 570 years of runtime to overflow this counter, which is currently
- believed to be an acceptably long time.
-<li> Linux systems can have thousands of CPUs running a single
- Linux kernel in a single shared-memory environment.
- RCU must therefore pay close attention to high-end scalability.
-</ol>
-
-<p>
-This last parallelism fact of life means that RCU must pay special
-attention to the preceding facts of life.
-The idea that Linux might scale to systems with thousands of CPUs would
-have been met with some skepticism in the 1990s, but these requirements
-would have otherwise have been unsurprising, even in the early 1990s.
-
-<h2><a name="Quality-of-Implementation Requirements">Quality-of-Implementation Requirements</a></h2>
-
-<p>
-These sections list quality-of-implementation requirements.
-Although an RCU implementation that ignores these requirements could
-still be used, it would likely be subject to limitations that would
-make it inappropriate for industrial-strength production use.
-Classes of quality-of-implementation requirements are as follows:
-
-<ol>
-<li> <a href="#Specialization">Specialization</a>
-<li> <a href="#Performance and Scalability">Performance and Scalability</a>
-<li> <a href="#Composability">Composability</a>
-<li> <a href="#Corner Cases">Corner Cases</a>
-</ol>
-
-<p>
-These classes is covered in the following sections.
-
-<h3><a name="Specialization">Specialization</a></h3>
-
-<p>
-RCU is and always has been intended primarily for read-mostly situations,
-which means that RCU's read-side primitives are optimized, often at the
-expense of its update-side primitives.
-Experience thus far is captured by the following list of situations:
-
-<ol>
-<li> Read-mostly data, where stale and inconsistent data is not
- a problem: RCU works great!
-<li> Read-mostly data, where data must be consistent:
- RCU works well.
-<li> Read-write data, where data must be consistent:
- RCU <i>might</i> work OK.
- Or not.
-<li> Write-mostly data, where data must be consistent:
- RCU is very unlikely to be the right tool for the job,
- with the following exceptions, where RCU can provide:
- <ol type=a>
- <li> Existence guarantees for update-friendly mechanisms.
- <li> Wait-free read-side primitives for real-time use.
- </ol>
-</ol>
-
-<p>
-This focus on read-mostly situations means that RCU must interoperate
-with other synchronization primitives.
-For example, the <tt>add_gp()</tt> and <tt>remove_gp_synchronous()</tt>
-examples discussed earlier use RCU to protect readers and locking to
-coordinate updaters.
-However, the need extends much farther, requiring that a variety of
-synchronization primitives be legal within RCU read-side critical sections,
-including spinlocks, sequence locks, atomic operations, reference
-counters, and memory barriers.
-
-<p>@@QQ@@
-What about sleeping locks?
-<p>@@QQA@@
-These are forbidden within Linux-kernel RCU read-side critical sections
-because it is not legal to place a quiescent state (in this case,
-voluntary context switch) within an RCU read-side critical section.
-However, sleeping locks may be used within userspace RCU read-side critical
-sections, and also within Linux-kernel sleepable RCU
-<a href="#Sleepable RCU">(SRCU)</a>
-read-side critical sections.
-In addition, the -rt patchset turns spinlocks into a sleeping locks so
-that the corresponding critical sections can be preempted, which
-also means that these sleeplockified spinlocks (but not other sleeping locks!)
-may be acquire within -rt-Linux-kernel RCU read-side critical sections.
-
-<p>
-Note that it <i>is</i> legal for a normal RCU read-side critical section
-to conditionally acquire a sleeping locks (as in <tt>mutex_trylock()</tt>),
-but only as long as it does not loop indefinitely attempting to
-conditionally acquire that sleeping locks.
-The key point is that things like <tt>mutex_trylock()</tt>
-either return with the mutex held, or return an error indication if
-the mutex was not immediately available.
-Either way, <tt>mutex_trylock()</tt> returns immediately without sleeping.
-<p>@@QQE@@
-
-<p>
-It often comes as a surprise that many algorithms do not require a
-consistent view of data, but many can function in that mode,
-with network routing being the poster child.
-Internet routing algorithms take significant time to propagate
-updates, so that by the time an update arrives at a given system,
-that system has been sending network traffic the wrong way for
-a considerable length of time.
-Having a few threads continue to send traffic the wrong way for a
-few more milliseconds is clearly not a problem: In the worst case,
-TCP retransmissions will eventually get the data where it needs to go.
-In general, when tracking the state of the universe outside of the
-computer, some level of inconsistency must be tolerated due to
-speed-of-light delays if nothing else.
-
-<p>
-Furthermore, uncertainty about external state is inherent in many cases.
-For example, a pair of veternarians might use heartbeat to determine
-whether or not a given cat was alive.
-But how long should they wait after the last heartbeat to decide that
-the cat is in fact dead?
-Waiting less than 400 milliseconds makes no sense because this would
-mean that a relaxed cat would be considered to cycle between death
-and life more than 100 times per minute.
-Moreover, just as with human beings, a cat's heart might stop for
-some period of time, so the exact wait period is a judgment call.
-One of our pair of veternarians might wait 30 seconds before pronouncing
-the cat dead, while the other might insist on waiting a full minute.
-The two veternarians would then disagree on the state of the cat during
-the final 30 seconds of the minute following the last heartbeat.
-
-<p>
-Interestingly enough, this same situation applies to hardware.
-When push comes to shove, how do we tell whether or not some
-external server has failed?
-We send messages to it periodically, and declare it failed if we
-don't receive a response within a given period of time.
-Policy decisions can usually tolerate short
-periods of inconsistency.
-The policy was decided some time ago, and is only now being put into
-effect, so a few milliseconds of delay is normally inconsequential.
-
-<p>
-However, there are algorithms that absolutely must see consistent data.
-For example, the translation between a user-level SystemV semaphore
-ID to the corresponding in-kernel data structure is protected by RCU,
-but it is absolutely forbidden to update a semaphore that has just been
-removed.
-In the Linux kernel, this need for consistency is accommodated by acquiring
-spinlocks located in the in-kernel data structure from within
-the RCU read-side critical section, and this is indicated by the
-green box in the figure above.
-Many other techniques may be used, and are in fact used within the
-Linux kernel.
-
-<p>
-In short, RCU is not required to maintain consistency, and other
-mechanisms may be used in concert with RCU when consistency is required.
-RCU's specialization allows it to do its job extremely well, and its
-ability to interoperate with other synchronization mechanisms allows
-the right mix of synchronization tools to be used for a given job.
-
-<h3><a name="Performance and Scalability">Performance and Scalability</a></h3>
-
-<p>
-Energy efficiency is a critical component of performance today,
-and Linux-kernel RCU implementations must therefore avoid unnecessarily
-awakening idle CPUs.
-I cannot claim that this requirement was premeditated.
-In fact, I learned of it during a telephone conversation in which I
-was given “frank and open” feedback on the importance
-of energy efficiency in battery-powered systems and on specific
-energy-efficiency shortcomings of the Linux-kernel RCU implementation.
-In my experience, the battery-powered embedded community will consider
-any unnecessary wakeups to be extremely unfriendly acts.
-So much so that mere Linux-kernel-mailing-list posts are
-insufficient to vent their ire.
-
-<p>
-Memory consumption is not particularly important for in most
-situations, and has become decreasingly
-so as memory sizes have expanded and memory
-costs have plummeted.
-However, as I learned from Matt Mackall's
-<a href="http://elinux.org/Linux_Tiny-FAQ">bloatwatch</a>
-efforts, memory footprint is critically important on single-CPU systems with
-non-preemptible (<tt>CONFIG_PREEMPT=n</tt>) kernels, and thus
-<a href="https://lkml.kernel.org/g/20090113221724.GA15307@linux.vnet.ibm.com">tiny RCU</a>
-was born.
-Josh Triplett has since taken over the small-memory banner with his
-<a href="https://tiny.wiki.kernel.org/">Linux kernel tinification</a>
-project, which resulted in
-<a href="#Sleepable RCU">SRCU</a>
-becoming optional for those kernels not needing it.
-
-<p>
-The remaining performance requirements are, for the most part,
-unsurprising.
-For example, in keeping with RCU's read-side specialization,
-<tt>rcu_dereference()</tt> should have negligible overhead (for
-example, suppression of a few minor compiler optimizations).
-Similarly, in non-preemptible environments, <tt>rcu_read_lock()</tt> and
-<tt>rcu_read_unlock()</tt> should have exactly zero overhead.
-
-<p>
-In preemptible environments, in the case where the RCU read-side
-critical section was not preempted (as will be the case for the
-highest-priority real-time process), <tt>rcu_read_lock()</tt> and
-<tt>rcu_read_unlock()</tt> should have minimal overhead.
-In particular, they should not contain atomic read-modify-write
-operations, memory-barrier instructions, preemption disabling,
-interrupt disabling, or backwards branches.
-However, in the case where the RCU read-side critical section was preempted,
-<tt>rcu_read_unlock()</tt> may acquire spinlocks and disable interrupts.
-This is why it is better to nest an RCU read-side critical section
-within a preempt-disable region than vice versa, at least in cases
-where that critical section is short enough to avoid unduly degrading
-real-time latencies.
-
-<p>
-The <tt>synchronize_rcu()</tt> grace-period-wait primitive is
-optimized for throughput.
-It may therefore incur several milliseconds of latency in addition to
-the duration of the longest RCU read-side critical section.
-On the other hand, multiple concurrent invocations of
-<tt>synchronize_rcu()</tt> are required to use batching optimizations
-so that they can be satisfied by a single underlying grace-period-wait
-operation.
