\documentclass[../thesis.tex]{subfiles}

\input{subfilepreamble}

\begin{document}
\input{subfileprefix}
\chapter{Implementation}%
\label{chp:implementation}
\begin{chapterabstract}
	\noindent This chapter shows the implementation of the \gls{MTASK} system by:
	\begin{itemize}
		\item shows the implementation of the byte code compiler for \gls{MTASK}'s \gls{TOP} language;
		\item gives details of the implementation of \gls{MTASK}'s \gls{TOP} engine that executes the \gls{MTASK} tasks on the microcontroller;
		\item and explains how the integration of \gls{MTASK} tasks and \glspl{SDS} is implemented both on the server and on the device.
	\end{itemize}
\end{chapterabstract}

Microcontrollers usually have flash-based program memory which wears out fairly quick.
For example, the atmega328p in the \gls{ARDUINO} UNO is rated for 10000 write cycles.
While this sounds like a lot, if new tasks are sent to the device every minute or so, a lifetime of not even seven days is guaranteed.
Hence, for dynamic applications, generating code at run-time for interpretation on the device is necessary.
This byte code is then interpreted on MCUs with very little memory and processing power and thus save precious write cycles of the program memory.
precious write cycles of the program memory.

In order to provide the device with the tools to interpret the byte code, it is programmed with a \gls{RTS}, a customisable domain-specific \gls{OS} that takes care of the execution of tasks but also low-level mechanisms such as the communication, multi tasking, and memory management.
Once the device is programmed with the \gls{MTASK} \gls{RTS}, it can continuously receive new tasks.

\subsection{Instruction set}
The instruction set is a fairly standard stack machine instruction set extended with special \gls{TOP} instructions.
\Cref{lst:instruction_type} shows the \gls{CLEAN} type representing the instruction set of which \cref{tbl:instr_task} gives detailed semantics.
Type synonyms are used to provide insight on the arguments of the instructions.
One notable instruction is the \cleaninline{MkTask} instruction, it allocates and initialises a task tree node and pushes a pointer to it on the stack.

\begin{lstClean}[caption={The type housing the instruction set.},label={lst:instruction_type}]
:: ArgWidth    :== UInt8         :: ReturnWidth :== UInt8
:: Depth       :== UInt8         :: Num         :== UInt8
:: SdsId       :== UInt8         :: JumpLabel   =: JL UInt16

//** Datatype housing all instructions
:: BCInstr
	//Return instructions
	//Jumps
	= BCJumpF JumpLabel | BCJump JumpLabel | BCLabel JumpLabel | BCJumpSR ArgWidth JumpLabel
	| BCReturn ReturnWidth ArgWidth | BCTailcall ArgWidth ArgWidth JumpLabel
	//Arguments
	| BCArgs ArgWidth ArgWidth
	//Task node creation and refinement
	| BCMkTask BCTaskType | BCTuneRateMs | BCTuneRateSec
	//Task value ops
	| BCIsStable | BCIsUnstable | BCIsNoValue | BCIsValue
	//Stack ops
	| BCPush String255 | BCPop Num | BCRot Depth Num | BCDup | BCPushPtrs
	//Casting
	| BCItoR | BCItoL | BCRtoI | ...
	// arith
	| BCAddI | BCSubI | ...
	...

//** Datatype housing all task types
:: BCTaskType
	= BCStableNode ArgWidth | ArgWidth
	// Pin io
	| BCReadD | BCWriteD | BCReadA | BCWriteA | BCPinMode
	// Interrupts
	| BCInterrupt
	// Repeat
	| BCRepeat
	// Delay
	| BCDelay | BCDelayUntil //* Only for internal use
	// Parallel
	| BCTAnd | BCTOr
	//Step
	| BCStep ArgWidth JumpLabel
	//Sds ops
	| BCSdsGet SdsId | BCSdsSet SdsId | BCSdsUpd SdsId JumpLabel
	// Rate limiter
	| BCRateLimit
	////Peripherals
	//DHT
	| BCDHTTemp UInt8 | BCDHTHumid UInt8
	...
\end{lstClean}

