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Dump abstract machine text.

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Daniel Hillerström 5 years ago
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  1. 25
      macros.tex
  2. 687
      thesis.tex

25
macros.tex
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@ -84,6 +84,7 @@
\newcommand{\reducesto}[0]{\ensuremath{\leadsto}}
\newcommand{\areducesto}{\ensuremath{\reducesto_{\textrm{a}}}}
\newcommand{\stepsto}[0]{\ensuremath{\longrightarrow}}
\newcommand{\Stepsto}{\Longrightarrow}
\newcommand{\EC}{\ensuremath{\mathcal{E}}}
\newcommand{\BL}{\ensuremath{\mathsf{BL}}}
@ -283,4 +284,26 @@
%
\renewcommand{\hid}{V_\mathrm{forward}}
\newcommand{\kid}{V_\mathrm{id}}
\newcommand{\rid}{V_\mathrm{ret}}
\newcommand{\rid}{V_\mathrm{ret}}
% Examples
\newcommand{\Pipe}{\dec{pipe}}
\newcommand{\Copipe}{\dec{copipe}}
\newcommand{\Ones}{\dec{ones}}
\newcommand{\Yield}{\dec{Yield}}
\newcommand{\Await}{\dec{Await}}
\newcommand{\AddTwo}{\ensuremath{\dec{add}_2}}
% Abstract machine
\newcommand{\cek}[1]{\ensuremath{\langle #1 \rangle}}
% Environments
\newcommand{\env}{\ensuremath{\gamma}}
% restrict an environment
\newcommand{\res}{\backslash}
% abstract machine translations
\newcommand{\val}[2]{\llbracket #1 \rrbracket #2}
\newcommand{\inv}[1]{\llparenthesis #1 \rrparenthesis}
% configurations
\newcommand{\conf}{\mathcal{C}}

687
thesis.tex
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@ -1077,8 +1077,11 @@ that the binder $x : A$.
\EC[M] &\reducesto& \EC[N], \hfill\quad \text{if } M \reducesto N \\
\end{reductions}
\begin{syntax}
\slab{Evaluation contexts} & \mathcal{E} \in \EvalCat &::=& [\,] \mid \Let \; x \revto \mathcal{E} \; \In \; N
\slab{Evaluation contexts} & \mathcal{E} \in \EvalCat &::=& [~] \mid \Let \; x \revto \mathcal{E} \; \In \; N
\end{syntax}
%
\dhil{Describe evaluation contexts as functions: decompose and plug.}
%
%%\[
% Evaluation context lift
%% \inferrule*[Lab=\semlab{Lift}]
@ -1090,13 +1093,14 @@ that the binder $x : A$.
\label{fig:base-language-small-step}
\end{figure}
%
In this section we present the dynamic semantics of \BCalc{}. We opt
to use a \citet{Felleisen87}-style contextual small-step semantics,
since in conjunction with fine-grain call-by-value (FGCBV), it yields
a considerably simpler semantics than the traditional structural
operational semantics (SOS)~\cite{Plotkin04a}, because only the rule
for let bindings admits a continuation wheres in ordinary
call-by-value SOS each congruence rule admits a continuation.
In this section I will present the dynamic semantics of \BCalc{}. I
have chosen opt to use a \citet{Felleisen87}-style contextual
small-step semantics, since in conjunction with fine-grain
call-by-value (FGCBV), it yields a considerably simpler semantics than
the traditional structural operational semantics
(SOS)~\cite{Plotkin04a}, because only the rule for let bindings admits
a continuation wheres in ordinary call-by-value SOS each congruence
rule admits a continuation.
%
The simpler semantics comes at the expense of a more verbose syntax,
which is not a concern as one can readily convert between fine-grain
@ -1247,7 +1251,7 @@ computation term, $M$, in some larger context, $\EC$, and reduce it in
the said context to another computation $N$ if $M$ reduces outside out
the context to that particular $N$. In our formalism, we call this
rule \semlab{Lift}. Evaluation contexts are generated from the empty
context $[~]$ and let expressions $\Let\;x \revto \EC \;\In\;N$.
context ($[~]$) and let expressions ($\Let\;x \revto \EC \;\In\;N$).
The choices of using fine-grain call-by-value and evaluation contexts
may seem odd, if not arbitrary at this point; the reader may wonder
@ -1298,8 +1302,8 @@ contexts.
By structural induction on $M$.
\begin{description}
\item[Base step] $M = N$ where $N$ is either $\Return\;V$,
$\Absurd^C\;V$, $V\,W$, or $V\,T$. In either case take
$\EC = [\,]$ such that $M = \EC[N]$.
$\Absurd^C\;V$, $V\,W$, or $V\,T$. In each case take $\EC = [\,]$
such that $M = \EC[N]$.
\item[Inductive step]
%
There are several cases to consider. In each case we must find an
@ -1313,10 +1317,8 @@ contexts.
\item[Case] $M = \Let\;x \revto M' \;\In\;N$: By the induction
hypothesis it follows that $M'$ is either stuck or it
decomposes (uniquely) into an evaluation context $\EC'$ and a
redex $N'$. If $M$ is stuck, then take
$\EC = \Let\;x \revto [\,] \;\In\;N$ such that $M =
\EC[M']$. Otherwise take $\EC = \Let\;x \revto \EC'\;\In\;N$
such that $M = \EC[N']$.
redex $N'$. If $M'$ is stuck, then so is $M$. Otherwise take
$\EC = \Let\;x \revto \EC'\;\In\;N$ such that $M = \EC[N']$.
\end{itemize}
\end{description}
\end{proof}
@ -1333,23 +1335,23 @@ realisable function in \BCalc{} is effect-free and total.