-For example, in the Linux kernel, it is not unusual for a single
-grace-period-wait operation to serve more than
-<a href="https://www.usenix.org/conference/2004-usenix-annual-technical-conference/making-rcu-safe-deep-sub-millisecond-response">1,000 separate invocations</a>
-of <tt>synchronize_rcu()</tt>, thus amortizing the per-invocation
-overhead down to nearly zero.
-However, the grace-period optimization is also required to avoid
-measurable degradation of real-time scheduling and interrupt latencies.
-
-<p>
-In some cases, the multi-millisecond <tt>synchronize_rcu()</tt>
-latencies are unacceptable.
-In these cases, <tt>synchronize_rcu_expedited()</tt> may be used
-instead, reducing the grace-period latency down to a few tens of
-microseconds on small systems, at least in cases where the RCU read-side
-critical sections are short.
-There are currently no special latency requirements for
-<tt>synchronize_rcu_expedited()</tt> on large systems, but,
-consistent with the empirical nature of the RCU specification,
-that is subject to change.
-However, there most definitely are scalability requirements:
-A storm of <tt>synchronize_rcu_expedited()</tt> invocations on 4096
-CPUs should at least make reasonable forward progress.
-In return for its shorter latencies, <tt>synchronize_rcu_expedited()</tt>
-is permitted to impose modest degradation of real-time latency
-on non-idle online CPUs.
-That said, it will likely be necessary to take further steps to reduce this
-degradation, hopefully to roughly that of a scheduling-clock interrupt.
-
-<p>
-There are a number of situations where even
-<tt>synchronize_rcu_expedited()</tt>'s reduced grace-period
-latency is unacceptable.
-In these situations, the asynchronous <tt>call_rcu()</tt> can be
-used in place of <tt>synchronize_rcu()</tt> as follows:
-
-<blockquote>
-<pre>
- 1 struct foo {
- 2 int a;
- 3 int b;
- 4 struct rcu_head rh;
- 5 };
- 6
- 7 static void remove_gp_cb(struct rcu_head *rhp)
- 8 {
- 9 struct foo *p = container_of(rhp, struct foo, rh);
-10
-11 kfree(p);
-12 }
-13
-14 bool remove_gp_asynchronous(void)
-15 {
-16 struct foo *p;
-17
-18 spin_lock(&gp_lock);
-19 p = rcu_dereference(gp);
-20 if (!p) {
-21 spin_unlock(&gp_lock);
-22 return false;
-23 }
-24 rcu_assign_pointer(gp, NULL);
-25 call_rcu(&p->rh, remove_gp_cb);
-26 spin_unlock(&gp_lock);
-27 return true;
-28 }
-</pre>
-</blockquote>
-
-<p>
-A definition of <tt>struct foo</tt> is finally needed, and appears
-on lines 1-5.
-The function <tt>remove_gp_cb()</tt> is passed to <tt>call_rcu()</tt>
-on line 25, and will be invoked after the end of a subsequent
-grace period.
-This gets the same effect as <tt>remove_gp_synchronous()</tt>,
-but without forcing the updater to wait for a grace period to elapse.
-The <tt>call_rcu()</tt> function may be used in a number of
-situations where neither <tt>synchronize_rcu()</tt> nor
-<tt>synchronize_rcu_expedited()</tt> would be legal,
-including within preempt-disable code, <tt>local_bh_disable()</tt> code,
-interrupt-disable code, and interrupt handlers.
-However, even <tt>call_rcu()</tt> is illegal within NMI handlers
-and from offline CPUs.
-The callback function (<tt>remove_gp_cb()</tt> in this case) will be
-executed within softirq (software interrupt) environment within the
-Linux kernel,
-either within a real softirq handler or under the protection
-of <tt>local_bh_disable()</tt>.
-In both the Linux kernel and in userspace, it is bad practice to
-write an RCU callback function that takes too long.
-Long-running operations should be relegated to separate threads or
-(in the Linux kernel) workqueues.
-
-<p>@@QQ@@
-Why does line 19 use <tt>rcu_access_pointer()</tt>?
-After all, <tt>call_rcu()</tt> on line 25 stores into the
-structure, which would interact badly with concurrent insertions.
-Doesn't this mean that <tt>rcu_dereference()</tt> is required?
-<p>@@QQA@@
-Presumably the <tt>->gp_lock</tt> acquired on line 18 excludes
-any changes, including any insertions that <tt>rcu_dereference()</tt>
-would protect against.
-Therefore, any insertions will be delayed until after <tt>->gp_lock</tt>
-is released on line 25, which in turn means that
-<tt>rcu_access_pointer()</tt> suffices.
-<p>@@QQE@@
-
-<p>
-However, all that <tt>remove_gp_cb()</tt> is doing is
-invoking <tt>kfree()</tt> on the data element.
-This is a common idiom, and is supported by <tt>kfree_rcu()</tt>,
-which allows “fire and forget” operation as shown below:
-
-<blockquote>
-<pre>
- 1 struct foo {
- 2 int a;
- 3 int b;
- 4 struct rcu_head rh;
- 5 };
- 6
- 7 bool remove_gp_faf(void)
- 8 {
- 9 struct foo *p;
-10
-11 spin_lock(&gp_lock);
-12 p = rcu_dereference(gp);
-13 if (!p) {
-14 spin_unlock(&gp_lock);
-15 return false;
-16 }
-17 rcu_assign_pointer(gp, NULL);
-18 kfree_rcu(p, rh);
-19 spin_unlock(&gp_lock);
-20 return true;
-21 }
-</pre>
-</blockquote>
-
-<p>
-Note that <tt>remove_gp_faf()</tt> simply invokes
-<tt>kfree_rcu()</tt> and proceeds, without any need to pay any
-further attention to the subsequent grace period and <tt>kfree()</tt>.
-It is permissible to invoke <tt>kfree_rcu()</tt> from the same
-environments as for <tt>call_rcu()</tt>.
-Interestingly enough, DYNIX/ptx had the equivalents of
-<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt>, but not
-<tt>synchronize_rcu()</tt>.
-This was due to the fact that RCU was not heavily used within DYNIX/ptx,
-so the very few places that needed something like
-<tt>synchronize_rcu()</tt> simply open-coded it.
-
-<p>@@QQ@@
-Earlier it was claimed that <tt>call_rcu()</tt> and
-<tt>kfree_rcu()</tt> allowed updaters to avoid being blocked
-by readers.
-But how can that be correct, given that the invocation of the callback
-and the freeing of the memory (respectively) must still wait for
-a grace period to elapse?
-<p>@@QQA@@
-We could define things this way, but keep in mind that this sort of
-definition would say that updates in garbage-collected languages
-cannot complete until the next time the garbage collector runs,
-which does not seem at all reasonable.
-The key point is that in most cases, an updater using either
-<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> can proceed to the
-next update as soon as it has invoked <tt>call_rcu()</tt> or
-<tt>kfree_rcu()</tt>, without having to wait for a subsequent
-grace period.
-<p>@@QQE@@
-
-<p>
-But what if the updater must wait for the completion of code to be
-executed after the end of the grace period, but has other tasks
-that can be carried out in the meantime?
-The polling-style <tt>get_state_synchronize_rcu()</tt> and
-<tt>cond_synchronize_rcu()</tt> functions may be used for this
-purpose, as shown below:
-
-<blockquote>
-<pre>
- 1 bool remove_gp_poll(void)
- 2 {
- 3 struct foo *p;
- 4 unsigned long s;
- 5
- 6 spin_lock(&gp_lock);
- 7 p = rcu_access_pointer(gp);
- 8 if (!p) {
- 9 spin_unlock(&gp_lock);
-10 return false;
-11 }
-12 rcu_assign_pointer(gp, NULL);
-13 spin_unlock(&gp_lock);
-14 s = get_state_synchronize_rcu();
-15 do_something_while_waiting();
-16 cond_synchronize_rcu(s);
-17 kfree(p);
-18 return true;
-19 }
-</pre>
-</blockquote>
-
-<p>
-On line 14, <tt>get_state_synchronize_rcu()</tt> obtains a
-“cookie” from RCU,
-then line 15 carries out other tasks,
-and finally, line 16 returns immediately if a grace period has
-elapsed in the meantime, but otherwise waits as required.
-The need for <tt>get_state_synchronize_rcu</tt> and
-<tt>cond_synchronize_rcu()</tt> has appeared quite recently,
-so it is too early to tell whether they will stand the test of time.
-
-<p>
-RCU thus provides a range of tools to allow updaters to strike the
-required tradeoff between latency, flexibility and CPU overhead.
-
-<h3><a name="Composability">Composability</a></h3>
-
-<p>
-Composability has received much attention in recent years, perhaps in part
-due to the collision of multicore hardware with object-oriented techniques
-designed in single-threaded environments for single-threaded use.
-And in theory, RCU read-side critical sections may be composed, and in
-fact may be nested arbitrarily deeply.
-In practice, as with all real-world implementations of composable
-constructs, there are limitations.
-
-<p>
-Implementations of RCU for which <tt>rcu_read_lock()</tt>
-and <tt>rcu_read_unlock()</tt> generate no code, such as
-Linux-kernel RCU when <tt>CONFIG_PREEMPT=n</tt>, can be
-nested arbitrarily deeply.
-After all, there is no overhead.
-Except that if all these instances of <tt>rcu_read_lock()</tt>
-and <tt>rcu_read_unlock()</tt> are visible to the compiler,
-compilation will eventually fail due to exhausting memory,
-mass storage, or user patience, whichever comes first.