\subsection{Compiler}
The bytecode compiler interpretation for the \gls{MTASK} language is implemented as a monad stack containing a writer monad and a state monad.
The writer monad is used to generate code snippets locally without having to store them in the monadic values.
The state monad accumulates the code, and stores the stateful data the compiler requires.
\Cref{lst:compiler_state} shows the data type for the state, storing:
function the compiler currently is in;
code of the main expression;
context (see \todo{insert ref to compilation rules step here});
code for the functions;
next fresh label;
a list of all the used \glspl{SDS}, either local \glspl{SDS} containing the initial value (\cleaninline{Left}) or lifted \glspl{SDS} (see \cref{sec:liftsds}) containing a reference to the associated \gls{ITASK} \gls{SDS};
and finally there is a list of peripherals used.

\begin{lstClean}[label={lst:compiler_state},caption={\Gls{MTASK}'s byte code compiler type}]
:: BCInterpret a :== StateT BCState (WriterT [BCInstr] Identity) a
:: BCState =
	{ bcs_infun        :: JumpLabel
	, bcs_mainexpr     :: [BCInstr]
	, bcs_context      :: [BCInstr]
	, bcs_functions    :: Map JumpLabel BCFunction
	, bcs_freshlabel   :: JumpLabel
	, bcs_sdses        :: [Either String255 MTLens]
	, bcs_hardware     :: [BCPeripheral]
	}
:: BCFunction =
	{ bcf_instructions :: [BCInstr]
	, bcf_argwidth     :: UInt8
	, bcf_returnwidth  :: UInt8
	}
\end{lstClean}

Executing the compiler is done by providing an initial state.
After compilation, several post-processing steps are applied to make the code suitable for the microprocessor.
First, in all tail call \cleaninline{BCReturn}'s are replaced by \cleaninline{BCTailCall} to implement tail call elimination.
Furthermore, all byte code is concatenated, resulting in one big program.
Many instructions have commonly used arguments so shorthands are introduced to reduce the program size.
For example, the \cleaninline{BCArg} instruction is often called with argument \qtyrange{0}{2} and can be replaced by the \cleaninline{BCArg0}--\cleaninline{BCArg2} shorthands.
Furthermore, redundant instructions (e.g.\ pop directly after push) are removed as well in order not to burden the code generation with these intricacies.
Finally the labels are resolved to represent actual program addresses instead of freshly generated identifiers.
After the byte code is ready, the lifted \glspl{SDS} are resolved to provide an initial value for them.
The result---byte code, \gls{SDS} specification and perpipheral specifications---are the result of the process, ready to be sent to the device. 

\section{Compilation rules}
This section describes the compilation rules, the translation from abstract syntax to byte code.
The compilation scheme consists of three schemes\slash{}functions.
When something is surrounded by double vertical bars, e.g.\ $\stacksize{a_i}$, it denotes the number of stack cells required to store it.

Some schemes have a \emph{context} $r$ as an argument which contains information about the location of the arguments in scope.
More information is given in the schemes requiring such arguments.

\newcommand{\cschemeE}[2]{\mathcal{E}\llbracket#1\rrbracket~#2}
\newcommand{\cschemeF}[1]{\mathcal{F}\llbracket#1\rrbracket}
\newcommand{\cschemeS}[3]{\mathcal{S}\llbracket#1\rrbracket~#2~#3}
\begin{table}
	\centering
	\begin{tabularx}{\linewidth}{l X}
		\toprule
		Scheme & Description\\
		\midrule
		$\cschemeE{e}{r}$ & Produces the value of expression $e$ given the context $r$ and pushes it on the stack.
			The result can be a basic value or a pointer to a task.\\
		$\cschemeF{e}$ & Generates the bytecode for functions.\\
		$\cschemeS{e}{r}{w} $ & Generates the function for the step continuation given the context $r$ and the width $w$ of the left-hand side task value.\\
		\bottomrule
	\end{tabularx}
\end{table}

