rules.tex

In this section we present the type checking algorithm for {\Core} with meta-variables.

First we extend
the syntax of signatures to include guarded constants and add a new syntactic
category for user expressions:

{\small
\[\begin{array}{lcll}
    C & ::= & \TypeConstr A B \Or \TermConstr M N A \Or
	      \TermConstr {\bar M} {\bar N} \Delta \\
    \Sigma & ::= & \ldots \Or
		\Ext \Sigma \ConstDecl p A M \Cs \\
    e & ::= & \LAM xe \Or x\,\bar e \Or c\,\bar e \Or \SET \Or \PI xee \Or ? \\
\end{array}\]
}

The input to the type checking algorithm is a user expression which could
represent either a type or a term. Apart from the usual constructions user
expressions can also contain $?$ representing a meta-variable. During type
checking user expressions are translated into {\Core} terms where
meta-variables are represented as fresh constants. Note that since we have
domain free lambda abstractions we cannot type check $\beta$-redexes. Hence the
syntax of user expressions disallows them.

A constraint $C$ is an equality constraint that has been postponed
because not enough information was available about the meta-variables. Since
our conversion checking algorithm is typed the constraints must also be typed. The
constraints show up in the signature as guards to guarded constants. We write
$\ConstDecl p A M \Cs$ for a guarded constant $p$ of type $A$ and value $M$
guarded by the set of constraints $\Cs$. We have the computation rule that $p$
computes to $M$ when $\Cs$ is the empty set.

We use the naming convention that lowercase greek letters $\alpha, \beta,
\ldots$ stand for constants representing meta-variables and $p$ and $q$ for
guarded constants.

\subsection{Operations on the signature}

All rules work on a signature $\Sigma$, containing previously defined
constants, meta-variables, and guarded constants.
%
In other words we can write all judgements on the form
\(\ExplicitJudgement \Sigma J {\Sigma'}\).
%
To make the rules easier to read we first define a set of operations reading
and modifying the signature and when presenting the algorithm simply write $J$
for the judgement above. In rules with multiple premisses the signature is
threaded top-down, left-to-right. % For instance,
% \[\begin{array}{ccc}
% {\small\begin{array}[c]{l}
% \infer{J}
%   {\begin{array}[b]{l}
%      P_1
%   \\ P_2
%   \end{array}
%   & P_3
%   }
% \end{array}}
% & ~\mbox{is short-hand for}~
% &
% {\small\begin{array}[c]{l}
% \infer{\ExplicitJudgement {\Sigma_1} J {\Sigma_4}}
%   {\begin{array}[b]{l}
%      \ExplicitJudgement {\Sigma_1} {P_1} {\Sigma_2}
%   \\ \ExplicitJudgement {\Sigma_2} {P_2} {\Sigma_3}
%   \end{array}
%   & \ExplicitJudgement {\Sigma_3} {P_3} {\Sigma_4}
%   }
% \end{array}}
% \end{array}\]

