(*<*)
theory Paper
imports "../Nominal/Nominal2"
"~~/src/HOL/Library/LaTeXsugar"
begin
consts
fv :: "'a \<Rightarrow> 'b"
abs_set :: "'a \<Rightarrow> 'b \<Rightarrow> 'c"
alpha_bn :: "'a \<Rightarrow> 'a \<Rightarrow> bool"
abs_set2 :: "'a \<Rightarrow> perm \<Rightarrow> 'b \<Rightarrow> 'c"
equ2 :: "'a \<Rightarrow> 'a \<Rightarrow> bool"
Abs_dist :: "'a \<Rightarrow> 'b \<Rightarrow> 'c"
Abs_print :: "'a \<Rightarrow> 'b \<Rightarrow> 'c"
definition
"equal \<equiv> (op =)"
fun alpha_set_ex where
"alpha_set_ex (bs, x) R f (cs, y) = (\<exists>pi. alpha_set (bs, x) R f pi (cs, y))"
fun alpha_res_ex where
"alpha_res_ex (bs, x) R f pi (cs, y) = (\<exists>pi. alpha_res (bs, x) R f pi (cs, y))"
fun alpha_lst_ex where
"alpha_lst_ex (bs, x) R f pi (cs, y) = (\<exists>pi. alpha_lst (bs, x) R f pi (cs, y))"
notation (latex output)
swap ("'(_ _')" [1000, 1000] 1000) and
fresh ("_ # _" [51, 51] 50) and
fresh_star ("_ #\<^sup>* _" [51, 51] 50) and
supp ("supp _" [78] 73) and
uminus ("-_" [78] 73) and
If ("if _ then _ else _" 10) and
alpha_set_ex ("_ \<approx>\<^raw:\,\raisebox{-1pt}{\makebox[0mm][l]{$_{\textit{set}}$}}>\<^bsup>_, _\<^esup> _") and
alpha_lst_ex ("_ \<approx>\<^raw:\,\raisebox{-1pt}{\makebox[0mm][l]{$_{\textit{list}}$}}>\<^bsup>_, _\<^esup> _") and
alpha_res_ex ("_ \<approx>\<^raw:\,\raisebox{-1pt}{\makebox[0mm][l]{$_{\textit{set+}}$}}>\<^bsup>_, _\<^esup> _") and
abs_set ("_ \<approx>\<^raw:{$\,_{\textit{abs\_set}}$}> _") and
abs_set2 ("_ \<approx>\<^raw:\raisebox{-1pt}{\makebox[0mm][l]{$\,_{\textit{list}}$}}>\<^bsup>_\<^esup> _") and
fv ("fa'(_')" [100] 100) and
equal ("=") and
alpha_abs_set ("_ \<approx>\<^raw:{$\,_{\textit{abs\_set}}$}> _") and
alpha_abs_lst ("_ \<approx>\<^raw:{$\,_{\textit{abs\_list}}$}> _") and
alpha_abs_res ("_ \<approx>\<^raw:{$\,_{\textit{abs\_set+}}$}> _") and
Abs_set ("[_]\<^bsub>set\<^esub>._" [20, 101] 999) and
Abs_lst ("[_]\<^bsub>list\<^esub>._" [20, 101] 999) and
Abs_dist ("[_]\<^bsub>#list\<^esub>._" [20, 101] 999) and
Abs_res ("[_]\<^bsub>set+\<^esub>._") and
Abs_print ("_\<^bsub>set\<^esub>._") and
Cons ("_::_" [78,77] 73) and
supp_set ("aux _" [1000] 10) and
alpha_bn ("_ \<approx>bn _")
consts alpha_trm ::'a
consts fa_trm :: 'a
consts fa_trm_al :: 'a
consts alpha_trm2 ::'a
consts fa_trm2 :: 'a
consts fa_trm2_al :: 'a
consts supp2 :: 'a
consts ast :: 'a
consts ast' :: 'a
consts bn_al :: "'b \<Rightarrow> 'a"
notation (latex output)
alpha_trm ("\<approx>\<^bsub>trm\<^esub>") and
fa_trm ("fa\<^bsub>trm\<^esub>") and
fa_trm_al ("fa\<AL>\<^bsub>trm\<^esub>") and
alpha_trm2 ("'(\<approx>\<^bsub>assn\<^esub>, \<approx>\<^bsub>trm\<^esub>')") and
fa_trm2 ("'(fa\<^bsub>assn\<^esub>, fa\<^bsub>trm\<^esub>')") and
fa_trm2_al ("'(fa\<AL>\<^bsub>assn\<^esub>, fa\<AL>\<^bsub>trm\<^esub>')") and
ast ("'(as, t')") and
ast' ("'(as', t\<PRIME> ')") and
equ2 ("'(=, =')") and
supp2 ("'(supp, supp')") and
bn_al ("bn\<^sup>\<alpha> _" [100] 100)
(*>*)
section {* Introduction *}
text {*
So far, Nominal Isabelle provided a mechanism for constructing alpha-equated
terms, for example lambda-terms
\[
@{text "t ::= x | t t | \<lambda>x. t"}
\]\smallskip
\noindent
where free and bound variables have names. For such alpha-equated terms,
Nominal Isabelle derives automatically a reasoning infrastructure that has
been used successfully in formalisations of an equivalence checking
algorithm for LF \cite{UrbanCheneyBerghofer08}, Typed
Scheme~\cite{TobinHochstadtFelleisen08}, several calculi for concurrency
\cite{BengtsonParow09} and a strong normalisation result for cut-elimination
in classical logic \cite{UrbanZhu08}. It has also been used by Pollack for
formalisations in the locally-nameless approach to binding
\cite{SatoPollack10}.
However, Nominal Isabelle has fared less well in a formalisation of the
algorithm W \cite{UrbanNipkow09}, where types and type-schemes are,
respectively, of the form
\begin{equation}\label{tysch}
\begin{array}{l}
@{text "T ::= x | T \<rightarrow> T"}\hspace{15mm}
@{text "S ::= \<forall>{x\<^isub>1,\<dots>, x\<^isub>n}. T"}
\end{array}
\end{equation}\smallskip
\noindent
and the @{text "\<forall>"}-quantification binds a finite (possibly empty) set of
type-variables. While it is possible to implement this kind of more general
binders by iterating single binders, this leads to a rather clumsy
formalisation of W.
{\bf add example about W}
The need of iterating single binders is also one reason
why Nominal Isabelle and similar theorem provers that only provide
mechanisms for binding single variables have not fared extremely well with
the more advanced tasks in the POPLmark challenge \cite{challenge05},
because also there one would like to bind multiple variables at once.
Binding multiple variables has interesting properties that cannot be captured
easily by iterating single binders. For example in the case of type-schemes we do not
want to make a distinction about the order of the bound variables. Therefore
we would like to regard in \eqref{ex1} below the first pair of type-schemes as alpha-equivalent,
but assuming that @{text x}, @{text y} and @{text z} are distinct variables,
the second pair should \emph{not} be alpha-equivalent:
\begin{equation}\label{ex1}
@{text "\<forall>{x, y}. x \<rightarrow> y \<approx>\<^isub>\<alpha> \<forall>{y, x}. y \<rightarrow> x"}\hspace{10mm}
@{text "\<forall>{x, y}. x \<rightarrow> y \<notapprox>\<^isub>\<alpha> \<forall>{z}. z \<rightarrow> z"}
\end{equation}\smallskip
\noindent
Moreover, we like to regard type-schemes as alpha-equivalent, if they differ
only on \emph{vacuous} binders, such as
\begin{equation}\label{ex3}
@{text "\<forall>{x}. x \<rightarrow> y \<approx>\<^isub>\<alpha> \<forall>{x, z}. x \<rightarrow> y"}
\end{equation}\smallskip
\noindent
where @{text z} does not occur freely in the type. In this paper we will
give a general binding mechanism and associated notion of alpha-equivalence
that can be used to faithfully represent this kind of binding in Nominal
Isabelle. The difficulty of finding the right notion for alpha-equivalence
can be appreciated in this case by considering that the definition given for
type-schemes by Leroy in \cite[Page 18--19]{Leroy92} is incorrect (it omits a side-condition).
However, the notion of alpha-equivalence that is preserved by vacuous
binders is not always wanted. For example in terms like
\begin{equation}\label{one}
@{text "\<LET> x = 3 \<AND> y = 2 \<IN> x - y \<END>"}
\end{equation}\smallskip
\noindent
we might not care in which order the assignments @{text "x = 3"} and
\mbox{@{text "y = 2"}} are given, but it would be often unusual (particularly
in strict languages) to regard \eqref{one} as alpha-equivalent with
\[
@{text "\<LET> x = 3 \<AND> y = 2 \<AND> z = foo \<IN> x - y \<END>"}
\]\smallskip
\noindent
Therefore we will also provide a separate binding mechanism for cases in
which the order of binders does not matter, but the `cardinality' of the
binders has to agree.
However, we found that this is still not sufficient for dealing with
language constructs frequently occurring in programming language
research. For example in @{text "\<LET>"}s containing patterns like
\begin{equation}\label{two}
@{text "\<LET> (x, y) = (3, 2) \<IN> x - y \<END>"}
\end{equation}\smallskip
\noindent
we want to bind all variables from the pattern inside the body of the
$\mathtt{let}$, but we also care about the order of these variables, since
we do not want to regard \eqref{two} as alpha-equivalent with
\[
@{text "\<LET> (y, x) = (3, 2) \<IN> x - y \<END>"}
\]\smallskip
\noindent
As a result, we provide three general binding mechanisms each of which binds
multiple variables at once, and let the user choose which one is intended
when formalising a term-calculus.
By providing these general binding mechanisms, however, we have to work
around a problem that has been pointed out by Pottier \cite{Pottier06} and
Cheney \cite{Cheney05}: in @{text "\<LET>"}-constructs of the form
\[
@{text "\<LET> x\<^isub>1 = t\<^isub>1 \<AND> \<dots> \<AND> x\<^isub>n = t\<^isub>n \<IN> s \<END>"}
\]\smallskip
\noindent
we care about the information that there are as many bound variables @{text
"x\<^isub>i"} as there are @{text "t\<^isub>i"}. We lose this information if
we represent the @{text "\<LET>"}-constructor by something like
\[
@{text "\<LET> (\<lambda>x\<^isub>1\<dots>x\<^isub>n . s) [t\<^isub>1,\<dots>,t\<^isub>n]"}
\]\smallskip
\noindent
where the notation @{text "\<lambda>_ . _"} indicates that the list of @{text
"x\<^isub>i"} becomes bound in @{text s}. In this representation the term
\mbox{@{text "\<LET> (\<lambda>x . s) [t\<^isub>1, t\<^isub>2]"}} is a perfectly
legal instance, but the lengths of the two lists do not agree. To exclude
such terms, additional predicates about well-formed terms are needed in
order to ensure that the two lists are of equal length. This can result in
very messy reasoning (see for example~\cite{BengtsonParow09}). To avoid
this, we will allow type specifications for @{text "\<LET>"}s as follows
\[
\mbox{\begin{tabular}{r@ {\hspace{2mm}}r@ {\hspace{2mm}}ll}
@{text trm} & @{text "::="} & @{text "\<dots>"} \\
& @{text "|"} & @{text "\<LET> as::assn s::trm"}\hspace{2mm}
\isacommand{binds} @{text "bn(as)"} \isacommand{in} @{text "s"}\\[1mm]
@{text assn} & @{text "::="} & @{text "\<ANIL>"}\\
& @{text "|"} & @{text "\<ACONS> name trm assn"}
\end{tabular}}
\]\smallskip
\noindent
where @{text assn} is an auxiliary type representing a list of assignments
and @{text bn} an auxiliary function identifying the variables to be bound
by the @{text "\<LET>"}. This function can be defined by recursion over @{text
assn} as follows
\[
@{text "bn(\<ANIL>) ="}~@{term "{}"} \hspace{10mm}
@{text "bn(\<ACONS> x t as) = {x} \<union> bn(as)"}
\]\smallskip
\noindent
The scope of the binding is indicated by labels given to the types, for
example @{text "s::trm"}, and a binding clause, in this case
\isacommand{binds} @{text "bn(as)"} \isacommand{in} @{text "s"}. This binding
clause states that all the names the function @{text "bn(as)"} returns
should be bound in @{text s}. This style of specifying terms and bindings
is heavily inspired by the syntax of the Ott-tool \cite{ott-jfp}. Our work
extends Ott in several aspects: one is that we support three binding
modes---Ott has only one, namely the one where the order of binders matters.
Another is that our reasoning infrastructure, like strong induction principles
and the notion of free variables, is derived from first principles within
the Isabelle/HOL theorem prover.
However, we will not be able to cope with all specifications that are
allowed by Ott. One reason is that Ott lets the user specify `empty' types
like \mbox{@{text "t ::= t t | \<lambda>x. t"}} where no clause for variables is
given. Arguably, such specifications make some sense in the context of Coq's
type theory (which Ott supports), but not at all in a HOL-based environment
where every datatype must have a non-empty set-theoretic model
\cite{Berghofer99}. Another reason is that we establish the reasoning
infrastructure for alpha-\emph{equated} terms. In contrast, Ott produces a
reasoning infrastructure in Isabelle/HOL for \emph{non}-alpha-equated, or
`raw', terms. While our alpha-equated terms and the `raw' terms produced by
Ott use names for bound variables, there is a key difference: working with
alpha-equated terms means, for example, that the two type-schemes
\[
@{text "\<forall>{x}. x \<rightarrow> y = \<forall>{x, z}. x \<rightarrow> y"}
\]\smallskip
\noindent
are not just alpha-equal, but actually \emph{equal}! As a result, we can
only support specifications that make sense on the level of alpha-equated
terms (offending specifications, which for example bind a variable according
to a variable bound somewhere else, are not excluded by Ott, but we have
to).
Our insistence on reasoning with alpha-equated terms comes from the
wealth of experience we gained with the older version of Nominal Isabelle:
for non-trivial properties, reasoning with alpha-equated terms is much
easier than reasoning with `raw' terms. The fundamental reason for this is
that the HOL-logic underlying Nominal Isabelle allows us to replace
`equals-by-equals'. In contrast, replacing
`alpha-equals-by-alpha-equals' in a representation based on `raw' terms
requires a lot of extra reasoning work.