\subsection{Expressions}
Almost all expression constructions are compiled using $\mathcal{E}$.
The argument of $\mathcal{E}$ is the context (see \cref{ssec:functions}).
Values are always placed on the stack; tuples and other compound data types are unpacked.
Function calls, function arguments and tasks are also compiled using $\mathcal{E}$ but their compilations is explained later.

\begin{align*}
	\cschemeE{\text{\cleaninline{lit}}~e}{r} & = \text{\cleaninline{BCPush (bytecode e)}};\\
	\cschemeE{e_1\mathbin{\text{\cleaninline{+.}}}e_2}{r} & = \cschemeE{e_1}{r};
			\cschemeE{e_2}{r};
			\text{\cleaninline{BCAdd}};\\
		{} & \text{\emph{Similar for other binary operators}}\\
	\cschemeE{\text{\cleaninline{Not}}~e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCNot}};\\
		{} & \text{\emph{Similar for other unary operators}}\\
	\cschemeE{\text{\cleaninline{If}}~e_1~e_2~e_3}{r} & =
			\cschemeE{e_1}{r};
			\text{\cleaninline{BCJmpF}}\enskip l_{else}; \mathbin{\phantom{=}} \cschemeE{e_2}{r}; \text{\cleaninline{BCJmp}}\enskip l_{endif};\\
		{} & \mathbin{\phantom{=}} \text{\cleaninline{BCLabel}}\enskip l_{else}; \cschemeE{e_3}{r}; \mathbin{\phantom{=}} \text{\cleaninline{BCLabel}}\enskip l_{endif};\\
		{} & \text{\emph{Where $l_{else}$ and $l_{endif}$ are fresh labels}}\\
	\cschemeE{\text{\cleaninline{tupl}}~e_1~e_2}{r} & =
			\cschemeE{e_1}{r};
			\cschemeE{e_2}{r};\\
		{} & \text{\emph{Similar for other unboxed compound data types}}\\
	\cschemeE{\text{\cleaninline{first}}~e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCPop}}\enskip w;\\
		{} & \text{\emph{Where $w$ is the width of the left value and}}\\
		{} & \text{\emph{similar for other unboxed compound data types}}\\
	\cschemeE{\text{\cleaninline{second}}\enskip e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCRot}}\enskip w_1\enskip (w_1+w_2);
			\text{\cleaninline{BCPop}}\enskip w_2;\\
		{} & \text{\emph{Where $w_1$ is the width of the left and, $w_2$ of the right value}}\\
		{} & \text{\emph{similar for other unboxed compound data types}}\\
\end{align*}

Translating $\mathcal{E}$ to \gls{CLEAN} code is very straightforward, it basically means executing the monad.
Almost always, the type of the interpretation is not used, i.e.\ it is a phantom type.
To still have the functions return the correct type, the \cleaninline{tell`}\footnote{\cleaninline{tell` :: [BCInstr] -> BCInterpret a}} helper is used.
This function is similar to the writer monad's \cleaninline{tell} function but is casted to the correct type.
\Cref{lst:imp_arith} shows the implementation for the arithmetic and conditional expressions.
Note that $r$, the context, is not an explicit argument but stored in the state.

\begin{lstClean}[caption={Interpretation implementation for the arithmetic and conditional classes.},label={lst:imp_arith}]
instance expr BCInterpret where
	lit   t   = tell` [BCPush (toByteCode{|*|} t)]
	(+.)  a b = a >>| b >>| tell` [BCAdd]
	...
	If c t e = freshlabel >>= \elselabel->freshlabel >>= \endiflabel->
		c >>| tell` [BCJumpF elselabel] >>|
		t >>| tell` [BCJump endiflabel,BCLabel elselabel] >>|
		e >>| tell` [BCLabel endiflabel]
\end{lstClean}

