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120 lines
19 KiB
120 lines
19 KiB
\chapter{Implementation}
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\label{chap:implementation}
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% Hier greift man einige wenige, interessante Gesichtspunkte der
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% Implementierung heraus. Das Kapitel darf nicht mit Dokumentation oder
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% sehr viele Gesichtspunkte aufgegriffen werden müssen, ist aber nicht
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% machen, daß man es bei der Arbeit nicht mit einem "Papiertiger"
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% auch ein sehr wichtiger Text für jemanden, der die Arbeit später
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% fortsetzt. Der dritte Gesichtspunkt dabei ist, einem Leser einen etwas
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In this chapter, we concentrate on specific implementation details, offering an in-depth view of how the design promises outlined in Chapter \ref{chap:design} are realized. Firstly, we delve into the usage of locking and atomics to achieve thread safety. Subsequently, we provide an example of the policy functions alluded to in Section \ref{sec:design:accel-usage}. Finally, we apply the cache to \glsentrylong{qdp}. \par
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\section{Locking and Usage of Atomics}
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The usage of locking and atomics has proven to be challenging. Their use is performance-critical, and mistakes may lead to deadlock. Consequently, these aspects constitute the most interesting part of the implementation, which is why this chapter will extensively focus on the details of their implementation. \par
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\subsection{Cache State Lock} \label{subsec:implementation:cache-state-lock}
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To keep track of the current cache state the \texttt{Cache} will hold a reference to each currently existing \texttt{CacheData} instance. The reason for this is twofold: In Section \ref{sec:design:cache} we decided to keep elements in the cache until forced by memory pressure to remove them. Secondly in Section \ref{subsec:design:cache-entry-reuse} we decided to reuse one cache entry for multiple consumers. The second part requires access to the structure holding this reference to be thread safe when accessing and modifying the cache state in \texttt{Cache::Access}, \texttt{Cache::Flush} and \texttt{Cache::Clear}. The latter two both require unique locking, preventing other calls to \texttt{Cache} from making progress while the operation is being processed. For \texttt{Cache::Access} the use of locking depends upon the caches state. At first, only a shared lock is acquired for checking whether the given address already resides in cache, allowing other \texttt{Cache::Access}-operations to also perform this check. If no entry for the region is present, a unique lock is required as well when adding the newly created entry to cache. \par
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A map-datastructure was chosen to represent the current cache state with the key being the memory address of the entry and as value the \texttt{CacheData} instance. As the caching policy is controlled by the user, one datum may be requested for caching in multiple locations. To accommodate this, one map is allocated for each available \glsentrylong{numa:node} of the system. This can be exploited to reduce lock contention by separately locking each \gls{numa:node}'s state instead of utilizing a global lock. This ensures that \texttt{Cache::Access} and the implicit \texttt{Cache::Flush} it may cause can not hinder progress of caching operations on other \gls{numa:node}s. Both \texttt{Cache::Clear} and a complete \texttt{Cache::Flush} as callable by the user will now iteratively perform their respective task per \gls{numa:node} state, also allowing other \gls{numa:node} to progress.\par
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Even with this optimization, in scenarios where the \texttt{Cache} is frequently tasked with flushing and re-caching by multiple threads from the same node, lock contention will negatively impact performance by delaying cache access. Due to passive waiting, this impact might be less noticeable when other threads on the system are able to make progress during the wait. \par
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\subsection{CacheData Atomicity}
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Throughout this section we will use the term \enquote{handler}, which was coined by \gls{intel:dml}, referring to an object associated with an operation on the accelerator. Through it, the state of a task may be queried, making the handler our connection to the asynchronously executed task. As we may split up one single copy into multiple distinct tasks for submission to multiple \gls{dsa}s, \texttt{CacheData} internally contains a vector of multiple of these handlers.