Although in informal settings a reasoning infrastructure for alpha-equated
terms is nearly always taken for granted, establishing it automatically in
Isabelle/HOL is a rather non-trivial task. For every
specification we will need to construct type(s) containing as elements the
alpha-equated terms. To do so, we use the standard HOL-technique of defining
a new type by identifying a non-empty subset of an existing type. The
construction we perform in Isabelle/HOL can be illustrated by the following picture:
\begin{equation}\label{picture}
\mbox{\begin{tikzpicture}[scale=1.1]
%\draw[step=2mm] (-4,-1) grid (4,1);
\draw[very thick] (0.7,0.4) circle (4.25mm);
\draw[rounded corners=1mm, very thick] ( 0.0,-0.8) rectangle ( 1.8, 0.9);
\draw[rounded corners=1mm, very thick] (-1.95,0.85) rectangle (-2.85,-0.05);
\draw (-2.0, 0.845) -- (0.7,0.845);
\draw (-2.0,-0.045) -- (0.7,-0.045);
\draw ( 0.7, 0.5) node {\footnotesize\begin{tabular}{@ {}c@ {}}$\alpha$-\\[-1mm]classes\end{tabular}};
\draw (-2.4, 0.5) node {\footnotesize\begin{tabular}{@ {}c@ {}}$\alpha$-eq.\\[-1mm]terms\end{tabular}};
\draw (1.8, 0.48) node[right=-0.1mm]
{\small\begin{tabular}{@ {}l@ {}}existing\\[-1mm] type\\ (sets of raw terms)\end{tabular}};
\draw (0.9, -0.35) node {\footnotesize\begin{tabular}{@ {}l@ {}}non-empty\\[-1mm]subset\end{tabular}};
\draw (-3.25, 0.55) node {\small\begin{tabular}{@ {}l@ {}}new\\[-1mm]type\end{tabular}};
\draw[<->, very thick] (-1.8, 0.3) -- (-0.1,0.3);
\draw (-0.95, 0.3) node[above=0mm] {\footnotesize{}isomorphism};
\end{tikzpicture}}
\end{equation}\smallskip
\noindent
We take as the starting point a definition of raw terms (defined as a
datatype in Isabelle/HOL); then identify the alpha-equivalence classes in
the type of sets of raw terms according to our alpha-equivalence relation,
and finally define the new type as these alpha-equivalence classes (the
non-emptiness requirement is always satisfied whenever the raw terms are
definable as datatype in Isabelle/HOL and our relation for alpha-equivalence
is an equivalence relation).
The fact that we obtain an isomorphism between the new type and the
non-empty subset shows that the new type is a faithful representation of
alpha-equated terms. That is not the case for example for terms using the
locally nameless representation of binders \cite{McKinnaPollack99}: in this
representation there are `junk' terms that need to be excluded by
reasoning about a well-formedness predicate.
The problem with introducing a new type in Isabelle/HOL is that in order to
be useful, a reasoning infrastructure needs to be `lifted' from the
underlying subset to the new type. This is usually a tricky and arduous
task. To ease it, we re-implemented in Isabelle/HOL \cite{KaliszykUrban11}
the quotient package described by Homeier \cite{Homeier05} for the HOL4
system. This package allows us to lift definitions and theorems involving
raw terms to definitions and theorems involving alpha-equated terms. For
example if we define the free-variable function over raw lambda-terms
as follows
\[
\mbox{\begin{tabular}{l@ {\hspace{1mm}}r@ {\hspace{1mm}}l}
@{text "fv(x)"} & @{text "\<equiv>"} & @{text "{x}"}\\
@{text "fv(t\<^isub>1 t\<^isub>2)"} & @{text "\<equiv>"} & @{text "fv(t\<^isub>1) \<union> fv(t\<^isub>2)"}\\
@{text "fv(\<lambda>x.t)"} & @{text "\<equiv>"} & @{text "fv(t) - {x}"}
\end{tabular}}
\]\smallskip
\noindent
then with the help of the quotient package we can obtain a function @{text "fv\<^sup>\<alpha>"}
operating on quotients, that is alpha-equivalence classes of lambda-terms. This
lifted function is characterised by the equations
\[
\mbox{\begin{tabular}{l@ {\hspace{1mm}}r@ {\hspace{1mm}}l}
@{text "fv\<^sup>\<alpha>(x)"} & @{text "="} & @{text "{x}"}\\
@{text "fv\<^sup>\<alpha>(t\<^isub>1 t\<^isub>2)"} & @{text "="} & @{text "fv\<^sup>\<alpha>(t\<^isub>1) \<union> fv\<^sup>\<alpha>(t\<^isub>2)"}\\
@{text "fv\<^sup>\<alpha>(\<lambda>x.t)"} & @{text "="} & @{text "fv\<^sup>\<alpha>(t) - {x}"}
\end{tabular}}
\]\smallskip
\noindent
(Note that this means also the term-constructors for variables, applications
and lambda are lifted to the quotient level.) This construction, of course,
only works if alpha-equivalence is indeed an equivalence relation, and the
`raw' definitions and theorems are respectful w.r.t.~alpha-equivalence.
For example, we will not be able to lift a bound-variable function. Although
this function can be defined for raw terms, it does not respect
alpha-equivalence and therefore cannot be lifted.
To sum up, every lifting
of theorems to the quotient level needs proofs of some respectfulness
properties (see \cite{Homeier05}). In the paper we show that we are able to
automate these proofs and as a result can automatically establish a reasoning
infrastructure for alpha-equated terms.\smallskip
The examples we have in mind where our reasoning infrastructure will be
helpful include the term language of Core-Haskell (see
Figure~\ref{corehas}). This term language involves patterns that have lists
of type-, coercion- and term-variables, all of which are bound in @{text
"\<CASE>"}-expressions. In these patterns we do not know in advance how many
variables need to be bound. Another example is the algorithm W,
which includes multiple binders in type-schemes.\medskip
\noindent
{\bf Contributions:} We provide three new definitions for when terms
involving general binders are alpha-equivalent. These definitions are
inspired by earlier work of Pitts \cite{Pitts04}. By means of automatic
proofs, we establish a reasoning infrastructure for alpha-equated terms,
including properties about support, freshness and equality conditions for
alpha-equated terms. We are also able to automatically derive strong
induction principles that have the variable convention already built in.
For this we simplify the earlier automated proofs by using the proving tools
from the function package~\cite{Krauss09} of Isabelle/HOL. The method
behind our specification of general binders is taken from the Ott-tool, but
we introduce crucial restrictions, and also extensions, so that our
specifications make sense for reasoning about alpha-equated terms. The main
improvement over Ott is that we introduce three binding modes (only one is
present in Ott), provide formalised definitions for alpha-equivalence and
for free variables of our terms, and also derive a reasoning infrastructure
for our specifications from `first principles' inside a theorem prover.
\begin{figure}[t]
\begin{boxedminipage}{\linewidth}
\begin{center}
\begin{tabular}{@ {\hspace{8mm}}r@ {\hspace{2mm}}r@ {\hspace{2mm}}l}
\multicolumn{3}{@ {}l}{Type Kinds}\\
@{text "\<kappa>"} & @{text "::="} & @{text "\<star> | \<kappa>\<^isub>1 \<rightarrow> \<kappa>\<^isub>2"}\smallskip\\
\multicolumn{3}{@ {}l}{Coercion Kinds}\\
@{text "\<iota>"} & @{text "::="} & @{text "\<sigma>\<^isub>1 \<sim> \<sigma>\<^isub>2"}\smallskip\\
\multicolumn{3}{@ {}l}{Types}\\
@{text "\<sigma>"} & @{text "::="} & @{text "a | T | \<sigma>\<^isub>1 \<sigma>\<^isub>2 | S\<^isub>n"}$\;\overline{@{text "\<sigma>"}}$@{text "\<^sup>n"}
@{text "| \<forall>a:\<kappa>. \<sigma> | \<iota> \<Rightarrow> \<sigma>"}\smallskip\\
\multicolumn{3}{@ {}l}{Coercion Types}\\
@{text "\<gamma>"} & @{text "::="} & @{text "c | C | \<gamma>\<^isub>1 \<gamma>\<^isub>2 | S\<^isub>n"}$\;\overline{@{text "\<gamma>"}}$@{text "\<^sup>n"}
@{text "| \<forall>c:\<iota>. \<gamma> | \<iota> \<Rightarrow> \<gamma> | refl \<sigma> | sym \<gamma> | \<gamma>\<^isub>1 \<circ> \<gamma>\<^isub>2"}\\
& @{text "|"} & @{text "\<gamma> @ \<sigma> | left \<gamma> | right \<gamma> | \<gamma>\<^isub>1 \<sim> \<gamma>\<^isub>2 | rightc \<gamma> | leftc \<gamma> | \<gamma>\<^isub>1 \<triangleright> \<gamma>\<^isub>2"}\smallskip\\
\multicolumn{3}{@ {}l}{Terms}\\
@{text "e"} & @{text "::="} & @{text "x | K | \<Lambda>a:\<kappa>. e | \<Lambda>c:\<iota>. e | e \<sigma> | e \<gamma> | \<lambda>x:\<sigma>. e | e\<^isub>1 e\<^isub>2"}\\
& @{text "|"} & @{text "\<LET> x:\<sigma> = e\<^isub>1 \<IN> e\<^isub>2 | \<CASE> e\<^isub>1 \<OF>"}$\;\overline{@{text "p \<rightarrow> e\<^isub>2"}}$ @{text "| e \<triangleright> \<gamma>"}\smallskip\\
\multicolumn{3}{@ {}l}{Patterns}\\
@{text "p"} & @{text "::="} & @{text "K"}$\;\overline{@{text "a:\<kappa>"}}\;\overline{@{text "c:\<iota>"}}\;\overline{@{text "x:\<sigma>"}}$\smallskip\\
\multicolumn{3}{@ {}l}{Constants}\\
& @{text C} & coercion constants\\
& @{text T} & value type constructors\\
& @{text "S\<^isub>n"} & n-ary type functions (which need to be fully applied)\\
& @{text K} & data constructors\smallskip\\
\multicolumn{3}{@ {}l}{Variables}\\
& @{text a} & type variables\\
& @{text c} & coercion variables\\
& @{text x} & term variables\\
\end{tabular}
\end{center}
\end{boxedminipage}
\caption{The System @{text "F\<^isub>C"}
\cite{CoreHaskell}, also often referred to as \emph{Core-Haskell}. In this
version of @{text "F\<^isub>C"} we made a modification by separating the
grammars for type kinds and coercion kinds, as well as for types and coercion
types. For this paper the interesting term-constructor is @{text "\<CASE>"},
which binds multiple type-, coercion- and term-variables (the overlines stand for lists).\label{corehas}}
\end{figure}
*}
section {* A Short Review of the Nominal Logic Work *}
text {*
At its core, Nominal Isabelle is an adaptation of the nominal logic work by
Pitts \cite{Pitts03}. This adaptation for Isabelle/HOL is described in
\cite{HuffmanUrban10} (including proofs). We shall briefly review this work
to aid the description of what follows.
Two central notions in the nominal logic work are sorted atoms and
sort-respecting permutations of atoms. We will use the letters @{text "a, b,
c, \<dots>"} to stand for atoms and @{text "\<pi>, \<pi>\<^isub>1, \<dots>"} to stand for permutations,
which in Nominal Isabelle have type @{typ perm}. The purpose of atoms is to
represent variables, be they bound or free. The sorts of atoms can be used
to represent different kinds of variables, such as the term-, coercion- and
type-variables in Core-Haskell. It is assumed that there is an infinite
supply of atoms for each sort. In the interest of brevity, we shall restrict
ourselves in what follows to only one sort of atoms.
Permutations are bijective functions from atoms to atoms that are
the identity everywhere except on a finite number of atoms. There is a
two-place permutation operation written
@{text "_ \<bullet> _ :: perm \<Rightarrow> \<beta> \<Rightarrow> \<beta>"}
where the generic type @{text "\<beta>"} is the type of the object
over which the permutation
acts. In Nominal Isabelle, the identity permutation is written as @{term "0::perm"},
the composition of two permutations @{term "\<pi>\<^isub>1"} and @{term "\<pi>\<^isub>2"} as \mbox{@{term "\<pi>\<^isub>1 + \<pi>\<^isub>2"}},
and the inverse permutation of @{term "\<pi>"} as @{text "- \<pi>"}. The permutation
operation is defined over Isabelle/HOL's type-hierarchy \cite{HuffmanUrban10};
for example permutations acting on atoms, products, lists, permutations, sets,
functions and booleans are given by:
\begin{equation}\label{permute}
\mbox{\begin{tabular}{@ {}c@ {\hspace{10mm}}c@ {}}
\begin{tabular}{@ {}l@ {}}
@{text "\<pi> \<bullet> a \<equiv> \<pi> a"}\\
@{thm permute_prod.simps[where p="\<pi>", no_vars, THEN eq_reflection]}\\[2mm]
@{thm permute_list.simps(1)[where p="\<pi>", no_vars, THEN eq_reflection]}\\
@{thm permute_list.simps(2)[where p="\<pi>", no_vars, THEN eq_reflection]}\\
\end{tabular} &
\begin{tabular}{@ {}l@ {}}
@{thm permute_perm_def[where p="\<pi>" and q="\<pi>'", no_vars, THEN eq_reflection]}\\
@{thm permute_set_eq[where p="\<pi>", no_vars, THEN eq_reflection]}\\
@{text "\<pi> \<bullet> f \<equiv> \<lambda>x. \<pi> \<bullet> (f (- \<pi> \<bullet> x))"}\\
@{thm permute_bool_def[where p="\<pi>", no_vars, THEN eq_reflection]}
\end{tabular}
\end{tabular}}
\end{equation}\smallskip
\noindent
Concrete permutations in Nominal Isabelle are built up from swappings,
written as \mbox{@{text "(a b)"}}, which are permutations that behave
as follows:
\[
@{text "(a b) = \<lambda>c. if a = c then b else if b = c then a else c"}
\]\smallskip
The most original aspect of the nominal logic work of Pitts is a general
definition for the notion of the `set of free variables of an object @{text
"x"}'. This notion, written @{term "supp x"}, is general in the sense that
it applies not only to lambda-terms (alpha-equated or not), but also to lists,
products, sets and even functions. Its definition depends only on the
permutation operation and on the notion of equality defined for the type of
@{text x}, namely:
\begin{equation}\label{suppdef}
@{thm supp_def[no_vars, THEN eq_reflection]}
\end{equation}\smallskip
\noindent
There is also the derived notion for when an atom @{text a} is \emph{fresh}
for an @{text x}, defined as
\[
@{thm fresh_def[no_vars]}
\]\smallskip
\noindent
We use for sets of atoms the abbreviation
@{thm (lhs) fresh_star_def[no_vars]}, defined as
@{thm (rhs) fresh_star_def[no_vars]}.