\subsection{Functions}
Compiling functions occurs in $\mathcal{F}$, which generates bytecode for the complete program by iterating over the functions and ending with the main expression.
When compiling the body of the function, the arguments of the function are added to the context so that the addresses can be determined when referencing arguments.
The main expression is a special case of $\mathcal{F}$ since it neither has arguments nor something to continue.
Therefore, it is just compiled using $\mathcal{E}$.

\begin{align*}
	\cschemeF{main=m} & =
		\cschemeE{m}{[]};\\
	\cschemeF{f~a_0 \ldots a_n = b~\text{\cleaninline{In}}~m} & =
		\text{\cleaninline{BCLabel}}~f; \cschemeE{b}{[\langle f, i\rangle, i\in \{(\Sigma^n_{i=0}\stacksize{a_i})..0\}]};\\
		{} & \mathbin{\phantom{=}} \text{\cleaninline{BCReturn}}~\stacksize{b}~n; \cschemeF{m};\\
\end{align*}

A function call starts by pushing the stack and frame pointer, and making space for the program counter (\cref{lst:funcall_pushptrs}) followed by evaluating the arguments in reverse order (\cref{lst:funcall_args}).
On executing \cleaninline{BCJumpSR}, the program counter is set and the interpreter jumps to the function (\cref{lst:funcall_jumpsr}).
When the function returns, the return value overwrites the old pointers and the arguments.
This occurs right after a \cleaninline{BCReturn} (\cref{lst:funcall_ret}).
Putting the arguments on top of pointers and not reserving space for the return value uses little space and facilitates tail call optimization.

\begin{figure}
	\begin{subfigure}{.24\linewidth}
		\centering
		\includestandalone{memory1}
		\caption{\cleaninline{BCPushPtrs}}\label{lst:funcall_pushptrs}
	\end{subfigure}
	\begin{subfigure}{.24\linewidth}
		\centering
		\includestandalone{memory2}
		\caption{Arguments}\label{lst:funcall_args}
	\end{subfigure}
	\begin{subfigure}{.24\linewidth}
		\centering
		\includestandalone{memory3}
		\caption{\cleaninline{BCJumpSR}}\label{lst:funcall_jumpsr}
	\end{subfigure}
	\begin{subfigure}{.24\linewidth}
		\centering
		\includestandalone{memory4}
		\caption{\cleaninline{BCReturn}}\label{lst:funcall_ret}
	\end{subfigure}
	\caption{The stack layout during function calls.}%
\end{figure}

Calling a function and referencing function arguments are an extension to $\mathcal{E}$ as shown below.
Arguments may be at different places on the stack at different times (see \cref{ssec:step}) and therefore the exact location always has to be determined from the context using \cleaninline{findarg}\footnote{\cleaninline{findarg [l`:r] l = if (l == l`) 0 (1 + findarg r l)}}.
Compiling argument $a_{f^i}$, the $i$th argument in function $f$, consists of traversing all positions in the current context.
Arguments wider than one stack cell are fetched in reverse to preserve the order.

\begin{align*}
	\cschemeE{f(a_0, \ldots, a_n)}{r} & =
		\text{\cleaninline{BCPushPtrs}}; \cschemeE{a_n}{r}; \cschemeE{a_{\ldots}}{r}; \cschemeE{a_0}{r}; \text{\cleaninline{BCJumpSR}}~n~f;\\
	\cschemeE{a_{f^i}}{r} & =
		\text{\cleaninline{BCArg}~findarg}(r, f, i)~\text{for all}~i\in\{w\ldots v\};\\
		{} & v = \Sigma^{i-1}_{j=0}\stacksize{a_{f^j}}~\text{ and }~ w = v + \stacksize{a_{f^i}}\\
\end{align*}

