Process and Port Tables — erts v15.2.7.5

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# Process and Port Tables

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## Problems

The process table is a mapping from process identifiers to process structure pointers. The process structure contains miscellaneous information about a process, as for example pointers to its heap, message queue, etc. When the runtime system needs to operate on a process, it looks up the process structure in the process table using the process identifier. An example of this is when passing a message to a process.

The process table has for a very long time just been an array of pointers to process structures. Since process identifiers internally in the runtime system are 28-bit integers it is quite easy to map a process identifier to index into the array. The 28-bits were divided into two sets. The least significant set of bits was used as index into the array. The most significant set of bits was only used to be able to distinguish between a number of identifiers with which map to the same index in the array. As long as process table sizes of a power of two was used we had 2^28 unique process identifiers.

When the first SMP support was implemented, the table still was kept more or less the same way, but protected by two types of locks. One lock that protected the whole table against modifications and an array of locks protecting different parts of the table. The exact locking strategy previously used isn't interesting. What is interesting is that it suffered from heavy lock contention especially when lots of modifications was being made, but also when only performing lookups.

In order to be able to detect when it is safe to deallocate a previously used process structure, reference counting of the structure was used. Also this was problematic, since simultaneous lookups needed to modify the reference counter which also caused contention on the cache line where the reference counter was located. This since all modifications needs to be communicated between all involved processors.

The port table is very similar to the process table. The major difference, at least in concept, is that it is a mapping from port identifiers to port structures. It had a similar implementation, but with some differences. Instead of being an array of pointers it was an array of structures, and instead of being protected by two types of locks it was only protected by one global lock. This table also suffered from lock contention in various situations.

## Solution

The process table was the major problem to address since processes are much more frequently used than ports. The first implementation only implemented this for processes, but since the port table is very similar and very similar problems occur on the port table, the process table implementation was later generalized so that it could also be used implementing the port table. For simplicity I will only talk about the process table in the following text, but the same will apply to the port table unless otherwise stated.

If we disregard the locking issues, the original solution is very appealing. The mapping from process identifier to index into the array is very fast, and this property is something we would like to keep. The vast majority of operations on these tables are lookups so optimizing for lookups is what we want to do.

### Lookup

Using a set of bits in the process identifier as index into an array seems hard to beat. By replacing the array of pointers with an array of our pointer sized atomic data type, a lookup will consist of the following:

1. Mapping the 28-bit integer to an index into the array.

   More about this mapping later.

2. Read the pointer using an atomic memory operation at determined index in array.

   On all platforms that we provide atomic memory operations, this is just a `volatile` read, preventing the compiler to use values in registers, forcing the a read from memory.

3. Depending on use, issue appropriate memory barrier.

   A common barrier used is a barrier with acquire semantics. On x86/x86_64 this maps to a compiler barrier preventing the compiler to reorder instructions, but on other hardware often some kind of light weight hardware memory barrier is also needed.

   When comparing with a locked approach, at least one heavy weight memory barrier will be issued when locking the lock on most, if not all, hardware architectures (including x86/x86_64), and often some kind of light weight memory barrier will be issued when unlocking the lock.

When looking at this very simple solution with very little overhead you might wonder why we didn't implement it this way from the beginning. It all boils down to the read operation of the pointer. We need some way to know that it is safe to access the memory pointed to. One way of doing this is to place a reference counter in the process structure. Increment of the reference counter at lookup needs to be done atomically with the lookup. A lock can typically provide this service for us, which was the approach we previously used. Another approach could be to co-locate the reference counter with the pointer in the table. The major problem with this approach is the modifications of the reference counter. This since these modification would have to be communicated between all involved processor cause contention on the cache line containing the reference counter. The new lookup approach above is possible since we can use the "thread progress" functionality in order to determine when it is safe to deallocate the process structure. We'll get back to this when describing deletion in the table.

Using this new lookup approach we won't modify any memory at all which is important. A lookup conceptually only read memory, now this is true in the implementation also which is important from a scalability perspective. The previous implementation modified the cache line containing the reference counter two times, and the cache line containing the corresponding lock two times at each lookup.

### Modifications of the Table

A lightweight lookup in the table was the most important feature, but we also wanted to improve modifications of the table. The process table is modified when a new process is spawned, i.e. a new pointer is inserted into the table, and when a process terminates, i.e. a pointer is deleted in the table.

