async_events.txt
Barret Rhoden

1. Overview
2. Async Syscalls and I/O
3. Event Delivery / Notification
4. Single-core Process (SCP) Events
5. Misc Things That Aren't Sorted Completely:

1. Overview
====================
1.1 Event Handling / Notifications / Async IO Issues:
------------------------------------------------------------------
Basically, syscalls use the ROS event delivery mechanisms, redefined and
described below.  Syscalls use the event delivery just like any other
subsystem would that wants to deliver messages to a process.  The only other
example we have right now are the "kernel notifications", which are the
one-sided, kernel-initiated messages that the kernel sends to a process.

Overall, there are several analogies from how vcores work to how the OS
handles interrupts.  This is a result of trying to make vcores run like
virtual multiprocessors, in control of their resources and aware of the lower
levels of the system.  This analogy has guided much of how the vcore layer
works.  Whenever we have issues with the 2-lsched, realize the amount of
control they want means using solutions that the OS must do too.

Note that there is some pointer chasing going on, though we try to keep it to
a minimum.  Any time the kernel chases a pointer, it needs to make sure it is
in the R/W section of userspace, though it doesn't need to check if the page
is present.  There's more info in the Page Fault sections of the
documentation.  (Briefly, if the kernel PFs on a user address, it will either
block and handle the PF, or if the address was unmapped, it will kill the
process).

1.2 Some Definitions:
---------------------------------------
ev_q, event_queue, event_q: all terms used interchangeably with each other.
They are the endpoint for communicating messages to a process, encapsulating
the method of delivery (such as IPI or not) with where to save the message.

Vcore context: the execution context of the virtual core on the "trampoline"
stack.  All executions start from the top of this stack, and no stack state is
saved between vcore_entry() calls.  All executions on here are non-blocking,
notifications (IPIs) are disabled, and there is a specific TLS loaded.  Vcore
context is used for running the second level scheduler (2LS), swapping between
threads, and handling notifications.  It is analagous to "interrupt context"
in the OS.  Any functions called from here should be brief.  Any memory
touched must be pinned.  In Lithe terms, vcore context might be called the
Hart / hard thread.  People often wonder if they can run out of vcore context
directly.  Technically, you can, but you lose the ability to take any fault
(page fault) or to get IPIs for notification.  In essence, you lose control,
analgous to running an application in the kernel with preemption/interrupts
disabled.  See the process documentation for more info.

2LS: is the second level scheduler/framework.  This code executes in vcore
context, and is Lithe / plugs in to Lithe (eventually).  Often used
interchangeably with "vcore context", usually when I want to emphasize the
scheduling nature of the code.

VCPD: "virtual core preemption data".  In procdata, there is an array of
struct preempt_data, one per vcore.  This is the default location to look for
all things related to the management of vcores, such as its event_mbox (queue
of incoming messages/notifications/events).  Both the kernel and the vcore
code know to look here for a variety of things.

Vcore-business: This is a term I use for a class of messages where the receiver
is the actual vcore, and not just using the vcore as a place to receive the
message.  Examples of vcore-business are INDIR events, preempt_pending events,
scheduling events (self-ipis by the 2LS from one vcore to another), and things
like that.  There are two types: public and private.  Private will only be
handled by that vcore.  Public might be handled by another vcore.

Notif_table: This is a list of event_q*s that correspond to certain
unexpected/"one-sided" events the kernel sends to the process.  It is similar
to an IRQ table in the kernel.  Each event_q tells the kernel how the process
wants to be told about the specific event type.

Notifications: used to be a generic event, but now used in terms of the verb
'notify' (do_notify()).  In older docs, passive notification is just writing a
message somewhere.  Active notification is an IPI delivered to a vcore.  I use
that term interchangeably with an IPI, and usually you can tell by context
that I'm talking about an IPI going to a process (and not just the kernel).
The details of it make it more complicated than just an IPI, but it's
analagous.  I've start referring to notification as the IPI, and "passive
notification" as just events, though older documentation has both meanings.

BCQ: "bounded concurrent queue".  It is a fixed size array of messages
(structs of notification events, or whatever).  It is non-blocking, supporting
multiple producers and consumers, where the producers do not trust the
consumers.  It is the primary mechanism for the kernel delivering message
payloads into a process's address space.  Note that producers don't trust each
other either (in the event of weirdness, the producers give up and say the
buffer is full).  This means that a process can produce for one of its ev_qs
(which is what they need to do to send message to itself).

UCQ: "unbounded concurrent queue".  This is a data structure allowing the kernel
to produce an unbounded number of messages for the process to consume.  The main
limitation to the number of messages is RAM.  Check out its documentation.

2. Async Syscalls and I/O
====================
2.1 Basics
----------------------------------------------
The syscall struct is the contract for work with the kernel, including async
I/O.  Lots of current OS async packages use epoll or other polling systems.
Note the distinction between Polling and Async I/O.  Polling is about finding
out if a call will block.  It is primarily used for sockets and pipes.  It
does relatively nothing for disk I/O, which requires a separate async I/O
system.  By having all syscalls be async, we can make polling a bit easier and
more unified with the generic event code that we use for all syscalls.

For instance, we can have a sys_poll syscall, which is async just like any
other syscall.  The call can be a "one shot / non-blocking", like the current
systems polling code, or it can also notify on change (not requiring future
polls) via the event_q mechanisms.  If you don't want to be IPId, you can
"poll" the syscall struct - not requiring another kernel crossing/syscall.

Note that we do not tie syscalls and polling to FDs.  We do events on
syscalls, which can be used to check FDs.  I think a bunch of polling cases
will not be needed once we have async syscalls, but for those that remain,
we'll have sys_poll() (or whatever).