Translating the compilation schemes for functions to Clean is not as straightforward as other schemes due to the nature of shallow embedding.\todo{deze \P{} moet ge\-\"up\-da\-ted worden}
The \cleaninline{fun} class has a single function with a single argument.
This argument is a Clean function that---when given a callable Clean function representing the mTask function---will produce \cleaninline{main} and a callable function.
To compile this, the argument must be called with a function representing a function call in mTask.
\Cref{lst:fun_imp} shows the implementation for this as Clean code.
To uniquely identify the function, a fresh label is generated.
The function is then called with the \cleaninline{callFunction} helper function that generates the instructions that correspond to calling the function.
That is, it pushes the pointers, compiles the arguments, and writes the \cleaninline{JumpSR} instruction.
The resulting structure (\cleaninline{g In m}) contains a function representing the mTask function (\cleaninline{g}) and the \cleaninline{main} structure to continue with.
To get the actual function, \cleaninline{g} must be called with representations for the argument, i.e.\ using \cleaninline{findarg} for all arguments.
The arguments are added to the context and \cleaninline{liftFunction} is called with the label, the argument width and the compiler.
This function executes the compiler, decorates the instructions with a label and places them in the function dictionary together with the metadata such as the argument width.
After lifting the function, the context is cleared again and compilation continues with the rest of the program.

\begin{lstClean}[label={lst:fun_imp},caption={The backend implementation for functions.}]
instance fun (BCInterpret a) BCInterpret | type a where
	fun def = {main=freshlabel >>= \funlabel->
		let (g In m) = def \a->callFunction funlabel (toByteWidth a) [a]
		    argwidth = toByteWidth (argOf g)
		in  addToCtx funlabel zero argwidth
		>>| infun funlabel
			(liftFunction funlabel argwidth
				(g (retrieveArgs funlabel zero argwidth)
				) ?None)
		>>| clearCtx >>| m.main
		}

argOf :: ((m a) -> b) a -> UInt8 | toByteWidth a
callFunction :: JumpLabel UInt8 [BCInterpret b] -> BCInterpret c | ...
liftFunction :: JumpLabel UInt8 (BCInterpret a) (?UInt8) -> BCInterpret ()
\end{lstClean}

\subsection{Tasks}
Task trees are created with the \cleaninline{BCMkTask} instruction that allocates a node and pushes it to the stack.
It pops arguments from the stack according to the given task type.
The following extension of $\mathcal{E}$ shows this compilation scheme (except for the step combinator, explained in \cref{ssec:step}).

\begin{align*}
	\cschemeE{\text{\cleaninline{rtrn}}~e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCMkTask BCStable}}_{\stacksize{e}};\\
	\cschemeE{\text{\cleaninline{unstable}}~e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCMkTask BCUnstable}}_{\stacksize{e}};\\
	\cschemeE{\text{\cleaninline{readA}}~e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCMkTask BCReadA}};\\
	\cschemeE{\text{\cleaninline{writeA}}~e_1~e_2}{r} & =
			\cschemeE{e_1}{r};
			\cschemeE{e_2}{r};
			\text{\cleaninline{BCMkTask BCWriteA}};\\
	\cschemeE{\text{\cleaninline{readD}}~e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCMkTask BCReadD}};\\
	\cschemeE{\text{\cleaninline{writeD}}~e_1~e_2}{r} & =
			\cschemeE{e_1}{r};
			\cschemeE{e_2}{r};
			\text{\cleaninline{BCMkTask BCWriteD}};\\
	\cschemeE{\text{\cleaninline{delay}}~e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCMkTask BCDelay}};\\
	\cschemeE{\text{\cleaninline{rpeat}}~e}{r} & =
			\cschemeE{e}{r};
			\text{\cleaninline{BCMkTask BCRepeat}};\\
	\cschemeE{e_1\text{\cleaninline{.\|\|.}}e_2}{r} & =
			\cschemeE{e_1}{r};
			\cschemeE{e_2}{r};
			\text{\cleaninline{BCMkTask BCOr}};\\
	\cschemeE{e_1\text{\cleaninline{.&&.}}e_2}{r} & =
			\cschemeE{e_1}{r};
			\cschemeE{e_2}{r};
			\text{\cleaninline{BCMkTask BCAnd}};\\
\end{align*}