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The choice made in \ref{subsec:design:cache-entry-reuse} necessitates thread-safe shared access to the same resource. The C++ standard library provides \texttt{std::shared\_ptr<T>}, a reference-counted pointer that is thread-safe for the required operations \cite{cppreference:shared-ptr}, making it a suitable candidate for this task. Although an implementation using it was explored, it presented its own set of challenges. \par
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As we aim to minimize the time spent in a locked region, only the task is added to the \gls{numa:node}'s cache state when locked, with the submission taking place outside the locked region. We assume the handlers of \gls{intel:dml} to be unsafe for access from multiple threads. To achieve the safety for \texttt{CacheData::WaitOnCompletion}, outlined in \ref{subsec:design:cache-entry-reuse}, threads need to coordinate which one performs the actual waiting. To avoid queuing multiple copies of the same task, the task must be added to the cache state before submission. These two aspects necessitate modifying the handlers atomically. We therefore use an atomic pointer in \texttt{CacheData} to allow safe exchange and waiting on modification. \par
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Using \texttt{std::shared\_ptr<T>} also introduces uncertainty, relying on the implementation to be performant. The standard does not specify whether a lock-free algorithm is to be used, and \cite{shared-ptr-perf} suggests abysmal performance for some implementations, although the full article is in Korean. No further research was found on this topic. \par
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Therefore, the decision was made to implement atomic reference counting for \texttt{CacheData}. This involves providing a custom constructor and destructor wherein a shared atomic integer is either incremented or decremented using atomic fetch sub and add operations to modify the reference count. In the case of a decrease to zero, the destructor was called for the last reference and then performs the actual destruction. \par
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\begin{figure}[h!tb]
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\centering
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\includegraphics[width=0.9\textwidth]{images/seq-blocking-wait.pdf}
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\caption{Sequence for Blocking Scenario. Observable in first draft implementation. Scenario where \(T_1\) performed first access to a datum followed \(T_2\) and \(T_3\). Then \(T_1\) holds the handlers exclusively, leading to the other threads having to wait for \(T_1\) to perform the work submission and waiting before they can access the datum through the cache.}
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\label{fig:impl-cachedata-threadseq-waitoncompletion}
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\end{figure}
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Due to the possibility of access by multiple threads, the implementation of \texttt{CacheData::WaitOnCompletion} proved to be challenging. In the first implementation, a thread would check if the handlers are available and atomically wait \cite{cppreference:atomic-wait} on a value change from nullptr, if they are not. As the handlers are only available after submission, a situation could arise where only one copy of \texttt{CacheData} is capable of actually waiting on them. \par
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To illustrate this, an exemplary scenario is used, as seen in the sequence diagram Figure \ref{fig:impl-cachedata-threadseq-waitoncompletion}. Assume that three threads \(T_1\), \(T_2\) and \(T_3\) wish to access the same resource. \(T_1\) is the first to call \texttt{CacheData::Access} and therefore adds it to the cache state and will perform the work submission. Before \(T_1\) may submit the work, it is interrupted and \(T_2\) and \(T_3\) obtain access to the incomplete \texttt{CacheData} on which they wait, causing them to see a nullptr for the handlers but invalid cache pointer, leading to atomic wait on the cache pointer (marked blue lines in Figure \ref{fig:impl-cachedata-threadseq-waitoncompletion}). \(T_1\) submits the work and sets the handlers (marked red lines in Figure \ref{fig:impl-cachedata-threadseq-waitoncompletion}), while \(T_2\) and \(T_3\) continue to wait. Therefore, only \(T_1\) can trigger the waiting and is therefore capable of keeping \(T_2\) and \(T_3\) from progressing. This is undesirable as it can lead to deadlocking if by some reason \(T_1\) does not wait and at the very least may lead to unnecessary delay for \(T_2\) and \(T_3\) if \(T_1\) does not wait immediately. \par
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\begin{figure}[h!tb]
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\centering
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\includegraphics[width=1.0\textwidth]{images/nsd-cachedata-waitoncompletion.pdf}
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\caption{\texttt{CacheData::WaitOnCompletion} Pseudocode. Final rendition of the implementation for a fair wait function.}
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\label{fig:impl-cachedata-waitoncompletion}
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\end{figure}
