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 <h1>A Tour Through RCU's Requirements</h1>

 <p>Copyright IBM Corporation, 2015</p>
 <p>Author: Paul E.&nbsp;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>,
 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>

 <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 &ldquo;rclock&rdquo;) 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&nbsp;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 &amp;&amp; r2 == 1)
 </pre>
 </blockquote>
 cannot happen.

 <table>
 <tr><th>&nbsp;</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>&nbsp;</td></tr>
 </table>

 <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>.

 <table>
 <tr><th>&nbsp;</th></tr>
 <tr><th align="left">Quick Quiz:</th></tr>
 <tr><td>
 	Why is the <tt>synchronize_rcu()</tt> on line&nbsp;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>&nbsp;</td></tr>
 </table>

 <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>-&gt;a</tt> and <tt>-&gt;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(&amp;gp_lock);
  7   if (rcu_access_pointer(gp)) {
  8     spin_unlock(&amp;gp_lock);
  9     return false;
 10   }
 11   p-&gt;a = a;
 12   p-&gt;b = a;
 13   gp = p; /* ORDERING BUG */
 14   spin_unlock(&amp;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(&amp;gp_lock);
  7   if (rcu_access_pointer(gp)) {
  8     spin_unlock(&amp;gp_lock);
  9     return false;
 10   }
 <b>11   gp = p; /* ORDERING BUG */
 12   p-&gt;a = a;
 13   p-&gt;b = a;</b>
 14   spin_unlock(&amp;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&nbsp;11,
 it will see garbage in the <tt>-&gt;a</tt> and <tt>-&gt;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(&amp;gp_lock);
  7   if (rcu_access_pointer(gp)) {
  8     spin_unlock(&amp;gp_lock);
  9     return false;
 10   }
 11   p-&gt;a = a;
 12   p-&gt;b = a;
 13   rcu_assign_pointer(gp, p);
 14   spin_unlock(&amp;gp_lock);
 15   return true;
 16 }
 </pre>
 </blockquote>

 <p>
 The <tt>rcu_assign_pointer()</tt> on line&nbsp;13 is conceptually
 equivalent to a simple assignment statement, but also guarantees
 that its assignment will
 happen after the two assignments in lines&nbsp;11 and&nbsp;12,
 similar to the C11 <tt>memory_order_release</tt> store operation.
 It also prevents any number of &ldquo;interesting&rdquo; compiler
 optimizations, for example, the use of <tt>gp</tt> as a scratch
 location immediately preceding the assignment.

 <table>
 <tr><th>&nbsp;</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-&gt;a</tt> and <tt>p-&gt;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-&gt;a</tt> and <tt>p-&gt;b</tt> cannot possibly
 	cause any problems.
 </font></td></tr>
 <tr><td>&nbsp;</td></tr>
 </table>

 <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-&gt;a, p-&gt;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-&gt;a, gp-&gt;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-&gt;a</tt>
 and <tt>gp-&gt;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-&gt;a, p-&gt;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(&amp;gp_lock);
  6   p = rcu_access_pointer(gp);
  7   if (!p) {
  8     spin_unlock(&amp;gp_lock);
  9     return false;
 10   }
 11   rcu_assign_pointer(gp, NULL);
 12   spin_unlock(&amp;gp_lock);
 13   synchronize_rcu();
 14   kfree(p);
 15   return true;
 16 }
 </pre>
 </blockquote>

 <p>
 This function is straightforward, with line&nbsp;13 waiting for a grace
 period before line&nbsp;14 frees the old data element.
 This waiting ensures that readers will reach line&nbsp;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&nbsp;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>

 <table>
 <tr><th>&nbsp;</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 distinct 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, <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>&nbsp;</td></tr>
 </table>

 <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&nbsp;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&nbsp;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&nbsp;14 of <tt>remove_gp_synchronous()</tt> really does
 	execute after the removal on line&nbsp;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&nbsp;14 of <tt>remove_gp_synchronous()</tt> really does
 	execute after the removal on
 	line&nbsp;11, but also in the case where the thread executing the
 	<tt>synchronize_rcu()</tt> migrates in the meantime.
 </ol>