This simply translates to Clean code by writing the correct \cleaninline{BCMkTask} instruction as exemplified in \cref{lst:imp_ret}.

\begin{lstClean}[caption={The backend implementation for \cleaninline{rtrn}.},label={lst:imp_ret}]
instance rtrn BCInterpret
where
	rtrn m = m >>| tell` [BCMkTask (bcstable m)]
\end{lstClean}

\subsection{Step combinator}\label{ssec:step}
The \cleaninline{step} construct is a special type of task because the task value of the left-hand side may change over time.
Therefore, the continuation tasks on the right-hand side are \emph{observing} this task value and acting upon it.
In the compilation scheme, all continuations are first converted to a single function that has two arguments: the stability of the task and its value.
This function either returns a pointer to a task tree or fails (denoted by $\bot$).
It is special because in the generated function, the task value of a task can actually be inspected.
Furthermore, it is a lazy node in the task tree: the right-hand side may yield a new task tree after several rewrite steps (i.e.\ it is allowed to create infinite task trees using step combinators).
The function is generated using the $\mathcal{S}$ scheme that requires two arguments: the context $r$ and the width of the left-hand side so that it can determine the position of the stability which is added as an argument to the function.
The resulting function is basically a list of if-then-else constructions to check all predicates one by one.
Some optimization is possible here but has currently not been implemented.

\begin{align*}
	\cschemeE{t_1\text{\cleaninline{>>*.}}t_2}{r} & =
		\cschemeE{a_{f^i}}{r}, \langle f, i\rangle\in r;
		\text{\cleaninline{BCMkTask}}~\text{\cleaninline{BCStable}}_{\stacksize{r}}; \cschemeE{t_1}{r};\\
	{} & \mathbin{\phantom{=}} \text{\cleaninline{BCMkTask}}~\text{\cleaninline{BCAnd}}; \text{\cleaninline{BCMkTask}}~(\text{\cleaninline{BCStep}}~(\cschemeS{t_2}{(r + [\langle l_s, i\rangle])}{\stacksize{t_1}}));\\
\end{align*}

\begin{align*}
	\cschemeS{[]}{r}{w} & =
		\text{\cleaninline{BCPush}}~\bot;\\
	\cschemeS{\text{\cleaninline{IfValue}}~f~t:cs}{r}{w} & =
		\text{\cleaninline{BCArg}} (\stacksize{r} + w);
		\text{\cleaninline{BCIsNoValue}};\\
	{} & \mathbin{\phantom{=}} \cschemeE{f}{r};
		\text{\cleaninline{BCAnd}};\\
	{} & \mathbin{\phantom{=}} \text{\cleaninline{BCJmpF}}~l_1;\\
	{} & \mathbin{\phantom{=}} \cschemeE{t}{r};
		\text{\cleaninline{BCJmp}}~l_2;\\
	{} & \mathbin{\phantom{=}} \text{\cleaninline{BCLabel}}~l_1;
		\cschemeS{cs}{r}{w};\\
	{} & \mathbin{\phantom{=}} \text{\cleaninline{BCLabel}}~l_2;\\
	{} & \text{\emph{Where $l_1$ and $l_2$ are fresh labels}}\\
	{} & \text{\emph{Similar for \cleaninline{IfStable} and \cleaninline{IfUnstable}}}\\
\end{align*}

First the context is evaluated.
The context contains arguments from functions and steps that need to be preserved after rewriting.
The evaluated context is combined with the left-hand side task value by means of a \cleaninline{.&&.} combinator to store it in the task tree so that it is available after a rewrite.
This means that the task tree is be transformed as follows:

\begin{lstClean}
t1 >>= \v1->t2 >>= \v2->t3 >>= ...
//is transformed to
t1 >>= \v1->rtrn v1 .&&. t2 >>= \v2->rtrn (v1, v2) .&&. t3 >>= ...
\end{lstClean}

The translation to \gls{CLEAN} is given in \cref{lst:imp_seq}.