\end{definition}
%
\begin{theorem}[Progress]\label{thm:base-language-progress}
Suppose $\typ{}{M : C}$, then there exists $\typ{}{N : C}$, such
that $M \reducesto^\ast N$ and $N$ is normal.
Suppose $\typ{}{M : C}$, then $M$ is normal or there exists
$\typ{}{N : C}$ such that $M \reducesto^\ast N$.
\end{theorem}
%
\begin{proof}
Proof by induction on typing derivations.
\end{proof}
%
\begin{corollary}
Every closed computation term in \BCalc{} is terminating.
\end{corollary}
% \begin{corollary}
% Every closed computation term in \BCalc{} is terminating.
% \end{corollary}
%
The calculus also satisfies the \emph{preservation} property,
which states that if some computation $M$ is well-typed and reduces to
some other computation $M'$, then $M'$ is also well-typed.
The calculus also satisfies the \emph{subject reduction} property,
which states that if some computation $M$ is well typed and reduces to
some other computation $M'$, then $M'$ is also well typed.
%
\begin{theorem}[Preservation]\label{thm:base-language-preservation}
\begin{theorem}[Subject reduction]\label{thm:base-language-preservation}
Suppose $\typ{\Gamma}{M : C}$ and $M \reducesto M'$, then
$\typ{\Gamma}{M' : C}$.
\end{theorem}
@ -1383,8 +1385,7 @@ abstraction form for recursive functions.
\end{syntax}
%
The $\Rec$ construct binds the function name $f$ and its argument $x$
in the body $M$. Typing of recursive functions is standard and
entirely straightforward.
in the body $M$. Typing of recursive functions is standard.
%
\begin{mathpar}
\inferrule*[Lab=\tylab{Rec}]
@ -1392,7 +1393,7 @@ entirely straightforward.
{\typ{\Delta;\Gamma}{(\Rec \; f^{A \to C} \, x . M) : A \to C}}
\end{mathpar}
%
The reduction semantics are similarly simple.
The reduction semantics are also standard.
%
\begin{reductions}
\semlab{Rec} &
@ -1407,9 +1408,8 @@ progress theorem as some programs may now diverge.
%
\begin{theorem}[Progress]
\label{thm:base-rec-language-progress}
Suppose $\typ{}{M : C}$, then there exists $\typ{}{N : C}$, such
that $M \reducesto^\ast N$ either diverges or $N\not\reducesto$ and
$N$ is normal.
Suppose $\typ{}{M : C}$, then $M$ is normal or there exists
$\typ{}{N : C}$ such that $M \reducesto^\ast N$.
\end{theorem}
%
\begin{proof}
@ -1420,7 +1420,7 @@ progress theorem as some programs may now diverge.
\label{sec:tracking-div}
%
With the $\Rec$-operator in \BCalcRec{} we can now implement the
customary factorial function.
factorial function.
%
\[
\bl
@ -1451,7 +1451,7 @@ Section~\ref{sec:notions-of-purity} we shall make clear the notion of
purity that we have mind, however, first let us briefly illustrate how
we might utilise the effect system to track divergence.
The key to track divergence is to modify the \tylab{Rec} to inject
The key to tracking divergence is to modify the \tylab{Rec} to inject
some primitive operation into the effect row.
%
\begin{mathpar}
@ -1542,11 +1542,11 @@ The term `pure' is heavily overloaded in the programming literature.
\label{sec:row-polymorphism}
\dhil{A discussion of alternative row systems}
\section{Type and effect inference}
\dhil{While I would like to detail the type and effect inference, it
may not be worth the effort. The reason I would like to do this goes
back to 2016 when Richard Eisenberg asked me about how we do effect
inference in Links.}
% \section{Type and effect inference}
% \dhil{While I would like to detail the type and effect inference, it
% may not be worth the effort. The reason I would like to do this goes
% back to 2016 when Richard Eisenberg asked me about how we do effect
% inference in Links.}
\chapter{Unary handlers}
\label{ch:unary-handlers}
@ -3545,6 +3545,619 @@ then in the image $\cps{M}\,k \reducesto^\ast k\,\cps{V}$.
\paragraph{Partial evaluation}
\chapter{Abstract machine semantics}
\dhil{The text is this chapter needs to be reworked}
In this chapter we develop an abstract machine that supports deep and
shallow handlers \emph{simultaneously}, using the generalised
continuation structure we identified in the previous section for the
CPS translation. We also build upon prior work~\citep{HillerstromL16}
that developed an abstract machine for deep handlers by generalising
the continuation structure of a CEK machine (Control, Environment,
Kontinuation)~\citep{FelleisenF86}.
%
% \citet{HillerstromL16} sketched an adaptation for shallow handlers. It
% turns out that this adaptation had a subtle flaw, similar to the flaw
% in the sketched implementation of a CPS translation for shallow
% handlers given by \citet{HillerstromLAS17}. We fix the flaw here with
% a full development of shallow handlers along with a statement of the
% correctness property.
\section{Syntax and semantics}
The abstract machine syntax is given in
Figure~\ref{fig:abstract-machine-syntax}.
% A machine continuation is a list of handler frames. A handler frame is
% a pair of a \emph{handler closure} (handler definition) and a
% \emph{pure continuation} (a sequence of let bindings), analogous to
% the structured frames used in the CPS translation in
% \Sec\ref{sec:higher-order-uncurried-cps}.
% %
% Handling an operation amounts to searching through the continuation
% for a matching handler.
% %
% The resumption is constructed during the search by reifying each
% handler frame. As in the CPS translation, the resumption is assembled
% in one of two ways depending on whether the matching handler is deep
% or shallow.
% %
% For a deep handler, the current handler closure is included, and a deep
% resumption is a reified continuation.