 <table>
 <tr><th>&nbsp;</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>

 	<p><font color="ffffff">
 	A related question is &ldquo;When <tt>rcu_read_lock()</tt>
 	doesn't generate any code, why does it matter how it relates
 	to a grace period?&rdquo;
 	The answer is that it is not the relationship of
 	<tt>rcu_read_lock()</tt> itself that is important, but rather
 	the relationship of the code within the enclosed RCU read-side
 	critical section to the code preceding and following the
 	grace period.
 	If we take this viewpoint, then a given RCU read-side critical
 	section begins before a given grace period when some access
 	preceding the grace period observes the effect of some access
 	within the critical section, in which case none of the accesses
 	within the critical section may observe the effects of any
 	access following the grace period.
 	</font>

 	<p><font color="ffffff">
 	As of late 2016, mathematical models of RCU take this
 	viewpoint, for example, see slides&nbsp;62 and&nbsp;63
 	of the
 	<a href="http://www2.rdrop.com/users/paulmck/scalability/paper/LinuxMM.2016.10.04c.LCE.pdf">2016 LinuxCon EU</a>
 	presentation.
 </font></td></tr>
 <tr><td>&nbsp;</td></tr>
 </table>

 <table>
 <tr><th>&nbsp;</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-&gt;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-&gt;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&nbsp;1 might end up accessing the freelist.
 	</font>

 	<p><font color="ffffff">
 	The &ldquo;as if&rdquo; 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>&nbsp;</td></tr>
 </table>

 <table>
 <tr><th>&nbsp;</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>&nbsp;</td></tr>
 </table>

 <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.

 <table>
 <tr><th>&nbsp;</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>&nbsp;</td></tr>
 </table>

 <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 &amp;&amp; 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.

 <table>
 <tr><th>&nbsp;</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>&nbsp;</td></tr>
 </table>

 <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(&amp;my_lock);
 16   WRITE_ONCE(x, 1);
 17   WRITE_ONCE(y, 1);
 18   spin_unlock(&amp;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.

 <table>
 <tr><th>&nbsp;</th></tr>
 <tr><th align="left">Quick Quiz:</th></tr>
 <tr><td>
 	Suppose that synchronize_rcu() did wait until <i>all</i>
 	readers had completed instead of waiting only on
 	pre-existing readers.
 	For how long would the updater be able to rely on there
 	being no readers?
 </td></tr>
 <tr><th align="left">Answer:</th></tr>
 <tr><td bgcolor="#ffffff"><font color="ffffff">
 	For no time at all.
 	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> can <i>never</i> rely on there being
 	no readers.
 </font></td></tr>
 <tr><td>&nbsp;</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>

 <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 &amp;&amp; r2 == 0 &amp;&amp; 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 &amp;&amp; r2 == 1 &amp;&amp; r3 == 0 &amp;&amp; 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 &amp;&amp; r2 == 1 &amp;&amp; r3 == 1 &amp;&amp; r4 == 0 &amp&amp; 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.

 <table>
 <tr><th>&nbsp;</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>&nbsp;</td></tr>
 </table>

 <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(&amp;mylock);
  9 rcu_read_lock();
 10 do_something();
 11 rcu_read_unlock();
 12 spin_unlock(&amp;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
 	&ldquo;extremely long&rdquo; 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.

 <table>
 <tr><th>&nbsp;</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>&nbsp;</td></tr>
 </table>

 <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 veterinarians 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 veterinarians might wait 30 seconds before pronouncing
 the cat dead, while the other might insist on waiting a full minute.
 The two veterinarians 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 &ldquo;frank and open&rdquo; 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.
 Here, &ldquo;modest&rdquo; means roughly the same latency
 degradation as 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(&amp;gp_lock);
 19   p = rcu_dereference(gp);
 20   if (!p) {
 21     spin_unlock(&amp;gp_lock);
 22     return false;
 23   }
 24   rcu_assign_pointer(gp, NULL);
 25   call_rcu(&amp;p-&gt;rh, remove_gp_cb);
 26   spin_unlock(&amp;gp_lock);
 27   return true;
 28 }
 </pre>
 </blockquote>

 <p>
 A definition of <tt>struct foo</tt> is finally needed, and appears
 on lines&nbsp;1-5.
 The function <tt>remove_gp_cb()</tt> is passed to <tt>call_rcu()</tt>
 on line&nbsp;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 idle and 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.