-If the nesting is not visible to the compiler, as is the case with
-mutually recursive functions each in its own translation unit,
-stack overflow will result.
-If the nesting takes the form of loops, either the control variable
-will overflow or (in the Linux kernel) you will get an RCU CPU stall warning.
-Nevertheless, this class of RCU implementations is one
-of the most composable constructs in existence.
-
-<p>
-RCU implementations that explicitly track nesting depth
-are limited by the nesting-depth counter.
-For example, the Linux kernel's preemptible RCU limits nesting to
-<tt>INT_MAX</tt>.
-This should suffice for almost all practical purposes.
-That said, a consecutive pair of RCU read-side critical sections
-between which there is an operation that waits for a grace period
-cannot be enclosed in another RCU read-side critical section.
-This is because it is not legal to wait for a grace period within
-an RCU read-side critical section: To do so would result either
-in deadlock or
-in RCU implicitly splitting the enclosing RCU read-side critical
-section, neither of which is conducive to a long-lived and prosperous
-kernel.
-
-<p>
-It is worth noting that RCU is not alone in limiting composability.
-For example, many transactional-memory implementations prohibit
-composing a pair of transactions separated by an irrevocable
-operation (for example, a network receive operation).
-For another example, lock-based critical sections can be composed
-surprisingly freely, but only if deadlock is avoided.
-
-<p>
-In short, although RCU read-side critical sections are highly composable,
-care is required in some situations, just as is the case for any other
-composable synchronization mechanism.
-
-<h3><a name="Corner Cases">Corner Cases</a></h3>
-
-<p>
-A given RCU workload might have an endless and intense stream of
-RCU read-side critical sections, perhaps even so intense that there
-was never a point in time during which there was not at least one
-RCU read-side critical section in flight.
-RCU cannot allow this situation to block grace periods: As long as
-all the RCU read-side critical sections are finite, grace periods
-must also be finite.
-
-<p>
-That said, preemptible RCU implementations could potentially result
-in RCU read-side critical sections being preempted for long durations,
-which has the effect of creating a long-duration RCU read-side
-critical section.
-This situation can arise only in heavily loaded systems, but systems using
-real-time priorities are of course more vulnerable.
-Therefore, RCU priority boosting is provided to help deal with this
-case.
-That said, the exact requirements on RCU priority boosting will likely
-evolve as more experience accumulates.
-
-<p>
-Other workloads might have very high update rates.
-Although one can argue that such workloads should instead use
-something other than RCU, the fact remains that RCU must
-handle such workloads gracefully.
-This requirement is another factor driving batching of grace periods,
-but it is also the driving force behind the checks for large numbers
-of queued RCU callbacks in the <tt>call_rcu()</tt> code path.
-Finally, high update rates should not delay RCU read-side critical
-sections, although some read-side delays can occur when using
-<tt>synchronize_rcu_expedited()</tt>, courtesy of this function's use
-of <tt>try_stop_cpus()</tt>.
-(In the future, <tt>synchronize_rcu_expedited()</tt> will be
-converted to use lighter-weight inter-processor interrupts (IPIs),
-but this will still disturb readers, though to a much smaller degree.)
-
-<p>
-Although all three of these corner cases were understood in the early
-1990s, a simple user-level test consisting of <tt>close(open(path))</tt>
-in a tight loop
-in the early 2000s suddenly provided a much deeper appreciation of the
-high-update-rate corner case.
-This test also motivated addition of some RCU code to react to high update
-rates, for example, if a given CPU finds itself with more than 10,000
-RCU callbacks queued, it will cause RCU to take evasive action by
-more aggressively starting grace periods and more aggressively forcing
-completion of grace-period processing.
-This evasive action causes the grace period to complete more quickly,
-but at the cost of restricting RCU's batching optimizations, thus
-increasing the CPU overhead incurred by that grace period.
-
-<h2><a name="Software-Engineering Requirements">
-Software-Engineering Requirements</a></h2>
-
-<p>
-Between Murphy's Law and “To err is human”, it is necessary to
-guard against mishaps and misuse:
-
-<ol>
-<li> It is all too easy to forget to use <tt>rcu_read_lock()</tt>
- everywhere that it is needed, so kernels built with
- <tt>CONFIG_PROVE_RCU=y</tt> will spat if
- <tt>rcu_dereference()</tt> is used outside of an
- RCU read-side critical section.
- Update-side code can use <tt>rcu_dereference_protected()</tt>,
- which takes a
- <a href="https://lwn.net/Articles/371986/">lockdep expression</a>
- to indicate what is providing the protection.
- If the indicated protection is not provided, a lockdep splat
- is emitted.
-
- <p>
- Code shared between readers and updaters can use
- <tt>rcu_dereference_check()</tt>, which also takes a
- lockdep expression, and emits a lockdep splat if neither
- <tt>rcu_read_lock()</tt> nor the indicated protection
- is in place.
- In addition, <tt>rcu_dereference_raw()</tt> is used in those
- (hopefully rare) cases where the required protection cannot
- be easily described.
- Finally, <tt>rcu_read_lock_held()</tt> is provided to
- allow a function to verify that it has been invoked within
- an RCU read-side critical section.
- I was made aware of this set of requirements shortly after Thomas
- Gleixner audited a number of RCU uses.
-<li> A given function might wish to check for RCU-related preconditions
- upon entry, before using any other RCU API.
- The <tt>rcu_lockdep_assert()</tt> does this job,
- asserting the expression in kernels having lockdep enabled
- and doing nothing otherwise.
-<li> It is also easy to forget to use <tt>rcu_assign_pointer()</tt>
- and <tt>rcu_dereference()</tt>, perhaps (incorrectly)
- substituting a simple assignment.
- To catch this sort of error, a given RCU-protected pointer may be
- tagged with <tt>__rcu</tt>, after which running sparse
- with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt> will complain
- about simple-assignment accesses to that pointer.
- Arnd Bergmann made me aware of this requirement, and also
- supplied the needed
- <a href="https://lwn.net/Articles/376011/">patch series</a>.
-<li> Kernels built with <tt>CONFIG_DEBUG_OBJECTS_RCU_HEAD=y</tt>
- will splat if a data element is passed to <tt>call_rcu()</tt>
- twice in a row, without a grace period in between.
- (This error is similar to a double free.)
- The corresponding <tt>rcu_head</tt> structures that are
- dynamically allocated are automatically tracked, but
- <tt>rcu_head</tt> structures allocated on the stack
- must be initialized with <tt>init_rcu_head_on_stack()</tt>
- and cleaned up with <tt>destroy_rcu_head_on_stack()</tt>.
- Similarly, statically allocated non-stack <tt>rcu_head</tt>
- structures must be initialized with <tt>init_rcu_head()</tt>
- and cleaned up with <tt>destroy_rcu_head()</tt>.
- Mathieu Desnoyers made me aware of this requirement, and also
- supplied the needed
- <a href="https://lkml.kernel.org/g/20100319013024.GA28456@Krystal">patch</a>.
-<li> An infinite loop in an RCU read-side critical section will
- eventually trigger an RCU CPU stall warning splat, with
- the duration of “eventually” being controlled by the
- <tt>RCU_CPU_STALL_TIMEOUT</tt> <tt>Kconfig</tt> option, or,
- alternatively, by the
- <tt>rcupdate.rcu_cpu_stall_timeout</tt> boot/sysfs
- parameter.
- However, RCU is not obligated to produce this splat
- unless there is a grace period waiting on that particular
- RCU read-side critical section.
- <p>
- Some extreme workloads might intentionally delay
- RCU grace periods, and systems running those workloads can
- be booted with <tt>rcupdate.rcu_cpu_stall_suppress</tt>
- to suppress the splats.
- This kernel parameter may also be set via <tt>sysfs</tt>.
- Furthermore, RCU CPU stall warnings are counter-productive
- during sysrq dumps and during panics.
- RCU therefore supplies the <tt>rcu_sysrq_start()</tt> and
- <tt>rcu_sysrq_end()</tt> API members to be called before
- and after long sysrq dumps.
- RCU also supplies the <tt>rcu_panic()</tt> notifier that is
- automatically invoked at the beginning of a panic to suppress
- further RCU CPU stall warnings.
-
- <p>
- This requirement made itself known in the early 1990s, pretty
- much the first time that it was necessary to debug a CPU stall.
- That said, the initial implementation in DYNIX/ptx was quite
- generic in comparison with that of Linux.
-<li> Although it would be very good to detect pointers leaking out
- of RCU read-side critical sections, there is currently no
- good way of doing this.
- One complication is the need to distinguish between pointers
- leaking and pointers that have been handed off from RCU to
- some other synchronization mechanism, for example, reference
- counting.
-<li> In kernels built with <tt>CONFIG_RCU_TRACE=y</tt>, RCU-related
- information is provided via both debugfs and event tracing.
-<li> Open-coded use of <tt>rcu_assign_pointer()</tt> and
- <tt>rcu_dereference()</tt> to create typical linked
- data structures can be surprisingly error-prone.
- Therefore, RCU-protected
- <a href="https://lwn.net/Articles/609973/#RCU List APIs">linked lists</a>
- and, more recently, RCU-protected
- <a href="https://lwn.net/Articles/612100/">hash tables</a>
- are available.
- Many other special-purpose RCU-protected data structures are
- available in the Linux kernel and the userspace RCU library.
-<li> Some linked structures are created at compile time, but still
- require <tt>__rcu</tt> checking.