\begin{figure}
\begin{tabular}{llll}
%    \multicolumn2l{Operations on meta-variables} \\
    & $\ExplicitJudgement \Sigma
	{\AddMeta \alpha A}
	{\Ext \Sigma \MetaDecl \alpha A}
    $ & if & $\alpha \notin \Sigma$
    \\
    & $\ExplicitJudgement
	    \Sigma
	    {\InstMeta \alpha M}
	    {\Ext {\Sigma_1} \Ext {\IMetaDecl \alpha A M} \Sigma_2}
    $ & if & $\Sigma ~=~ \Ext {\Sigma_1} \Ext {\MetaDecl \alpha A} \Sigma_2$
%     \\
%     & $\ExplicitJudgement
% 	    \Sigma
% 	    {\LookupMeta \alpha M}
% 	    \Sigma
%     $ & if & $\IMetaDecl \alpha A M ~\in~ \Sigma$
%     \\
%     & $\ExplicitJudgement
% 	    \Sigma
% 	    {\Uninstantiated \alpha}
% 	    \Sigma
%     $ & if & $\MetaDecl \alpha A ~\in~ \Sigma$
    \\ {} \\
%    \multicolumn2l{Operations on guarded constants} \\
    & \multicolumn3l{$\ExplicitJudgement
	    \Sigma
	    {\AddConst p A M \Cs}
	    {\Ext \Sigma \ConstDecl p A M \Cs}
    $} \\
    & & if & $p \notin \Sigma$
%     \\
%     & \multicolumn3l{$\ExplicitJudgement
% 	    \Sigma
% 	    {\UpdateGuard p \Cs}
% 	    {\Ext {\Sigma_1} \Ext {\ConstDecl p A M \Cs} \Sigma_2}
%     $} \\
%     & & if & $\Sigma ~=~ \Ext {\Sigma_1} \Ext {\ConstDecl p A M \Cs^\prime} \Sigma_2$
%     \\
%     & $\ExplicitJudgement
% 	    \Sigma
% 	    {\LookupConst p M}
% 	    \Sigma
%     $ & if & $\ConstDecl p A M \emptyset ~\in~ \Sigma$
%     \\
%     & $\ExplicitJudgement
% 	    \Sigma
% 	    {\Guarded p}
% 	    \Sigma
%     $ & if & $\ConstDecl p A M \Cs ~\in~ \Sigma$ and $\Cs \neq \emptyset$
    \\ {} \\
%    \multicolumn2l{General operations} \\
%     & $\ExplicitJudgement
% 	    \Sigma
% 	    {\LookupType c A}
% 	    \Sigma
%     $ & if & $c : A \in \Sigma$
%     \\
    & $\ExplicitJudgement
	    \Sigma
	    {\InScope \alpha M}
	    \Sigma
    $ & if & $\begin{array}[t]{ll}
	\Sigma ~=~ \Ext {\Sigma_1} \Ext {\MetaDecl \alpha A} \Sigma_2 ~~ \mbox{and}
	\\ c \in M ~~ \mbox{implies} ~~ c \in \Sigma_1
%     \\	\IsTypeCS {\Sigma_1} {} A ~~ \mbox{and} ~~ \HasTypeCS \Sigma {} M A ~~ \mbox{implies}
%     \\	\HasTypeCS {\Sigma_1} {} M A
    \end{array}$
    \\
%     & $\ExplicitJudgement
% 	    \Sigma
% 	    {\WithSig {\Sigma'} J}
% 	    {\Sigma''}
%     $ & if & $\ExplicitJudgement
% 	    {\Sigma'}
% 	    J
% 	    {\Sigma''}
%     $ \\
\end{tabular}
\caption{Operations on the signature}
\label{figOperations}
\end{figure}

We introduce two operations to manipulate meta-variables: \\ $\AddMeta \alpha A$
adds a new meta-variable $\alpha$ of type $A$ to the signature, and $\InstMeta
\alpha M$ instantiates $\alpha$ to $M$.
%
%% Never used
% $\LookupMeta \alpha M$ looks up the
% value of an instantiated meta-variable, and $\Uninstantiated \alpha$ verifies
% that $\alpha$ is uninstantiated.
%
For guarded constants we just add the operation $\AddConst p A M \Cs$ to add a
new guarded constant to the signature.  In Section~\ref{secAlgorithm} we
explain the rules for solving the constraints of a guarded constant.
%
% , $\UpdateGuard p \Cs$ to update the guard of $p$,
% $\LookupConst p M$ to get the value of a guarded constant whose guard has been
% solved, and $\Guarded p$ to verify that the guard of $p$ has not been solved.
%
We also introduce an operation $\InScope \alpha M$ to check that $M$ is in
scope at the definition site of $\alpha$ (to ensure that $\alpha$ can be
instantiated to $M$).
%
%, and $\WithSig \Sigma J$ which checks $J$ in the signature $\Sigma$.
%
Detailed definitions of the operations can be found in
Figure~\ref{figOperations}.

\subsection{The algorithm} \label{secAlgorithm}

Next we present the type checking algorithm.  We use a bidirectional algorithm,
consisting of the following main judgement forms.

{\small
\[\begin{array}{lcl}
    \IsType e A && \mbox{well-formed types} \\
    \CheckType e A M && \mbox{type checking} \\
    \InferType e A M && \mbox{type inference} \\
    \EqualType A B \Cs && \mbox{type conversion} \\
    \Equal M N A \Cs &~& \mbox{term conversion} \\
\end{array}\]
}

The rules for well-formed types and type checking and inference take a user
expression and produce a type or term in {\Core} which is a
well-typed approximation of the user expression. Conversion checking produces
a set of unsolved constraints which needs to be solved for the judgement to be
true in {\Core}.

We use typed conversion for two reasons: it is a nice way to implement
$\eta$-equality, and perhaps more importantly to prove the correctness of the
algorithm we need the invariant that when checking $\Equal M N A \Cs$ we have
$\HasTypeC \Gamma M A$ and $\HasTypeC \Gamma N A$, so we need to record the
type to make sure the invariant is preserved.