To receive an event on a syscall completion or status change, just fill in the
event_q pointer.  If it is 0, the kernel will assume you poll the actual
syscall struct.

struct syscall {
	current stuff 		/* arguments, retvals */
	struct ev_queue * 	/* struct used for messaging, including IPIs*/
	void * 			/* used by 2LS, usually a struct u_thread * */
}

One issue with async syscalls is that there can be too many outstanding IOs
(normally sync calls provide feedback / don't allow you to over-request).
Eventually, processes can exhaust kernel memory (the kthreads, specifically).
We need a way to limit the kthreads per proc, etc.  Shouldn't be a big deal.

Normally, we talk about changing the flag in a syscall to SC_DONE.  Async
syscalls can be SC_PROGRESS (new stuff happened on it), which can trigger a
notification event.  Some calls, like AIO or bulk accept, exist for a while
and slowly get filled in / completed.  In the future, we'll also want a way to
abort the in-progress syscalls (possibly any syscall!).

2.2 Uthreads Blocking on Syscalls
----------------------------------------------
Many threading libraries will want some notion of a synchronous, blocking
thread.  These threads use regular I/O calls, which are async under the hood,
but don't want to bother with call backs or other details of async I/O.  In
this section, I'll talk a bit about how this works, esp regarding
uthreads/pthreads.

'Blocking' refers to user threads, and has nothing to do with an actual
process blocking/waiting on some kernel event.  The kernel does not know
anything about what goes on here.  While a bit confusing, this allows
applications to do whatever they want on top of an async interface, and is a
consequence of decoupling cores from user-threads from kthreads.

2.2.1 Basics of Uthread Blocking
---------------
When a thread calls a glibc function that makes a system call, if the syscall
is not yet complete when the kernel returns to userspace, glibc will check for
the existence of a second level scheduler and attempt to use it to yield its
uthread.  If there is no 2LS, the code just spins for now.  Eventually, it
will try to suspend/yield the process for a while (til the call is done), aka,
block in the kernel.

If there is a 2LS, the current thread will yield, and call out to the 2LS's
blockon_sysc() method, which needs a way to stop the thread and be able to
restart it when the syscall completes.  Specifically, the pthread 2LS registers
the syscall to respond to an event (described in detail elsewhere in this doc).
When the event comes in, meaning the syscall is complete, the thread is put on
the runnable list.

Details:
- A pointer to the struct pthread is stored in the syscall's void*.  When the
  syscall is done, we normally get a message from the kernel, and the payload
  tells us the syscall is done, which tells us which thread to unblock. 
- The pthread code also always asks for an IPI and event message for every
  syscall that completes.  This is far from ideal.  Still, the basics are the
  same for any threading library.  Once you know a thread is done, you need to
  do something about it.
- The pthread code does syscall blocking and event notification on a per-core
  basis.  Using the default (VCPD) ev_mbox for this is a bad idea (which we did
  at some point).
- There's a race between the 2LS trying to sign up for events and the kernel
  finishing the event.  We handle this in uthread code, so use the helper to
  register_evq(), which does the the right thing (atomics, careful ordering
  with writes, etc).

2.2.1 Recovering from Event Overflow
---------------
Event overflow recovery is unnecessary, since syscall ev_qs use UCQs now.  this
section is kept around for some useful tidbits, such as details about
deregistering ev_qs for a syscall:

---------------------------
The pthread code expects to receive an event somehow to unblock a thread
once its syscall is done.  One limitation to our messaging systems is that you
can't send an infinite amount of event messages.  (By messages, I mean a chunk
of memory with a payload, in this case consisting of a struct syscall *).
Event delivery degrades to a bit in the case of the message queue being full
(more details on that later).

The pthread code (and any similar 2LS) needs to handle waking up syscalls when
the event message was lost and all we know is that some syscall that was meant
to have a message sent to a particular event queue (per-core in the case of
pthread stuff (actually the VCPD for now)).  The basic idea is to poll all
outstanding system calls and unblock whoever is done.

The key problem is due to a race: for a given syscall we don't know if we're
going to get a message for a syscall or not.  There could be a completion
message in the queue for the syscall while we are going through the list of
blocked threads.  If we assume we already got the message (or it was lost in
the overflow), but didn't really, then if we finish as SC and free its memory
(free or return up the stack), we could later get a message for it, and all
sorts of things would go wrong (like trying to unblock a pointer that is
gibberish).

Here's what we do:
1) Set a "handling overflow" flag so we don't recurse.
2) Turn off event delivery for all syscalls on our list
3) Handle any event messages.  This is how we make a distinction between
finished syscalls that had a message sent and those that didn't.  We're doing
the message-sent ones here.
4) For any left on the list, check to see if they are done.  We actually do
this by attempting to turn on event delivery for them.  Turning on event
delivery can fail if the call is already done.  So if it fails, they are done
and we unblock them (similar to how we block the threads in the first place).
If it doesn't fail, they are now ready to receive messages.  This can be
tweaked a bit.
5) Unset the overflow-handling flag.

One thing to be careful of is that when we turn off event delivery, you need to
be sure the kernel isn't in the process of sending an event.  This is why we
have the SC_K_LOCK syscall flag.  Uthread code will not consider deregistration
complete while that flag is set, since the kernel is still mucking with the
syscall (and sending an event).  Once the flag is clear, the event has been
delivered (the ev_msg is in the ev_mbox), and our assumptions remain true.

There are a couple implications of this style.  If you have a shared event
queue (with other event sources), those events can get mixed in with the
recovery.  Don't leave the vcore context due to other events.  This'll
probably need work.  The other thing is that completed syscalls can get
handled in a different order than they were signaled.  Shouldn't be a big
deal.