 <table>
 <tr><th>&nbsp;</th></tr>
 <tr><th align="left">Quick Quiz:</th></tr>
 <tr><td>
 	Why does line&nbsp;19 use <tt>rcu_access_pointer()</tt>?
 	After all, <tt>call_rcu()</tt> on line&nbsp;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>-&gt;gp_lock</tt> acquired on line&nbsp;18 excludes
 	any changes, including any insertions that <tt>rcu_dereference()</tt>
 	would protect against.
 	Therefore, any insertions will be delayed until after
 	<tt>-&gt;gp_lock</tt>
 	is released on line&nbsp;25, which in turn means that
 	<tt>rcu_access_pointer()</tt> suffices.
 </font></td></tr>
 <tr><td>&nbsp;</td></tr>
 </table>

 <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 &ldquo;fire and forget&rdquo; 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(&amp;gp_lock);
 12   p = rcu_dereference(gp);
 13   if (!p) {
 14     spin_unlock(&amp;gp_lock);
 15     return false;
 16   }
 17   rcu_assign_pointer(gp, NULL);
 18   kfree_rcu(p, rh);
 19   spin_unlock(&amp;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.

 <table>
 <tr><th>&nbsp;</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>&nbsp;</td></tr>
 </table>

 <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(&amp;gp_lock);
  7   p = rcu_access_pointer(gp);
  8   if (!p) {
  9     spin_unlock(&amp;gp_lock);
 10     return false;
 11   }
 12   rcu_assign_pointer(gp, NULL);
 13   spin_unlock(&amp;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&nbsp;14, <tt>get_state_synchronize_rcu()</tt> obtains a
 &ldquo;cookie&rdquo; from RCU,
 then line&nbsp;15 carries out other tasks,
 and finally, line&nbsp;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, perhaps in the guise of tail
 recursion, 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 small read-side delays can occur when using
 <tt>synchronize_rcu_expedited()</tt>, courtesy of this function's use
 of <tt>smp_call_function_single()</tt>.

 <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 &ldquo;To err is human&rdquo;, 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 splat 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 sparse
 	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 &ldquo;eventually&rdquo; 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 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
 &ldquo;out of the box.&rdquo;

 <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 &ldquo;ghost CPUs&rdquo; (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
 all of RCU's kthreads have been spawned, which occurs at
 <tt>early_initcall()</tt> time.
 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>),
 <a href="#Sched Flavor"><tt>synchronize_sched()</tt></a>,
 <tt>synchronize_rcu_expedited()</tt>,
 <tt>synchronize_rcu_bh_expedited()</tt>, and
 <tt>synchronize_sched_expedited()</tt>
 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>
 However, once the scheduler has spawned its first kthread, this early
 boot trick fails for <tt>synchronize_rcu()</tt> (as well as for
 <tt>synchronize_rcu_expedited()</tt>) in <tt>CONFIG_PREEMPT=y</tt>
 kernels.
 The reason is that an RCU read-side critical section might be preempted,
 which means that a subsequent <tt>synchronize_rcu()</tt> really does have
 to wait for something, as opposed to simply returning immediately.
 Unfortunately, <tt>synchronize_rcu()</tt> can't do this until all of
 its kthreads are spawned, which doesn't happen until some time during
 <tt>early_initcalls()</tt> time.
 But this is no excuse:  RCU is nevertheless required to correctly handle
 synchronous grace periods during this time period.
 Once all of its kthreads are up and running, RCU starts running
 normally.

 <table>
 <tr><th>&nbsp;</th></tr>
 <tr><th align="left">Quick Quiz:</th></tr>
 <tr><td>
 	How can RCU possibly handle grace periods before all of its
 	kthreads have been spawned???
 </td></tr>
 <tr><th align="left">Answer:</th></tr>
 <tr><td bgcolor="#ffffff"><font color="ffffff">
 	Very carefully!
 	</font>

 	<p><font color="ffffff">
 	During the &ldquo;dead zone&rdquo; between the time that the
 	scheduler spawns the first task and the time that all of RCU's
 	kthreads have been spawned, all synchronous grace periods are
 	handled by the expedited grace-period mechanism.
 	At runtime, this expedited mechanism relies on workqueues, but
 	during the dead zone the requesting task itself drives the
 	desired expedited grace period.
 	Because dead-zone execution takes place within task context,
 	everything works.
 	Once the dead zone ends, expedited grace periods go back to
 	using workqueues, as is required to avoid problems that would
 	otherwise occur when a user task received a POSIX signal while
 	driving an expedited grace period.
 	</font>

 	<p><font color="ffffff">
 	And yes, this does mean that it is unhelpful to send POSIX
 	signals to random tasks between the time that the scheduler
 	spawns its first kthread and the time that RCU's kthreads
 	have all been spawned.
 	If there ever turns out to be a good reason for sending POSIX
 	signals during that time, appropriate adjustments will be made.
 	(If it turns out that POSIX signals are sent during this time for
 	no good reason, other adjustments will be made, appropriate
 	or otherwise.)
 </font></td></tr>
 <tr><td>&nbsp;</td></tr>
 </table>

 <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 &ldquo;half-interrupts&rdquo; 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.