% %
% An invocation of a deep resumption amounts to concatenating it with
% the current machine continuation.
% %
% For a shallow handler, the current handler closure must be discarded
% leaving behind a dangling pure continuation, and a shallow resumption
% is a pair of this pure continuation and the remaining reified
% continuation.
% %
% (By contrast, the prior flawed adaptation prematurely precomposed the
% pure continuation with the outer handler in the current resumption.)
% %
% An invocation of a shallow resumption again amounts to concatenating
% it with the current machine continuation, but taking care to
% concatenate the dangling pure continuation with that of the next
% frame.
%
\begin{figure}[t]
\flushleft
\begin{syntax}
\slab{Configurations} & \conf &::= & \cek{M \mid \env \mid \shk \circ \shk'} \\
\slab{Value environments} &\env &::= & \emptyset \mid \env[x \mapsto v] \\
\slab{Values} &v, w &::= & (\env, \lambda x^A . M) \mid (\env, \Lambda \alpha^K . M) \\
& &\mid& \Record{} \mid \Record{\ell = v; w} \mid (\ell\, v)^R \mid \shk^A \mid (\shk, \slk)^A \medskip\\
% \end{syntax}
% \begin{displaymath}
% \ba{@{}l@{~~}r@{~}c@{~}l@{\quad}l@{~~}r@{~}c@{~}l@{}}
% \slab{Continuations} &\shk &::= & \nil \mid \shf \cons \shk & \slab{Continuation frames} &\shf &::= & (\slk, \chi) \\
% & & & & \slab{Handler closures} &\chi &::= & (\env, H)^\depth \smallskip \\
% \slab{Pure continuations} &\slk &::= & \nil \mid \slf \cons \slk & \slab{Pure continuation frames} &\slf &::= & (\env, x, N) \\
% \ea
% \end{displaymath}
% \begin{syntax}
\slab{Continuations} &\shk &::= & \nil \mid \shf \cons \shk \\
\slab{Continuation frames} &\shf &::= & (\slk, \chi) \\
\slab{Pure continuations} &\slk &::= & \nil \mid \slf \cons \slk \\
\slab{Pure continuation frames} &\slf &::= & (\env, x, N) \\
\slab{Handler closures} &\chi &::= & (\env, H) \mid (\env, H)^\dagger \medskip \\
\end{syntax}
\caption{Abstract machine syntax.}
\label{fig:abstract-machine-syntax}
\end{figure}
%% A CEK machine~\citep{FelleisenF86} operates on configurations, which
%% are (Control, Environment, Continuation) triples.
%% %
% Our machine, like Hillerström and Lindley's, generalises the usual
% notion of continuation to accommodate handlers.
%
%
\begin{figure}[p]
\dhil{Fix figure formatting}
\begin{minipage}{0.90\textheight}%
%% Identity continuation
%% \[
%% \shk_0 = [(\nil, (\emptyset, \{\Return\;x \mapsto x\}))]
%% \]
\textbf{Transition function}
\begin{displaymath}
\bl
\ba{@{}l@{~}r@{~}c@{~}l@{\quad}l@{}}
\mlab{Init} & \multicolumn{4}{@{}c@{}}{M \stepsto \cek{M \mid \emptyset \mid [(\nil, (\emptyset, \{\Return\;x \mapsto \Return\;x\}))]}} \\[1ex]
% App
\mlab{AppClosure} & \cek{ V\;W \mid \env \mid \shk}
&\stepsto& \cek{ M \mid \env'[x \mapsto \val{W}{\env}] \mid \shk},
&\text{if }\val{V}{\env} = (\env', \lambda x^A.M) \\
\mlab{AppRec} & \cek{ V\;W \mid \env \mid \shk}
&\stepsto& \cek{ M \mid \env'[g \mapsto (\env', \Rec\,g^{A \to C}\,x.M), x \mapsto \val{W}{\env}] \mid \shk},
&\text{if }\val{V}{\env} = (\env', \Rec\,g^{A \to C}\,x.M) \\
% App - continuation
\mlab{AppCont} & \cek{ V\;W \mid \env \mid \shk}
&\stepsto& \cek{ \Return \; W \mid \env \mid \shk' \concat \shk},
&\text{if }\val{V}{\env} = (\shk')^A \\
\mlab{AppCont^\dagger} & \cek{ V\,W \mid \env \mid (\slk, \chi) \cons \shk}
&\stepsto&
\cek{\Return\; W \mid \env \mid \shk' \concat ((\slk' \concat \slk, \chi) \cons \shk)},
&\text{if } \val{V}{\env} = (\shk', \slk')^A \\
% TyApp
\mlab{AppType} & \cek{ V\,T \mid \env \mid \shk}
&\stepsto& \cek{ M[T/\alpha] \mid \env' \mid \shk},
&\text{if }\val{V}{\env} = (\env', \Lambda \alpha^K . \, M) \\[1ex]
\ea \\
\ba{@{}l@{}r@{~}c@{~}l@{\quad}l@{}}
\mlab{Split} & \cek{ \Let \; \Record{\ell = x;y} = V \; \In \; N \mid \env \mid \shk}
&\stepsto& \cek{ N \mid \env[x \mapsto v, y \mapsto w] \mid \shk},
& \text {if }\val{V}{\env} = \Record{\ell=v; w} \\
% Case
\mlab{Case} & \cek{ \Case\; V\, \{ \ell~x \mapsto M; y \mapsto N\} \mid \env \mid \shk}
&\stepsto& \left\{\ba{@{}l@{}}
\cek{ M \mid \env[x \mapsto v] \mid \shk}, \\
\cek{ N \mid \env[y \mapsto \ell'\, v] \mid \shk}, \\
\ea \right.