Note on the overflow handling flag and unsetting it.  There should not be any
races with this.  The flag prevented us from handling overflows on the event
queue.  Other than when we checked for events that had been succesfully sent,
we didn't try to handle events.  We can unset the flag, and at that point we
can start handling missed events.  If there was an overflow after we last
checked the list, but before we cleared the overflow-handling flag, we'll
still catch it since we haven't tried handling events in between checking the
list and clearing the flag.  That flag doesn't even matter until we want to
handle_events, so we aren't missing anything.  the next handle_events() will
deal with everything from scratch.

For blocking threads that block concurrently with the overflow handling: in
the pthread case, this can't happen since everything is per-vcore.  If you do
have process-wide thread blocking/syscall management, we can add new ones, but
they must have event delivery turned off when they are added to the list.  And
you'll need to lock the list, etc.  This should work in part due to new
syscalls being added to the end of the list, and the overflow-handler
proceeding linearly through the list.

Also note that we shouldn't handle the event for unblocking a syscall on a
different core than the one it was submitted to.  This could result in
concurrent modifications to the original core's TAILQ (bad).  This restriction
is dependent on how a 2LS does its thread handling/blocking.

Eventually, we'll want a way to detect and handle excessive overflow, since
it's probably quite expensive.  Perhaps turn it off and periodically poll the
syscalls for completion (but don't bother turning on the ev_q).
---------------------------

3. Event Delivery / Notification
====================
3.1 Basics
----------------------------------------------
The mbox (mailbox) is where the actual messages go.

struct ev_mbox {
	bcq of notif_events 	/* bounded buffer, multi-consumer/producer */
	msg_bitmap
}
struct ev_queue {		/* aka, event_q, ev_q, etc. */
	struct ev_mbox *
	void handler(struct event_q *)
	vcore_to_be_told
	flags 			/* IPI_WANTED, RR, 2L-handle-it, etc */
}
struct ev_queue_big {
	struct ev_mbox *	/* pointing to the internal storage */
	vcore_to_be_told
	flags 			/* IPI_WANTED, RR, 2L-handle-it, etc */
	struct ev_mbox { }	/* never access this directly */
}

The purpose of the big one is to simply embed some storage.  Still, only
access the mbox via the pointer.  The big one can be casted (and stored as)
the regular, so long as you know to dealloc a big one (free() knows, custom
styles or slabs would need some help).

The ev_mbox says where to put the actual message, and the flags handle things
such as whether or not an IPI is wanted.

Using pointers for the ev_q like this allows multiple event queues to use the
same mbox.  For example, we could use the vcpd queue for both kernel-generated
events as well as async syscall responses.  The notification table is actually
a bunch of ev_qs, many of which could be pointing to the same vcore/vcpd-mbox,
albeit with different flags.

3.2 Kernel Notification Using Event Queues
----------------------------------------------
The notif_tbl/notif_methods (kernel-generated 'one-sided' events) is just an
array of struct ev_queue*s.  Handling a notification is like any other time
when we want to send an event.  Follow a pointer, send a message, etc.  As
with all ev_qs, ev_mbox* points to where you want the message for the event,
which usually is the vcpd's mbox.  If the ev_q pointer is 0, then we know the
process doesn't want the event (equivalent to the older 'NOTIF_WANTED' flag).
Theoretically, we can send kernel notifs to user threads.  While it isn't
clear that anyone will ever want this, it is possible (barring other issues),
since they are just events.

Also note the flag EVENT_VCORE_APPRO.  Processes should set this for certain
types of events where they want the kernel to send the event/IPI to the
'appropriate' vcore.  For example, when sending a message about a preemption
coming in, it makes sense for the kernel to send it to the vcore that is going
to get preempted, but the application could choose to ignore the notification.
When this flag is set, the kernel will also use the vcore's ev_mbox, ignoring
the process's choice.  We can change this later, but it doesn't really make
sense for a process to pick an mbox and also say VCORE_APPRO.

There are also interfaces in the kernel to put a message in an ev_mbox
regardless of the process's wishes (post_vcore_event()), and can send an IPI
at any time (proc_notify()).

3.3 IPIs, Indirection Events, and Fallback (Spamming Indirs)
----------------------------------------------
An ev_q can ask for an IPI, for an indirection event, and for an indirection
event to be spammed in case a vcore is offline (sometimes called the 'fallback'
option.  Or any combination of these.  Note that these have little to do with
the actual message being sent.  The actual message is dropped in the ev_mbox
pointed to by the ev_q.

The main use for all of this is for syscalls.  If you want to receive an event
when a syscall completes or has a change in status, simply allocate an event_q,
and point the syscall at it.  syscall: ev_q* -> "vcore for IPI, syscall message
in the ev_q mbox", etc.  You can also point it to an existing ev_q.  Pthread
code has examples of two ways to do this.  Both have per vcore ev_qs, requesting
IPIs, INDIRS, and SPAM_INDIR.  One way is to have an ev_mbox per vcore, and
another is to have a global ev_mbox that all ev_qs point to.  As a side note, if
you do the latter, you don't need to worry about a vcore's ev_q if it gets
preempted: just check the global ev_mbox (which is done by checking your own
vcore's syscall ev_q).

3.3.1: IPIs and INDIRs
---------------
An EVENT_IPI simply means we'll send an IPI to the given vcore.  Nothing else.
This will usually be paired with an Indirection event (EVENT_INDIR).  An INDIR
is a message of type EV_EVENT with an ev_q* payload.  It means "check this
ev_q".  Most ev_qs that ask for an IPI will also want an INDIR so that the vcore
knows why it was IPIed.  You don't have to do this: for instance, your 2LS might
poll its own ev_q, so you won't need the indirection event.