 <p>
 <b>Important note</b>: The <tt>rcu_barrier()</tt> function is not,
 repeat, <i>not</i>, obligated to wait for a grace period.
 It is instead only required to wait for RCU callbacks that have
 already been posted.
 Therefore, if there are no RCU callbacks posted anywhere in the system,
 <tt>rcu_barrier()</tt> is within its rights to return immediately.
 Even if there are callbacks posted, <tt>rcu_barrier()</tt> does not
 necessarily need to wait for a grace period.

 <table>
 <tr><th>&nbsp;</th></tr>
 <tr><th align="left">Quick Quiz:</th></tr>
 <tr><td>
 	Wait a minute!
 	Each RCU callbacks must wait for a grace period to complete,
 	and <tt>rcu_barrier()</tt> must wait for each pre-existing
 	callback to be invoked.
 	Doesn't <tt>rcu_barrier()</tt> therefore need to wait for
 	a full grace period if there is even one callback posted anywhere
 	in the system?
 </td></tr>
 <tr><th align="left">Answer:</th></tr>
 <tr><td bgcolor="#ffffff"><font color="ffffff">
 	Absolutely not!!!
 	</font>

 	<p><font color="ffffff">
 	Yes, each RCU callbacks must wait for a grace period to complete,
 	but it might well be partly (or even completely) finished waiting
 	by the time <tt>rcu_barrier()</tt> is invoked.
 	In that case, <tt>rcu_barrier()</tt> need only wait for the
 	remaining portion of the grace period to elapse.
 	So even if there are quite a few callbacks posted,
 	<tt>rcu_barrier()</tt> might well return quite quickly.
 	</font>

 	<p><font color="ffffff">
 	So if you need to wait for a grace period as well as for all
 	pre-existing callbacks, you will need to invoke both
 	<tt>synchronize_rcu()</tt> and <tt>rcu_barrier()</tt>.
 	If latency is a concern, you can always use workqueues
 	to invoke them concurrently.
 </font></td></tr>
 <tr><td>&nbsp;</td></tr>
 </table>

 <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,
 with the exception of <a href="#Sleepable RCU">SRCU</a> read-side
 critical sections.
 This requirement was present from day one in DYNIX/ptx, but
 on the other hand, the Linux kernel's CPU-hotplug implementation
 is &ldquo;interesting.&rdquo;

 <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 synchronous grace-period operations such as
 <tt>synchronize_rcu()</tt> and <tt>synchronize_rcu_expedited()</tt>.

 <p>
 However, 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.
 Furthermore, <tt>rcu_barrier()</tt> blocks CPU-hotplug operations
 during its execution, which results in another type of deadlock
 when invoked from a CPU-hotplug notifier.

 <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)&nbsp;it releases them before exiting that same
 RCU read-side critical section, or
 (2)&nbsp;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.
 However, there are some restrictions on the code placed within
 <tt>RCU_NONIDLE()</tt>:

 <ol>
 <li>	Blocking is prohibited.
 	In practice, this is not a serious restriction given that idle
 	tasks are prohibited from blocking to begin with.
 <li>	Although nesting <tt>RCU_NONIDLE()</tt> is permitted, they cannot
 	nest indefinitely deeply.
 	However, given that they can be nested on the order of a million
 	deep, even on 32-bit systems, this should not be a serious
 	restriction.
 	This nesting limit would probably be reached long after the
 	compiler OOMed or the stack overflowed.
 <li>	Any code path that enters <tt>RCU_NONIDLE()</tt> must sequence
 	out of that same <tt>RCU_NONIDLE()</tt>.
 	For example, the following is grossly illegal:

 	<blockquote>
 	<pre>
  1     RCU_NONIDLE({
  2       do_something();
  3       goto bad_idea;  /* BUG!!! */
  4       do_something_else();});
  5   bad_idea:
 	</pre>
 	</blockquote>

 	<p>
 	It is just as illegal to transfer control into the middle of
 	<tt>RCU_NONIDLE()</tt>'s argument.
 	Yes, in theory, you could transfer in as long as you also
 	transferred out, but in practice you could also expect to get sharply
 	worded review comments.
 </ol>

 <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.
 RCU must therefore 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>-&gt;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>-&gt;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 &ldquo;lazy&rdquo;
 variant of <tt>call_rcu()</tt> that might one day be created for
 energy-efficiency purposes.