\begin{lstClean}[caption={Backend implementation for the step class.},label={lst:imp_seq}]
instance step BCInterpret where
	(>>*.) lhs cont
		//Fetch a fresh label and fetch the context
		=   freshlabel >>= \funlab->gets (\s->s.bcs_context)
		//Generate code for lhs
		>>= \ctx->lhs
		//Possibly add the context
		>>| tell` (if (ctx =: []) []
				//The context is just the arguments up till now in reverse
				(  [BCArg (UInt8 i)\\i<-reverse (indexList ctx)]
				++ map BCMkTask (bcstable (UInt8 (length ctx)))
				++ [BCMkTask BCTAnd]
				))
		//Increase the context
		>>| addToCtx funlab zero lhswidth
		//Lift the step function
		>>| liftFunction funlab
				//Width of the arguments is the width of the lhs plus the
				//stability plus the context
				(one + lhswidth + (UInt8 (length ctx)))
				//Body     label  ctx width            continuations
				(contfun funlab (UInt8 (length ctx)))
				//Return width (always 1, a task pointer)
				(Just one)
		>>| modify (\s->{s & bcs_context=ctx})
		>>| tell` [BCMkTask (instr rhswidth funlab)]

toContFun :: JumpLabel UInt8 -> BCInterpret a
toContFun steplabel contextwidth
	= foldr tcf (tell` [BCPush fail]) cont
where
	tcf (IfStable f t)
		= If ((stability >>| tell` [BCIsStable]) &. f val)
			(t val >>| tell` [])
	...
	stability = tell` [BCArg (lhswidth + contextwidth)]
	val = retrieveArgs steplabel zero lhswidth
\end{lstClean}

\subsection{\texorpdfstring{\Glspl{SDS}}{Shared data sources}}
The compilation scheme for \gls{SDS} definitions is a trivial extension to $\mathcal{F}$ since there is no code generated as seen below.

\begin{align*}
	\cschemeF{\text{\cleaninline{sds}}~x=i~\text{\cleaninline{In}}~m} & =
		\cschemeF{m};\\
\end{align*}

The \gls{SDS} access tasks have a compilation scheme similar to other tasks (see \cref{ssec:scheme_tasks}).
The \cleaninline{getSds} task just pushes a task tree node with the \gls{SDS} identifier embedded.
The \cleaninline{setSds} task evaluates the value, lifts that value to a task tree node and creates an \gls{SDS} set node.

\begin{align*}
	\cschemeE{\text{\cleaninline{getSds}}~s}{r} & =
		\text{\cleaninline{BCMkTask}} (\text{\cleaninline{BCSdsGet}} s);\\
	\cschemeE{\text{\cleaninline{setSds}}~s~e}{r} & =
		\cschemeE{e}{r};
		\text{\cleaninline{BCMkTask BCStable}}_{\stacksize{e}};\\
	{} & \mathbin{\phantom{=}} \text{\cleaninline{BCMkTask}} (\text{\cleaninline{BCSdsSet}} s);\\
\end{align*}

While there is no code generated in the definition, the byte code compiler is storing the \gls{SDS} data in the \cleaninline{bcs_sdses} field in the compilation state.
The \glspl{SDS} are typed as functions in the host language so an argument for this function must be created that represents the \gls{SDS} on evaluation.
For this, an \cleaninline{BCInterpret} is created that emits this identifier.
When passing it to the function, the initial value of the \gls{SDS} is returned.
This initial value is stored as a byte code encoded value in the state and the compiler continues with the rest of the program.

Compiling \cleaninline{getSds} is a matter of executing the \cleaninline{BCInterpret} representing the \gls{SDS}, which yields the identifier that can be embedded in the instruction.
Setting the \gls{SDS} is similar: the identifier is retrieved and the value is written to put in a task tree so that the resulting task can remember the value it has written.
Lifted SDSs are compiled in a very similar way.\todo{deze \P{} moet naar integration?}
The only difference is that there is no initial value but an iTasks SDS when executing the Clean function.
A lens on this SDS converting \cleaninline{a} from the \cleaninline{Shared a} to a \cleaninline{String255}---a bytecode encoded version---is stored in the state.
The encoding and decoding itself is unsafe when used directly but the type system of the language and the abstractions make it safe.
Upon sending the mTask task to the device, the initial values of the lifted SDSs are fetched to complete the SDS specification.