Additionally, note that IPIs and INDIRs can be spurious.  It's not a big deal to
receive and IPI and have nothing to do, or to be told to check an empty ev_q.
All of the event handling code can deal with this.

INDIR events are sent to the VCPD public mbox, which means they will get handled
if the vcore gets preempted.  Any other messages sent here will also get handled
during a preemption.  However, the only type of messages you should use this for
are ones that can handle spurious messages.  The completion of a syscall is an
example of a message that cannot be spurious.  Since INDIRs can be spurious, we
can use the public mbox.  (Side note: the kernel may spam INDIRs in attempting
to make sure you get the message on a vcore that didn't yield.)

Never use a VCPD mbox (public or private) for messages you might want to receive
if that vcore is offline.  If you want to be sure to get a message, create your
own ev_q and set flags for INDIR, SPAM_INDIR, and IPI.  There's no guarantee a
*specific* message will get looked at.  In cases where it won't, the kernel will
send that message to another vcore.  For example, if the kernel posts an INDIR
to a VCPD mbox (the public one btw) and it loses a race with the vcore yielding,
the vcore might never see that message.  However, the kernel knows it lost the
race, and will find another vcore to send it to.

3.3.2: Spamming Indirs / Fallback
---------------
Both IPI and INDIR need an actual vcore.  If that vcore is unavailable and if
EVENT_SPAM_INDIR is set, the kernel will pick another vcore and send the
messages there.  This allows an ev_q to be set up to handle work when the vcore
is online, while allowing the program to handle events when that core yields,
without having to reset all of its ev_qs to point to "known" available vcores
(and avoiding those races).  Note 'online' is synonymous with 'mapped', when
talking about vcores.  A vcore technically isn't always online, only destined
to be online, when it is mapped to a pcore (kmsg on the way, etc).  It's
easiest to think of it being online for the sake of this discussion.

One question is whether or not 2LSs need a SPAM_INDIR flag for their ev_qs.
The main use for SPAM_INDIR is so that vcores can yield.  (Note that fallback
won't help you *miss* INDIR messages in the event of a preemption; you can
always lose that race due to it taking too long to process the messages).  An
alternative would be for vcores to pick another vcore and change all of its
ev_qs to that vcore.  There are a couple problems with this.  One is that it'll
be a pain to get those ev_qs back when the vcore comes back online (if ever).
Another issue is that other vcores will build up a list of ev_qs that they
aren't aware of, which will be hard to deal with when *they* yield.  SPAM_INDIR
avoids all of those problems.

An important aspect of spamming indirs is that it works with yielded vcores,
not preempted vcores.  It could be that there are no cores that are online, but
there should always be at least one core that *will* be online in the future, a
core that the process didn't want to lose and will deal with in the future.  If
not for this distinction, SPAM_INDIR could fail.  An older idea would be to have
fallback send the msg to the desired vcore if there were no others.  This would
not work if the vcore yielded and then the entire process was preempted or
otherwise not running.  Another way to put this is that we need a field to
determine whether a vcore is offline temporarily or permanently.

This is why we have the VCPD flag 'VC_CAN_RCV_MSG'.  It tells the kernel's event
delivery code that the vcore will check the messages: it is an acceptable
destination for a spammed indir.  There are two reasons to put this in VCPD:
1) Userspace can remotely turn off a vcore's msg reception.  This is necessary
for handling preemption of a vcore that was in uthread context, so that we can
remotely 'yield' the core without having to sys_change_vcore() (which I discuss
below, and is meant to 'unstick' a vcore).
2) Yield is simplified.  The kernel no longer races with itself nor has to worry
about turning off that flag - userspace can do it when it wants to yield.  (turn
off the flag, check messages, then yield).  This is less big of a deal now that
the kernel races with vcore membership in the online_vcs list.

Two aspects of the code make this work nicely.  The VC_CAN_RCV_MSG flag greatly
simplifies the kernel's job.  There are a lot of weird races we'd have to deal
with, such as process state (RUNNING_M), whether a mass preempt is going on, or
just one core, or a bunch of cores, mass yields, etc.  A flag that does one
thing well helps a lot - esp since preemption is not the same as yielding.  The
other useful thing is being able to handle spurious events.  Vcore code can
handle extra IPIs and INDIRs to non-VCPD ev_qs.  Any vcore can handle an ev_q
that is "non-VCPD business".

Worth mentioning is the difference between 'notif_pending' and VC_CAN_RCV_MSG.
VC_CAN_RCV_MSG is the process saying it will check for messages.
'notif_pending' is when the kernel says it *has* sent a message.
'notif_pending' is also used by the kernel in proc_yield() and the 2LS in
pop_user_ctx() to make sure the sent message is not missed.

Also, in case this comes up, there's a slight race on changing the mbox* and the
vcore number within the event_q.  The message could have gone to the wrong (old)
vcore, but not the IPI.  Not a big deal - IPIs can be spurious, and the other
vcore will eventually get it.  The real way around this is create a new ev_q and
change the pointer (thus atomically changing the entire ev_q's contents), though
this can be a bit tricky if you have multiple places pointing to the same ev_q
(can't change them all at once).

3.3.3: Fallback and Preemption
---------------
SPAM_INDIR doesn't protect you from preemptions.  A vcore can be preempted and
have INDIRs in its VCPD.

It is tempting to just use sys_change_vcore(), which will change the calling
vcore to the new one.  This should only be used to "unstick" a vcore.  A vcore
is stuck when it was preempted while it had notifications disabled.  This is
usually when it is vcore context, but also in any lock holding code for locks
shared with vcore context (the userspace equivalent of irqsave locks).  With
this syscall, you could change to the offline vcore and process its INDIRs.