A striking consequence of these definitions is that we can prove
without knowing anything about the structure of @{term x} that
swapping two fresh atoms, say @{text a} and @{text b}, leaves
@{text x} unchanged, namely
\begin{prop}\label{swapfreshfresh}
If @{thm (prem 1) swap_fresh_fresh[no_vars]} and @{thm (prem 2) swap_fresh_fresh[no_vars]}
then @{thm (concl) swap_fresh_fresh[no_vars]}.
\end{prop}
While often the support of an object can be relatively easily
described, for example for atoms, products, lists, function applications,
booleans and permutations as follows
\begin{equation}\label{supps}\mbox{
\begin{tabular}{c@ {\hspace{10mm}}c}
\begin{tabular}{rcl}
@{term "supp a"} & $=$ & @{term "{a}"}\\
@{term "supp (x, y)"} & $=$ & @{term "supp x \<union> supp y"}\\
@{term "supp []"} & $=$ & @{term "{}"}\\
@{term "supp (x#xs)"} & $=$ & @{term "supp x \<union> supp xs"}\\
\end{tabular}
&
\begin{tabular}{rcl}
@{text "supp (f x)"} & @{text "\<subseteq>"} & @{term "supp f \<union> supp x"}\\
@{term "supp b"} & $=$ & @{term "{}"}\\
@{term "supp \<pi>"} & $=$ & @{term "{a. \<pi> \<bullet> a \<noteq> a}"}
\end{tabular}
\end{tabular}}
\end{equation}\smallskip
\noindent
in some cases it can be difficult to characterise the support precisely, and
only an approximation can be established (as for function applications
above). Reasoning about such approximations can be simplified with the
notion \emph{supports}, defined as follows:
\begin{defi}
A set @{text S} \emph{supports} @{text x}, if for all atoms @{text a} and @{text b}
not in @{text S} we have @{term "(a \<rightleftharpoons> b) \<bullet> x = x"}.
\end{defi}
\noindent
The main point of @{text supports} is that we can establish the following
two properties.
\begin{prop}\label{supportsprop}
Given a set @{text "as"} of atoms.\\
{\it (i)} If @{thm (prem 1) supp_is_subset[where S="as", no_vars]}
and @{thm (prem 2) supp_is_subset[where S="as", no_vars]} then
@{thm (concl) supp_is_subset[where S="as", no_vars]}.\\
{\it (ii)} @{thm supp_supports[no_vars]}.
\end{prop}
Another important notion in the nominal logic work is \emph{equivariance}.
For a function @{text f}, say of type @{text "\<alpha> \<Rightarrow> \<beta>"}, to be equivariant
it is required that every permutation leaves @{text f} unchanged, that is
\begin{equation}\label{equivariancedef}
@{term "\<forall>\<pi>. \<pi> \<bullet> f = f"}
\end{equation}\smallskip
\noindent or equivalently that a permutation applied to the application
@{text "f x"} can be moved to the argument @{text x}. That means for equivariant
functions @{text f}, we have for all permutations @{text "\<pi>"}:
\begin{equation}\label{equivariance}
@{text "\<pi> \<bullet> f = f"} \;\;\;\;\textit{if and only if}\;\;\;\;
@{text "\<forall>x. \<pi> \<bullet> (f x) = f (\<pi> \<bullet> x)"}
\end{equation}\smallskip
\noindent
From property \eqref{equivariancedef} and the definition of @{text supp}, we
can easily deduce that equivariant functions have empty support. There is
also a similar notion for equivariant relations, say @{text R}, namely the property
that
\[
@{text "x R y"} \;\;\textit{implies}\;\; @{text "(\<pi> \<bullet> x) R (\<pi> \<bullet> y)"}
\]\smallskip
Using freshness, the nominal logic work provides us with general means for renaming
binders.
\noindent
While in the older version of Nominal Isabelle, we used extensively
Property~\ref{swapfreshfresh} to rename single binders, this property
proved too unwieldy for dealing with multiple binders. For such binders the
following generalisations turned out to be easier to use.
\begin{prop}\label{supppermeq}
@{thm[mode=IfThen] supp_perm_eq[where p="\<pi>", no_vars]}
\end{prop}
\begin{prop}\label{avoiding}
For a finite set @{text as} and a finitely supported @{text x} with
@{term "as \<sharp>* x"} and also a finitely supported @{text c}, there
exists a permutation @{text "\<pi>"} such that @{term "(\<pi> \<bullet> as) \<sharp>* c"} and
@{term "supp x \<sharp>* \<pi>"}.
\end{prop}
\noindent
The idea behind the second property is that given a finite set @{text as}
of binders (being bound, or fresh, in @{text x} is ensured by the
assumption @{term "as \<sharp>* x"}), then there exists a permutation @{text "\<pi>"} such that
the renamed binders @{term "\<pi> \<bullet> as"} avoid @{text c} (which can be arbitrarily chosen
as long as it is finitely supported) and also @{text "\<pi>"} does not affect anything
in the support of @{text x} (that is @{term "supp x \<sharp>* \<pi>"}). The last
fact and Property~\ref{supppermeq} allow us to `rename' just the binders
@{text as} in @{text x}, because @{term "\<pi> \<bullet> x = x"}.
Note that @{term "supp x \<sharp>* \<pi>"}
is equivalent with @{term "supp \<pi> \<sharp>* x"}, which means we could also formulate
Propositions \ref{supppermeq} and \ref{avoiding} in the other `direction'; however the
reasoning infrastructure of Nominal Isabelle is set up so that it provides more
automation for the formulation given above.
Most properties given in this section are described in detail in \cite{HuffmanUrban10}
and all are formalised in Isabelle/HOL. In the next sections we will make
use of these properties in order to define alpha-equivalence in
the presence of multiple binders.
*}
section {* Abstractions\label{sec:binders} *}
text {*
In Nominal Isabelle, the user is expected to write down a specification of a
term-calculus and then a reasoning infrastructure is automatically derived
from this specification (remember that Nominal Isabelle is a definitional
extension of Isabelle/HOL, which does not introduce any new axioms).
In order to keep our work with deriving the reasoning infrastructure
manageable, we will wherever possible state definitions and perform proofs
on the `user-level' of Isabelle/HOL, as opposed to writing custom ML-code that
generates them anew for each specification.
To that end, we will consider
first pairs @{text "(as, x)"} of type @{text "(atom set) \<times> \<beta>"}. These pairs
are intended to represent the abstraction, or binding, of the set of atoms @{text
"as"} in the body @{text "x"}.
The first question we have to answer is when two pairs @{text "(as, x)"} and
@{text "(bs, y)"} are alpha-equivalent? (For the moment we are interested in
the notion of alpha-equivalence that is \emph{not} preserved by adding
vacuous binders.) To answer this question, we identify four conditions: {\it (i)}
given a free-atom function @{text "fa"} of type \mbox{@{text "\<beta> \<Rightarrow> atom
set"}}, then @{text x} and @{text y} need to have the same set of free
atoms; moreover there must be a permutation @{text \<pi>} such that {\it
(ii)} @{text \<pi>} leaves the free atoms of @{text x} and @{text y} unchanged, but
{\it (iii)} `moves' their bound names so that we obtain modulo a relation,
say \mbox{@{text "_ R _"}}, two equivalent terms. We also require that {\it (iv)}
@{text \<pi>} makes the sets of abstracted atoms @{text as} and @{text bs} equal. The
requirements {\it (i)} to {\it (iv)} can be stated formally as:
\begin{defi}[Alpha-Equivalence for Set-Bindings]\label{alphaset}\mbox{}\\
\begin{tabular}{@ {\hspace{10mm}}l@ {\hspace{5mm}}rl}
@{term "alpha_set_ex (as, x) R fa (bs, y)"}\hspace{2mm}@{text "\<equiv>"} &
\multicolumn{2}{@ {}l}{if there exists a @{text "\<pi>"} such that:}\\
& \mbox{\it (i)} & @{term "fa(x) - as = fa(y) - bs"}\\
& \mbox{\it (ii)} & @{term "(fa(x) - as) \<sharp>* \<pi>"}\\
& \mbox{\it (iii)} & @{text "(\<pi> \<bullet> x) R y"} \\
& \mbox{\it (iv)} & @{term "(\<pi> \<bullet> as) = bs"} \\
\end{tabular}
\end{defi}
\noindent
Note that the relation is
dependent on a free-atom function @{text "fa"} and a relation @{text
"R"}. The reason for this extra generality is that we will use
$\approx_{\,\textit{set}}^{\textit{R}, \textit{fa}}$ for both raw terms and
alpha-equated terms. In
the latter case, @{text R} will be replaced by equality @{text "="} and we
will prove that @{text "fa"} is equal to @{text "supp"}.
Definition \ref{alphaset} does not make any distinction between the
order of abstracted atoms. If we want this, then we can define alpha-equivalence
for pairs of the form \mbox{@{text "(as, x)"}} with type @{text "(atom list) \<times> \<beta>"}
as follows
\begin{defi}[Alpha-Equivalence for List-Bindings]\label{alphalist}\mbox{}\\
\begin{tabular}{@ {\hspace{10mm}}l@ {\hspace{5mm}}rl}
@{term "alpha_lst_ex (as, x) R fa (bs, y)"}\hspace{2mm}@{text "\<equiv>"} &
\multicolumn{2}{@ {}l}{if there exists a @{text "\<pi>"} such that:}\\
& \mbox{\it (i)} & @{term "fa(x) - (set as) = fa(y) - (set bs)"}\\
& \mbox{\it (ii)} & @{term "(fa(x) - set as) \<sharp>* \<pi>"}\\
& \mbox{\it (iii)} & @{text "(\<pi> \<bullet> x) R y"}\\
& \mbox{\it (iv)} & @{term "(\<pi> \<bullet> as) = bs"}\\
\end{tabular}
\end{defi}
\noindent
where @{term set} is the function that coerces a list of atoms into a set of atoms.
Now the last clause ensures that the order of the binders matters (since @{text as}
and @{text bs} are lists of atoms).
If we do not want to make any difference between the order of binders \emph{and}
also allow vacuous binders, that means according to Pitts~\cite{Pitts04}
\emph{restrict} atoms, then we keep sets of binders, but drop
condition {\it (iv)} in Definition~\ref{alphaset}:
\begin{defi}[Alpha-Equivalence for Set+-Bindings]\label{alphares}\mbox{}\\
\begin{tabular}{@ {\hspace{10mm}}l@ {\hspace{5mm}}rl}
@{term "alpha_res_ex (as, x) R fa (bs, y)"}\hspace{2mm}@{text "\<equiv>"} &
\multicolumn{2}{@ {}l}{if there exists a @{text "\<pi>"} such that:}\\
& \mbox{\it (i)} & @{term "fa(x) - as = fa(y) - bs"}\\
& \mbox{\it (ii)} & @{term "(fa(x) - as) \<sharp>* \<pi>"}\\
& \mbox{\it (iii)} & @{text "(\<pi> \<bullet> x) R y"}\\
\end{tabular}
\end{defi}
It might be useful to consider first some examples how these definitions
of alpha-equivalence pan out in practice. For this consider the case of
abstracting a set of atoms over types (as in type-schemes). We set
@{text R} to be the usual equality @{text "="} and for @{text "fa(T)"} we
define
\[
@{text "fa(x) \<equiv> {x}"} \hspace{10mm} @{text "fa(T\<^isub>1 \<rightarrow> T\<^isub>2) \<equiv> fa(T\<^isub>1) \<union> fa(T\<^isub>2)"}
\]\smallskip
\noindent
Now recall the examples shown in \eqref{ex1} and
\eqref{ex3}. It can be easily checked that @{text "({x, y}, x \<rightarrow> y)"} and
@{text "({y, x}, y \<rightarrow> x)"} are alpha-equivalent according to
$\approx_{\,\textit{set}}$ and $\approx_{\,\textit{set+}}$ by taking @{text "\<pi>"} to
be the swapping @{term "(x \<rightleftharpoons> y)"}. In case of @{text "x \<noteq> y"}, then @{text
"([x, y], x \<rightarrow> y)"} $\not\approx_{\,\textit{list}}$ @{text "([y, x], x \<rightarrow> y)"}
since there is no permutation that makes the lists @{text "[x, y]"} and
@{text "[y, x]"} equal, and also leaves the type \mbox{@{text "x \<rightarrow> y"}}
unchanged. Another example is @{text "({x}, x)"} $\approx_{\,\textit{set+}}$
@{text "({x, y}, x)"} which holds by taking @{text "\<pi>"} to be the identity
permutation. However, if @{text "x \<noteq> y"}, then @{text "({x}, x)"}
$\not\approx_{\,\textit{set}}$ @{text "({x, y}, x)"} since there is no
permutation that makes the sets @{text "{x}"} and @{text "{x, y}"} equal
(similarly for $\approx_{\,\textit{list}}$). It can also relatively easily be
shown that all three notions of alpha-equivalence coincide, if we only
abstract a single atom. In this case they also agree with the alpha-equivalence
used in older versions of Nominal Isabelle \cite{Urban08}.\footnote{We omit a
proof of this fact since the details are hairy and not really important for the
purpose of this paper.}
In the rest of this section we are going to show that the alpha-equivalences
really lead to abstractions where some atoms are bound (or more precisely
removed from the support). For this we will consider three abstraction
types that are quotients of the relations
\begin{equation}
\begin{array}{r}
@{term "alpha_set_ex (as, x) equal supp (bs, y)"}\smallskip\\
@{term "alpha_res_ex (as, x) equal supp (bs, y)"}\smallskip\\
@{term "alpha_lst_ex (as, x) equal supp (bs, y)"}\\
\end{array}
\end{equation}\smallskip
\noindent
Note that in these relations we replaced the free-atom function @{text "fa"}
with @{term "supp"} and the relation @{text R} with equality. We can show
the following two properties:
\begin{lem}\label{alphaeq}
The relations $\approx_{\,\textit{set}}^{=, \textit{supp}}$,
$\approx_{\,\textit{set+}}^{=, \textit{supp}}$
and $\approx_{\,\textit{list}}^{=, \textit{supp}}$ are
equivalence relations and equivariant.