When checking conversion we also need the following judgement forms.

{\small
\[\begin{array}{lcl}
    \EqualWhnf M N A \Cs &~& \mbox{conversion of weak head normal forms} \\
    \Equal {\bar M} {\bar N} \Delta \Cs &~& \mbox{conversion of sequences of terms} \\
\end{array}\]
}

Type checking with dependent types involves normalising arbitrary (type
correct) terms, so we need to know how to normalise terms in a signature
containing meta-variables and guarded constants. We do this by translating the
signature to {\Core} and performing the normalisation in {\Core}.

\begin{definition}
    Given a signature $\Sigma$ containing meta-variables and guarded constants
    we define its {\Core} restriction $\CoreSig \Sigma$ by replacing
    guarded constants with normal constants, replacing $\ConstDecl p A M \Cs$
    by $p : A = M$ if $\Cs$ is empty, and $p : A$ otherwise.
\end{definition}

The correctness of the type checking algorithm relies on the invariant that
when $\ExplicitJudgement \Sigma {\CheckType e A M} {\Sigma'}$, we have
$\HasTypeCS {\CoreSig {\Sigma'}} \Gamma M A$ (see Theorem~\ref{thmSoundNoCs}).

We write $\ExplicitJudgement \Sigma {\whnf M {M'}} \Sigma$ if $M'$ is the weak
head normal form of $M$ in $\CoreSig \Sigma$. Similarly $\Normalise M {M'}$
means that $M'$ is the normal form of $M$.

\subsubsection{Type checking rules}

To save some space we omit the rules for checking well-formed types and most of
the rules for type checking and inference. The rules are simple extensions of
standard type checking algorithms to produce well-typed terms. The interesting
type checking rules are the rule for type checking meta-variables and the
conversion rules.

\URules{

\infer{ \CheckType {?} A {\alpha \, \Gamma }}
{\begin{array}{l}
    \AddMeta \alpha {\Gamma \to A}
\end{array}}

\quad

\infer{
    \CheckType e A M
}{\begin{array}{l}
    \InferType e B M
\\  \EqualType A B \emptyset
\end{array}}

\quad

\infer{
    \CheckType e A {p\,\Gamma}
}{\begin{array}{l}
    \InferType e B M
\\  \EqualType A B \Cs \neq \emptyset
\\  \AddConst p {\Gamma \to A} {\LAM \Gamma M} \Cs
\end{array}}

}

When type checking a user meta-variable we create a fresh meta-variable, add it
to the signature and return it. Since meta-variables are part of the signature
they have to be lifted to the top-level.

We have two versions of the conversion rule. The first corresponds to the
normal conversion rule and applies when no constraints are generated. The
interesting case is when we cannot safely conclude that $A = B$, in which case
we introduce a fresh guarded constant. As meta-variables, guarded constants
are lifted to the top-level.

\subsubsection{Conversion rules}

When checking conversion of two function types, an interesting question is what
to do when comparing the domains gives rise to constraints. The rule in question is

\URules{
% (x : A) -> B = (x : A') -> B'
\infer{
    \EqualType {\PI x {A_1} {B_1}} {\PI x {A_2} {B_2}} {\Cs \cup \Cs^\prime}
}{\begin{array}{l}
    \EqualType {A_1} {A_2} \Cs, ~~ \Cs \neq \emptyset
\\  \AddConst p {\Gamma \to A_1 \to A_2} {\LAM {\Gamma\,x} x} \Cs
\\  \EqualTypeCtx {\Ext \Gamma x : A_1} {B_1} {\SubstD {B_2} {p ~ \Gamma \, x}} {\Cs^\prime}
\end{array}}
}

To ensure the correctness of the algorithm we need to maintain the invariant
that when we check $\EqualTypeCtx {} A B \Cs$ we have $\IsTypeC {} A$ and
$\IsTypeC {} B$. Thus if we do not know whether $A_1 = A_2$ it is not
correct to check $\EqualTypeCtx {x : A_1} {B_1} {B_2} {\Cs^\prime}$
since $B_2$ is not well-formed in the context $x : A_1$. To solve the problem
we substitute a guarded constant $p \, x$ for $x$ in $B_2$, where $p \, x$
reduces to $x$ when $A_1$ and $A_2$ are convertible.