% VimTeX: SynIgnore on
\begin{lstClean}[caption={Backend implementation for the SDS classes.},label={lst:comp_sds}]
:: Sds a = Sds Int
instance sds BCInterpret where
	sds def = {main = freshsds >>= \sdsi->
			let sds = modify (\s->{s & bcs_sdses=put sdsi
						(Left (toByteCode t)) s.bcs_sdses})
					>>| pure (Sds sdsi)
			    (t In e) = def sds
			in e.main}
	getSds f   = f >>= \(Sds i)-> tell` [BCMkTask (BCSdsGet (fromInt i))]
	setSds f v = f >>= \(Sds i)->v >>| tell`
		(  map BCMkTask (bcstable (byteWidth v))
		++ [BCMkTask (BCSdsSet (fromInt i))])
\end{lstClean}
% VimTeX: SynIgnore off

\section{\texorpdfstring{\Glsxtrlong{RTS}}{Run time system}}

The \gls{RTS} is designed to run on systems with as little as \qty{2}{\kibi\byte} of \gls{RAM}.
Aggressive memory management is therefore vital.
Not all firmwares for microprocessors support heaps and---when they do---allocation often leaves holes when not used in a \emph{last in first} out strategy.
Therefore the \gls{RTS} uses a chunk of memory in the global data segment with its own memory manager tailored to the needs of \gls{MTASK}.
The size of this block can be changed in the configuration of the \gls{RTS} if necessary.
On an \gls{ARDUINO} UNO ---equipped with \qty{2}{\kibi\byte} of \gls{RAM}--- this size is about \qty{1500}{\byte}.

In memory, task data grows from the bottom up and an interpreter stack is located directly on top of it growing in the same direction.
As a consequence, the stack moves when a new task is received.
This never happens within execution because communication is always processed before execution.
Values in the interpreter are always stored on the stack, even tuples.
Task trees grow from the top down as in a heap.
This approach allows for flexible ratios, i.e.\ many tasks and small trees or few tasks and big trees.

The event loop of the \gls{RTS} is executed repeatedly and consists of three distinct phases.
\todo{plaa\-tje van me\-mo\-ry hier}
\todo{pseu\-do\-code hier van de ex\-e\-cu\-tie}

%TODO evt subsubsections verwijderen
\subsection{Communication}
In the first phase, the communication channels are processed.
The messages announcing \gls{SDS} updates are applied immediately, the initialization of new tasks is delayed until the next phase.

\subsection{Execution}
The second phase consists of executing tasks.
The \gls{RTS} executes all tasks in a round robin fashion.
If a task is not initialized yet, the bytecode of the main function is interpreted to produce the initial task tree.
The rewriting engine uses the interpreter when needed, e.g.\ to calculate the step continuations.
The rewriter and the interpreter use the same stack to store intermediate values.
Rewriting steps are small so that interleaving results in seemingly parallel execution.
In this phase new task tree nodes may be allocated.
Both rewriting and initialization are atomic operations in the sense that no processing on SDSs is done other than SDS operations from the task itself.
The host is notified if a task value is changed after a rewrite step.

\subsection{Memory management}
The third and final phase is memory management.
Stable tasks, and unreachable task tree nodes are removed.
If a task is to be removed, tasks with higher memory addresses are moved down.
For task trees---stored in the heap---the \gls{RTS} already marks tasks and task trees as trash during rewriting so the heap can be compacted in a single pass.
This is possible because there is no sharing or cycles in task trees and nodes contain pointers pointers to their parent.

\input{subfilepostamble}
\end{document}