- The <tt>RCU_POINTER_INITIALIZER()</tt> macro serves this
- purpose.
-<li> It is not necessary to use <tt>rcu_assign_pointer()</tt>
- when creating linked structures that are to be published via
- a single external pointer.
- The <tt>RCU_INIT_POINTER()</tt> macro is provided for
- this task and also for assigning <tt>NULL</tt> pointers
- at runtime.
-</ol>
-
-<p>
-This not a hard-and-fast list: RCU's diagnostic capabilities will
-continue to be guided by the number and type of usage bugs found
-in real-world RCU usage.
-
-<h2><a name="Linux Kernel Complications">Linux Kernel Complications</a></h2>
-
-<p>
-The Linux kernel provides an interesting environment for all kinds of
-software, including RCU.
-Some of the relevant points of interest are as follows:
-
-<ol>
-<li> <a href="#Configuration">Configuration</a>.
-<li> <a href="#Firmware Interface">Firmware Interface</a>.
-<li> <a href="#Early Boot">Early Boot</a>.
-<li> <a href="#Interrupts and NMIs">
- Interrupts and non-maskable interrupts (NMIs)</a>.
-<li> <a href="#Loadable Modules">Loadable Modules</a>.
-<li> <a href="#Hotplug CPU">Hotplug CPU</a>.
-<li> <a href="#Scheduler and RCU">Scheduler and RCU</a>.
-<li> <a href="#Tracing and RCU">Tracing and RCU</a>.
-<li> <a href="#Energy Efficiency">Energy Efficiency</a>.
-<li> <a href="#Memory Efficiency">Memory Efficiency</a>.
-<li> <a href="#Performance, Scalability, Response Time, and Reliability">
- Performance, Scalability, Response Time, and Reliability</a>.
-</ol>
-
-<p>
-This list is probably incomplete, but it does give a feel for the
-most notable Linux-kernel complications.
-Each of the following sections covers one of the above topics.
-
-<h3><a name="Configuration">Configuration</a></h3>
-
-<p>
-RCU's goal is automatic configuration, so that almost nobody
-needs to worry about RCU's <tt>Kconfig</tt> options.
-And for almost all users, RCU does in fact work well
-“out of the box.”
-
-<p>
-However, there are specialized use cases that are handled by
-kernel boot parameters and <tt>Kconfig</tt> options.
-Unfortunately, the <tt>Kconfig</tt> system will explicitly ask users
-about new <tt>Kconfig</tt> options, which requires almost all of them
-be hidden behind a <tt>CONFIG_RCU_EXPERT</tt> <tt>Kconfig</tt> option.
-
-<p>
-This all should be quite obvious, but the fact remains that
-Linus Torvalds recently had to
-<a href="https://lkml.kernel.org/g/CA+55aFy4wcCwaL4okTs8wXhGZ5h-ibecy_Meg9C4MNQrUnwMcg@mail.gmail.com">remind</a>
-me of this requirement.
-
-<h3><a name="Firmware Interface">Firmware Interface</a></h3>
-
-<p>
-In many cases, kernel obtains information about the system from the
-firmware, and sometimes things are lost in translation.
-Or the translation is accurate, but the original message is bogus.
-
-<p>
-For example, some systems' firmware overreports the number of CPUs,
-sometimes by a large factor.
-If RCU naively believed the firmware, as it used to do,
-it would create too many per-CPU kthreads.
-Although the resulting system will still run correctly, the extra
-kthreads needlessly consume memory and can cause confusion
-when they show up in <tt>ps</tt> listings.
-
-<p>
-RCU must therefore wait for a given CPU to actually come online before
-it can allow itself to believe that the CPU actually exists.
-The resulting “ghost CPUs” (which are never going to
-come online) cause a number of
-<a href="https://paulmck.livejournal.com/37494.html">interesting complications</a>.
-
-<h3><a name="Early Boot">Early Boot</a></h3>
-
-<p>
-The Linux kernel's boot sequence is an interesting process,
-and RCU is used early, even before <tt>rcu_init()</tt>
-is invoked.
-In fact, a number of RCU's primitives can be used as soon as the
-initial task's <tt>task_struct</tt> is available and the
-boot CPU's per-CPU variables are set up.
-The read-side primitives (<tt>rcu_read_lock()</tt>,
-<tt>rcu_read_unlock()</tt>, <tt>rcu_dereference()</tt>,
-and <tt>rcu_access_pointer()</tt>) will operate normally very early on,
-as will <tt>rcu_assign_pointer()</tt>.
-
-<p>
-Although <tt>call_rcu()</tt> may be invoked at any
-time during boot, callbacks are not guaranteed to be invoked until after
-the scheduler is fully up and running.
-This delay in callback invocation is due to the fact that RCU does not
-invoke callbacks until it is fully initialized, and this full initialization
-cannot occur until after the scheduler has initialized itself to the
-point where RCU can spawn and run its kthreads.
-In theory, it would be possible to invoke callbacks earlier,
-however, this is not a panacea because there would be severe restrictions
-on what operations those callbacks could invoke.
-
-<p>
-Perhaps surprisingly, <tt>synchronize_rcu()</tt>,
-<a href="#Bottom-Half Flavor"><tt>synchronize_rcu_bh()</tt></a>
-(<a href="#Bottom-Half Flavor">discussed below</a>),
-and
-<a href="#Sched Flavor"><tt>synchronize_sched()</tt></a>
-will all operate normally
-during very early boot, the reason being that there is only one CPU
-and preemption is disabled.
-This means that the call <tt>synchronize_rcu()</tt> (or friends)
-itself is a quiescent
-state and thus a grace period, so the early-boot implementation can
-be a no-op.
-
-<p>
-Both <tt>synchronize_rcu_bh()</tt> and <tt>synchronize_sched()</tt>
-continue to operate normally through the remainder of boot, courtesy
-of the fact that preemption is disabled across their RCU read-side
-critical sections and also courtesy of the fact that there is still
-only one CPU.
-However, once the scheduler starts initializing, preemption is enabled.
-There is still only a single CPU, but the fact that preemption is enabled
-means that the no-op implementation of <tt>synchronize_rcu()</tt> no
-longer works in <tt>CONFIG_PREEMPT=y</tt> kernels.
-Therefore, as soon as the scheduler starts initializing, the early-boot
-fastpath is disabled.
-This means that <tt>synchronize_rcu()</tt> switches to its runtime
-mode of operation where it posts callbacks, which in turn means that
-any call to <tt>synchronize_rcu()</tt> will block until the corresponding
-callback is invoked.
-Unfortunately, the callback cannot be invoked until RCU's runtime
-grace-period machinery is up and running, which cannot happen until
-the scheduler has initialized itself sufficiently to allow RCU's
-kthreads to be spawned.
-Therefore, invoking <tt>synchronize_rcu()</tt> during scheduler
-initialization can result in deadlock.
-
-<p>@@QQ@@
-So what happens with <tt>synchronize_rcu()</tt> during
-scheduler initialization for <tt>CONFIG_PREEMPT=n</tt>
-kernels?
-<p>@@QQA@@
-In <tt>CONFIG_PREEMPT=n</tt> kernel, <tt>synchronize_rcu()</tt>
-maps directly to <tt>synchronize_sched()</tt>.
-Therefore, <tt>synchronize_rcu()</tt> works normally throughout
-boot in <tt>CONFIG_PREEMPT=n</tt> kernels.
-However, your code must also work in <tt>CONFIG_PREEMPT=y</tt> kernels,
-so it is still necessary to avoid invoking <tt>synchronize_rcu()</tt>
-during scheduler initialization.
-<p>@@QQE@@
-
-<p>
-I learned of these boot-time requirements as a result of a series of
-system hangs.
-
-<h3><a name="Interrupts and NMIs">Interrupts and NMIs</a></h3>
-
-<p>
-The Linux kernel has interrupts, and RCU read-side critical sections are
-legal within interrupt handlers and within interrupt-disabled regions
-of code, as are invocations of <tt>call_rcu()</tt>.
-
-<p>
-Some Linux-kernel architectures can enter an interrupt handler from
-non-idle process context, and then just never leave it, instead stealthily
-transitioning back to process context.
-This trick is sometimes used to invoke system calls from inside the kernel.
-These “half-interrupts” mean that RCU has to be very careful
-about how it counts interrupt nesting levels.
-I learned of this requirement the hard way during a rewrite
-of RCU's dyntick-idle code.
-
-<p>
-The Linux kernel has non-maskable interrupts (NMIs), and
-RCU read-side critical sections are legal within NMI handlers.
-Thankfully, RCU update-side primitives, including
-<tt>call_rcu()</tt>, are prohibited within NMI handlers.
-
-<p>
-The name notwithstanding, some Linux-kernel architectures
-can have nested NMIs, which RCU must handle correctly.
-Andy Lutomirski
-<a href="https://lkml.kernel.org/g/CALCETrXLq1y7e_dKFPgou-FKHB6Pu-r8+t-6Ds+8=va7anBWDA@mail.gmail.com">surprised me</a>
-with this requirement;
-he also kindly surprised me with
-<a href="https://lkml.kernel.org/g/CALCETrXSY9JpW3uE6H8WYk81sg56qasA2aqmjMPsq5dOtzso=g@mail.gmail.com">an algorithm</a>
-that meets this requirement.
-
-<h3><a name="Loadable Modules">Loadable Modules</a></h3>
-
-<p>
-The Linux kernel has loadable modules, and these modules can
-also be unloaded.
-After a given module has been unloaded, any attempt to call
-one of its functions results in a segmentation fault.