The problem with that plan is the calling core (that is trying to save the
other) may have extra messages, and that sys_change_vcore does not return.  We
need a way to deal with our other messages.  We're back to the same problem we
had before, just with different vcores.  The only thing we really accomplished
is that we unstuck the other vcore.  We could tell the restarted vcore (via an
event) to switch back to us, but by the time it does that, it may have other
events that got lost.  So we're back to polling the ev_qs that it might have
received INDIRs about.  Note that we still want to send an event with
sys_change_vcore().  We want the new vcore to know the old vcore was put
offline: a preemption (albeit one that it chose to do, and one that isn't stuck
in vcore context).

One older way to deal with this was to force the 2LS to deal with this. The 2LS
would check the ev_mboxes/ev_qs of all ev_qs that could send INDIRS to the
offline vcore.  There could be INDIRS in the VCPD that are just lying there.
The 2LS knows which ev_qs these are (such as for completed syscalls), and for
many things, this will be a common ev_q (such as for 'vcore-x-was-preempted').
However, this is a huge pain in the ass, since a preempted vcore could have the
spammed INDIR for an ev_q associated with another vcore.  To deal with this,
the 2LS would need to check *every* ev_q that requests INDIRs.  We don't do
this.

Instead, we simply have the remote core check the VCPD public mbox of the
preempted vcore.  INDIRs (and other vcore business that other vcores can handle)
will get sorted here.

3.3.5: Lists to Find Vcores
---------------
A process has three lists: online, bulk_preempt, and inactive.  These not only
are good for process management, but also for helping alert_vcore() find
potentially alertable vcores.  alert_vcore() and its associated helpers are
failry complicated and heavily commented.  I've set things up so both the
online_vcs and the bulk_preempted_vcs lists can be handled the same way: post to
the first element, then see if it still VC_CAN_RCV_MSG.  If not, if it is still
the first on the list, then it hasn't proc_yield()ed yet, and it will eventually
restart when it tries to yield.  And this all works without locking the
proc_lock.  There are a bunch more details and races avoided.  Check the code
out.

3.3.6: Vcore Business and the VCPD mboxs
---------------
There are two types of VCPD mboxes: public and private.  Public ones will get
handled during preemption recovery.  Messages sent here need to be handle-able
by any vcore.  Private messages are for that specific vcore.  In the common
case, the public mbox will usually only get looked at by its vcore.  Only during
recovery and some corner cases will we deal with it remotely.

Here's some guidelines: if you message is spammy and the handler can deal with
spurious events and it doesn't need to be on a specific vcore, then go with
public.  Examples of public mbox events are ones that need to be spammed:
preemption recovery, INDIRs, etc.  Note that you won't need to worry about
these: uthread code and the kernel handle them.  But if you have something
similar, then that's where it would go.  You can also send non-spammy things,
but there's no guarantee they'll be looked at.

Some messages should only be sent to the private mbox.  These include ones that
make no sense for other vcores to handle.  Examples: 2LS IPIs/preemptions (like
"change your scheduling policy vcore 3", preemption-pending notifs from the
kernel, timer interrupts, etc.

An example of something that shouldn't be sent to either is syscall completions.
They can't be spammed, so you can't send them around like INDIRs.  And they need
to be dealt with.  Other than carefully-spammed public messages, there's no
guarantee of getting a message for certain scenarios (yields).  Instead, use an
ev_q with INDIR set.

Also note that a 2LS could set up a big ev_q with EVENT_IPI and not EVENT_INDIR,
and then poll for that in their vcore_entry().  This is equivalent to setting up
a small ev_q with EVENT_IPI and pointing it at the private mbox.

3.4 Application-specific Event Handling
---------------------------------------
So what happens when the vcore/2LS isn't handling an event queue, but has been
"told" about it?  This "telling" is in the form of an IPI.  The vcore was
prodded, but is not supposed to handle the event.  This is actually what
happens now in Linux when you send signals for AIO.  It's all about who (which
thread, in their world) is being interrupted to process the work in an
application specific way.  The app sets the handler, with the option to have a
thread spawned (instead of a sighandler), etc.

This is not exactly the same as the case above where the ev_mbox* pointed to
the vcore's default mbox.  That issue was just about avoiding extra messages
(and messages in weird orders).  A vcore won't handle an ev_q if the
message/contents of the queue aren't meant for the vcore/2LS.  For example, a
thread can want to run its own handler, perhaps because it performs its own
asynchronous I/O (compared to relying on the 2LS to schedule synchronous
blocking u_threads).

There are a couple ways to handle this.  Ultimately, the application is supposed
to handle the event.  If it asked for an IPI, it is because something ought to
be done, which really means running a handler.  We used to support the
application setting EVENT_THREAD in the ev_q's flags, and the 2LS would spawn a
thread to run the ev_q's handler.  Now we just have the application block a
uthread on the evq.  If an ev_handler is set, the vcore will execute the
handler itself.  Careful with this, since the only memory it touches must be
pinned, the function must not block (this is only true for the handlers called
directly out of vcore context), and it should return quickly.

Note that in either case, vcore-written code (library code) does not look at
the contents of the notification event.  Also note the handler takes the whole
event_queue, and not a specific message.  It is more flexible, can handle
multiple specific events, and doesn't require the vcore code to dequeue the
event and either pass by value or allocate more memory.

These ev_q handlers are different than ev_handlers.  The former handles an
event_queue.  The latter is the 2LS's way to handle specific types of messages.
If an app wants to process specific messages, have them sent to an ev_q under
its control; don't mess with ev_handlers unless you're the 2LS (or example
code).

Continuing the analogy between vcores getting IPIs and the OS getting HW
interrupts, what goes on in vcore context is like what goes on in interrupt
context, and the threaded handler is like running a threaded interrupt handler
(in Linux).  In the ROS world, it is like having the interrupt handler kick
off a kernel message to defer the work out of interrupt context.