&
\ba{@{}l@{}}
\text{if }\val{V}{\env} = \ell\, v \\
\text{if }\val{V}{\env} = \ell'\, v \text{ and } \ell \neq \ell' \\
\ea \\[2ex]
% Let - eval M
\mlab{Let} & \cek{ \Let \; x \revto M \; \In \; N \mid \env \mid (\slk, \chi) \cons \shk}
&\stepsto& \cek{ M \mid \env \mid ((\env,x,N) \cons \slk, \chi) \cons \shk} \\
% Handle
\mlab{Handle} & \cek{ \Handle^\depth \, M \; \With \; H \mid \env \mid \shk}
&\stepsto& \cek{ M \mid \env \mid (\nil, (\env, H)^\depth) \cons \shk} \\[1ex]
% Return - let binding
\mlab{RetCont} &\cek{ \Return \; V \mid \env \mid ((\env',x,N) \cons \slk, \chi) \cons \shk}
&\stepsto& \cek{ N \mid \env'[x \mapsto \val{V}{\env}] \mid (\slk, \chi) \cons \shk} \\
% Return - handler
\mlab{RetHandler} & \cek{ \Return \; V \mid \env \mid (\nil, (\env',H)) \cons \shk}
&\stepsto& \cek{ M \mid \env'[x \mapsto \val{V}{\env}] \mid \shk},
& \text{if } \hret = \{\Return\; x \mapsto M\} \\
\mlab{RetTop} & \cek{\Return\;V \mid \env \mid \nil} &\stepsto& \val{V}{\env} \\[1ex]
% Deep
\mlab{Do} & \cek{ (\Do \; \ell \; V)^E \mid \env \mid ((\slk, (\env', H)) \cons \shk) \circ \shk'}
&\stepsto& \multicolumn{2}{@{}l@{}}{\cek{M \mid \env'[x \mapsto \val{V}{\env},
r \mapsto (\shk' \concat [(\slk, (\env', H))])^B] \mid \shk},} \\
&&& \multicolumn{2}{@{}r@{}}{\text{if } \ell : A \to B \in E \text{ and } \hell = \{\ell\; x \; r \mapsto M\}} \\
% Shallow
\mlab{Do^\dagger} & \cek{ (\Do \; \ell \; V)^E \mid \env \mid ((\slk, (\gamma', H)^\dagger) \cons \shk) \circ \shk'} &\stepsto&
\multicolumn{2}{@{}l@{}}{\cek{M \mid \env'[x \mapsto \val{V}{\env}, r \mapsto (\shk', \slk)^B] \mid \shk},}\\
&&&\multicolumn{2}{@{}r@{}}{\text{if } \ell : A \to B \in E \text{ and } \hell = \{\ell\; x \; r \mapsto M\}} \\
% Forward
\mlab{Forward} & \cek{ (\Do \; \ell \; V)^E \mid \env \mid ((\slk, (\env', H)^\depth) \cons \shk) \circ \shk'}
&\stepsto& \cek{ (\Do \; \ell \; V)^E \mid \env \mid \shk \circ (\shk' \concat [(\slk, (\env', H)^\depth)])},
& \text{if } \hell = \emptyset \\
\ea \\
\el
\end{displaymath}
\textbf{Value interpretation}
\begin{displaymath}
\ba{@{}r@{~}c@{~}l@{}}
\val{x}{\env} &=& \env(x) \\
\val{\Record{}}{\env} &=& \Record{} \\
\ea
\qquad
\ba{@{}r@{~}c@{~}l@{}}
\val{\lambda x^A.M}{\env} &=& (\env, \lambda x^A.M) \\
\val{\Record{\ell = V; W}}{\env} &=& \Record{\ell = \val{V}{\env}; \val{W}{\env}} \\
\ea
\qquad
\ba{@{}r@{~}c@{~}l@{}}
\val{\Lambda \alpha^K.M}{\env} &=& (\env, \Lambda \alpha^K.M) \\
\val{(\ell\, V)^R}{\env} &=& (\ell\; \val{V}{\env})^R \\
\ea
\qquad
\val{\Rec\,g^{A \to C}\,x.M}{\env} = (\env, \Rec\,g^{A \to C}\,x.M) \\
\end{displaymath}
\caption{Abstract machine semantics.}
\label{fig:abstract-machine-semantics}
\end{minipage}
\end{figure}
%
%
A configuration $\conf = \cek{M \mid \env \mid \shk \circ \shk'}$
of our abstract machine is a quadruple of a computation term ($M$), an
environment ($\env$) mapping free variables to values, and two
continuations ($\shk$) and ($\shk'$).
%
The latter continuation is always the identity, except when forwarding
an operation, in which case it is used to keep track of the extent to
which the operation has been forwarded.
%
We write $\cek{M \mid \env \mid \shk}$ as syntactic sugar for $\cek{M
\mid \env \mid \shk \circ []}$ where $[]$ is the identity
continuation.
%
%% Our continuations differ from the standard machine. On the one hand,
%% they are somewhat simplified, due to our strict separation between
%% computations and values. On the other hand, they have considerably
%% more structure in order to accommodate effects and handlers.
Values consist of function closures, type function closures, records,
variants, and captured continuations.
%
A continuation $\shk$ is a stack of frames $[\shf_1, \dots,
\shf_n]$. We annotate captured continuations with input types in order
to make the results of Section~\ref{subsec:machine-correctness} easier
to state. Each frame $\shf = (\slk, \chi)$ represents pure
continuation $\slk$, corresponding to a sequence of let bindings,
inside handler closure $\chi$.
%
A pure continuation is a stack of pure frames. A pure frame $(\env, x,
N)$ closes a let-binding $\Let \;x=[~] \;\In\;N$ over environment
$\env$. A handler closure $(\env, H)$ closes a handler definition $H$
over environment $\env$.