Assuming that we spawn fewer processes than the maximum amount of unique process identifiers in the system, one has always been able to determine the order of process creation just by comparing process identifiers. If PidX is larger than PidY, then PidX was created after PidY assuming both identifiers originates from the same node. However, since we have a quite limited amount of unique identifiers today (2^28), this property cannot be relied upon if we create large amount of processes. But never the less, this is a property the system always have had.

If we would have had a huge amount of unique identifiers available, it would have tempting to drop or modify this ordering property as described above. The ordering property could for example be based on the scheduler performing the spawn operation. It would have been possible to reserve large ranges of identifiers exclusive for each scheduler thread which could be used minimizing the need for communication when allocating identifiers. The amount of identifiers we got to work with today is, however, not even close to be enough for such an approach.

Since we have a limited amount of unique identifiers, we need to be careful not to waste them. If previously used identifiers are reused too quick, identifiers originating from terminated processes will refer to newly created processes, and mixups will occur. The previously used approach was quite good at not wasting identifiers. Using a modified version of the same approach also lets us keep the ordering property that we have always had.

#### Insert

The original approach is more or less to search for next free index or slot in the array. The search starts from the last slot allocated. If we reach the end of the array we increase a "wrapped counter" and then continue the search. The process identifier is constructed by writing the index to the least significant set of bits, and the "wrapped counter" to the most significant set of bits. The amount of bits in each set of bits is decided at boot time, so that maximum index will just fit into the least significant set of bits.

In the modified lock free version of this approach we more or less do it the same way, but with some important modifications trying to avoid unnecessary contention when multiple schedulers create processes simultaneously. Since multiple threads might be trying to search for the next free slot at the same time from the same starting point we want subsequent slots to be located in different cache lines. Multiple schedulers simultaneously writing new pointers into the table are therefore very likely to write into adjacent slots. If adjacent slots are located in the same cache line all modification of this cache line needs to be communicated between all involved processors which will be very expensive and scale very poor. By locating adjacent slots in different cache lines only true conflicts will trigger communication between involved processors, i.e., avoiding false sharing.

A cache line is larger than a pointer, typically 8 or 16 times larger, so using one cache line for each slot only containing one pointer would be a waste of space. Each cache line will be able to hold a fixed amount of slots. The first slot of the table will be the first slot of the first cache line, the second slot of the table will be the first slot of the second cache line until we reach the end of the array. The next slot after that will be the second slot of the first cache line, etc, moving forward one cache line internal slot each time we wrap. This way we will be able to fit the same amount of pointers into an array of the same size while always keeping adjacent slots in different cache lines.

The mapping from identifier to slot or index into the array gets a bit more complicated than before. Instead of a `shift` and a bitwise `and`, we get two `shift`s, two bitwise `and`s, and an `add` (see implementation of `erts_ptab_data2pix()` in `erl_ptab.h`). However, by storing this information optimized for lookup we only need a `shift` and a bitwise `and` on 32-bit platforms. On 64-bit platforms we got enough room for the 28-bit identifier in the least significant halfword, and the index in the most significant halfword, in other words, we just need to read the most significant halfword to get the index. That is, this operation is as fast, or faster than before. The downside is that on 32-bit platforms we need to convert this information into the 28-bit identifier number when printing, or when ordering identifiers from the same node. These operations are, however, extremely infrequent compared to lookups.

When we insert a new element in the table we do the following:

1. We begin by reserving space in the table by atomically incrementing a counter of processes in the table. If our increment brings the counter above the maximum size of the table, the operation fail and a `system_limit` exception is raised.

2. The table contains a 64-bit atomic variable of the last identifier used. Only the least significant bits will be used when actually creating the identifier. This identifier is where the search begin.

3. We increment last identifier value used. In order determine the slot that corresponds to this identifier we call `erts_ptab_data2pix()` that maps identifier to slot. We read the content of the slot. If the slot is free we try to write a reservation marker using an atomic compare and swap. If this fails we repeat this step until it succeeds.

4. Change the table variable of last identifier used. Since multiple writes might occur at the same time this value may already have been changed by to an identifier larger that the one we got. In this case we can continue; otherwise, we need to change it to the identifier we got.

5. We now do some initializations of the process structure that cannot be done before we know the process identifier, and have to be done before we publish the structure in the table. This, for example, includes storing the identifier in the process structure.

6. Now we can publish the structure in the table by writing the pointer to the process structure in the slot previously reserved in 3.