If neither of the application-specific handling flags are set, the vcore will
respond to the IPI by attempting to handle the event on its own (lookup table
based on the type of event (like "syscall complete")).  If you didn't want the
vcore to handle it, then you shouldn't have asked for an IPI.  Those flags are
the means by which the vcore can distinguish between its event_qs and the
applications.  It does not make sense otherwise to send the vcore an IPI and
an event_q, but not tell give the code the info it needs to handle it.

In the future, we might have the ability to block a u_thread on an event_q, so
we'll have other EV_ flags to express this, and probably a void*.  This may
end up being redudant, since u_threads will be able to block on syscalls (and
not necessarily IPIs sent to vcores).

As a side note, a vcore can turn off the IPI wanted flag at any time.  For
instance, when it spawns a thread to handle an ev_q, the vcore can turn off
IPI wanted on that event_q, and the thread handler can turn it back on when it
is done processing and wants to be re-IPId.  The reason for this is to avoid
taking future IPIs (once we leave vcore context, IPIs are enabled) to let us
know about an event for which a handler is already running.

3.5 Overflowed/Missed Messages in the VCPD 
---------------------------------------
This too is no longer necessary.  It's useful in that it shows what we don't
have to put up with.  Missing messages requires potentially painful
infrastructure to handle it:

-----------------------------
All event_q's requesting IPIs ought to register with the 2LS.  This is for
recovering in case the vcpd's mbox overflowed, and the vcore knows it missed a
NE_EVENT type message.  At that point, it would have to check all of its
IPI-based queues.  To do so, it could check to see if the mbox has any
messages, though in all likelihood, we'll just act as if there was a message
on each of the queues (all such handlers should be able to handle spurious
IPIs anyways).  This is analagous to how the OS's block drivers don't solely
rely on receiving an interrupt (they deal with it via timeouts).  Any user
code requiring an IPI must do this.  Any code that runs better due to getting
the IPI ought to do this.

We could imagine having a thread spawned to handle an ev_q, and the vcore
never has to touch the ev_q (which might make it easier for memory
allocation).  This isn't a great idea, but I'll still explain it.  In the
notif_ev message sent to the vcore, it has the event_q*.  We could also send a
flag with the same info as in the event_q's flags, and also send the handler.
The problem with this is that it isn't resilient to failure.  If there was a
message overflow, it would have the check the event_q (which was registered
before) anyway, and could potentially page fault there.  Also the kernel would
have faulted on it (and read it in) back when it tried to read those values.
It's somewhat moot, since we're going to have an allocator that pins event_qs.
-----------------------------

3.6 Round-Robin or Other IPI-delivery styles
---------------------------------------
In the same way that the IOAPIC can deliver interrupts to a group of cores,
round-robinning between them, so can we imagine processes wanting to
distribute the IPI/active notification of events across its vcores.  This is
only meaningful is the NOTIF_IPI_WANTED flag is set.

Eventually we'll support this, via a flag in the event_q.  When
NE_ROUND_ROBIN, or whatever, is set a couple things will happen.  First, the
vcore field will be used in a "delivery-specific" manner.  In the case of RR,
it will probably be the most recent destination.  Perhaps it will be a bitmask
of vcores available to receive.  More important is the event_mbox*.  If it is
set, then the event message will be sent there.  Whichever vcore gets selected
will receive an IPI, and its vcpd mbox will get a NE_EVENT message.  If the
event_mbox* is 0, then the actual message will get delivered to the vcore's
vcpd mbox (the default location).

3.7 Event_q-less Notifications
---------------------------------------
Some events needs to be delivered directly to the vcore, regardless of any
event_qs.  This happens currently when we bypass the notification table (e.g.,
sys_self_notify(), preemptions, etc).  These notifs will just use the vcore's
default mbox.  In essence, the ev_q is being generated/sent with the call.
The implied/fake ev_q points to the vcpd's mbox, with the given vcore set, and
with IPI_WANTED set.  It is tempting to make those functions take a
dynamically generated ev_q, though more likely we'll just use the lower level
functions in the kernel, much like the Round Robin set will need to do.  No
need to force things to fit just for the sake of using a 'solution'.  We want
tools to make solutions, not packaged solutions.

3.8 UTHREAD_DONT_MIGRATE
---------------------------------------
DONT_MIGRATE exists to allow uthreads to disable notifications/IPIs and enter
vcore context.  It is needed since you need to read vcoreid to disable notifs,
but once you read it, you need to not move to another vcore.  Here are a few
rules/guidelines.

We turn off the flag so that we can disable notifs, but turn the flag back on
before enabling.  The thread won't get migrated in that instant since notifs are
off.  But if it was the other way, we could miss a message (because we skipped
an opportunity to be dropped into vcore context to read a message).

Don't check messages/handle events when you have a DONT_MIGRATE uthread.  There
are issues with preemption recovery if you do.  In short, if two uthreads are
both DONT_MIGRATE with notifs enabled on two different vcores, and one vcore
gets preempted while the other gets an IPI telling it to recover the other one,
both could keep bouncing back and forth if they handle their preemption
*messages* without dealing with their own DONT_MIGRATEs first.  Note that the
preemption recovery code can handle having a DONT_MIGRATE thread on the vcore.
This is a special case, and it is very careful about how cur_uthread works.

All uses of DONT_MIGRATE must reenable notifs (and check messages) at some
point.  One such case is uthread_yield().  Another is mcs_unlock_notifsafe().
Note that mcs_notif_safe locks have uthreads that can't migrate for a
potentially long time.  notifs are also disabled, so it's not a big deal.  It's
basically just the same as if you were in vcore context (though technically you
aren't) when it comes to preemption recovery: we'll just need to restart the
vcore via a syscall.  Also note that it would be a real pain in the ass to
migrate a notif_safe locking uthread.  The whole point of it is in case it grabs
a lock that would be held by vcore context, and there's no way to know it isn't
a lock on the restart-path.