\end{lem}
\begin{proof}
Reflexivity is by taking @{text "\<pi>"} to be @{text "0"}. For symmetry we have
a permutation @{text "\<pi>"} and for the proof obligation take @{term "-
\<pi>"}. In case of transitivity, we have two permutations @{text "\<pi>\<^isub>1"}
and @{text "\<pi>\<^isub>2"}, and for the proof obligation use @{text
"\<pi>\<^isub>1 + \<pi>\<^isub>2"}. Equivariance means @{term "alpha_set_ex (\<pi> \<bullet> as,
\<pi> \<bullet> x) equal supp (\<pi> \<bullet> bs, \<pi> \<bullet> y)"} holds provided \mbox{@{term
"alpha_set_ex (as, x) equal supp(bs, y)"}} holds. From the assumption we
have a permutation @{text "\<pi>'"} and for the proof obligation use @{text "\<pi> \<bullet>
\<pi>'"}. To show equivariance, we need to `pull out' the permutations,
which is possible since all operators, namely as @{text "#\<^sup>*, -, =, \<bullet>,
set"} and @{text "supp"}, are equivariant (see
\cite{HuffmanUrban10}). Finally, we apply the permutation operation on
booleans.
\end{proof}
\noindent
Recall the picture shown in \eqref{picture} about new types in HOL.
The lemma above allows us to use our quotient package for introducing
new types @{text "\<beta> abs\<^bsub>set\<^esub>"}, @{text "\<beta> abs\<^bsub>set+\<^esub>"} and @{text "\<beta> abs\<^bsub>list\<^esub>"}
representing alpha-equivalence classes of pairs of type
@{text "(atom set) \<times> \<beta>"} (in the first two cases) and of type @{text "(atom list) \<times> \<beta>"}
(in the third case).
The elements in these types will be, respectively, written as
\[
@{term "Abs_set as x"} \hspace{10mm}
@{term "Abs_res as x"} \hspace{10mm}
@{term "Abs_lst as x"}
\]\smallskip
\noindent
indicating that a set (or list) of atoms @{text as} is abstracted in @{text x}. We will
call the types \emph{abstraction types} and their elements
\emph{abstractions}. The important property we need to derive is the support of
abstractions, namely:
\begin{thm}[Support of Abstractions]\label{suppabs}
Assuming @{text x} has finite support, then
\[
\begin{array}{l@ {\;=\;}l}
@{thm (lhs) supp_Abs(1)[no_vars]} & @{thm (rhs) supp_Abs(1)[no_vars]}\\
@{thm (lhs) supp_Abs(2)[no_vars]} & @{thm (rhs) supp_Abs(2)[no_vars]}\\
@{thm (lhs) supp_Abs(3)[where bs="as", no_vars]} &
@{thm (rhs) supp_Abs(3)[where bs="as", no_vars]}\\
\end{array}
\]\smallskip
\end{thm}
\noindent
In effect, this theorem states that the atoms @{text "as"} are bound in the
abstraction. As stated earlier, this can be seen as a litmus test that our
Definitions \ref{alphaset}, \ref{alphalist} and \ref{alphares} capture the
idea of alpha-equivalence relations. Below we will give the proof for the
first equation of Theorem \ref{suppabs}. The others follow by similar
arguments. By definition of the abstraction type @{text
"abs\<^bsub>set\<^esub>"} we have
\begin{equation}\label{abseqiff}
@{thm (lhs) Abs_eq_iff(1)[where bs="as" and bs'="bs", no_vars]} \;\;\;\text{if and only if}\;\;\;
@{term "alpha_set_ex (as, x) equal supp (bs, y)"}
\end{equation}\smallskip
\noindent
and also
\begin{equation}\label{absperm}
@{thm permute_Abs(1)[where p="\<pi>", no_vars]}
\end{equation}\smallskip
\noindent
The second fact derives from the definition of permutations acting on pairs
\eqref{permute} and alpha-equivalence being equivariant
(see Lemma~\ref{alphaeq}). With these two facts at our disposal, we can show
the following lemma about swapping two atoms in an abstraction.
\begin{lem}
If @{thm (prem 1) Abs_swap1(1)[where bs="as", no_vars]} and
@{thm (prem 2) Abs_swap1(1)[where bs="as", no_vars]} then
@{thm (concl) Abs_swap1(1)[where bs="as", no_vars]}
\end{lem}
\begin{proof}
If @{term "a = b"} the lemma is immediate, since @{term "(a \<rightleftharpoons> b)"} is then
the identity permutation.
Also in the other case the lemma is straightforward using \eqref{abseqiff}
and observing that the assumptions give us @{term "(a \<rightleftharpoons> b) \<bullet> (supp x - as) =
(supp x - as)"}. We therefore can use the swapping @{term "(a \<rightleftharpoons> b)"} as
the permutation for the proof obligation.
\end{proof}
\noindent
Assuming that @{text "x"} has finite support, this lemma together
with \eqref{absperm} allows us to show
\begin{equation}\label{halfone}
@{thm Abs_supports(1)[no_vars]}
\end{equation}\smallskip
\noindent
which by Property~\ref{supportsprop} gives us `one half' of
Theorem~\ref{suppabs}. The `other half' is a bit more involved. To establish
it, we use a trick from \cite{Pitts04} and first define an auxiliary
function @{text aux}, taking an abstraction as argument
\[
@{thm supp_set.simps[THEN eq_reflection, no_vars]}
\]\smallskip
\noindent
Using the second equation in \eqref{equivariance}, we can show that
@{text "aux"} is equivariant (since @{term "\<pi> \<bullet> (supp x - as) = (supp (\<pi> \<bullet> x)) - (\<pi> \<bullet> as)"})
and therefore has empty support.
This in turn means
\[
@{term "supp (supp_set (Abs_set as x)) \<subseteq> supp (Abs_set as x)"}
\]\smallskip
\noindent
using the fact about the support of function applications in \eqref{supps}. Assuming
@{term "supp x - as"} is a finite set, we further obtain
\begin{equation}\label{halftwo}
@{thm (concl) Abs_supp_subset1(1)[no_vars]}
\end{equation}\smallskip
\noindent
This is because for every finite set of atoms, say @{text "bs"}, we have
@{thm (concl) supp_finite_atom_set[where S="bs", no_vars]}.
Finally, taking \eqref{halfone} and \eqref{halftwo} together establishes
the first equation of Theorem~\ref{suppabs}. The others are similar.
Recall the definition of support given in \eqref{suppdef}, and note the difference between
the support of a raw pair and an abstraction
\[
@{term "supp (as, x) = supp as \<union> supp x"}\hspace{15mm}
@{term "supp (Abs_set as x) = supp x - as"}
\]\smallskip
\noindent
While the permutation operations behave in both cases the same (a permutation
is just moved to the arguments), the notion of equality is different for pairs and
abstractions. Therefore we have different supports. In case of abstractions,
we have established in Theorem~\ref{suppabs} that bound atoms are removed from
the support of the abstractions' bodies.
The method of first considering abstractions of the form @{term "Abs_set as
x"} etc is motivated by the fact that we can conveniently establish at the
Isabelle/HOL level properties about them. It would be extremely laborious
to write custom ML-code that derives automatically such properties for every
term-constructor that binds some atoms. Also the generality of the
definitions for alpha-equivalence will help us in the next sections.
*}
section {* Specifying General Bindings\label{sec:spec} *}
text {*
Our choice of syntax for specifications is influenced by the existing
datatype package of Isabelle/HOL \cite{Berghofer99}
and by the syntax of the
Ott-tool \cite{ott-jfp}. For us a specification of a term-calculus is a
collection of (possibly mutually recursive) type declarations, say @{text
"ty\<AL>\<^isub>1, \<dots>, ty\<AL>\<^isub>n"}, and an associated collection of
binding functions, say @{text "bn\<AL>\<^isub>1, \<dots>, bn\<AL>\<^isub>m"}. The
syntax in Nominal Isabelle for such specifications is schematically as follows:
\begin{equation}\label{scheme}
\mbox{\begin{tabular}{@ {}p{2.5cm}l}
type \mbox{declaration part} &
$\begin{cases}
\mbox{\begin{tabular}{l}
\isacommand{nominal\_datatype} @{text "ty\<AL>\<^isub>1 = \<dots>"}\\
\isacommand{and} @{text "ty\<AL>\<^isub>2 = \<dots>"}\\
\raisebox{2mm}{$\ldots$}\\[-2mm]
\isacommand{and} @{text "ty\<AL>\<^isub>n = \<dots>"}\\
\end{tabular}}
\end{cases}$\\[2mm]
binding \mbox{function part} &
$\begin{cases}
\mbox{\begin{tabular}{l}
\isacommand{binder} @{text "bn\<AL>\<^isub>1"} \isacommand{and} \ldots \isacommand{and} @{text "bn\<AL>\<^isub>m"}\\
\isacommand{where}\\
\raisebox{2mm}{$\ldots$}\\[-2mm]
\end{tabular}}
\end{cases}$\\
\end{tabular}}
\end{equation}\smallskip
\noindent
Every type declaration @{text ty}$^\alpha_{1..n}$ consists of a collection
of term-constructors, each of which comes with a list of labelled types that
stand for the types of the arguments of the term-constructor. For example a
term-constructor @{text "C\<^sup>\<alpha>"} might be specified with
\[
@{text "C\<^sup>\<alpha> label\<^isub>1::ty"}\mbox{$'_1$} @{text "\<dots> label\<^isub>l::ty"}\mbox{$'_l\;\;\;\;\;$}
@{text "binding_clauses"}
\]\smallskip
\noindent
whereby some of the @{text ty}$'_{1..l}$ (or their components) can be
contained in the collection of @{text ty}$^\alpha_{1..n}$ declared in
\eqref{scheme}. In this case we will call the corresponding argument a
\emph{recursive argument} of @{text "C\<^sup>\<alpha>"}. The types of such
recursive arguments need to satisfy a `positivity' restriction, which
ensures that the type has a set-theoretic semantics (see
\cite{Berghofer99}). The labels annotated on the types are optional. Their
purpose is to be used in the (possibly empty) list of \emph{binding
clauses}, which indicate the binders and their scope in a term-constructor.
They come in three \emph{modes}:
\[\mbox{
\begin{tabular}{@ {}l@ {}}
\isacommand{binds} {\it binders} \isacommand{in} {\it bodies}\\
\isacommand{binds (set)} {\it binders} \isacommand{in} {\it bodies}\\
\isacommand{binds (set+)} {\it binders} \isacommand{in} {\it bodies}
\end{tabular}}
\]\smallskip
\noindent
The first mode is for binding lists of atoms (the order of bound atoms
matters); the second is for sets of binders (the order does not matter, but
the cardinality does) and the last is for sets of binders (with vacuous
binders preserving alpha-equivalence). As indicated, the labels in the
`\isacommand{in}-part' of a binding clause will be called \emph{bodies};
the `\isacommand{binds}-part' will be called \emph{binders}. In contrast to
Ott, we allow multiple labels in binders and bodies. For example we allow
binding clauses of the form:
\[\mbox{
\begin{tabular}{@ {}ll@ {}}
@{text "Foo\<^isub>1 x::name y::name t::trm s::trm"} &
\isacommand{binds} @{text "x y"} \isacommand{in} @{text "t s"}\\
@{text "Foo\<^isub>2 x::name y::name t::trm s::trm"} &
\isacommand{binds} @{text "x y"} \isacommand{in} @{text "t"},
\isacommand{binds} @{text "x y"} \isacommand{in} @{text "s"}\\
\end{tabular}}
\]\smallskip
\noindent
Similarly for the other binding modes. Interestingly, in case of
\isacommand{binds (set)} and \isacommand{binds (set+)} the binding clauses
above will make a difference to the semantics of the specifications (the
corresponding alpha-equivalence will differ). We will show this later with
an example.
There are also some restrictions we need to impose on our binding clauses in
comparison to Ott. The main idea behind these restrictions is
that we obtain a notion of alpha-equivalence where it is ensured
that within a given scope an atom occurrence cannot be both bound and free
at the same time. The first restriction is that a body can only occur in
\emph{one} binding clause of a term constructor. So for example
\[\mbox{
@{text "Foo x::name y::name t::trm"}\hspace{3mm}
\isacommand{binds} @{text "x"} \isacommand{in} @{text "t"},
\isacommand{binds} @{text "y"} \isacommand{in} @{text "t"}}
\]\smallskip
\noindent
is not allowed. This ensures that the bound atoms of a body cannot be free
at the same time by specifying an alternative binder for the same body.
For binders we distinguish between \emph{shallow} and \emph{deep} binders.
Shallow binders are just labels. The restriction we need to impose on them
is that in case of \isacommand{binds (set)} and \isacommand{binds (set+)} the
labels must either refer to atom types or to sets of atom types; in case of
\isacommand{binds} the labels must refer to atom types or to lists of atom
types. Two examples for the use of shallow binders are the specification of
lambda-terms, where a single name is bound, and type-schemes, where a finite
set of names is bound:
\[\mbox{
\begin{tabular}{@ {}c@ {\hspace{8mm}}c@ {}}
\begin{tabular}{@ {}l}
\isacommand{nominal\_datatype} @{text lam} $=$\\
\hspace{2mm}\phantom{$\mid$}~@{text "Var name"}\\
\hspace{2mm}$\mid$~@{text "App lam lam"}\\
\hspace{2mm}$\mid$~@{text "Lam x::name t::lam"}\hspace{3mm}%
\isacommand{binds} @{text x} \isacommand{in} @{text t}\\
\\
\end{tabular} &
\begin{tabular}{@ {}l@ {}}
\isacommand{nominal\_datatype}~@{text ty} $=$\\
\hspace{2mm}\phantom{$\mid$}~@{text "TVar name"}\\
\hspace{2mm}$\mid$~@{text "TFun ty ty"}\\
\isacommand{and}~@{text "tsc ="}\\
\hspace{2mm}\phantom{$\mid$}~@{text "TAll xs::(name fset) T::ty"}\hspace{3mm}%
\isacommand{binds (set+)} @{text xs} \isacommand{in} @{text T}\\
\end{tabular}
\end{tabular}}
\]\smallskip
\noindent
In these specifications @{text "name"} refers to a (concrete) atom type, and @{text
"fset"} to the type of finite sets. Note that for @{text Lam} it does not
matter which binding mode we use. The reason is that we bind only a single
@{text name}, in which case all three binding modes coincide. However, having
\isacommand{binds (set)} or just \isacommand{binds}
in the second case makes a difference to the semantics of the specification
(which we will define in the next section).