%
We write $\nil$ for an empty stack, $x \cons s$ for the result of
pushing $x$ on top of stack $s$, and $s \concat s'$ for the
concatenation of stack $s$ on top of $s'$. We use pattern matching to
deconstruct stacks.
The abstract machine semantics defining the transition function $\stepsto$ is given in
Fig.~\ref{fig:abstract-machine-semantics}.
%
It depends on an interpretation function $\val{-}$ for values.
%
The machine is initialised (\mlab{Init}) by placing a term in a
configuration alongside the empty environment and identity
continuation.
%
The rules (\mlab{AppClosure}), (\mlab{AppRec}), (\mlab{AppCont}),
(\mlab{AppCont^\dagger}), (\mlab{AppType}), (\mlab{Split}), and
(\mlab{Case}) enact the elimination of values.
%
The rules (\mlab{Let}) and (\mlab{Handle}) extend the current
continuation with let bindings and handlers respectively.
%
The rule (\mlab{RetCont}) binds a returned value if there is a pure
continuation in the current continuation frame;
%
(\mlab{RetHandler}) invokes the return clause of a handler if the pure
continuation is empty; and (\mlab{RetTop}) returns a final value if
the continuation is empty.
%
%% The rule (\mlab{Op}) switches to a special four place configuration in
%% order to handle an operation. The fourth component of the
%% configuration is an auxiliary forwarding continuation, which keeps
%% track of the continuation frames through which the operation has been
%% forwarded. It is initialised to be empty.
%% %
The rule (\mlab{Do}) applies the current handler to an operation if
the label matches one of the operation clauses. The captured
continuation is assigned the forwarding continuation with the current
frame appended to the end of it.
%
The rule (\mlab{Do^\dagger}) is much like (\mlab{Do}), except it
constructs a shallow resumption, discarding the current handler but
keeping the current pure continuation.
%
The rule (\mlab{Forward}) appends the current continuation
frame onto the end of the forwarding continuation.
%
The (\mlab{Init}) rule provides a canonical way to map a computation
term onto a configuration.
\newcommand{\Chiid}{\ensuremath{\Chi_{\text{id}}}}
\newcommand{\kappaid}{\ensuremath{\kappa_{\text{id}}}}
\paragraph{Example} To make the transition rules in
Figure~\ref{fig:abstract-machine-semantics} concrete we give an
example of the abstract machine in action. We reuse the small producer
and consumer from Section~\ref{sec:shallow-handlers-tutorial}. We
reproduce their definitions here in ANF.
%
\[
\bl
\Ones \defas \Rec\;ones~\Unit. \Do\; \Yield~1; ones~\Unit \\
\AddTwo \defas
\lambda \Unit.
\Let\; x \revto \Do\;\Await~\Unit \;\In\;
\Let\; y \revto \Do\;\Await~\Unit \;\In\;
x + y \\
\el
\]%
%
Let $N_x$ denote the term $\Let\;x \revto \Do\;\Await~\Unit\;\In\;N_y$
and $N_y$ the term $\Let\;y \revto \Do\;\Await~\Unit\;\In\;x+y$.
%
%% For clarity, we annotate each bound variable $resume$ with a subscript
%% $\Await$ or $\Yield$ according to whether it was captured by a
%% consumer or a producer.
%
Suppose $\Ones$, $\AddTwo$, $\Pipe$, and $\Copipe$ are bound in
$\env_\top$. Furthermore, let $H_\Pipe$ and $H_\Copipe$ denote the
pipe and copipe handler definitions. The machine begins by applying
$\Pipe$.
%
\begin{derivation}
&\cek{\Pipe~\Record{\Ones, \AddTwo} \mid \env_\top \mid \kappaid}\\
\stepsto& \reason{apply $\Pipe$}\\
&\cek{\ShallowHandle\;c~\Unit\;\With\;H_\Pipe \mid \env_\top[c \mapsto (\emptyset, \AddTwo), p \mapsto (\emptyset, \Ones)] \mid \kappaid}\\
\stepsto& \reason{install $H_\Pipe$ with $\env_\Pipe = \env_\top[c \mapsto (\emptyset, \AddTwo), p \mapsto (\emptyset, \Ones)]$}\\
&\cek{c~\Unit \mid \env_\Pipe \mid (\nil, (\env_\Pipe, H_\Pipe)) \cons \kappaid}\\
\stepsto& \reason{apply $c$ and $\val{c}\env_\Pipe = (\emptyset, \AddTwo)$}\\
&\cek{N_x \mid \emptyset \mid (\nil, (\env_\Pipe, H_\Pipe)) \cons \kappaid}\\
\stepsto& \reason{focus on left operand}\\
&\cek{\Do\;\Await~\Unit \mid \emptyset \mid ([(\emptyset, x, N_y)], (\env_\Pipe, H_\Pipe)) \cons \kappaid}\\
\stepsto& \reason{shallow continuation capture $v_\Await = (\nil, [(\emptyset, x, N_y)])$}\\
&\cek{\Copipe~\Record{resume, p} \mid \env_\Pipe[resume \mapsto v_\Await] \mid \kappaid}\\
\end{derivation}
%
The invocation of $\Await$ begins a search through the machine
continuation to locate a matching handler. In this instance the
top-most handler $H_\Pipe$ handles $\Await$. The complete shallow
resumption consists of an empty continuation and a singleton pure
continuation. The former is empty as $H_\Pipe$ is a shallow handler,
meaning that it is discarded.
Evaluation continues by applying $\Copipe$.