3.9 Why Preemption Handling Doesn't Lock Up (probably)
---------------------------------------
One of the concerns of preemption handling is that we don't get into some form
of livelock, where we ping-pong back and forth between vcores (or a set of
vcores), all of which are trying to handle each other's preemptions.  Part of
the concern is that when a vcore sys_changes to another, it can result in
another preemption message being sent.  We want to be sure that we're making
progress, and not just livelocked doing sys_change_vcore()s.

A few notes first:
1) If a vcore is holding locks or otherwise isn't handling events and is
preempted, it will let go of its locks before it gets to the point of
attempting to handle any other vcore preemption events.  Event handling is only
done when it is okay to never return (meaning no locks are held).  If this is
the situation, eventually it'll work itself out or get to a potential ping-pong
scenario.

2) When you change_to while handling preemption, once you start back up, you
will leave change_to and eventually fetch a new event.  This means any
potential ping-pong needs to happen on a fresh event.

3) If there are enough pcores for the vcores to all run, we won't issue any
change_tos, since the vcores are no longer preempted.  This means we only are
worried about situations with insufficient vcores.  We'll mostly talk about 1
pcore and 2 vcores.

4) Preemption handlers will not call change_to on their target vcore if they
are also the one STEALING from that vcore.  The handler will stop STEALING
first.

So the only way to get stuck permanently is if both cores are stuck doing a
sys_change_to(FALSE).  This means we want to become the other vcore, *and* we
need to restart our vcore where it left off.  This is due to some invariant
that keeps us from abandoning vcore context.  If we were to abandon vcore
context (with a sys_change_to(TRUE)), we basically don't need to be
preempt-recovered.  We already packaged up our cur_uthread, and we know we
aren't holding any locks or otherwise breaking any invariants.  The system will
work fine if we never run again.  (Someone just needs to check our messages).

Now, there are only two cases where we will do a sys_change_to(FALSE) *while*
handling preemptions.  Again, we aren't concerned about things like MCS-PDR
locks; those all work because the change_tos are done where we'd normally just
busy loop.  We are only concerned about change_tos during handle_vc_preempt.
These two cases are when the changing/handling vcore has a DONT_MIGRATE uthread
or when someone else is STEALING its uthread.  Note that both of these cases
are about the calling vcore, not its target.

If a vcore (referred to as "us") has a DONT_MIGRATE uthread and it is handling
events, it is because someone else is STEALING from our vcore, and we are in
the short one-shot event handling loop at the beginning of
uthread_vcore_entry().  Whichever vcore is STEALING will quickly realize it
can't steal (it sees the DONT_MIGRATE), and bail out.  If that vcore isn't
running now, we will change_to it (which is the purpose of our handling their
preemption).  Once that vcore realizes it can't steal, it will stop STEALING
and change to us.  At this point, no one is STEALING from us, and we move along
in the code.  Specifically, we do *not* handle events (we now have an event
about the other vcore being preempted when it changed_to us), and instead we
start up the DONT_MIGRATE uthread and let it run until it is migratable, at
which point we handle events and will deal with the other vcore.  

So DONT_MIGRATE will be sorted out.  Likewise, STEALING gets sorted out too,
quite easily.  If someone is STEALING from us, they will quickly stop STEALING
and change to us.  There are only two ways this could even happen: they are
running concurrently with us, and somehow saw us out of vcore context before
deciding to STEAL, or they were in the process of STEALING and got preempted by
the kernel.  They would not have willingly stopped running while STEALING our
cur_uthread.  So if we are running and someone is stealing, after a round of
change_tos, eventually they run, and stop STEALING.

Note that once someone stops STEALING from us, they will not start again,
unless we leave vcore context.  If that happened, we basically broke out of the
ping-pong, and now we're onto another set of preemptions.  We wouldn't leave
vcore context if we still had preemption events to deal with.

Finally, note that we needed to only check for one message at a time at the
beginning of uthread_vcore_entry().  If we just handled the entire mbox without
checking STEALING, then we might not break out of that loop if there is a
constant supply of messages (perhaps from a vcore in a similar loop).

Anyway, that's the basic plan behind the preemption handler and how we avoid
the ping-ponging.  change_to_vcore() is built so that we handle our own
preemption before changing (pack up our current uthread), so that we make
progress.  The two cases where we can't do that get sorted out after everyone
gets to run once, and since you can't steal or have other uthread's turn on
DONT_MIGRATE while we're in vcore context, eventually we clear everything up.
There might be other bugs or weird corner cases, possibly involving multiple
vcores, but I think we're okay for now.

3.10: Handling Messages for Other Vcores
---------------------------------------
First, remember that when a vcore handles an event, there's no guarantee that
the vcore will return from the handler.  It may start fresh in vcore_entry().

The issue is that when you handle another vcore's INDIRs, you may handle
preemption messages.  If you have to do a change_to, the kernel will make sure
a message goes out about your demise.  Thus someone who recovers that will
check your public mbox.  However, the recoverer won't know that you were
working on another vcore's mbox, so those messages might never be checked.

The way around it is to send yourself a "check the other guy's messages" event.
When we might change_to and never return, if we were dealing with another
vcores mbox, we'll send ourselves a message to finish up that mbox (if there
are any messages left).  Whoever reads our messages will eventually get that
message, and deal with it.