A \emph{deep} binder uses an auxiliary binding function that `picks' out
the atoms in one argument of the term-constructor, which can be bound in
other arguments and also in the same argument (we will call such binders
\emph{recursive}, see below). The binding functions are
expected to return either a set of atoms (for \isacommand{binds (set)} and
\isacommand{binds (set+)}) or a list of atoms (for \isacommand{binds}). They need
to be defined by recursion over the corresponding type; the equations
must be given in the binding function part of the scheme shown in
\eqref{scheme}. For example a term-calculus containing @{text "Let"}s with
tuple patterns may be specified as:
\begin{equation}\label{letpat}
\mbox{%
\begin{tabular}{l}
\isacommand{nominal\_datatype} @{text trm} $=$\\
\hspace{5mm}\phantom{$\mid$}~@{term "Var name"}\\
\hspace{5mm}$\mid$~@{term "App trm trm"}\\
\hspace{5mm}$\mid$~@{text "Lam x::name t::trm"}
\;\;\isacommand{binds} @{text x} \isacommand{in} @{text t}\\
\hspace{5mm}$\mid$~@{text "Let_pat p::pat trm t::trm"}
\;\;\isacommand{binds} @{text "bn(p)"} \isacommand{in} @{text t}\\
\isacommand{and} @{text pat} $=$\\
\hspace{5mm}\phantom{$\mid$}~@{text "PVar name"}\\
\hspace{5mm}$\mid$~@{text "PTup pat pat"}\\
\isacommand{binder}~@{text "bn::pat \<Rightarrow> atom list"}\\
\isacommand{where}~@{text "bn(PVar x) = [atom x]"}\\
\hspace{5mm}$\mid$~@{text "bn(PTup p\<^isub>1 p\<^isub>2) = bn(p\<^isub>1) @ bn(p\<^isub>2)"}\smallskip\\
\end{tabular}}
\end{equation}\smallskip
\noindent
In this specification the function @{text "bn"} determines which atoms of
the pattern @{text p} (fifth line) are bound in the argument @{text "t"}. Note that in the
second-last @{text bn}-clause the function @{text "atom"} coerces a name
into the generic atom type of Nominal Isabelle \cite{HuffmanUrban10}. This
allows us to treat binders of different atom type uniformly.
For deep binders we allow binding clauses such as
\[\mbox{
\begin{tabular}{ll}
@{text "Bar p::pat t::trm"} &
\isacommand{binds} @{text "bn(p)"} \isacommand{in} @{text "p t"} \\
\end{tabular}}
\]\smallskip
\noindent
where the argument of the deep binder also occurs in the body. We call such
binders \emph{recursive}. To see the purpose of such recursive binders,
compare `plain' @{text "Let"}s and @{text "Let_rec"}s in the following
specification:
\begin{equation}\label{letrecs}
\mbox{%
\begin{tabular}{@ {}l@ {}l}
\isacommand{nominal\_datatype}~@{text "trm ="}~\ldots\\
\hspace{5mm}$\mid$~@{text "Let as::assn t::trm"}
& \hspace{-19mm}\isacommand{binds} @{text "bn(as)"} \isacommand{in} @{text t}\\
\hspace{5mm}$\mid$~@{text "Let_rec as::assn t::trm"}
& \hspace{-19mm}\isacommand{binds} @{text "bn(as)"} \isacommand{in} @{text "as t"}\\
\isacommand{and} @{text "assn"} $=$\\
\hspace{5mm}\phantom{$\mid$}~@{text "ANil"}\\
\hspace{5mm}$\mid$~@{text "ACons name trm assn"}\\
\isacommand{binder} @{text "bn::assn \<Rightarrow> atom list"}\\
\isacommand{where}~@{text "bn(ANil) = []"}\\
\hspace{5mm}$\mid$~@{text "bn(ACons a t as) = [atom a] @ bn(as)"}\\
\end{tabular}}
\end{equation}\smallskip
\noindent
The difference is that with @{text Let} we only want to bind the atoms @{text
"bn(as)"} in the term @{text t}, but with @{text "Let_rec"} we also want to bind the atoms
inside the assignment. This difference has consequences for the associated
notions of free-atoms and alpha-equivalence.
To make sure that atoms bound by deep binders cannot be free at the
same time, we cannot have more than one binding function for a deep binder.
Consequently we exclude specifications such as
\[\mbox{
\begin{tabular}{@ {}l@ {\hspace{2mm}}l@ {}}
@{text "Baz\<^isub>1 p::pat t::trm"} &
\isacommand{binds} @{text "bn\<^isub>1(p) bn\<^isub>2(p)"} \isacommand{in} @{text t}\\
@{text "Baz\<^isub>2 p::pat t\<^isub>1::trm t\<^isub>2::trm"} &
\isacommand{binds} @{text "bn\<^isub>1(p)"} \isacommand{in} @{text "t\<^isub>1"},
\isacommand{binds} @{text "bn\<^isub>2(p)"} \isacommand{in} @{text "t\<^isub>2"}\\
\end{tabular}}
\]\smallskip
\noindent
Otherwise it is possible that @{text "bn\<^isub>1"} and @{text "bn\<^isub>2"} pick
out different atoms to become bound, respectively be free, in @{text "p"}.
(Since the Ott-tool does not derive a reasoning infrastructure for
alpha-equated terms with deep binders, it can permit such specifications.)
We also need to restrict the form of the binding functions in order to
ensure the @{text "bn"}-functions can be defined for alpha-equated
terms. The main restriction is that we cannot return an atom in a binding
function that is also bound in the corresponding term-constructor.
Consider again the specification for @{text "trm"} and a contrived
version for assignments @{text "assn"}:
\begin{equation}\label{bnexp}
\mbox{%
\begin{tabular}{@ {}l@ {}}
\isacommand{nominal\_datatype}~@{text "trm ="}~\ldots\\
\isacommand{and} @{text "assn"} $=$\\
\hspace{5mm}\phantom{$\mid$}~@{text "ANil'"}\\
\hspace{5mm}$\mid$~@{text "ACons' x::name y::name t::trm assn"}
\;\;\isacommand{binds} @{text "y"} \isacommand{in} @{text t}\\
\isacommand{binder} @{text "bn::assn \<Rightarrow> atom list"}\\
\isacommand{where}~@{text "bn(ANil') = []"}\\
\hspace{5mm}$\mid$~@{text "bn(ACons' x y t as) = [atom x] @ bn(as)"}\\
\end{tabular}}
\end{equation}\smallskip
\noindent
In this example the term constructor @{text "ACons'"} has four arguments with
a binding clause involving two of them. This constructor is also used in the definition
of the binding function. The restriction we have to impose is that the
binding function can only return free atoms, that is the ones that are \emph{not}
mentioned in a binding clause. Therefore @{text "y"} cannot be used in the
binding function @{text "bn"} (since it is bound in @{text "ACons'"} by the
binding clause), but @{text x} can (since it is a free atom). This
restriction is sufficient for lifting the binding function to alpha-equated
terms. If we would permit @{text "bn"} to return @{text "y"},
then it would not be respectful and therefore cannot be lifted to
alpha-equated lambda-terms.
In the version of Nominal Isabelle described here, we also adopted the
restriction from the Ott-tool that binding functions can only return: the
empty set or empty list (as in case @{text ANil'}), a singleton set or
singleton list containing an atom (case @{text PVar} in \eqref{letpat}), or
unions of atom sets or appended atom lists (case @{text ACons'}). This
restriction will simplify some automatic definitions and proofs later on.
In order to simplify our definitions of free atoms and alpha-equivalence,
we shall assume specifications
of term-calculi are implicitly \emph{completed}. By this we mean that
for every argument of a term-constructor that is \emph{not}
already part of a binding clause given by the user, we add implicitly a special \emph{empty} binding
clause, written \isacommand{binds}~@{term "{}"}~\isacommand{in}~@{text "labels"}. In case
of the lambda-terms, the completion produces
\[\mbox{
\begin{tabular}{@ {}l@ {\hspace{-1mm}}}
\isacommand{nominal\_datatype} @{text lam} =\\
\hspace{5mm}\phantom{$\mid$}~@{text "Var x::name"}
\;\;\isacommand{binds}~@{term "{}"}~\isacommand{in}~@{text "x"}\\
\hspace{5mm}$\mid$~@{text "App t\<^isub>1::lam t\<^isub>2::lam"}
\;\;\isacommand{binds}~@{term "{}"}~\isacommand{in}~@{text "t\<^isub>1 t\<^isub>2"}\\
\hspace{5mm}$\mid$~@{text "Lam x::name t::lam"}
\;\;\isacommand{binds}~@{text x} \isacommand{in} @{text t}\\
\end{tabular}}
\]\smallskip
\noindent
The point of completion is that we can make definitions over the binding
clauses and be sure to have captured all arguments of a term constructor.
*}
section {* Alpha-Equivalence and Free Atoms\label{sec:alpha} *}
text {*
Having dealt with all syntax matters, the problem now is how we can turn
specifications into actual type definitions in Isabelle/HOL and then
establish a reasoning infrastructure for them. As Pottier and Cheney pointed
out \cite{Cheney05,Pottier06}, just re-arranging the arguments of
term-constructors so that binders and their bodies are next to each other
will result in inadequate representations in cases like \mbox{@{text "Let
x\<^isub>1 = t\<^isub>1\<dots>x\<^isub>n = t\<^isub>n in s"}}. Therefore we will
first extract `raw' datatype definitions from the specification and then
define explicitly an alpha-equivalence relation over them. We subsequently
construct the quotient of the datatypes according to our alpha-equivalence.
The `raw' datatype definition can be obtained by stripping off the
binding clauses and the labels from the types given by the user. We also have to invent
new names for the types @{text "ty\<^sup>\<alpha>"} and the term-constructors @{text "C\<^sup>\<alpha>"}.
In our implementation we just use the affix ``@{text "_raw"}''.
But for the purpose of this paper, we use the superscript @{text "_\<^sup>\<alpha>"} to indicate
that a notion is given for alpha-equivalence classes and leave it out
for the corresponding notion given on the raw level. So for example
we have @{text "ty\<^sup>\<alpha> / ty"} and @{text "C\<^sup>\<alpha> / C"}
where @{term ty} is the type used in the quotient construction for
@{text "ty\<^sup>\<alpha>"} and @{text "C"} is the term-constructor of the raw type @{text "ty"},
respectively @{text "C\<^sup>\<alpha>"} is the corresponding term-constructor of @{text "ty\<^sup>\<alpha>"}.
The resulting datatype definition is legal in Isabelle/HOL provided the datatypes are
non-empty and the types in the constructors only occur in positive
position (see \cite{Berghofer99} for an in-depth description of the datatype package
in Isabelle/HOL).
We subsequently define each of the user-specified binding
functions @{term "bn"}$_{1..m}$ by recursion over the corresponding
raw datatype. We also define permutation operations by
recursion so that for each term constructor @{text "C"} we have that
\begin{equation}\label{ceqvt}
@{text "\<pi> \<bullet> (C z\<^isub>1 \<dots> z\<^isub>n) = C (\<pi> \<bullet> z\<^isub>1) \<dots> (\<pi> \<bullet> z\<^isub>n)"}
\end{equation}\smallskip
\noindent
We will need this operation later when we define the notion of alpha-equivalence.
The first non-trivial step we have to perform is the generation of
\emph{free-atom functions} from the specifications.\footnote{Admittedly, the
details of our definitions will be somewhat involved. However they are still
conceptually simple in comparison with the `positional' approach taken in
Ott \cite[Pages 88--95]{ott-jfp}, which uses the notions of \emph{occurrences} and
\emph{partial equivalence relations} over sets of occurrences.} For the
\emph{raw} types @{text "ty"}$_{1..n}$ we define the free-atom functions
\begin{equation}\label{fvars}
\mbox{@{text "fa_ty"}$_{1..n}$}
\end{equation}\smallskip
\noindent
by recursion.
We define these functions together with auxiliary free-atom functions for
the binding functions. Given raw binding functions @{text "bn"}$_{1..m}$
we define
\[
@{text "fa_bn"}\mbox{$_{1..m}$}.
\]\smallskip
\noindent
The reason for this setup is that in a deep binder not all atoms have to be
bound, as we saw in \eqref{letrecs} with the example of `plain' @{text Let}s. We need
therefore functions that calculate those free atoms in deep binders.
While the idea behind these free-atom functions is simple (they just
collect all atoms that are not bound), because of our rather complicated
binding mechanisms their definitions are somewhat involved. Given
a raw term-constructor @{text "C"} of type @{text ty} and some associated
binding clauses @{text "bc\<^isub>1\<dots>bc\<^isub>k"}, the result of @{text
"fa_ty (C z\<^isub>1 \<dots> z\<^isub>n)"} will be the union @{text
"fa(bc\<^isub>1) \<union> \<dots> \<union> fa(bc\<^isub>k)"} where we will define below what @{text "fa"} for a binding
clause means. We only show the details for the mode \isacommand{binds (set)} (the other modes are similar).