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As a solution for this, a more intricate implementation is required. When waiting, the threads now immediately check whether the cache pointer contains a valid value and return if it does, as nothing has to be waited for in this case. We will use the same example as before to illustrate the second part of the waiting procedure. Both \(T_2\) and \(T_3\) arrive in this latter section as the cache was invalid at the point in time when waiting was called for. They now atomically wait on the handlers-pointer to change, instead of doing it the other way around as before. Now when \(T_1\) supplies the handlers, it also uses \texttt{std::atomic<T>::notify\_one} \cite{cppreference:atomic-notify-one} to wake at least one thread waiting on value change of the handlers-pointer, if there are any. Through this the exclusion that was observable in the first implementation is already avoided. If nobody is waiting, then the handlers will be set to a valid pointer and a thread may pass the atomic wait instruction later on. Following this wait, the handlers-pointer is atomically exchanged \cite{cppreference:atomic-exchange} with nullptr, invalidating it. Each thread again checks whether it has received a valid local pointer to the handlers from the exchange. If it has then the atomic operation guarantees that is now in sole possession of the pointer. The owning thread is tasked with actually waiting. All other threads will now regress and call \texttt{CacheData::WaitOnCompletion} again. The solo thread may proceed to wait on the handlers and should update the cache pointer. \par
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Additional cases must be considered for the latter implementation to be safe and free of deadlocks. We will now discuss these edge cases and their resolution. \par
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\subsubsection{Initial Invalid State}
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\label{subsubsec:impl:cdatomicity:initial-invalid-state}
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We previously mentioned the possibly problematic situation where both the cache pointer and the handlers are not yet available for an instance in \texttt{CacheData}. This situation is avoided explicitly by the implementation due to waiting on the handlers being atomically updated from nullptr to valid. When the handlers will be set in the future by the thread calling \texttt{Cache::Access} first, progress is guaranteed. \par
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\subsubsection{Invalid State on Immediate Destruction}
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The previous Section discussed the initial invalid state and noted that, as long as the handlers will be set in the future, progress is guaranteed. We now discuss the situation where handlers will not be set. This situation is encountered when a memory region is accessed by threads \(T_1\) and \(T_2\) concurrently. One will win the data race to add the entry to the cache state, we choose \(T_1\). \(T_2\) then must follow Section \ref{subsec:design:cache-entry-reuse} and return the entry already present in cache state. Therefore, \(T_2\) has to destroy the \texttt{CacheData} instance it created previously. \par
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The destructor of \texttt{CacheData} waits on operation completion in order to ensure that no running jobs require the cache memory region, before deallocating it. This necessitates usability of \texttt{CacheData::WaitOnCompletion} for the case of immediate destruction. As the instance of \texttt{CacheData} is destroyed immediately, no tasks will be submitted to the \gls{dsa} and therefore handlers never become available, leading to deadlock on destruction. \par
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To circumvent this deadlock, the initial state of \texttt{CacheData} was modified to be safe for deletion. An initialization function was added to \texttt{CacheData}, which is required to be called when the instance is to be used. \par
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\subsubsection{Invalid State on Operation Failure}
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\texttt{CacheData::WaitOnCompletion} first checks for a valid cache pointer and then waits on the handlers becoming valid. To process the handlers, the global atomic pointer is read into a local copy and then set to nullptr using \texttt{std::atomic<T>::exchange}. During evaluation of the handlers completion states, an unsuccessful operation may be found. In this case, the cache memory region remains invalid and may therefore not be used. In this case, both the handlers and the cache pointer will be nullptr. This results in an invalid state, like the one discussed in Section \ref{subsubsec:impl:cdatomicity:initial-invalid-state}. \par
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In this invalid state, progress is not guaranteed by the measures set forth to handle the initial invalidity. The cache is still nullptr and as the handlers have already been set and processed, they will also be nullptr without the chance of them ever becoming valid. \par
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Edge case handling is introduced and the cache pointer is set to the source address, providing validity. \par
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\subsubsection{Locally Invalid State due to Race Condition}
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The guarantee of \texttt{std::atomic<T>::wait} to only wake up when the value has changed \cite{cppreference:atomic-wait} was found to be stronger than the promise of waking up all waiting threads with \texttt{std::atomic<T>::notify\_all} \cite{cppreference:atomic-notify-all}. \par