One thing that is a little ugly is that the way you deal with messages two
layers deep is to send yourself the message.  So if VC1 is handling VC2's
messages, and then wants to change_to VC3, VC1 sends a message to VC1 to check
VC2.  Later, when VC3 is checking VC1's messages, it'll handle the "check VC2's messages"
message.  VC3 can't directly handle VC2's messages, since it could run a
handler that doesn't return.  Nor can we just forget about VC2.  So VC3 sends
itself a message to check VC2 later.  Alternatively, VC3 could send itself a
message to continue checking VC1, and then move on to VC2.  Both seem
equivalent.  In either case, we ought to check to make sure the mbox has
something before bothering sending the message.

So for either a "change_to that might not return" or for a "check INDIRs on yet
another vcore", we send messages to ourself so that we or someone else will
deal with it.

Note that we use TLS to track whether or not we are handling another vcore's
messages, and if we do plan to change_to that might not return, we clear the
bool so that when our vcore starts over at vcore_entry(), it starts over and
isn't still checking someone elses message.

As a reminder of why this is important: these messages we are hunting down
include INDIRs, specifically ones to ev_qs such as the "syscall completed
ev_q".  If we never get that message, a uthread will block forever.  If we
accidentally yield a vcore instead of checking that message, we would end up
yielding the process forever since that uthread will eventually be the last
one, but our main thread is probably blocked on a join call.  Our process is
blocked on a message that already came, but we just missed it. 

4. Single-core Process (SCP) Events:
====================
4.1 Basics:
---------------------------------------
Event delivery is important for SCP's blocking syscalls.  It can also be used
(in the future) to deliver POSIX signals, which would just be another kernel
event.

SCPs can receive events just like MCPs.  For the most part, the code paths are
the same on both sides of the K/U interface.  The kernel sends events (which
can detect an SCP and will send it to vcore0), the kernel will make sure you
can't yield/miss an event, etc.  Userspace preps vcore context in advance, and
can do all the things vcore context does: handle events, select thread to run.
For an SCP, there is only one thread to run.

4.2 Degenerate Event Delivery:
---------------------------------------
That being said, there are a few tricky things.  First, there is a time before
the SCP is ready to fully receive events.  Specifically, before
vcore_event_init(), which is called out of glibc's _start.  More importantly,
the runtime linker never calls that function, yet it wants to block.

The important thing to note is that there are a few parts to event delivery:
registration (user), sending the event (kernel), making sure the proc wakes up
(kernel), and actually handling the event (user).  For syscalls, the only thing
the process (even rtld) needs is the first three.  Registration is easy - can be
done with nothing more than kernel headers (no need for parlib) for NO_MSG ev_qs
(no need to init the UCQ).  Event handling is trickier, and requires parlib
(which rtld can't link against).  To support processes that could register for
events, but not handle them (or even enter vcore context), the kernel needed a
few changes (checking the VC_SCP_NOVCCTX flag) so that it would wake the
process, but never put it in vcore context.  

This degenerate event handling just wakes the process up, at which point it can
check on its syscall.  Very early in the process's life, it'll init vcore0's
UCQ and be able to handle full events, enter vcore context, etc.

Once the SCP is up and running, it can receive events like normal.  One thing to
note is that the SCPs are not using a handle_syscall() event handler, like the
MCPs do.  They are only using the event to get the process restarted, at which
point their vcore 0 restarts thread0.  One consequence of this is that if a
process receives an unrelated event while blocking on a syscall, it'll handle
that event, then restart thread0.  Thread0 will see its syscall isn't complete,
and then re-block.  (It also re-registers its ev_q, which is harmless).  When
that syscall is finally done, the kernel will send an event and wake it up
again.

4.3 Extra Tidbits:
---------------------------------------
If we receive an event right as we transition from SCP to MCP, vcore0 could get
spammed with a message that is never received.  Right now, it's not a problem,
since vcore0 is the first vcore that will get woken up as an MCP.  This could be
an issue if we ever allow transitions from MCP back to SCP.

On a related note, it's now wrong for SCPs to sys_yield(FALSE) (not being nice,
meaning they are waiting for an event) in a loop that does not check events or
otherwise allow them to break out of that loop.  This should be fairly obvious.
A little more subtle is that these loops also need to sort out notif_pending.
If you are trying to yield and still have an old notif_pending set, the kernel
won't let you yield (it thinks you are missing the notif).  For the degenerate
mode, (VC_SCP_NOVCCTX is set on vcore0), the kernel will handle dealing with
this flag.

Finally, note that while the SCP is in vcore context, it has none of the
guarantees of an MCP.  It's somewhat meaningless to talk about being gang
scheduled or knowing about the state of other vcores.  If you're running, you're
on a physical core.  You may get unexpected interrupts, descheduled, etc.  Aside
from the guarantees and being the only vcore, the main differences are really up
to the kernel scheduler.  In that sense, we have somewhat of a new state for
processes - SCPs that can enter vcore context.  From the user's perspective,
they look a lot like an MCP, and the degenerate/early mode SCPs are like the
old, dumb SCPs.  The big difference for userspace is that there isn't a 2LS yet
(will need to reinit things slightly).  The kernel treats SCPs and MCPs very
differently too, but that may not always be the case.

5. Misc Things That Aren't Sorted Completely:
====================
5.1 What about short handlers?
---------------------------------------
Once we sort the other issues, we can ask for them via a flag in the event_q,
and run the handler in the event_q struct.

5.2 What about blocking on a syscall?
---------------------------------------
The current plan is to set a flag, and let the kernel go from there.  The
kernel knows which process it is, since that info is saved in the kthread that
blocked.  One issue is that the process could muck with that flag and then go
to sleep forever.  To deal with that, maybe we'd have a long running timer to
reap those.  Arguably, it's like having a process while(1).  You can screw
yourself, etc.  Killing the process would still work.