%
\begin{derivation}
\stepsto& \reason{apply $\Copipe$}\\
&\cek{\ShallowHandle\; p~\Unit \;\With\;H_\Copipe \mid \env_\top[c \mapsto v_\Await, p \mapsto (\emptyset, \Ones)] \mid \kappaid}\\
\stepsto& \reason{install $H_\Copipe$ with $\env_\Copipe = \env_\top[c \mapsto v_\Await, p \mapsto (\emptyset, \Ones)]$}\\
&\cek{p~\Unit \mid \env_\Copipe \mid (\emptyset, (\env_\Copipe, H_\Copipe)) \cons \kappaid}\\
\stepsto^2& \reason{apply $p$, $\val{p}\env_\Copipe = (\emptyset, \Ones)$, and $\env_{\Ones} = \emptyset[ones \mapsto (\emptyset, \Ones)]$}\\
% &\cek{\Do\;\Yield~1;~ones~\Unit \mid \env_{ones} \mid (\nil, (\env_\Copipe, H_\Copipe)) \cons \kappaid}\\
% \stepsto^2& \reason{focus on $\Yield$}\\
&\cek{\Do\;\Yield~1 \mid \env_{\Ones} \mid ([(\env_{\Ones},\_,ones~\Unit)], (\env_\Copipe, H_\Copipe)) \cons \kappaid}\\
\stepsto& \reason{shallow continuation capture $v_\Yield = (\nil, [(\env_{\Ones}, \_, ones~\Unit)])$}\\
&\cek{\Pipe~\Record{resume, \lambda\Unit.c~s} \mid \env_\Copipe[s \mapsto 1,resume \mapsto v_\Yield] \mid \kappaid}
\end{derivation}
%
At this point the situation is similar as before: the invocation of
$\Yield$ causes the continuation to be unwound in order to find an
appropriate handler, which happens to be $H_\Copipe$. Next $\Pipe$ is
applied.
%
\begin{derivation}
\stepsto& \reason{apply $\Pipe$ and $\env_\Pipe' = \env_\top[c \mapsto (\env_\Copipe[c \mapsto v_\Await, s\mapsto 1]), p \mapsto v_\Yield])]$}\\
&\cek{\ShallowHandle\;c~\Unit\;\With\; H_\Pipe \mid \env_\Pipe' \mid \kappaid}\\
\stepsto& \reason{install $H_\Pipe$}\\
&\cek{c~\Unit \mid \env_\Pipe' \mid (\nil, (\env_\Pipe', H_\Pipe)) \cons \kappaid}\\
\stepsto& \reason{apply $c$ and $\val{c}\env_\Pipe' = (\env_\Copipe[c \mapsto v_\Await, s\mapsto 1])$}\\
&\cek{c~s \mid \env_\Copipe[c \mapsto v_\Await, s\mapsto 1] \mid (\nil, (\env_\Pipe', H_\Pipe)) \cons \kappaid}\\
\stepsto& \reason{shallow resume with $v_\Await = (\nil, [(\emptyset,x,N_y)])$}\\
&\cek{\Return\;1 \mid \env_\Pipe' \mid ([(\emptyset,x,N_y)], (\env_\Pipe', H_\Pipe)) \cons \kappaid}
\end{derivation}
%
Applying the resumption concatenates the first component (the
continuation) with the machine continuation. The second component
(pure continuation) gets concatenated with the pure continuation of
the top-most frame of the machine continuation. Thus in this
particular instance, the machine continuation is manipulated as
follows.
%
\[
\ba{@{~}l@{~}l}
&\nil \concat ([(\emptyset,x,N_y)] \concat \nil, (\env_\Pipe', H_\Pipe)) \cons \kappaid\\
=& ([(\emptyset,x,N_y)], (\env_\Pipe', H_\Pipe)) \cons \kappaid
\ea
\]
%
Because the control component contains the expression $\Return\;1$ and
the pure continuation is nonempty, the machine applies the pure
continuation.
\begin{derivation}
\stepsto& \reason{apply pure continuation, $\env_{\AddTwo} = \emptyset[x \mapsto 1]$}\\
&\cek{N_y \mid \env_{\AddTwo} \mid (\nil, (\env_\Pipe', H_\Pipe)) \cons \kappaid}\\
\stepsto& \reason{focus on right operand}\\
&\cek{\Do\;\Await~\Unit \mid \env_{\AddTwo} \mid ([(\env_{\AddTwo}, y, x + y)], (\env_\Pipe', H_\Pipe)) \cons \kappaid}\\
\stepsto^2& \reason{shallow continuation capture $w_\Await = (\nil, [(\env_{\AddTwo}, y, x + y)])$, apply $\Copipe$}\\
&\cek{\ShallowHandle\;p~\Unit \;\With\; \env_\Copipe \mid \env_\top[c \mapsto w_\Await, p \mapsto v_\Yield] \mid \kappaid}\\
\reducesto& \reason{install $H_\Copipe$ with $\env_\Copipe = \env_\top[c \mapsto w_\Await, p \mapsto v_\Yield]$}\\
&\cek{p~\Unit \mid \env_\Copipe \mid (\nil, (\env_\Copipe, H_\Copipe)) \cons \kappaid}
\end{derivation}
%
The variable $p$ is bound to the shallow resumption $v_\Yield$, thus
invoking it will transfer control back to the $\Ones$ computation.