% \quad
% 
% \infer{
%     \EqualType {\PI x {A_1} {B_1}} {\PI x {A_2} {B_2}} \Cs
% }{\begin{array}{l}
%     \EqualType {A_1} {A_2} \emptyset
% \\  \EqualTypeCtx {\Ext \Gamma x : A_1} {B_1} {B_2} \Cs
% \end{array}}
% 
%  \\{}\\
% 
% % El M = El N
% 
% \infer{
%     \EqualType {\EL M} {\EL N} \Cs
% }{\begin{array}{l}
%     \Equal M N \SET \Cs
% \end{array}}
% 
% \end{array}\]}

\subsubsection{Term conversion rules}

Checking conversion of terms is done on weak head normal forms. The only rule
that is applied before weak head normalisation is the $\eta$-rule.

\URules{

% Eta

\infer{
    \Equal M N {\PI x A B} \Cs
}{\begin{array}{l}
    \EqualCtx {\Ext \Gamma x : A} {M \, x} {N \, x} B \Cs
\end{array}}

\qquad

% Weak head normalisation

\infer{
    \Equal M N A \Cs
}{\begin{array}[b]{lcl}
    \whnf M {M'}
\\  \whnf N {N'}
&&  \EqualWhnf {M'} {N'} A \Cs
\end{array}
}

}

In {\Core} function types are not terms so a meta-variable can never be
instantiated to a function type. If this was the case we would have to check if
the type was a meta-variable, and if so postpone the constraint, since we would
not know whether or not the $\eta$-rule should be applied.

The weak head normal forms we compare will be of atomic type and so they are of
the form $h\,\bar M$ where the head $h$ is a variable, constant, meta-variable,
or guarded constant. If both terms have the same variable or constant head $h :
\Delta \to A$ we compare the arguments in $\Delta$.

\URules{

\infer{
    \EqualWhnf {h ~ \bar M} {h ~ \bar N} A \Cs
}{\begin{array}{lcl}
    h : \Delta \to B
&&   \Equal {\bar M} {\bar N} \Delta \Cs
\end{array}}
}

If the heads are different constants or variables conversion checking fails.
If one of the heads is a guarded constant we give up and return the problem as
a constraint.

\URules{
\infer{
    \EqualWhnf {p ~ \bar M} N A {\left\{\TermConstr {p ~ \bar M} N A\right\}}
}{}
}

If one of the heads is a meta variable we use a restricted form of pattern
unification, but we believe that our correctness proof can be extended to more
powerful unification algorithms, for example
\cite{dowek:matching,dowek:unification,miller:pattern,Nipkow-LICS-93,pfenning:unification}. The crucial step is
to prove that meta-variable instantiations are well-typed.
In the examples we have studied, using meta-variables for
implicit arguments, this simpler form of unification seems to be sufficient.
The rule for meta-variable instantiation is

\URules{

% Instantiation

\infer{
    \EqualWhnf {\alpha ~ \bar x} M A \emptyset
}{
\begin{array}[b]{l}
    \bar x~\mathit{distinct}
\\  \Normalise M {M'}
\\  \FV {M'} \subseteq \bar x
\end{array}
& \begin{array}[b]{l}
    \InScope \alpha {\LAM {\bar x} M'}
\\  \InstMeta \alpha {\LAM {\bar x} M'}
\end{array}
}
}

Given the problem $\alpha \, \bar x = M$ we would like to instantiate $\alpha$ to
$\LAM {\bar x} M$. This is only correct if $\bar x$ are distinct variables, $M$
does not contain any variables other than $\bar x$, and any constants refered
to by $M$ are in scope at the declaration site of $\alpha$\footnote{Note that
scope checking subsumes the usual occurs check, since constants are non-recursive.}.
Now $M$ might refer to meta-variables introduced after $\alpha$ but which have
been instantiated.  For this reason we normalise $M$ to $M'$ and try to
instantiate $\alpha$ to $\LAM {\bar x} M'$. A possible optimisation might be to
only normalise if $M$ contains out-of-scope constants or variables.
If any of the premisses in this rule fail or $\alpha$ is not applied only to
variables, we return the constraint as it is.

When checking conversion of argument lists, the interesting case is when
comparing the first arguments results in some unsolved constraints.