Suppose the binding clause @{text bc\<^isub>i} is of the form
\[
\mbox{\isacommand{binds (set)} @{text "b\<^isub>1\<dots>b\<^isub>p"} \isacommand{in} @{text "d\<^isub>1\<dots>d\<^isub>q"}}
\]\smallskip
\noindent
in which the body-labels @{text "d"}$_{1..q}$ refer to types @{text
ty}$_{1..q}$, and the binders @{text b}$_{1..p}$ either refer to labels of
atom types (in case of shallow binders) or to binding functions taking a
single label as argument (in case of deep binders). Assuming @{text "D"}
stands for the set of free atoms of the bodies, @{text B} for the set of
binding atoms in the binders and @{text "B'"} for the set of free atoms in
non-recursive deep binders, then the free atoms of the binding clause @{text
bc\<^isub>i} are
\begin{equation}\label{fadef}
\mbox{@{text "fa(bc\<^isub>i) \<equiv> (D - B) \<union> B'"}}.
\end{equation}\smallskip
\noindent
The set @{text D} is formally defined as
\[
@{text "D \<equiv> fa_ty\<^isub>1 d\<^isub>1 \<union> ... \<union> fa_ty\<^isub>q d\<^isub>q"}
\]\smallskip
\noindent
where in case @{text "d\<^isub>i"} refers to one of the raw types @{text "ty"}$_{1..n}$ from the
specification, the function @{text "fa_ty\<^isub>i"} is the corresponding free-atom function
we are defining by recursion; otherwise we set \mbox{@{text "fa_ty\<^isub>i \<equiv> supp"}}. The reason
for the latter is that @{text "ty"}$_i$ is not a type that is part of the specification, and
we assume @{text supp} is the generic function that characterises the free variables of
a type (in fact in the next section we will show that the free-variable functions we
define here, are equal to the support once lifted to alpha-equivalence classes).
In order to formally define the set @{text B} we use the following auxiliary @{text "bn"}-functions
for atom types to which shallow binders may refer\\[-4mm]
\begin{equation}\label{bnaux}\mbox{
\begin{tabular}{r@ {\hspace{2mm}}c@ {\hspace{2mm}}l}
@{text "bn\<^bsub>atom\<^esup> a"} & @{text "\<equiv>"} & @{text "{atom a}"}\\
@{text "bn\<^bsub>atom_set\<^esup> as"} & @{text "\<equiv>"} & @{text "atoms as"}\\
@{text "bn\<^bsub>atom_list\<^esub> as"} & @{text "\<equiv>"} & @{text "atoms (set as)"}
\end{tabular}}
\end{equation}\smallskip
\noindent
Like the function @{text atom}, the function @{text "atoms"} coerces
a set of atoms to a set of the generic atom type.
It is defined as @{text "atoms as \<equiv> {atom a | a \<in> as}"}.
The set @{text B} in \eqref{fadef} is then formally defined as
\begin{equation}\label{bdef}
@{text "B \<equiv> bn_ty\<^isub>1 b\<^isub>1 \<union> ... \<union> bn_ty\<^isub>p b\<^isub>p"}
\end{equation}\smallskip
\noindent
where we use the auxiliary binding functions from \eqref{bnaux} for shallow
binders (that means when @{text "ty"}$_i$ is of type @{text "atom"}, @{text "atom set"} or
@{text "atom list"}).
The set @{text "B'"} in \eqref{fadef} collects all free atoms in
non-recursive deep binders. Let us assume these binders in the binding
clause @{text "bc\<^isub>i"} are
\[
\mbox{@{text "bn\<^isub>1 l\<^isub>1, \<dots>, bn\<^isub>r l\<^isub>r"}}
\]\smallskip
\noindent
with @{text "l"}$_{1..r}$ $\subseteq$ @{text "b"}$_{1..p}$ and
none of the @{text "l"}$_{1..r}$ being among the bodies
@{text "d"}$_{1..q}$. The set @{text "B'"} is defined as
\begin{equation}\label{bprimedef}
@{text "B' \<equiv> fa_bn\<^isub>1 l\<^isub>1 \<union> ... \<union> fa_bn\<^isub>r l\<^isub>r"}
\end{equation}\smallskip
\noindent
This completes all clauses for the free-atom functions @{text "fa_ty"}$_{1..n}$.
Note that for non-recursive deep binders, we have to add in \eqref{fadef}
the set of atoms that are left unbound by the binding functions @{text
"bn"}$_{1..m}$. We used for
the definition of this set the functions @{text "fa_bn"}$_{1..m}$. The
definition for those functions needs to be extracted from the clauses the
user provided for @{text "bn"}$_{1..m}$ Assume the user specified a @{text
bn}-clause of the form
\[
@{text "bn (C z\<^isub>1 \<dots> z\<^isub>s) = rhs"}
\]\smallskip
\noindent
where the @{text "z"}$_{1..s}$ are of types @{text "ty"}$_{1..s}$. For
each of the arguments we calculate the free atoms as follows:
\[\mbox{
\begin{tabular}{c@ {\hspace{2mm}}p{0.9\textwidth}}
$\bullet$ & @{term "fa_ty\<^isub>i z\<^isub>i"} provided @{text "z\<^isub>i"} does not occur in @{text "rhs"}\\
& (that means nothing is bound in @{text "z\<^isub>i"} by the binding function),\smallskip\\
$\bullet$ & @{term "fa_bn\<^isub>i z\<^isub>i"} provided @{text "z\<^isub>i"} occurs in @{text "rhs"}
with the recursive call @{text "bn\<^isub>i z\<^isub>i"}\\
& (that means whatever is `left over' from the @{text "bn"}-function is free)\smallskip\\
$\bullet$ & @{term "{}"} provided @{text "z\<^isub>i"} occurs in @{text "rhs"},
but without a recursive call\\
& (that means @{text "z\<^isub>i"} is supposed to become bound by the binding function)\\
\end{tabular}}
\]\smallskip
\noindent
For defining @{text "fa_bn (C z\<^isub>1 \<dots> z\<^isub>n)"} we just union up all these sets.
To see how these definitions work in practice, let us reconsider the
term-constructors @{text "Let"} and @{text "Let_rec"} shown in
\eqref{letrecs} together with the term-constructors for assignments @{text
"ANil"} and @{text "ACons"}. Since there is a binding function defined for
assignments, we have three free-atom functions, namely @{text
"fa\<^bsub>trm\<^esub>"}, @{text "fa\<^bsub>assn\<^esub>"} and @{text
"fa\<^bsub>bn\<^esub>"} as follows:
\[\mbox{
\begin{tabular}{@ {}l@ {\hspace{1mm}}c@ {\hspace{1mm}}l@ {}}
@{text "fa\<^bsub>trm\<^esub> (Let as t)"} & @{text "\<equiv>"} & @{text "(fa\<^bsub>trm\<^esub> t - set (bn as)) \<union> fa\<^bsub>bn\<^esub> as"}\\
@{text "fa\<^bsub>trm\<^esub> (Let_rec as t)"} & @{text "\<equiv>"} & @{text "(fa\<^bsub>assn\<^esub> as \<union> fa\<^bsub>trm\<^esub> t) - set (bn as)"}\smallskip\\
@{text "fa\<^bsub>assn\<^esub> (ANil)"} & @{text "\<equiv>"} & @{term "{}"}\\
@{text "fa\<^bsub>assn\<^esub> (ACons a t as)"} & @{text "\<equiv>"} & @{text "(supp a) \<union> (fa\<^bsub>trm\<^esub> t) \<union> (fa\<^bsub>assn\<^esub> as)"}\smallskip\\
@{text "fa\<^bsub>bn\<^esub> (ANil)"} & @{text "\<equiv>"} & @{term "{}"}\\
@{text "fa\<^bsub>bn\<^esub> (ACons a t as)"} & @{text "\<equiv>"} & @{text "(fa\<^bsub>trm\<^esub> t) \<union> (fa\<^bsub>bn\<^esub> as)"}
\end{tabular}}
\]\smallskip
\noindent
Recall that @{text ANil} and @{text "ACons"} have no binding clause in the
specification. The corresponding free-atom function @{text
"fa\<^bsub>assn\<^esub>"} therefore returns all free atoms of an assignment
(in case of @{text "ACons"}, they are given in terms of @{text supp}, @{text
"fa\<^bsub>trm\<^esub>"} and @{text "fa\<^bsub>assn\<^esub>"}). The binding
only takes place in @{text Let} and @{text "Let_rec"}. In case of @{text
"Let"}, the binding clause specifies that all atoms given by @{text "set (bn
as)"} have to be bound in @{text t}. Therefore we have to subtract @{text
"set (bn as)"} from @{text "fa\<^bsub>trm\<^esub> t"}. However, we also need
to add all atoms that are free in @{text "as"}. This is in contrast with
@{text "Let_rec"} where we have a recursive binder to bind all occurrences
of the atoms in @{text "set (bn as)"} also inside @{text "as"}. Therefore we
have to subtract @{text "set (bn as)"} from both @{text
"fa\<^bsub>trm\<^esub> t"} and @{text "fa\<^bsub>assn\<^esub> as"}. Like the
function @{text "bn"}, the function @{text "fa\<^bsub>bn\<^esub>"} traverses
the list of assignments, but instead returns the free atoms, which means in
this example the free atoms in the argument @{text "t"}.
An interesting point in this example is that a `naked' assignment (@{text
"ANil"} or @{text "ACons"}) does not bind any atoms, even if the binding
function is specified over assignments. Only in the context of a @{text Let}
or @{text "Let_rec"}, where the binding clauses are given, will some atoms
actually become bound. This is a phenomenon that has also been pointed out
in \cite{ott-jfp}. For us this observation is crucial, because we would not
be able to lift the @{text "bn"}-functions to alpha-equated terms if they
act on atoms that are bound. In that case, these functions would \emph{not}
respect alpha-equivalence.
Having the free-atom functions at our disposal, we can next define the
alpha-equivalence relations for the raw types @{text
"ty"}$_{1..n}$. We write them as
\[
\mbox{@{text "\<approx>ty"}$_{1..n}$}.
\]\smallskip
\noindent
Like with the free-atom functions, we also need to
define auxiliary alpha-equivalence relations
\[
\mbox{@{text "\<approx>bn\<^isub>"}$_{1..m}$}
\]\smallskip
\noindent
for the binding functions @{text "bn"}$_{1..m}$,
To simplify our definitions we will use the following abbreviations for
\emph{compound equivalence relations} and \emph{compound free-atom functions} acting on tuples.
\[\mbox{
\begin{tabular}{r@ {\hspace{2mm}}c@ {\hspace{2mm}}l}
@{text "(x\<^isub>1,\<dots>, x\<^isub>n) (R\<^isub>1,\<dots>, R\<^isub>n) (y\<^isub>1,\<dots>, y\<^isub>n)"} & @{text "\<equiv>"} &
@{text "x\<^isub>1 R\<^isub>1 y\<^isub>1 \<and> \<dots> \<and> x\<^isub>n R\<^isub>n y\<^isub>n"}\\
@{text "(fa\<^isub>1,\<dots>, fa\<^isub>n) (x\<^isub>1,\<dots>, x\<^isub>n)"} & @{text "\<equiv>"} & @{text "fa\<^isub>1 x\<^isub>1 \<union> \<dots> \<union> fa\<^isub>n x\<^isub>n"}\\
\end{tabular}}
\]\smallskip
The alpha-equivalence relations are defined as inductive predicates
having a single clause for each term-constructor. Assuming a
term-constructor @{text C} is of type @{text ty} and has the binding clauses
@{term "bc"}$_{1..k}$, then the alpha-equivalence clause has the form
\[
\mbox{\infer{@{text "C z\<^isub>1 \<dots> z\<^isub>n \<approx>ty C z\<PRIME>\<^isub>1 \<dots> z\<PRIME>\<^isub>n"}}
{@{text "prems(bc\<^isub>1) \<dots> prems(bc\<^isub>k)"}}}
\]\smallskip
\noindent
The task below is to specify what the premises corresponding to a binding
clause are. To understand better what the general pattern is, let us first
treat the special instance where @{text "bc\<^isub>i"} is the empty binding clause
of the form
\[
\mbox{\isacommand{binds (set)} @{term "{}"} \isacommand{in} @{text "d\<^isub>1\<dots>d\<^isub>q"}.}
\]\smallskip
\noindent
In this binding clause no atom is bound and we only have to `alpha-relate'
the bodies. For this we build first the tuples @{text "D \<equiv> (d\<^isub>1,\<dots>,
d\<^isub>q)"} and @{text "D' \<equiv> (d\<PRIME>\<^isub>1,\<dots>, d\<PRIME>\<^isub>q)"}
whereby the labels @{text "d"}$_{1..q}$ refer to some of the arguments @{text
"z"}$_{1..n}$ and respectively @{text "d\<PRIME>"}$_{1..q}$ to some of @{text
"z\<PRIME>"}$_{1..n}$ of the term-constructor. In order to relate two such
tuples we define the compound alpha-equivalence relation @{text "R"} as
follows
\begin{equation}\label{rempty}
\mbox{@{text "R \<equiv> (R\<^isub>1,\<dots>, R\<^isub>q)"}}
\end{equation}\smallskip
\noindent
with @{text "R\<^isub>i"} being @{text "\<approx>ty\<^isub>i"} if the corresponding
labels @{text "d\<^isub>i"} and @{text "d\<PRIME>\<^isub>i"} refer to a
recursive argument of @{text C} and have type @{text "ty\<^isub>i"}; otherwise
we take @{text "R\<^isub>i"} to be the equality @{text "="}. Again the
latter is because @{text "ty\<^isub>i"} is not part of the specified types
and alpha-equivalence of any previously defined type is supposed to coincide
with equality. This lets us now define the premise for an empty binding
clause succinctly as @{text "prems(bc\<^isub>i) \<equiv> D R D'"}, which can be
unfolded to the series of premises
\[
@{text "d\<^isub>1 R\<^isub>1 d\<PRIME>\<^isub>1 \<dots> d\<^isub>q R\<^isub>q d\<PRIME>\<^isub>q"}.
\]\smallskip
\noindent
We will use the unfolded version in the examples below.