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As visible in Figure \ref{fig:impl-cachedata-waitoncompletion}, we wait while the handlers-pointer is nullptr, if the cache pointer is invalid. To exemplify we use the following scenario. Both \(T_1\) and \(T_2\) call \texttt{CacheData::WaitOnCompletion}, with \(T_1\) preceding \(T_2\). \(T_1\) exchanges the global handlers-pointer with nullptr, invalidating it. Before \(T_1\) can check the status of the handlers and update the cache pointer, \(T_2\) sees an invalid cache pointer and then waits for the handlers becoming available. \par
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This has again caused a similar state of invalidity as the previous two Sections handled. As the handlers will not become available again due to being cleared by \(T_1\), the second consumer, \(T_2\), will now wait indefinitely. This missed update is commonly referred to as \enquote{ABA-Problem} for which multiple solutions exist. \par
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One could use double-width atomic operations and introduce a counter which would allow resetting the pointer back to null while setting a flag indicating the exchange took place. The handlers-pointer would then be contained in a struct with this flag, allowing exchange with a composite of nullptr and flag-set. Other threads then would then wait on the struct changing from nullptr and flag-unset, allowing them to pass if either the flag is set or the handlers have become non-null. As standard C++ does not yet support the required operations, we chose to avoid the missed update differently. \cite{dwcas-cpp} \par
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The chosen solution for this is to not exchange the handlers-pointer with nullptr but with a second invalid value. We must determine a secondary invalid pointer for use in the exchange. Therefore, we introduce a new attribute, of the same type as the one pointed to by the handlers-pointer, to \texttt{Cache}. The \texttt{Cache} then shares it with each instance of \texttt{CacheData}, where it is then used in \texttt{CacheData::WaitOnCompletion}. \par
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This secondary value allows \(T_2\) to pass the wait, then perform the exchange of handlers itself. \(T_2\) then checks the local copy of the handlers-pointer for validity. The invalid state now includes both nullptr and the secondary invalid pointer chosen. With this, the deadlock is avoided and \(T_2\) will wait for \(T_1\) completing the processing of the handlers. \par
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\section{Accelerator Usage}
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After \ref{sec:design:accel-usage} the implementation of \texttt{Cache} provided leaves it up to the user to choose a caching and copy method policy which is accomplished through submitting function pointers at initialization of the \texttt{Cache}. In \ref{sec:state:setup-and-config} we configured our system to have separate \gls{numa:node}s for accessing \gls{hbm} which are assigned a \gls{numa:node}-ID by adding eight to the \gls{numa:node}s ID of the \gls{numa:node} that physically contains the \gls{hbm}. Therefore, given \gls{numa:node} 3 accesses some datum, the most efficient placement for the copy would be on \gls{numa:node} \(3 + 8 = 11\). As the \texttt{Cache} is intended for multithreaded usage, conserving accelerator resources is important, so that concurrent cache requests complete quickly. To get high per-copy performance while maintaining low usage, the Push-Pull method is selected as described in \ref{subsec:perf:datacopy} for larger copies, while small copies will be handled exclusively by the current node. We introduce this distinction by datum size due to the observations in Section \ref{subsec:perf:submitmethod} that with lower transfer sizes, the effect of submission cost is amplified. \par
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\section{Application to \glsentrylong{qdp}}
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Applying the \texttt{Cache} to \gls{qdp} is straightforward. We adapted the benchmarking code developed by Anna Bartuschka and André Berthold \cite{dimes-prefetching}, calling \texttt{Cache::Access} for both prefetching and cache access. \par
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During performance analysis of the developed \texttt{Cache}, we found that \gls{intel:dml} does not utilize interrupt-based completion signalling (Section \ref{subsubsec:state:completion-signal}), but instead busy-waits on the completion descriptor being updated. As this busy waiting costs CPU cycles, waiting on task completion is out of question, requiring code modifications which we now detail. \par
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We extended \texttt{CacheData} and \texttt{Cache} to provide support for weak waiting. Through a flag configurable in \texttt{Cache}, all instances of \texttt{CacheData} created through \texttt{Cache::Access} will only check once whether the \gls{dsa} has completed processing the operation, and otherwise return without updating the cache-pointer. Calls to \texttt{CacheData::GetDataLocation} can therefore return nullptr, even after waiting. The user is then responsible to access the data through its source location. For latency-critical applications, \texttt{Cache::Access} provides the option for weak access. When set, the function will only return an existing instance of \texttt{CacheData} and therefore does not cause work submission to the \gls{dsa}. \par
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