-The module-unload functions must therefore cancel any
-delayed calls to loadable-module functions, for example,
-any outstanding <tt>mod_timer()</tt> must be dealt with
-via <tt>del_timer_sync()</tt> or similar.
-
-<p>
-Unfortunately, there is no way to cancel an RCU callback;
-once you invoke <tt>call_rcu()</tt>, the callback function is
-going to eventually be invoked, unless the system goes down first.
-Because it is normally considered socially irresponsible to crash the system
-in response to a module unload request, we need some other way
-to deal with in-flight RCU callbacks.
-
-<p>
-RCU therefore provides
-<tt><a href="https://lwn.net/Articles/217484/">rcu_barrier()</a></tt>,
-which waits until all in-flight RCU callbacks have been invoked.
-If a module uses <tt>call_rcu()</tt>, its exit function should therefore
-prevent any future invocation of <tt>call_rcu()</tt>, then invoke
-<tt>rcu_barrier()</tt>.
-In theory, the underlying module-unload code could invoke
-<tt>rcu_barrier()</tt> unconditionally, but in practice this would
-incur unacceptable latencies.
-
-<p>
-Nikita Danilov noted this requirement for an analogous filesystem-unmount
-situation, and Dipankar Sarma incorporated <tt>rcu_barrier()</tt> into RCU.
-The need for <tt>rcu_barrier()</tt> for module unloading became
-apparent later.
-
-<h3><a name="Hotplug CPU">Hotplug CPU</a></h3>
-
-<p>
-The Linux kernel supports CPU hotplug, which means that CPUs
-can come and go.
-It is of course illegal to use any RCU API member from an offline CPU.
-This requirement was present from day one in DYNIX/ptx, but
-on the other hand, the Linux kernel's CPU-hotplug implementation
-is “interesting.”
-
-<p>
-The Linux-kernel CPU-hotplug implementation has notifiers that
-are used to allow the various kernel subsystems (including RCU)
-to respond appropriately to a given CPU-hotplug operation.
-Most RCU operations may be invoked from CPU-hotplug notifiers,
-including even normal synchronous grace-period operations
-such as <tt>synchronize_rcu()</tt>.
-However, expedited grace-period operations such as
-<tt>synchronize_rcu_expedited()</tt> are not supported,
-due to the fact that current implementations block CPU-hotplug
-operations, which could result in deadlock.
-
-<p>
-In addition, all-callback-wait operations such as
-<tt>rcu_barrier()</tt> are also not supported, due to the
-fact that there are phases of CPU-hotplug operations where
-the outgoing CPU's callbacks will not be invoked until after
-the CPU-hotplug operation ends, which could also result in deadlock.
-
-<h3><a name="Scheduler and RCU">Scheduler and RCU</a></h3>
-
-<p>
-RCU depends on the scheduler, and the scheduler uses RCU to
-protect some of its data structures.
-This means the scheduler is forbidden from acquiring
-the runqueue locks and the priority-inheritance locks
-in the middle of an outermost RCU read-side critical section unless either
-(1) it releases them before exiting that same
-RCU read-side critical section, or
-(2) interrupts are disabled across
-that entire RCU read-side critical section.
-This same prohibition also applies (recursively!) to any lock that is acquired
-while holding any lock to which this prohibition applies.
-Adhering to this rule prevents preemptible RCU from invoking
-<tt>rcu_read_unlock_special()</tt> while either runqueue or
-priority-inheritance locks are held, thus avoiding deadlock.
-
-<p>
-Prior to v4.4, it was only necessary to disable preemption across
-RCU read-side critical sections that acquired scheduler locks.
-In v4.4, expedited grace periods started using IPIs, and these
-IPIs could force a <tt>rcu_read_unlock()</tt> to take the slowpath.
-Therefore, this expedited-grace-period change required disabling of
-interrupts, not just preemption.
-
-<p>
-For RCU's part, the preemptible-RCU <tt>rcu_read_unlock()</tt>
-implementation must be written carefully to avoid similar deadlocks.
-In particular, <tt>rcu_read_unlock()</tt> must tolerate an
-interrupt where the interrupt handler invokes both
-<tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>.
-This possibility requires <tt>rcu_read_unlock()</tt> to use
-negative nesting levels to avoid destructive recursion via
-interrupt handler's use of RCU.
-
-<p>
-This pair of mutual scheduler-RCU requirements came as a
-<a href="https://lwn.net/Articles/453002/">complete surprise</a>.
-
-<p>
-As noted above, RCU makes use of kthreads, and it is necessary to
-avoid excessive CPU-time accumulation by these kthreads.
-This requirement was no surprise, but RCU's violation of it
-when running context-switch-heavy workloads when built with
-<tt>CONFIG_NO_HZ_FULL=y</tt>
-<a href="http://www.rdrop.com/users/paulmck/scalability/paper/BareMetal.2015.01.15b.pdf">did come as a surprise [PDF]</a>.
-RCU has made good progress towards meeting this requirement, even
-for context-switch-have <tt>CONFIG_NO_HZ_FULL=y</tt> workloads,
-but there is room for further improvement.
-
-<h3><a name="Tracing and RCU">Tracing and RCU</a></h3>
-
-<p>
-It is possible to use tracing on RCU code, but tracing itself
-uses RCU.
-For this reason, <tt>rcu_dereference_raw_notrace()</tt>
-is provided for use by tracing, which avoids the destructive
-recursion that could otherwise ensue.
-This API is also used by virtualization in some architectures,
-where RCU readers execute in environments in which tracing
-cannot be used.
-The tracing folks both located the requirement and provided the
-needed fix, so this surprise requirement was relatively painless.
-
-<h3><a name="Energy Efficiency">Energy Efficiency</a></h3>
-
-<p>
-Interrupting idle CPUs is considered socially unacceptable,
-especially by people with battery-powered embedded systems.
-RCU therefore conserves energy by detecting which CPUs are
-idle, including tracking CPUs that have been interrupted from idle.
-This is a large part of the energy-efficiency requirement,
-so I learned of this via an irate phone call.
-
-<p>
-Because RCU avoids interrupting idle CPUs, it is illegal to
-execute an RCU read-side critical section on an idle CPU.
-(Kernels built with <tt>CONFIG_PROVE_RCU=y</tt> will splat
-if you try it.)
-The <tt>RCU_NONIDLE()</tt> macro and <tt>_rcuidle</tt>
-event tracing is provided to work around this restriction.
-In addition, <tt>rcu_is_watching()</tt> may be used to
-test whether or not it is currently legal to run RCU read-side
-critical sections on this CPU.
-I learned of the need for diagnostics on the one hand
-and <tt>RCU_NONIDLE()</tt> on the other while inspecting
-idle-loop code.
-Steven Rostedt supplied <tt>_rcuidle</tt> event tracing,
-which is used quite heavily in the idle loop.
-
-<p>
-It is similarly socially unacceptable to interrupt an
-<tt>nohz_full</tt> CPU running in userspace.
-RCU must therefore track <tt>nohz_full</tt> userspace
-execution.
-And in
-<a href="https://lwn.net/Articles/558284/"><tt>CONFIG_NO_HZ_FULL_SYSIDLE=y</tt></a>
-kernels, RCU must separately track idle CPUs on the one hand and
-CPUs that are either idle or executing in userspace on the other.
-In both cases, RCU must be able to sample state at two points in
-time, and be able to determine whether or not some other CPU spent
-any time idle and/or executing in userspace.
-
-<p>
-These energy-efficiency requirements have proven quite difficult to
-understand and to meet, for example, there have been more than five
-clean-sheet rewrites of RCU's energy-efficiency code, the last of
-which was finally able to demonstrate
-<a href="http://www.rdrop.com/users/paulmck/realtime/paper/AMPenergy.2013.04.19a.pdf">real energy savings running on real hardware [PDF]</a>.
-As noted earlier,
-I learned of many of these requirements via angry phone calls:
-Flaming me on the Linux-kernel mailing list was apparently not
-sufficient to fully vent their ire at RCU's energy-efficiency bugs!
-
-<h3><a name="Memory Efficiency">Memory Efficiency</a></h3>
-
-<p>
-Although small-memory non-realtime systems can simply use Tiny RCU,
-code size is only one aspect of memory efficiency.
-Another aspect is the size of the <tt>rcu_head</tt> structure
-used by <tt>call_rcu()</tt> and <tt>kfree_rcu()</tt>.
-Although this structure contains nothing more than a pair of pointers,
-it does appear in many RCU-protected data structures, including
-some that are size critical.
-The <tt>page</tt> structure is a case in point, as evidenced by
-the many occurrences of the <tt>union</tt> keyword within that structure.
-
-<p>
-This need for memory efficiency is one reason that RCU uses hand-crafted
-singly linked lists to track the <tt>rcu_head</tt> structures that
-are waiting for a grace period to elapse.
-It is also the reason why <tt>rcu_head</tt> structures do not contain
-debug information, such as fields tracking the file and line of the
-<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> that posted them.
-Although this information might appear in debug-only kernel builds at some
-point, in the meantime, the <tt>->func</tt> field will often provide
-the needed debug information.
-
-<p>
-However, in some cases, the need for memory efficiency leads to even
-more extreme measures.
-Returning to the <tt>page</tt> structure, the <tt>rcu_head</tt> field
-shares storage with a great many other structures that are used at
-various points in the corresponding page's lifetime.