%
\begin{derivation}
\stepsto & \reason{shallow resume with $v_\Yield = (\nil, [(\env_{\Ones}, \_, ones~\Unit)])$}\\
&\cek{\Return\;\Unit \mid \env_\Copipe \mid ([(\env_{\Ones}, \_, ones~\Unit)], (\env_\Copipe, H_\Copipe)) \cons \kappaid}\\
\stepsto^3& \reason{apply pure continuation, apply $ones$, focus on $\Yield$}\\
&\cek{\Do\;\Yield~1 \mid \env_{\Ones} \mid ([(\env_{\Ones}, \_, ones~\Unit)], (\env_\Copipe, H_\Copipe)) \cons \kappaid}\\
\end{derivation}
%
At this stage the machine repeats the transitions from before: the
shallow continuation of $\Do\;\Yield~1$ is captured, control passes to
the $\Yield$ clause in $H_\Copipe$, which again invokes $\Pipe$ and
subsequently installs the $H_\Pipe$ handler with an environment
$\env_\Pipe''$. The handler runs the computation $c~\Unit$, where $c$
is an abstraction over the resumption $w_\Await$ applied to the
yielded value $1$.
%
\begin{derivation}
\stepsto^6 & \reason{by the above reasoning, shallow resume with $w_\Await = (\nil, [(\env_{\AddTwo}, y, x + y)])$}\\
&\cek{x + y \mid \env_{\AddTwo}[y \mapsto 1] \mid (\nil, (\env_\Pipe'', H_\Pipe)) \cons \kappaid}\\
\stepsto& \reason{$\val{x}\env_{\AddTwo}[y\mapsto 1] = 1$ and $\val{y}\env_{\AddTwo}[y\mapsto 1] = 1$}\\
&\cek{\Return\;2 \mid \env_{\AddTwo}[y \mapsto 1] \mid (\nil, (\env_\Pipe'', H_\Pipe)) \cons \kappaid}
\end{derivation}
%
Since the pure continuation is empty the $\Return$ clause of $H_\Pipe$
gets invoked with the value $2$. Afterwards the $\Return$ clause of the
identity continuation in $\kappaid$ is invoked, ultimately
transitioning to the following final configuration.
%
\begin{derivation}
\stepsto^2& \reason{by the above reasoning}\\
&\cek{\Return\;2 \mid \emptyset \mid \nil}
\end{derivation}
%
%\env_\top[c \mapsto (\env_\Copipe, \lambda\Unit.w_\Await~1),p \mapsto (\nil, [(\env_{\Ones}, \_, ones~\Unit)])]
\begin{figure}[t]
\flushleft
\newcommand{\contapp}[2]{#1 #2}
\newcommand{\contappp}[2]{#1(#2)}
%% \newcommand{\contapp}[2]{#1[#2]}
%% \newcommand{\contapp}[2]{#1\mathbin{@}#2}
%% \newcommand{\contappp}[2]{#1\mathbin{@}(#2)}
%
Configurations
\begin{displaymath}
\inv{\cek{M \mid \env \mid \shk \circ \shk'}} = \contappp{\inv{\shk' \concat \shk}}{\inv{M}\env}
= \contappp{\inv{\shk'}}{\contapp{\inv{\shk}}{\inv{M}\env}}
\end{displaymath}
Pure continuations
\begin{displaymath}
\contapp{\inv{[]}}{M} = M \qquad \contapp{\inv{((\env, x, N) \cons \slk)}}{M}
= \contappp{\inv{\slk}}{\Let\; x \revto M \;\In\; \inv{N}(\env \res \{x\})}
\end{displaymath}
%% \begin{equations}
%% \contapp{\inv{[]}}{M}
%% &=& M \\
%% \contapp{\inv{((\env, x, N) \cons \slk)}}{M}
%% &=& \contappp{\inv{\slk}}{\Let\; x \revto M \;\In\; \inv{N}(\env \res \{x\})} \\
%% \end{equations}
Continuations
\begin{displaymath}
\contapp{\inv{[]}}{M}
= M \qquad
\contapp{\inv{(\slk, \chi) \cons \shk}}{M}
= \contapp{\inv{\shk}}{(\contappp{\inv{\chi}}{\contappp{\inv{\slk}}{M}})}
\end{displaymath}
%% \begin{equations}
%% \contapp{\inv{[]}}{M}
%% &=& M \\
%% \contapp{\inv{(\slk, \chi) \cons \shk}}{M}
%% &=& \contapp{\inv{\shk}}{(\contappp{\inv{\chi}}{\contappp{\inv{\slk}}{M}})} \\
%% \end{equations}
Handler closures
\begin{displaymath}
\contapp{\inv{(\env, H)}^\depth}{M}
= \Handle^\depth\;M\;\With\;\inv{H}\env
\end{displaymath}
%% \begin{equations}
%% \contapp{\inv{(\env, H)}}{M}
%% &=& \Handle\;M\;\With\;\inv{H}\env \\
%% \contapp{\inv{(\env, H)^\dagger}}{M}
%% &=& \ShallowHandle\;M\;\With\;\inv{H}\env \\
%% \end{equations}
Computation terms
\begin{equations}
\inv{V\,W}\env &=& \inv{V}\env\,\inv{W}{\env} \\
\inv{V\,T}\env &=& \inv{V}\env\,T \\
\inv{\Let\;\Record{\ell = x; y} = V \;\In\;N}\env
&=& \Let\;\Record{\ell = x; y} =\inv{V}\env \;\In\; \inv{N}(\env \res \{x, y\}) \\
\inv{\Case\;V\,\{\ell\;x \mapsto M; y \mapsto N\}}\env
&=& \Case\;\inv{V}\env \,\{\ell\;x \mapsto \inv{M}(\env \res \{x\}); y \mapsto \inv{N}(\env \res \{y\})\} \\
\inv{\Return\;V}\env &=& \Return\;\inv{V}\env \\
\inv{\Let\;x \revto M \;\In\;N}\env
&=& \Let\;x \revto\inv{M}\env \;\In\; \inv{N}(\env \res \{x\}) \\
\inv{\Do\;\ell\;V}\env
&=& \Do\;\ell\;\inv{V}\env \\
\inv{\Handle^\depth\;M\;\With\;H}\env
&=& \Handle^\depth\;\inv{M}\env\;\With\;\inv{H}\env \\
\end{equations}
Handler definitions
\begin{equations}
\inv{\{\Return\;x \mapsto M\}}\env
&=& \{\Return\;x \mapsto \inv{M}(\env \res \{x\})\} \\