Now suppose the binding clause @{text "bc\<^isub>i"} is of the general form
\begin{equation}\label{nonempty}
\mbox{\isacommand{binds (set)} @{text "b\<^isub>1\<dots>b\<^isub>p"} \isacommand{in} @{text "d\<^isub>1\<dots>d\<^isub>q"}.}
\end{equation}\smallskip
\noindent
In this case we define a premise @{text P} using the relation
$\approx_{\,\textit{set}}^{\textit{R}, \textit{fa}}$ given in Section~\ref{sec:binders} (similarly
$\approx_{\,\textit{set+}}^{\textit{R}, \textit{fa}}$ and
$\approx_{\,\textit{list}}^{\textit{R}, \textit{fa}}$ for the other
binding modes). As above, we first build the tuples @{text "D"} and
@{text "D'"} for the bodies @{text "d"}$_{1..q}$, and the corresponding
compound alpha-relation @{text "R"} (shown in \eqref{rempty}).
For $\approx_{\,\textit{set}}^{\textit{R}, \textit{fa}}$ we also need
a compound free-atom function for the bodies defined as
\[
\mbox{@{text "fa \<equiv> (fa_ty\<^isub>1,\<dots>, fa_ty\<^isub>q)"}}
\]\smallskip
\noindent
with the assumption that the @{text "d"}$_{1..q}$ refer to arguments of types @{text "ty"}$_{1..q}$.
The last ingredient we need are the sets of atoms bound in the bodies.
For this we take
\[
@{text "B \<equiv> bn_ty\<^isub>1 b\<^isub>1 \<union> \<dots> \<union> bn_ty\<^isub>p b\<^isub>p"}\;.\\
\]\smallskip
\noindent
Similarly for @{text "B'"} using the labels @{text "b\<PRIME>"}$_{1..p}$. This
lets us formally define the premise @{text P} for a non-empty binding clause as:
\[
\mbox{@{term "P \<equiv> alpha_set_ex (B, D) R fa (B', D')"}}\;.
\]\smallskip
\noindent
This premise accounts for alpha-equivalence of the bodies of the binding
clause. However, in case the binders have non-recursive deep binders, this
premise is not enough: we also have to `propagate' alpha-equivalence
inside the structure of these binders. An example is @{text "Let"} where we
have to make sure the right-hand sides of assignments are
alpha-equivalent. For this we use relations @{text "\<approx>bn"}$_{1..m}$ (which we
will define shortly). Let us assume the non-recursive deep binders
in @{text "bc\<^isub>i"} are
\[
@{text "bn\<^isub>1 l\<^isub>1, \<dots>, bn\<^isub>r l\<^isub>r"}.
\]\smallskip
\noindent
The tuple @{text L} consists then of all these binders @{text "(l\<^isub>1,\<dots>,l\<^isub>r)"}
(similarly @{text "L'"}) and the compound equivalence relation @{text "R'"}
is @{text "(\<approx>bn\<^isub>1,\<dots>,\<approx>bn\<^isub>r)"}. All premises for @{text "bc\<^isub>i"} are then given by
\[
@{text "prems(bc\<^isub>i) \<equiv> P \<and> L R' L'"}
\]\smallskip
\noindent
The auxiliary alpha-equivalence relations @{text "\<approx>bn"}$_{1..m}$
in @{text "R'"} are defined as follows: assuming a @{text bn}-clause is of the form
\[
@{text "bn (C z\<^isub>1 \<dots> z\<^isub>s) = rhs"}
\]\smallskip
\noindent
where the @{text "z"}$_{1..s}$ are of types @{text "ty"}$_{1..s}$,
then the corresponding alpha-equivalence clause for @{text "\<approx>bn"} has the form
\[
\mbox{\infer{@{text "C z\<^isub>1 \<dots> z\<^isub>s \<approx>bn C z\<PRIME>\<^isub>1 \<dots> z\<PRIME>\<^isub>s"}}
{@{text "z\<^isub>1 R\<^isub>1 z\<PRIME>\<^isub>1 \<dots> z\<^isub>s R\<^isub>s z\<PRIME>\<^isub>s"}}}
\]\smallskip
\noindent
In this clause the relations @{text "R"}$_{1..s}$ are given by
\[\mbox{
\begin{tabular}{c@ {\hspace{2mm}}p{0.9\textwidth}}
$\bullet$ & @{text "z\<^isub>i \<approx>ty z\<PRIME>\<^isub>i"} provided @{text "z\<^isub>i"} does not occur in @{text rhs} and
is a recursive argument of @{text C},\smallskip\\
$\bullet$ & @{text "z\<^isub>i = z\<PRIME>\<^isub>i"} provided @{text "z\<^isub>i"} does not occur in @{text rhs}
and is a non-recursive argument of @{text C},\smallskip\\
$\bullet$ & @{text "z\<^isub>i \<approx>bn\<^isub>i z\<PRIME>\<^isub>i"} provided @{text "z\<^isub>i"} occurs in @{text rhs}
with the recursive call @{text "bn\<^isub>i x\<^isub>i"} and\smallskip\\
$\bullet$ & @{text True} provided @{text "z\<^isub>i"} occurs in @{text rhs} but without a
recursive call.
\end{tabular}}
\]\smallskip
\noindent
This completes the definition of alpha-equivalence. As a sanity check, we can show
that the premises of empty binding clauses are a special case of the clauses for
non-empty ones (we just have to unfold the definition of
$\approx_{\,\textit{set}}^{\textit{R}, \textit{fa}}$ and take @{text "0"}
for the existentially quantified permutation).
Again let us take a look at a concrete example for these definitions. For
the specification shown in \eqref{letrecs}
we have three relations $\approx_{\textit{trm}}$, $\approx_{\textit{assn}}$ and
$\approx_{\textit{bn}}$ with the following rules:
\begin{equation}\label{rawalpha}\mbox{
\begin{tabular}{@ {}c @ {}}
\infer{@{text "Let as t \<approx>\<^bsub>trm\<^esub> Let as' t'"}}
{@{term "alpha_lst_ex (bn as, t) alpha_trm fa_trm (bn as', t')"} &
\hspace{5mm}@{text "as \<approx>\<^bsub>bn\<^esub> as'"}}\\
\\
\makebox[0mm]{\infer{@{text "Let_rec as t \<approx>\<^bsub>trm\<^esub> Let_rec as' t'"}}
{@{term "alpha_lst_ex (bn as, ast) alpha_trm2 fa_trm2 (bn as', ast')"}}}\\
\\
\begin{tabular}{@ {}c @ {}}
\infer{@{text "ANil \<approx>\<^bsub>assn\<^esub> ANil"}}{}\hspace{9mm}
\infer{@{text "ACons a t as \<approx>\<^bsub>assn\<^esub> ACons a' t' as"}}
{@{text "a = a'"} & \hspace{5mm}@{text "t \<approx>\<^bsub>trm\<^esub> t'"} & \hspace{5mm}@{text "as \<approx>\<^bsub>assn\<^esub> as'"}}
\end{tabular}\\
\\
\begin{tabular}{@ {}c @ {}}
\infer{@{text "ANil \<approx>\<^bsub>bn\<^esub> ANil"}}{}\hspace{9mm}
\infer{@{text "ACons a t as \<approx>\<^bsub>bn\<^esub> ACons a' t' as"}}
{@{text "t \<approx>\<^bsub>trm\<^esub> t'"} & \hspace{5mm}@{text "as \<approx>\<^bsub>bn\<^esub> as'"}}
\end{tabular}
\end{tabular}}
\end{equation}\smallskip
\noindent
Notice the difference between $\approx_{\textit{assn}}$ and
$\approx_{\textit{bn}}$: the latter only `tracks' alpha-equivalence of
the components in an assignment that are \emph{not} bound. This is needed in the
clause for @{text "Let"} (which has
a non-recursive binder).
The underlying reason is that the terms inside an assignment are not meant
to be `under' the binder. Such a premise is \emph{not} needed in @{text "Let_rec"},
because there all components of an assignment are `under' the binder.
Note also that in case of more than one body (that is in the @{text "Let_rec"}-case above)
we need to parametrise the relation $\approx_{\textit{list}}$ with a compound
equivalence relation and a compound free-atom function. This is because the
corresponding binding clause specifies a binder with two bodies, namely
@{text "as"} and @{text "t"}.
*}
section {* Establishing the Reasoning Infrastructure *}
text {*
Having made all necessary definitions for raw terms, we can start with
establishing the reasoning infrastructure for the alpha-equated types @{text
"ty\<AL>"}$_{1..n}$, that is the types the user originally specified. We
give in this section and the next the proofs we need for establishing this
infrastructure. One point of our work is that we have completely
automated these proofs in Isabelle/HOL.
First we establish that the free-variable functions, the binding functions and the
alpha-equi\-va\-lences are equivariant.
\begin{lem}\mbox{}\\
@{text "(i)"} The functions @{text "fa_ty"}$_{1..n}$, @{text "fa_bn"}$_{1..m}$ and
@{text "bn"}$_{1..m}$ are equivariant.\\
@{text "(ii)"} The relations @{text "\<approx>ty"}$_{1..n}$ and
@{text "\<approx>bn"}$_{1..m}$ are equivariant.
\end{lem}
\begin{proof}
The function package of Isabelle/HOL allows us to prove the first part by
mutual induction over the definitions of the functions.\footnote{We have
that the free-atom functions are terminating. From this the function
package derives an induction principle~\cite{Krauss09}.} The second is by a
straightforward induction over the rules of @{text "\<approx>ty"}$_{1..n}$ and
@{text "\<approx>bn"}$_{1..m}$ using the first part.
\end{proof}
\noindent
Next we establish that the alpha-equivalence relations defined in the
previous section are indeed equivalence relations.
\begin{lem}\label{equiv}
The relations @{text "\<approx>ty"}$_{1..n}$ and @{text "\<approx>bn"}$_{1..m}$ are
equivalence relations.
\end{lem}
\begin{proof}
The proofs are by induction. The non-trivial
cases involve premises built up by $\approx_{\textit{set}}$,
$\approx_{\textit{set+}}$ and $\approx_{\textit{list}}$. They
can be dealt with as in Lemma~\ref{alphaeq}. However, the transitivity
case needs in addition the fact that the relations are equivariant.
\end{proof}
\noindent
We can feed the last lemma into our quotient package and obtain new types
@{text "ty"}$^\alpha_{1..n}$ representing alpha-equated terms of types
@{text "ty"}$_{1..n}$. We also obtain definitions for the term-constructors
@{text "C"}$^\alpha_{1..k}$ from the raw term-constructors @{text
"C"}$_{1..k}$, and similar definitions for the free-atom functions @{text
"fa_ty"}$^\alpha_{1..n}$ and @{text "fa_bn"}$^\alpha_{1..m}$ as well as the
binding functions @{text "bn"}$^\alpha_{1..m}$. However, these definitions
are not really useful to the user, since they are given in terms of the
isomorphisms we obtained by creating new types in Isabelle/HOL (recall the
picture shown in the Introduction).
The first useful property for the user is the fact that distinct
term-constructors are not equal, that is the property
\begin{equation}\label{distinctalpha}
\mbox{@{text "C"}$^\alpha$~@{text "x\<^isub>1 \<dots> x\<^isub>r"}~@{text "\<noteq>"}~%
@{text "D"}$^\alpha$~@{text "y\<^isub>1 \<dots> y\<^isub>s"}}
\end{equation}\smallskip
\noindent
whenever @{text "C"}$^\alpha$~@{text "\<noteq>"}~@{text "D"}$^\alpha$.
In order to derive this property, we use the definition of alpha-equivalence
and establish that
\begin{equation}\label{distinctraw}
\mbox{@{text "C x\<^isub>1 \<dots> x\<^isub>r"}\;$\not\approx$@{text ty}\;@{text "D y\<^isub>1 \<dots> y\<^isub>s"}}
\end{equation}\smallskip
\noindent
holds for the corresponding raw term-constructors.
In order to deduce \eqref{distinctalpha} from \eqref{distinctraw}, our quotient
package needs to know that the raw term-constructors @{text "C"} and @{text "D"}
are \emph{respectful} w.r.t.~the alpha-equivalence relations (see \cite{Homeier05}).
Given, for example, @{text "C"} is of type @{text "ty"} with argument types
@{text "ty"}$_{1..r}$, respectfulness amounts to showing that
\[\mbox{
@{text "C x\<^isub>1 \<dots> x\<^isub>r \<approx>ty C x\<PRIME>\<^isub>1 \<dots> x\<PRIME>\<^isub>r"}
}\]\smallskip
\noindent
holds under the assumptions \mbox{@{text
"x\<^isub>i \<approx>ty\<^isub>i x\<PRIME>\<^isub>i"}} whenever @{text "x\<^isub>i"}
and @{text "x\<PRIME>\<^isub>i"} are recursive arguments of @{text C}, and
@{text "x\<^isub>i = x\<PRIME>\<^isub>i"} whenever they are non-recursive arguments
(similarly for @{text "D"}). For this we have to show
by induction over the definitions of alpha-equivalences the following
auxiliary implications
\begin{equation}\label{fnresp}\mbox{
\begin{tabular}{lll}
@{text "x \<approx>ty\<^isub>i x'"} & implies & @{text "fa_ty\<^isub>i x = fa_ty\<^isub>i x'"}\\
@{text "x \<approx>ty\<^isub>l x'"} & implies & @{text "fa_bn\<^isub>j x = fa_bn\<^isub>j x'"}\\
@{text "x \<approx>ty\<^isub>l x'"} & implies & @{text "bn\<^isub>j x = bn\<^isub>j x'"}\\
@{text "x \<approx>ty\<^isub>l x'"} & implies & @{text "x \<approx>bn\<^isub>j x'"}\\
\end{tabular}
}\end{equation}\smallskip
\noindent
whereby @{text "ty\<^isub>l"} is the type over which @{text "bn\<^isub>j"}
is defined. Whereas the first, second and last implication are true by
how we stated our definitions, the third \emph{only} holds because of our
restriction imposed on the form of the binding functions---namely \emph{not}
to return any bound atoms. In Ott, in contrast, the user may define @{text
"bn"}$_{1..m}$ so that they return bound atoms and in this case the third
implication is \emph{not} true. A result is that in general the lifting of the
corresponding binding functions in Ott to alpha-equated terms is impossible.
Having established respectfulness for the raw term-constructors, the
quotient package is able to automatically deduce \eqref{distinctalpha} from
\eqref{distinctraw}.