-In order to correctly resolve certain
-<a href="https://lkml.kernel.org/g/1439976106-137226-1-git-send-email-kirill.shutemov@linux.intel.com">race conditions</a>,
-the Linux kernel's memory-management subsystem needs a particular bit
-to remain zero during all phases of grace-period processing,
-and that bit happens to map to the bottom bit of the
-<tt>rcu_head</tt> structure's <tt>->next</tt> field.
-RCU makes this guarantee as long as <tt>call_rcu()</tt>
-is used to post the callback, as opposed to <tt>kfree_rcu()</tt>
-or some future “lazy”
-variant of <tt>call_rcu()</tt> that might one day be created for
-energy-efficiency purposes.
-
-<h3><a name="Performance, Scalability, Response Time, and Reliability">
-Performance, Scalability, Response Time, and Reliability</a></h3>
-
-<p>
-Expanding on the
-<a href="#Performance and Scalability">earlier discussion</a>,
-RCU is used heavily by hot code paths in performance-critical
-portions of the Linux kernel's networking, security, virtualization,
-and scheduling code paths.
-RCU must therefore use efficient implementations, especially in its
-read-side primitives.
-To that end, it would be good if preemptible RCU's implementation
-of <tt>rcu_read_lock()</tt> could be inlined, however, doing
-this requires resolving <tt>#include</tt> issues with the
-<tt>task_struct</tt> structure.
-
-<p>
-The Linux kernel supports hardware configurations with up to
-4096 CPUs, which means that RCU must be extremely scalable.
-Algorithms that involve frequent acquisitions of global locks or
-frequent atomic operations on global variables simply cannot be
-tolerated within the RCU implementation.
-RCU therefore makes heavy use of a combining tree based on the
-<tt>rcu_node</tt> structure.
-RCU is required to tolerate all CPUs continuously invoking any
-combination of RCU's runtime primitives with minimal per-operation
-overhead.
-In fact, in many cases, increasing load must <i>decrease</i> the
-per-operation overhead, witness the batching optimizations for
-<tt>synchronize_rcu()</tt>, <tt>call_rcu()</tt>,
-<tt>synchronize_rcu_expedited()</tt>, and <tt>rcu_barrier()</tt>.
-As a general rule, RCU must cheerfully accept whatever the
-rest of the Linux kernel decides to throw at it.
-
-<p>
-The Linux kernel is used for real-time workloads, especially
-in conjunction with the
-<a href="https://rt.wiki.kernel.org/index.php/Main_Page">-rt patchset</a>.
-The real-time-latency response requirements are such that the
-traditional approach of disabling preemption across RCU
-read-side critical sections is inappropriate.
-Kernels built with <tt>CONFIG_PREEMPT=y</tt> therefore
-use an RCU implementation that allows RCU read-side critical
-sections to be preempted.
-This requirement made its presence known after users made it
-clear that an earlier
-<a href="https://lwn.net/Articles/107930/">real-time patch</a>
-did not meet their needs, in conjunction with some
-<a href="https://lkml.kernel.org/g/20050318002026.GA2693@us.ibm.com">RCU issues</a>
-encountered by a very early version of the -rt patchset.
-
-<p>
-In addition, RCU must make do with a sub-100-microsecond real-time latency
-budget.
-In fact, on smaller systems with the -rt patchset, the Linux kernel
-provides sub-20-microsecond real-time latencies for the whole kernel,
-including RCU.
-RCU's scalability and latency must therefore be sufficient for
-these sorts of configurations.
-To my surprise, the sub-100-microsecond real-time latency budget
-<a href="http://www.rdrop.com/users/paulmck/realtime/paper/bigrt.2013.01.31a.LCA.pdf">
-applies to even the largest systems [PDF]</a>,
-up to and including systems with 4096 CPUs.
-This real-time requirement motivated the grace-period kthread, which
-also simplified handling of a number of race conditions.
-
-<p>
-RCU must avoid degrading real-time response for CPU-bound threads, whether
-executing in usermode (which is one use case for
-<tt>CONFIG_NO_HZ_FULL=y</tt>) or in the kernel.
-That said, CPU-bound loops in the kernel must execute
-<tt>cond_resched_rcu_qs()</tt> at least once per few tens of milliseconds
-in order to avoid receiving an IPI from RCU.
-
-<p>
-Finally, RCU's status as a synchronization primitive means that
-any RCU failure can result in arbitrary memory corruption that can be
-extremely difficult to debug.
-This means that RCU must be extremely reliable, which in
-practice also means that RCU must have an aggressive stress-test
-suite.
-This stress-test suite is called <tt>rcutorture</tt>.
-
-<p>
-Although the need for <tt>rcutorture</tt> was no surprise,
-the current immense popularity of the Linux kernel is posing
-interesting—and perhaps unprecedented—validation
-challenges.
-To see this, keep in mind that there are well over one billion
-instances of the Linux kernel running today, given Android
-smartphones, Linux-powered televisions, and servers.
-This number can be expected to increase sharply with the advent of
-the celebrated Internet of Things.
-
-<p>
-Suppose that RCU contains a race condition that manifests on average
-once per million years of runtime.
-This bug will be occurring about three times per <i>day</i> across
-the installed base.
-RCU could simply hide behind hardware error rates, given that no one
-should really expect their smartphone to last for a million years.
-However, anyone taking too much comfort from this thought should
-consider the fact that in most jurisdictions, a successful multi-year
-test of a given mechanism, which might include a Linux kernel,
-suffices for a number of types of safety-critical certifications.
-In fact, rumor has it that the Linux kernel is already being used
-in production for safety-critical applications.
-I don't know about you, but I would feel quite bad if a bug in RCU
-killed someone.
-Which might explain my recent focus on validation and verification.
-
-<h2><a name="Other RCU Flavors">Other RCU Flavors</a></h2>
-
-<p>
-One of the more surprising things about RCU is that there are now
-no fewer than five <i>flavors</i>, or API families.
-In addition, the primary flavor that has been the sole focus up to
-this point has two different implementations, non-preemptible and
-preemptible.
-The other four flavors are listed below, with requirements for each
-described in a separate section.
-
-<ol>
-<li> <a href="#Bottom-Half Flavor">Bottom-Half Flavor</a>
-<li> <a href="#Sched Flavor">Sched Flavor</a>
-<li> <a href="#Sleepable RCU">Sleepable RCU</a>
-<li> <a href="#Tasks RCU">Tasks RCU</a>
-<li> <a href="#Waiting for Multiple Grace Periods">
- Waiting for Multiple Grace Periods</a>
-</ol>
-
-<h3><a name="Bottom-Half Flavor">Bottom-Half Flavor</a></h3>
-
-<p>
-The softirq-disable (AKA “bottom-half”,
-hence the “_bh” abbreviations)
-flavor of RCU, or <i>RCU-bh</i>, was developed by
-Dipankar Sarma to provide a flavor of RCU that could withstand the
-network-based denial-of-service attacks researched by Robert
-Olsson.
-These attacks placed so much networking load on the system
-that some of the CPUs never exited softirq execution,
-which in turn prevented those CPUs from ever executing a context switch,
-which, in the RCU implementation of that time, prevented grace periods
-from ever ending.
-The result was an out-of-memory condition and a system hang.
-
-<p>
-The solution was the creation of RCU-bh, which does
-<tt>local_bh_disable()</tt>
-across its read-side critical sections, and which uses the transition
-from one type of softirq processing to another as a quiescent state
-in addition to context switch, idle, user mode, and offline.
-This means that RCU-bh grace periods can complete even when some of
-the CPUs execute in softirq indefinitely, thus allowing algorithms
-based on RCU-bh to withstand network-based denial-of-service attacks.
-
-<p>
-Because
-<tt>rcu_read_lock_bh()</tt> and <tt>rcu_read_unlock_bh()</tt>
-disable and re-enable softirq handlers, any attempt to start a softirq
-handlers during the
-RCU-bh read-side critical section will be deferred.
-In this case, <tt>rcu_read_unlock_bh()</tt>
-will invoke softirq processing, which can take considerable time.
-One can of course argue that this softirq overhead should be associated
-with the code following the RCU-bh read-side critical section rather
-than <tt>rcu_read_unlock_bh()</tt>, but the fact
-is that most profiling tools cannot be expected to make this sort
-of fine distinction.
-For example, suppose that a three-millisecond-long RCU-bh read-side
-critical section executes during a time of heavy networking load.
-There will very likely be an attempt to invoke at least one softirq
-handler during that three milliseconds, but any such invocation will
-be delayed until the time of the <tt>rcu_read_unlock_bh()</tt>.
-This can of course make it appear at first glance as if
-<tt>rcu_read_unlock_bh()</tt> was executing very slowly.
-
-<p>
-The
-<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">RCU-bh API</a>
-includes
-<tt>rcu_read_lock_bh()</tt>,
-<tt>rcu_read_unlock_bh()</tt>,
-<tt>rcu_dereference_bh()</tt>,
-<tt>rcu_dereference_bh_check()</tt>,
-<tt>synchronize_rcu_bh()</tt>,
-<tt>synchronize_rcu_bh_expedited()</tt>,
-<tt>call_rcu_bh()</tt>,
-<tt>rcu_barrier_bh()</tt>, and
-<tt>rcu_read_lock_bh_held()</tt>.
-
-<h3><a name="Sched Flavor">Sched Flavor</a></h3>
-
-<p>
-Before preemptible RCU, waiting for an RCU grace period had the
-side effect of also waiting for all pre-existing interrupt
-and NMI handlers.
-However, there are legitimate preemptible-RCU implementations that
-do not have this property, given that any point in the code outside
-of an RCU read-side critical section can be a quiescent state.