 <p>
 That said, there are limits.
 RCU requires that the <tt>rcu_head</tt> structure be aligned to a
 two-byte boundary, and passing a misaligned <tt>rcu_head</tt>
 structure to one of the <tt>call_rcu()</tt> family of functions
 will result in a splat.
 It is therefore necessary to exercise caution when packing
 structures containing fields of type <tt>rcu_head</tt>.
 Why not a four-byte or even eight-byte alignment requirement?
 Because the m68k architecture provides only two-byte alignment,
 and thus acts as alignment's least common denominator.

 <p>
 The reason for reserving the bottom bit of pointers to
 <tt>rcu_head</tt> structures is to leave the door open to
 &ldquo;lazy&rdquo; callbacks whose invocations can safely be deferred.
 Deferring invocation could potentially have energy-efficiency
 benefits, but only if the rate of non-lazy callbacks decreases
 significantly for some important workload.
 In the meantime, reserving the bottom bit keeps this option open
 in case it one day becomes useful.

 <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&mdash;and perhaps unprecedented&mdash;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 &ldquo;bottom-half&rdquo;,
 hence the &ldquo;_bh&rdquo; 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 &ldquo;classic&rdquo;
 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 &ldquo;I need to block within
 an RCU read-side critical section&rdquo; 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(&amp;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 &ldquo;cookie&rdquo; 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(&amp;ss);
  4 do_something();
  5 srcu_read_unlock(&amp;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(&amp;ss);
  4 do_something();
  5 synchronize_srcu(&amp;ss);
  6 srcu_read_unlock(&amp;ss, idx);
 </pre>
 </blockquote>

 <p>
 However, if line&nbsp;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&nbsp;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>
 Also unlike other RCU flavors, SRCU's callbacks-wait function
 <tt>srcu_barrier()</tt> may be invoked from CPU-hotplug notifiers,
 though this is not necessarily a good idea.
 The reason that this is possible is that SRCU is insensitive
 to whether or not a CPU is online, which means that <tt>srcu_barrier()</tt>
 need not exclude CPU-hotplug operations.

 <p>
 SRCU also differs from other RCU flavors in that SRCU's expedited and
 non-expedited grace periods are implemented by the same mechanism.
 This means that in the current SRCU implementation, expediting a
 future grace period has the side effect of expediting all prior
 grace periods that have not yet completed.
 (But please note that this is a property of the current implementation,
 not necessarily of future implementations.)
 In addition, if SRCU has been idle for longer than the interval
 specified by the <tt>srcutree.exp_holdoff</tt> kernel boot parameter
 (25&nbsp;microseconds by default),
 and if a <tt>synchronize_srcu()</tt> invocation ends this idle period,
 that invocation will be automatically expedited.

 <p>
 As of v4.12, SRCU's callbacks are maintained per-CPU, eliminating
 a locking bottleneck present in prior kernel versions.
 Although this will allow users to put much heavier stress on
 <tt>call_srcu()</tt>, it is important to note that SRCU does not
 yet take any special steps to deal with callback flooding.
 So if you are posting (say) 10,000 SRCU callbacks per second per CPU,
 you are probably totally OK, but if you intend to post (say) 1,000,000
 SRCU callbacks per second per CPU, please run some tests first.
 SRCU just might need a few adjustment to deal with that sort of load.
 Of course, your mileage may vary based on the speed of your CPUs and
 the size of your memory.

 <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 &ldquo;trampolines&rdquo; 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 &ldquo;just say no!&rdquo; 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(&amp;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.

 <table>
 <tr><th>&nbsp;</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>&nbsp;</td></tr>
 </table>

 <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
 <tt>rcu_barrier()</tt> operations.
 If there is a strong reason to use <tt>rcu_barrier()</tt> 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 &ldquo;interesting&rdquo; 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.

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