Using this approach we keep the properties like identifier ordering, and identifier reuse while improving performance and scalability. It has one flaw, though. There is no guarantee that the operation will terminate. This can quite easily be fixed though, and will be fixed in the next release. We will get back to this below.

#### Delete

When a process terminates, we mark the process as terminated in the process structure, the counter of number of processes in the table is decreased, and the reference to the process structure is removed by writing a `NULL` pointer into the corresponding slot. The scheduler thread performing this then schedule a thread progress later job which will do the final cleanup and deallocate the process structure. The thread progress functionality will make sure that this job will not execute until it is certain that all managed threads have dropped all references to the process structure.

### BIF Iterating Over the Table

The `erlang:processes/0` and `erlang:ports/0` BIFs iterate over the tables and return corresponding identifiers. These BIF should return a consistent snapshot of the table content during some time when the BIF is executing. In order to implement this we use locking in a strange way. We use an "inverted rwlock".

When performing lookups in the table we do not need to bother about the locking at all, but when modifying the table we read lock the rwlock protecting the table which allows for multiple writers during normal operation. When the BIF that iterates over the table need access to the table it write locks the rwlock and reads content of the table. The BIF do not read the whole table in one go but instead read small chunks at time only write locking while reading. The actual implementation of the BIFs is out of the scope of this document.

An out of the box rwlock will typically suffer from contention on the single cache line containing the state of the rwlock even in the case we are only read locking. Instead of using such an rwlock, we have our own implementation of reader optimized rwlocks which keeps track of reader threads in separate thread specific cache lines. This in order to avoid contention on a single cache line. As long as we only do read lock operations, threads only need to read a global cache line and modify its own cache line, and by this minimize communication between involved processors. The iterating BIFs are normally very infrequently used, so in the normal case we will only do read lock operations on the table global rwlock.

### Future Improvements

The first improvement is to fix the guarantee so that insert operations will be guaranteed to terminate. When the operation starts we verify that there actually exist a free slot that we can use. The problem is that we might not find it since it may move when multiple threads modify the table at the same time as we are trying to find the slot. The easy fix is to abort the operation if an empty slot could not be found in a finite number operation, and then restart the operation under a write lock. This will be implemented in next release, but further work should be made trying to find a better solution.

This and also previous implementation do not work well when the table is nearly full. We will both get long search times for free slots, and we will reuse identifiers more frequently since we more frequently wrap during the search. These tables works best when the table is much larger than the amount of simultaneous existing processes. One easy improvement is to always have room for more processes than we allow in the table. This will also be implemented in the next release, but this should probably also be worked more on trying to find an even better solution.

It would also be nice to get rid of the rwlock all together. The use of a reader optimized rwlock makes sure we do not any contention on the lock, but unnecessary memory barriers will be issued due to the lock. The main issue here is to modify iterating BIFs so that they do not require exclusive access to the table while reading a sequence of slots. In principle this should be rather easy, the code can handle sequences of variable sizes, so shrinking the sequence size of slots to one would solv the problem. This will, however, need some tweeks and modifications of not trivial code, but is something that should be looked at in the future.

By increasing the size of identifiers, at least on 64-bit machines (which isn't as easy as it first might seem) we get further room for improvement. Besides the obvious improvement of not reusing identifiers as fast as we currently do, it makes it possible to further avoid contention when inserting elements in the table. At least if we drop this ordering property, which isn't that useful anyway.

### Some Benchmark Results

In order to test modifications of the process table we ran a couple of benchmarks where lots of processes are spawned and terminated simultaneously, and got a speedup of between 150-200%. Running a similar benchmark but with ports we got a speedup of about 130%.

The BIF `erlang:is_process_alive/1` is the closest you can get to a process table lookup only. The BIF looks up the process corresponding to the process identifier passed as argument, and then checks if it is alive. By running multiple processes looping over this BIF checking the same process, we get a speedup between 20000-23000%. Conceptually this operation only involve read operations. In the implementation used in R16B also only read operation are performed, while the previous implementation need to lock structures in order to read the data, suffering from both lock contention and contention due to modifications of cache lines used by lock internal data structures and the reference counter on the process being looked up.

The benchmarks were run on a relatively new machine with an Intel i7 quad core processor with hyper-threading using 8 schedulers. On a machine with more communication overhead and/or larger amount of logical processors the speedups are expected to be even larger.

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