-Therefore, <i>RCU-sched</i> was created, which follows “classic”
-RCU in that an RCU-sched grace period waits for for pre-existing
-interrupt and NMI handlers.
-In kernels built with <tt>CONFIG_PREEMPT=n</tt>, the RCU and RCU-sched
-APIs have identical implementations, while kernels built with
-<tt>CONFIG_PREEMPT=y</tt> provide a separate implementation for each.
-
-<p>
-Note well that in <tt>CONFIG_PREEMPT=y</tt> kernels,
-<tt>rcu_read_lock_sched()</tt> and <tt>rcu_read_unlock_sched()</tt>
-disable and re-enable preemption, respectively.
-This means that if there was a preemption attempt during the
-RCU-sched read-side critical section, <tt>rcu_read_unlock_sched()</tt>
-will enter the scheduler, with all the latency and overhead entailed.
-Just as with <tt>rcu_read_unlock_bh()</tt>, this can make it look
-as if <tt>rcu_read_unlock_sched()</tt> was executing very slowly.
-However, the highest-priority task won't be preempted, so that task
-will enjoy low-overhead <tt>rcu_read_unlock_sched()</tt> invocations.
-
-<p>
-The
-<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">RCU-sched API</a>
-includes
-<tt>rcu_read_lock_sched()</tt>,
-<tt>rcu_read_unlock_sched()</tt>,
-<tt>rcu_read_lock_sched_notrace()</tt>,
-<tt>rcu_read_unlock_sched_notrace()</tt>,
-<tt>rcu_dereference_sched()</tt>,
-<tt>rcu_dereference_sched_check()</tt>,
-<tt>synchronize_sched()</tt>,
-<tt>synchronize_rcu_sched_expedited()</tt>,
-<tt>call_rcu_sched()</tt>,
-<tt>rcu_barrier_sched()</tt>, and
-<tt>rcu_read_lock_sched_held()</tt>.
-However, anything that disables preemption also marks an RCU-sched
-read-side critical section, including
-<tt>preempt_disable()</tt> and <tt>preempt_enable()</tt>,
-<tt>local_irq_save()</tt> and <tt>local_irq_restore()</tt>,
-and so on.
-
-<h3><a name="Sleepable RCU">Sleepable RCU</a></h3>
-
-<p>
-For well over a decade, someone saying “I need to block within
-an RCU read-side critical section” was a reliable indication
-that this someone did not understand RCU.
-After all, if you are always blocking in an RCU read-side critical
-section, you can probably afford to use a higher-overhead synchronization
-mechanism.
-However, that changed with the advent of the Linux kernel's notifiers,
-whose RCU read-side critical
-sections almost never sleep, but sometimes need to.
-This resulted in the introduction of
-<a href="https://lwn.net/Articles/202847/">sleepable RCU</a>,
-or <i>SRCU</i>.
-
-<p>
-SRCU allows different domains to be defined, with each such domain
-defined by an instance of an <tt>srcu_struct</tt> structure.
-A pointer to this structure must be passed in to each SRCU function,
-for example, <tt>synchronize_srcu(&ss)</tt>, where
-<tt>ss</tt> is the <tt>srcu_struct</tt> structure.
-The key benefit of these domains is that a slow SRCU reader in one
-domain does not delay an SRCU grace period in some other domain.
-That said, one consequence of these domains is that read-side code
-must pass a “cookie” from <tt>srcu_read_lock()</tt>
-to <tt>srcu_read_unlock()</tt>, for example, as follows:
-
-<blockquote>
-<pre>
- 1 int idx;
- 2
- 3 idx = srcu_read_lock(&ss);
- 4 do_something();
- 5 srcu_read_unlock(&ss, idx);
-</pre>
-</blockquote>
-
-<p>
-As noted above, it is legal to block within SRCU read-side critical sections,
-however, with great power comes great responsibility.
-If you block forever in one of a given domain's SRCU read-side critical
-sections, then that domain's grace periods will also be blocked forever.
-Of course, one good way to block forever is to deadlock, which can
-happen if any operation in a given domain's SRCU read-side critical
-section can block waiting, either directly or indirectly, for that domain's
-grace period to elapse.
-For example, this results in a self-deadlock:
-
-<blockquote>
-<pre>
- 1 int idx;
- 2
- 3 idx = srcu_read_lock(&ss);
- 4 do_something();
- 5 synchronize_srcu(&ss);
- 6 srcu_read_unlock(&ss, idx);
-</pre>
-</blockquote>
-
-<p>
-However, if line 5 acquired a mutex that was held across
-a <tt>synchronize_srcu()</tt> for domain <tt>ss</tt>,
-deadlock would still be possible.
-Furthermore, if line 5 acquired a mutex that was held across
-a <tt>synchronize_srcu()</tt> for some other domain <tt>ss1</tt>,
-and if an <tt>ss1</tt>-domain SRCU read-side critical section
-acquired another mutex that was held across as <tt>ss</tt>-domain
-<tt>synchronize_srcu()</tt>,
-deadlock would again be possible.
-Such a deadlock cycle could extend across an arbitrarily large number
-of different SRCU domains.
-Again, with great power comes great responsibility.
-
-<p>
-Unlike the other RCU flavors, SRCU read-side critical sections can
-run on idle and even offline CPUs.
-This ability requires that <tt>srcu_read_lock()</tt> and
-<tt>srcu_read_unlock()</tt> contain memory barriers, which means
-that SRCU readers will run a bit slower than would RCU readers.
-It also motivates the <tt>smp_mb__after_srcu_read_unlock()</tt>
-API, which, in combination with <tt>srcu_read_unlock()</tt>,
-guarantees a full memory barrier.
-
-<p>
-The
-<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">SRCU API</a>
-includes
-<tt>srcu_read_lock()</tt>,
-<tt>srcu_read_unlock()</tt>,
-<tt>srcu_dereference()</tt>,
-<tt>srcu_dereference_check()</tt>,
-<tt>synchronize_srcu()</tt>,
-<tt>synchronize_srcu_expedited()</tt>,
-<tt>call_srcu()</tt>,
-<tt>srcu_barrier()</tt>, and
-<tt>srcu_read_lock_held()</tt>.
-It also includes
-<tt>DEFINE_SRCU()</tt>,
-<tt>DEFINE_STATIC_SRCU()</tt>, and
-<tt>init_srcu_struct()</tt>
-APIs for defining and initializing <tt>srcu_struct</tt> structures.
-
-<h3><a name="Tasks RCU">Tasks RCU</a></h3>
-
-<p>
-Some forms of tracing use “tramopolines” to handle the
-binary rewriting required to install different types of probes.
-It would be good to be able to free old trampolines, which sounds
-like a job for some form of RCU.
-However, because it is necessary to be able to install a trace
-anywhere in the code, it is not possible to use read-side markers
-such as <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>.
-In addition, it does not work to have these markers in the trampoline
-itself, because there would need to be instructions following
-<tt>rcu_read_unlock()</tt>.
-Although <tt>synchronize_rcu()</tt> would guarantee that execution
-reached the <tt>rcu_read_unlock()</tt>, it would not be able to
-guarantee that execution had completely left the trampoline.
-
-<p>
-The solution, in the form of
-<a href="https://lwn.net/Articles/607117/"><i>Tasks RCU</i></a>,
-is to have implicit
-read-side critical sections that are delimited by voluntary context
-switches, that is, calls to <tt>schedule()</tt>,
-<tt>cond_resched_rcu_qs()</tt>, and
-<tt>synchronize_rcu_tasks()</tt>.
-In addition, transitions to and from userspace execution also delimit
-tasks-RCU read-side critical sections.
-
-<p>
-The tasks-RCU API is quite compact, consisting only of
-<tt>call_rcu_tasks()</tt>,
-<tt>synchronize_rcu_tasks()</tt>, and
-<tt>rcu_barrier_tasks()</tt>.
-
-<h3><a name="Waiting for Multiple Grace Periods">
-Waiting for Multiple Grace Periods</a></h3>
-
-<p>
-Perhaps you have an RCU protected data structure that is accessed from
-RCU read-side critical sections, from softirq handlers, and from
-hardware interrupt handlers.
-That is three flavors of RCU, the normal flavor, the bottom-half flavor,
-and the sched flavor.
-How to wait for a compound grace period?
-
-<p>
-The best approach is usually to “just say no!” and
-insert <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>
-around each RCU read-side critical section, regardless of what
-environment it happens to be in.
-But suppose that some of the RCU read-side critical sections are
-on extremely hot code paths, and that use of <tt>CONFIG_PREEMPT=n</tt>
-is not a viable option, so that <tt>rcu_read_lock()</tt> and
-<tt>rcu_read_unlock()</tt> are not free.
-What then?
-
-<p>
-You <i>could</i> wait on all three grace periods in succession, as follows:
-
-<blockquote>
-<pre>
- 1 synchronize_rcu();
- 2 synchronize_rcu_bh();
- 3 synchronize_sched();
-</pre>
-</blockquote>
-
-<p>
-This works, but triples the update-side latency penalty.
-In cases where this is not acceptable, <tt>synchronize_rcu_mult()</tt>
-may be used to wait on all three flavors of grace period concurrently:
-
-<blockquote>
-<pre>
- 1 synchronize_rcu_mult(call_rcu, call_rcu_bh, call_rcu_sched);
-</pre>
-</blockquote>
-
-<p>
-But what if it is necessary to also wait on SRCU?