\inv{\{\ell\;x\;k \mapsto M\} \uplus H}\env
&=& \{\ell\;x\;k \mapsto \inv{M}(\env \res \{x, k\}\} \uplus \inv{H}\env \\
\end{equations}
Value terms and values
\begin{displaymath}
\ba{@{}c@{}}
\begin{eqs}
\inv{x}\env &=& \inv{v}, \quad \text{ if }\env(x) = v \\
\inv{x}\env &=& x, \quad \text{ if }x \notin dom(\env) \\
\inv{\lambda x^A.M}\env &=& \lambda x^A.\inv{M}(\env \res \{x\}) \\
\inv{\Lambda \alpha^K.M}\env &=& \Lambda \alpha^K.\inv{M}\env \\
\inv{\Record{}}\env &=& \Record{} \\
\inv{\Record{\ell=V; W}}\env &=& \Record{\ell=\inv{V}\env; \inv{W}\env} \\
\inv{(\ell\;V)^R}\env &=& (\ell\;\inv{V}\env)^R \\
\end{eqs}
\quad
\begin{eqs}
\inv{\shk^A} &=& \lambda x^A.\inv{\shk}(\Return\;x) \\
\inv{(\shk, \slk)^A} &=& \lambda x^A.\inv{\slk}(\inv{\shk}(\Return\;x)) \\
\inv{(\env, \lambda x^A.M)} &=& \lambda x^A.\inv{M}(\env \res \{x\}) \\
\inv{(\env, \Lambda \alpha^K.M)} &=& \Lambda \alpha^K.\inv{M}\env \\
\inv{\Record{}} &=& \Record{} \\
\inv{\Record{\ell=v; w}} &=& \Record{\ell=\inv{v}; \inv{w}} \\
\inv{(\ell\;v)^R} &=& (\ell\;\inv{v})^R \\
\end{eqs} \smallskip\\
\inv{\Rec\,g^{A \to C}\,x.M}\env = \Rec\,g^{A \to C}\,x.\inv{M}(\env \res \{g, x\})
= \inv{(\env, \Rec\,g^{A \to C}\,x.M)} \\
\ea
\end{displaymath}
\caption{Mapping from abstract machine configurations to terms.}
\label{fig:config-to-term}
\end{figure}
\paragraph{Remark} If the main continuation is empty then the machine gets
stuck. This occurs when an operation is unhandled, and the forwarding
continuation describes the succession of handlers that have failed to
handle the operation along with any pure continuations that were
encountered along the way.
%
Assuming the input is a well-typed closed computation term $\typc{}{M
: A}{E}$, the machine will either not terminate, return a value of
type $A$, or get stuck failing to handle an operation appearing in
$E$. We now make the correspondence between the operational semantics
and the abstract machine more precise.
\section{Correctness}
\label{subsec:machine-correctness}
Fig.~\ref{fig:config-to-term} defines an inverse mapping $\inv{-}$
from configurations to computation terms via a collection of mutually
recursive functions defined on configurations, continuations,
computation terms, handler definitions, value terms, and values.
%
We write $dom(\gamma)$ for the domain of $\gamma$ and $\gamma \res
\{x_1, \dots, x_n\}$ for the restriction of environment $\gamma$ to
$dom(\gamma) \res \{x_1, \dots, x_n\}$.
%
The $\inv{-}$ function enables us to classify the abstract machine
reduction rules by how they relate to the operational
semantics.
%
The rules (\mlab{Init}) and (\mlab{RetTop}) concern only initial
input and final output, neither a feature of the operational
semantics.
%
The rules (\mlab{AppCont^\depth}), (\mlab{Let}), (\mlab{Handle}),
and (\mlab{Forward}) are administrative in that $\inv{-}$ is invariant
under them.
%
This leaves $\beta$-rules (\mlab{AppClosure}), (\mlab{AppRec}),
(\mlab{AppType}), (\mlab{Split}), (\mlab{Case}), (\mlab{RetCont}),
(\mlab{RetHandler}), (\mlab{Do}), and (\mlab{Do^\dagger}),
each of which corresponds directly to performing a reduction in the
operational semantics.
%
We write $\stepsto_a$ for administrative steps, $\stepsto_\beta$ for
$\beta$-steps, and $\Stepsto$ for a sequence of steps of the form
$\stepsto_a^* \stepsto_\beta$.
Each reduction in the operational semantics is simulated by a sequence
of administrative steps followed by a single $\beta$-step in the
abstract machine. The $Id$ handler (\S\ref{subsec:terms})
implements the top-level identity continuation.
\\
\begin{theorem}[Simulation]
\label{lem:simulation}
If $M \reducesto N$, then for any $\conf$ such that $\inv{\conf} =
Id(M)$ there exists $\conf'$ such that $\conf \Stepsto \conf'$ and
$\inv{\conf'} = Id(N)$.
%% If $M \reducesto N$, then for any $\conf$ such that $\inv{\conf} = M$
%% there exists $\conf'$ such that $\conf \Stepsto \conf'$ and
%% $\inv{\conf'} = N$.
\end{theorem}
\begin{proof}
By induction on the derivation of $M \reducesto N$.
\end{proof}
\begin{corollary}
If $\typc{}{M : A}{E}$ and $M \reducesto^+ N \not\reducesto$, then $M
\stepsto^+ \conf$ with $\inv{\conf} = N$.
\end{corollary}
\part{Expressiveness}
\chapter{Computability, complexity, and expressivness}

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