\URules{

% No constraints

% \infer{
%     \Equal {M, \, \bar M} {N, \, \bar N} {(x : A) \Delta} \Cs
% }{\begin{array}{l}
%     \Equal M N A \emptyset
% \\  \Equal {\bar M} {\bar N} {\SubstD \Delta M} \Cs
% \end{array}}
% 
% \\{}\\

% Some constraints

\Rule{
    \Equal {M, \, \bar M} {N, \, \bar N} {(x : A) \Delta} {~~} \\
    \hfill \left\{ \TermConstr {M, \, \bar M} {N, \, \bar N} {(x : A) \Delta} \right\}
}{\begin{array}{lcl}
    \Equal M N A \Cs \neq \emptyset
&&  x \in \FV \Delta
\end{array}}

\quad

\Rule{
    \Equal {M, \, \bar M} {N, \, \bar N} {(x : A) \Delta} {\Cs_1 \cup \Cs_2}
}{\begin{array}{lcl}
    \Equal M N A {\Cs_1} \neq \emptyset
\\  \Equal {\bar M} {\bar N} \Delta {\Cs_2}
&& x \notin \FV \Delta
\end{array}}

}

If the value of the first argument is used in the types of later arguments ($x
\in \FV \Delta$) we have to stop and produce a constraint since the types of
$\bar M$ and $\bar N$ differ. If on the other hand the types of later arguments
are independent of the value of the first argument, we can proceed and compare
them without knowing whether the first arguments are convertible.

% There is a possible inefficiency in that the constraint produced in the first
% case does not remember the result of comparing $M$ and $N$. This could be
% remedied by giving more structure to the constraint sets, requiring that the
% constraints in $\Cs$ are solved before comparing $\bar M$ and $\bar N$.

\subsubsection{Constraint Solving}

So far, we have not looked at when or how the guards of a constant are
simplified or solved. In principle this can be done at any time, for instance
as a separate phase after type checking. In practise, however, it might be a
better idea to interleave constraint solving and type checking. In
Section~\ref{secProof} we prove that this can be done safely.
Constraint solving amounts to rechecking the guard of a constant and replacing
it by the resulting constraints.

% The notation
% $\CheckConstr \Cs {\Cs^\prime}$ means checking each of the constraints in $\Cs$
% and taking $\Cs^\prime$ to be the union of the results.
% 
% % Constraint solving
% \URules{
%     \infer{
% 	\ExplicitJudgement
% 	    \Sigma
% 	    \Simplify
% 	    {\Ext {\Sigma^\prime_1} \Ext {\ConstDecl p A M {\Cs^\prime}} \Sigma_2}
%     }{\begin{array}{ll}
% 	\multicolumn2l{\Sigma ~ = ~ \Ext {\Sigma_1} \Ext {\ConstDecl p A M \Cs} {\Sigma_2}} \\
% 	\ExplicitJudgement {\Sigma_1} {\CheckConstr \Cs {\Cs^\prime}} {\Sigma^\prime_1} 
% 	% & \Cs \neq {\Cs^\prime} \\
%     \end{array}
%     }
% }

% Not used
% \begin{definition}[Normal signature]
%     A signature $\Sigma$ is in {\em normal form} if it is not the case that
%     $\ExplicitJudgement \Sigma \Simplify {\Sigma^\prime}$ for some $\Sigma^\prime$.
% \end{definition}

\if \NoteOnPatternMatching 1
\subsection{Adding pattern matching} \label{secAddPatternMatching}

If we have definitions by pattern matching reduction to weak head normal form
might be blocked by an uninstantiated meta variable. For instance $\neg ~
\alpha$ cannot be reduced to weak head normal form if $\neg$ is defined by
$\neg ~\mathit{true} = \mathit{false}$ and $\neg ~\mathit{false} =
\mathit{true}$. Since conversion checking is done on weak head normal forms we
generate a constraint when encountering a blocked term.

% Blocked terms
% \URules{
%     \infer{
% 	\Blocked {c \, \bar M}
%     }{\begin{array}{l}
% 	c ~ \mbox{pattern matches on its $i^\mathrm{th}$ argument}
%     \\	M_i = \alpha \, \bar N ~ or ~ M_i = p \, \bar N ~ or ~ \Blocked {M_i}
%     \end{array}}
% 
% \qquad
% 
% \infer{
%     \Equal M N A {\left\{ \TermConstr {M'} {N'} A \right\}}
% }{\begin{array}[b]{l}
%     \whnf M {M'}
% \\  \whnf N {N'}
% \\  \Blocked {M'} ~ \mathit{or} ~ \Blocked {N'}
% \end{array}
% }
% }

\fi