-This can be done as follows:
-
-<blockquote>
-<pre>
- 1 static void call_my_srcu(struct rcu_head *head,
- 2 void (*func)(struct rcu_head *head))
- 3 {
- 4 call_srcu(&my_srcu, head, func);
- 5 }
- 6
- 7 synchronize_rcu_mult(call_rcu, call_rcu_bh, call_rcu_sched, call_my_srcu);
-</pre>
-</blockquote>
-
-<p>
-If you needed to wait on multiple different flavors of SRCU
-(but why???), you would need to create a wrapper function resembling
-<tt>call_my_srcu()</tt> for each SRCU flavor.
-
-<p>@@QQ@@
-But what if I need to wait for multiple RCU flavors, but I also need
-the grace periods to be expedited?
-<p>@@QQA@@
-If you are using expedited grace periods, there should be less penalty
-for waiting on them in succession.
-But if that is nevertheless a problem, you can use workqueues or multiple
-kthreads to wait on the various expedited grace periods concurrently.
-<p>@@QQE@@
-
-<p>
-Again, it is usually better to adjust the RCU read-side critical sections
-to use a single flavor of RCU, but when this is not feasible, you can use
-<tt>synchronize_rcu_mult()</tt>.
-
-<h2><a name="Possible Future Changes">Possible Future Changes</a></h2>
-
-<p>
-One of the tricks that RCU uses to attain update-side scalability is
-to increase grace-period latency with increasing numbers of CPUs.
-If this becomes a serious problem, it will be necessary to rework the
-grace-period state machine so as to avoid the need for the additional
-latency.
-
-<p>
-Expedited grace periods scan the CPUs, so their latency and overhead
-increases with increasing numbers of CPUs.
-If this becomes a serious problem on large systems, it will be necessary
-to do some redesign to avoid this scalability problem.
-
-<p>
-RCU disables CPU hotplug in a few places, perhaps most notably in the
-expedited grace-period and <tt>rcu_barrier()</tt> operations.
-If there is a strong reason to use expedited grace periods in CPU-hotplug
-notifiers, it will be necessary to avoid disabling CPU hotplug.
-This would introduce some complexity, so there had better be a <i>very</i>
-good reason.
-
-<p>
-The tradeoff between grace-period latency on the one hand and interruptions
-of other CPUs on the other hand may need to be re-examined.
-The desire is of course for zero grace-period latency as well as zero
-interprocessor interrupts undertaken during an expedited grace period
-operation.
-While this ideal is unlikely to be achievable, it is quite possible that
-further improvements can be made.
-
-<p>
-The multiprocessor implementations of RCU use a combining tree that
-groups CPUs so as to reduce lock contention and increase cache locality.
-However, this combining tree does not spread its memory across NUMA
-nodes nor does it align the CPU groups with hardware features such
-as sockets or cores.
-Such spreading and alignment is currently believed to be unnecessary
-because the hotpath read-side primitives do not access the combining
-tree, nor does <tt>call_rcu()</tt> in the common case.
-If you believe that your architecture needs such spreading and alignment,
-then your architecture should also benefit from the
-<tt>rcutree.rcu_fanout_leaf</tt> boot parameter, which can be set
-to the number of CPUs in a socket, NUMA node, or whatever.
-If the number of CPUs is too large, use a fraction of the number of
-CPUs.
-If the number of CPUs is a large prime number, well, that certainly
-is an “interesting” architectural choice!
-More flexible arrangements might be considered, but only if
-<tt>rcutree.rcu_fanout_leaf</tt> has proven inadequate, and only
-if the inadequacy has been demonstrated by a carefully run and
-realistic system-level workload.
-
-<p>
-Please note that arrangements that require RCU to remap CPU numbers will
-require extremely good demonstration of need and full exploration of
-alternatives.
-
-<p>
-There is an embarrassingly large number of flavors of RCU, and this
-number has been increasing over time.
-Perhaps it will be possible to combine some at some future date.
-
-<p>
-RCU's various kthreads are reasonably recent additions.
-It is quite likely that adjustments will be required to more gracefully
-handle extreme loads.
-It might also be necessary to be able to relate CPU utilization by
-RCU's kthreads and softirq handlers to the code that instigated this
-CPU utilization.
-For example, RCU callback overhead might be charged back to the
-originating <tt>call_rcu()</tt> instance, though probably not
-in production kernels.
-
-<h2><a name="Summary">Summary</a></h2>
-
-<p>
-This document has presented more than two decade's worth of RCU
-requirements.
-Given that the requirements keep changing, this will not be the last
-word on this subject, but at least it serves to get an important
-subset of the requirements set forth.
-
-<h2><a name="Acknowledgments">Acknowledgments</a></h2>
-
-I am grateful to Steven Rostedt, Lai Jiangshan, Ingo Molnar,
-Oleg Nesterov, Borislav Petkov, Peter Zijlstra, Boqun Feng, and
-Andy Lutomirski for their help in rendering
-this article human readable, and to Michelle Rankin for her support
-of this effort.
-Other contributions are acknowledged in the Linux kernel's git archive.
-The cartoon is copyright (c) 2013 by Melissa Broussard,
-and is provided
-under the terms of the Creative Commons Attribution-Share Alike 3.0
-United States license.
-
-<p>@@QQAL@@
-
-</body></html>
diff --git a/Documentation/RCU/Design/htmlqqz.sh b/Documentation/RCU/Design/htmlqqz.sh
deleted file mode 100755
index d354f069559b..000000000000
--- a/Documentation/RCU/Design/htmlqqz.sh
+++ /dev/null
@@ -1,108 +0,0 @@
-#!/bin/sh
-#
-# Usage: sh htmlqqz.sh file
-#
-# Extracts and converts quick quizzes in a proto-HTML document file.htmlx.
-# Commands, all of which must be on a line by themselves:
-#
-# "<p>@@QQ@@": Start of a quick quiz.
-# "<p>@@QQA@@": Start of a quick-quiz answer.
-# "<p>@@QQE@@": End of a quick-quiz answer, and thus of the quick quiz.
-# "<p>@@QQAL@@": Place to put quick-quiz answer list.
-#
-# Places the result in file.html.
-#
-# This program is free software; you can redistribute it and/or modify
-# it under the terms of the GNU General Public License as published by
-# the Free Software Foundation; either version 2 of the License, or
-# (at your option) any later version.
-#
-# This program is distributed in the hope that it will be useful,
-# but WITHOUT ANY WARRANTY; without even the implied warranty of
-# MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the
-# GNU General Public License for more details.
-#
-# You should have received a copy of the GNU General Public License
-# along with this program; if not, you can access it online at
-# http://www.gnu.org/licenses/gpl-2.0.html.
-#
-# Copyright (c) 2013 Paul E. McKenney, IBM Corporation.
-
-fn=$1
-if test ! -r $fn.htmlx
-then
- echo "Error: $fn.htmlx unreadable."
- exit 1
-fi
-
-echo "<!-- DO NOT HAND EDIT. -->" > $fn.html
-echo "<!-- Instead, edit $fn.htmlx and run 'sh htmlqqz.sh $fn' -->" >> $fn.html
-awk < $fn.htmlx >> $fn.html '
-
-state == "" && $1 != "<p>@@QQ@@" && $1 != "<p>@@QQAL@@" {
- print $0;
- if ($0 ~ /^<p>@@QQ/)
- print "Bad Quick Quiz command: " NR " (expected <p>@@QQ@@ or <p>@@QQAL@@)." > "/dev/stderr"
- next;
-}
-
-state == "" && $1 == "<p>@@QQ@@" {
- qqn++;
- qqlineno = NR;
- haveqq = 1;
- state = "qq";
- print "<p><a name=\"Quick Quiz " qqn "\"><b>Quick Quiz " qqn "</b>:</a>"
- next;
-}
-
-state == "qq" && $1 != "<p>@@QQA@@" {
- qq[qqn] = qq[qqn] $0 "\n";
- print $0
- if ($0 ~ /^<p>@@QQ/)
- print "Bad Quick Quiz command: " NR ". (expected <p>@@QQA@@)" > "/dev/stderr"
- next;
-}
-
-state == "qq" && $1 == "<p>@@QQA@@" {
- state = "qqa";
- print "<br><a href=\"#qq" qqn "answer\">Answer</a>"
- next;
-}
-
-state == "qqa" && $1 != "<p>@@QQE@@" {
- qqa[qqn] = qqa[qqn] $0 "\n";
- if ($0 ~ /^<p>@@QQ/)
- print "Bad Quick Quiz command: " NR " (expected <p>@@QQE@@)." > "/dev/stderr"
- next;
-}
-
-state == "qqa" && $1 == "<p>@@QQE@@" {
- state = "";
- next;
-}
-
-state == "" && $1 == "<p>@@QQAL@@" {
- haveqq = "";
- print "<h3><a name=\"Answers to Quick Quizzes\">"
- print "Answers to Quick Quizzes</a></h3>"
- print "";
- for (i = 1; i <= qqn; i++) {
- print "<a name=\"qq" i "answer\"></a>"
- print "<p><b>Quick Quiz " i "</b>:"
- print qq[i];
- print "";
- print "</p><p><b>Answer</b>:"
- print qqa[i];
- print "";
- print "</p><p><a href=\"#Quick%20Quiz%20" i "\"><b>Back to Quick Quiz " i "</b>.</a>"
- print "";
- }
- next;
-}
-
-END {
- if (state != "")
- print "Unterminated Quick Quiz: " qqlineno "." > "/dev/stderr"
- else if (haveqq)
- print "Missing \"<p>@@QQAL@@\", no Quick Quiz." > "/dev/stderr"
-}'
--
2.5.2
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