Next we can lift the permutation operations defined in \eqref{ceqvt}. In
order to make this lifting to go through, we have to show that the
permutation operations are respectful. This amounts to showing that the
alpha-equivalence relations are equivariant, which
we already established in Lemma~\ref{equiv}. As a result we can add the
equations
\begin{equation}\label{calphaeqvt}
@{text "\<pi> \<bullet> (C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>r) = C\<^sup>\<alpha> (\<pi> \<bullet> x\<^isub>1) \<dots> (\<pi> \<bullet> x\<^isub>r)"}
\end{equation}\smallskip
\noindent
to our infrastructure. In a similar fashion we can lift the defining equations
of the free-atom functions @{text "fa_ty\<AL>"}$_{1..n}$ and
@{text "fa_bn\<AL>"}$_{1..m}$ as well as of the binding functions @{text
"bn\<AL>"}$_{1..m}$ and size functions @{text "size_ty\<AL>"}$_{1..n}$.
The latter are defined automatically for the raw types @{text "ty"}$_{1..n}$
by the datatype package of Isabelle/HOL.
We also need to lift the properties that characterise when two raw terms of the form
\[
\mbox{@{text "C x\<^isub>1 \<dots> x\<^isub>r \<approx>ty C x\<PRIME>\<^isub>1 \<dots> x\<PRIME>\<^isub>r"}}
\]\smallskip
\noindent
are alpha-equivalent. This gives us conditions when the corresponding
alpha-equated terms are \emph{equal}, namely
\[
@{text "C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>r = C\<^sup>\<alpha> x\<PRIME>\<^isub>1 \<dots> x\<PRIME>\<^isub>r"}.
\]\smallskip
\noindent
We call these conditions \emph{quasi-injectivity}. They correspond to the
premises in our alpha-equiva\-lence relations, except that the
relations @{text "\<approx>ty"}$_{1..n}$ are all replaced by equality (and similarly
the free-atom and binding functions are replaced by their lifted
counterparts). Recall the alpha-equivalence rules for @{text "Let"} and
@{text "Let_rec"} shown in \eqref{rawalpha}. For @{text "Let\<^sup>\<alpha>"} and
@{text "Let_rec\<^sup>\<alpha>"} we have
\begin{equation}\label{alphalift}\mbox{
\begin{tabular}{@ {}c @ {}}
\infer{@{text "Let\<^sup>\<alpha> as t = Let\<^sup>\<alpha> as' t'"}}
{@{term "alpha_lst_ex (bn_al as, t) equal fa_trm_al (bn as', t')"} &
\hspace{5mm}@{text "as \<approx>\<AL>\<^bsub>bn\<^esub> as'"}}\\
\\
\makebox[0mm]{\infer{@{text "Let_rec\<^sup>\<alpha> as t = Let_rec\<^sup>\<alpha> as' t'"}}
{@{term "alpha_lst_ex (bn_al as, ast) equ2 fa_trm2_al (bn_al as', ast')"}}}\\
\end{tabular}}
\end{equation}\smallskip
We can also add to our infrastructure cases lemmas and a (mutual)
induction principle for the types @{text "ty\<AL>"}$_{1..n}$. The cases
lemmas allow the user to deduce a property @{text "P"} by exhaustively
analysing how an element of a type, say @{text "ty\<AL>"}$_i$, can be
constructed (that means one case for each of the term-constructors in @{text
"ty\<AL>"}$_i\,$). The lifted cases lemma for the type @{text
"ty\<AL>"}$_i\,$ looks as follows
\begin{equation}\label{cases}
\infer{P}
{\begin{array}{l}
@{text "\<forall>x\<^isub>1\<dots>x\<^isub>k. y = C\<AL>\<^isub>1 x\<^isub>1 \<dots> x\<^isub>k \<Rightarrow> P"}\\
\hspace{5mm}\vdots\\
@{text "\<forall>x\<^isub>1\<dots>x\<^isub>l. y = C\<AL>\<^isub>m x\<^isub>1 \<dots> x\<^isub>l \<Rightarrow> P"}\\
\end{array}}
\end{equation}\smallskip
\noindent
where @{text "y"} is a variable of type @{text "ty\<AL>"}$_i$ and @{text "P"} is the
property that is established by the case analysis. Similarly, we have a (mutual)
induction principle for the types @{text "ty\<AL>"}$_{1..n}$, which is of the
form
\begin{equation}\label{induct}
\infer{@{text "P\<^isub>1 y\<^isub>1 \<and> \<dots> \<and> P\<^isub>n y\<^isub>n "}}
{\begin{array}{l}
@{text "\<forall>x\<^isub>1\<dots>x\<^isub>k. P\<^isub>i x\<^isub>i \<and> \<dots> \<and> P\<^isub>j x\<^isub>j \<Rightarrow> P (C\<AL>\<^isub>1 x\<^isub>1 \<dots> x\<^isub>k)"}\\
\hspace{5mm}\vdots\\
@{text "\<forall>x\<^isub>1\<dots>x\<^isub>l. P\<^isub>r x\<^isub>r \<and> \<dots> \<and> P\<^isub>s x\<^isub>s \<Rightarrow> P (C\<AL>\<^isub>m x\<^isub>1 \<dots> x\<^isub>l)"}\\
\end{array}}
\end{equation}\smallskip
\noindent
whereby the @{text P}$_{1..n}$ are the properties established by the
induction, and the @{text y}$_{1..n}$ are of type @{text
"ty\<AL>"}$_{1..n}$. Note that for the term constructor @{text
"C"}$^\alpha_1$ the induction principle has a hypothesis of the form
\[
\mbox{@{text "\<forall>x\<^isub>1\<dots>x\<^isub>k. P\<^isub>i x\<^isub>i \<and> \<dots> \<and> P\<^isub>j x\<^isub>j \<Rightarrow> P (C\<AL>\<^sub>1 x\<^isub>1 \<dots> x\<^isub>k)"}}
\]\smallskip
\noindent
in which the @{text "x"}$_{i..j}$ @{text "\<subseteq>"} @{text "x"}$_{1..k}$ are the
recursive arguments of this term constructor (similarly for the other
term-constructors).
Recall the lambda-calculus with @{text "Let"}-patterns shown in
\eqref{letpat}. The cases lemmas and the induction principle shown in
\eqref{cases} and \eqref{induct} boil down in that example to the following three inference
rules:
\begin{equation}\label{inductex}\mbox{
\begin{tabular}{c}
\multicolumn{1}{@ {\hspace{-5mm}}l}{cases lemmas:}\smallskip\\
\infer{@{text "P\<^bsub>trm\<^esub>"}}
{\begin{array}{@ {}l@ {}}
@{text "\<forall>x. y = Var\<^sup>\<alpha> x \<Rightarrow> P\<^bsub>trm\<^esub>"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. y = App\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 \<Rightarrow> P\<^bsub>trm\<^esub>"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. y = Lam\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 \<Rightarrow> P\<^bsub>trm\<^esub>"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2 x\<^isub>3. y = Let_pat\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 x\<^isub>3 \<Rightarrow> P\<^bsub>trm\<^esub>"}
\end{array}}\hspace{10mm}
\infer{@{text "P\<^bsub>pat\<^esub>"}}
{\begin{array}{@ {}l@ {}}
@{text "\<forall>x. y = PVar\<^sup>\<alpha> x \<Rightarrow> P\<^bsub>pat\<^esub>"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. y = PTup\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 \<Rightarrow> P\<^bsub>pat\<^esub>"}
\end{array}}\medskip\\
\multicolumn{1}{@ {\hspace{-5mm}}l}{induction principle:}\smallskip\\
\infer{@{text "P\<^bsub>trm\<^esub> y\<^isub>1 \<and> P\<^bsub>pat\<^esub> y\<^isub>2"}}
{\begin{array}{@ {}l@ {}}
@{text "\<forall>x. P\<^bsub>trm\<^esub> (Var\<^sup>\<alpha> x)"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. P\<^bsub>trm\<^esub> x\<^isub>1 \<and> P\<^bsub>trm\<^esub> x\<^isub>2 \<Rightarrow> P\<^bsub>trm\<^esub> (App\<^sup>\<alpha> x\<^isub>1 x\<^isub>2)"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. P\<^bsub>trm\<^esub> x\<^isub>2 \<Rightarrow> P\<^bsub>trm\<^esub> (Lam\<^sup>\<alpha> x\<^isub>1 x\<^isub>2)"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2 x\<^isub>3. P\<^bsub>pat\<^esub> x\<^isub>1 \<and> P\<^bsub>trm\<^esub> x\<^isub>2 \<and> P\<^bsub>trm\<^esub> x\<^isub>3 \<Rightarrow> P\<^bsub>trm\<^esub> (Let_pat\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 x\<^isub>3)"}\\
@{text "\<forall>x. P\<^bsub>pat\<^esub> (PVar\<^sup>\<alpha> x)"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. P\<^bsub>pat\<^esub> x\<^isub>1 \<and> P\<^bsub>pat\<^esub> x\<^isub>2 \<Rightarrow> P\<^bsub>pat\<^esub> (PTup\<^sup>\<alpha> x\<^isub>1 x\<^isub>2)"}
\end{array}}
\end{tabular}}
\end{equation}\smallskip
By working now completely on the alpha-equated level, we
can first show using \eqref{calphaeqvt} and Property~\ref{swapfreshfresh} that the support of each term
constructor is included in the support of its arguments,
namely
\[
@{text "(supp x\<^isub>1 \<union> \<dots> \<union> supp x\<^isub>r) supports (C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>r)"}
\]\smallskip
\noindent
This allows us to prove using the induction principle for @{text "ty\<AL>"}$_{1..n}$
that every element of type @{text "ty\<AL>"}$_{1..n}$ is finitely supported
(using Proposition~\ref{supportsprop}{\it (i)}).
Similarly, we can establish by induction that the free-atom functions and binding
functions are equivariant, namely
\[\mbox{
\begin{tabular}{rcl}
@{text "\<pi> \<bullet> (fa_ty\<AL>\<^isub>i x)"} & $=$ & @{text "fa_ty\<AL>\<^isub>i (\<pi> \<bullet> x)"}\\
@{text "\<pi> \<bullet> (fa_bn\<AL>\<^isub>j x)"} & $=$ & @{text "fa_bn\<AL>\<^isub>j (\<pi> \<bullet> x)"}\\
@{text "\<pi> \<bullet> (bn\<AL>\<^isub>j x)"} & $=$ & @{text "bn\<AL>\<^isub>j (\<pi> \<bullet> x)"}\\
\end{tabular}}
\]\smallskip
\noindent
Lastly, we can show that the support of elements in @{text
"ty\<AL>"}$_{1..n}$ is the same as the free-atom functions @{text
"fa_ty\<AL>"}$_{1..n}$. This fact is important in the nominal setting where
the general theory is formulated in terms of support and freshness, but also
provides evidence that our notions of free-atoms and alpha-equivalence
`match up' correctly.
\begin{thm}\label{suppfa}
For @{text "x"}$_{1..n}$ with type @{text "ty\<AL>"}$_{1..n}$, we have
@{text "supp x\<^isub>i = fa_ty\<AL>\<^isub>i x\<^isub>i"}.
\end{thm}
\begin{proof}
The proof is by induction on @{text "x"}$_{1..n}$. In each case
we unfold the definition of @{text "supp"}, move the swapping inside the
term-constructors and then use the quasi-injectivity lemmas in order to complete the
proof. For the abstraction cases we use then the facts derived in Theorem~\ref{suppabs},
for which we have to know that every body of an abstraction is finitely supported.
This, we have proved earlier.
\end{proof}
\noindent
Consequently, we can replace the free-atom functions by @{text "supp"} in
our quasi-injection lemmas. In the examples shown in \eqref{alphalift}, for instance,
we obtain for @{text "Let\<^sup>\<alpha>"} and @{text "Let_rec\<^sup>\<alpha>"}
\[\mbox{
\begin{tabular}{@ {}c @ {}}
\infer{@{text "Let\<^sup>\<alpha> as t = Let\<^sup>\<alpha> as' t'"}}
{@{term "alpha_lst_ex (bn_al as, t) equal supp (bn_al as', t')"} &
\hspace{5mm}@{text "as \<approx>\<AL>\<^bsub>bn\<^esub> as'"}}\\
\\
\makebox[0mm]{\infer{@{text "Let_rec\<^sup>\<alpha> as t = Let_rec\<^sup>\<alpha> as' t'"}}
{@{term "alpha_lst_ex (bn_al as, ast) equ2 supp2 (bn_al as', ast')"}}}\\
\end{tabular}}
\]\smallskip
\noindent
Taking into account that the compound equivalence relation @{term
"equ2"} and the compound free-atom function @{term "supp2"} are by
definition equal to @{term "equal"} and @{term "supp"}, respectively, the
above rules simplify further to
\[\mbox{
\begin{tabular}{@ {}c @ {}}
\infer{@{text "Let\<^sup>\<alpha> as t = Let\<^sup>\<alpha> as' t'"}}
{@{term "Abs_lst (bn_al as) t = Abs_lst (bn_al as') t'"} &
\hspace{5mm}@{text "as \<approx>\<AL>\<^bsub>bn\<^esub> as'"}}\\
\\
\makebox[0mm]{\infer{@{text "Let_rec\<^sup>\<alpha> as t = Let_rec\<^sup>\<alpha> as' t'"}}
{@{term "Abs_lst (bn_al as) ast = Abs_lst (bn_al as') ast'"}}}\\
\end{tabular}}
\]\smallskip
\noindent
which means we can characterise equality between term-constructors (on the
alpha-equated level) in terms of equality between the abstractions defined
in Section~\ref{sec:binders}. From this we can deduce the support for @{text
"Let\<^sup>\<alpha>"} and @{text "Let_rec\<^sup>\<alpha>"}, namely
\[\mbox{
\begin{tabular}{l@ {\hspace{2mm}}l@ {\hspace{2mm}}l}
@{text "supp (Let\<^sup>\<alpha> as t)"} & @{text "="} & @{text "(supp t - set (bn\<^sup>\<alpha> as)) \<union> fa\<AL>\<^bsub>bn\<^esub> as"}\\
@{text "supp (Let_rec\<^sup>\<alpha> as t)"} & @{text "="} & @{text "(supp t \<union> supp as) - set (bn\<^sup>\<alpha> as)"}\\
\end{tabular}}
\]\smallskip
\noindent
using the support of abstractions derived in Theorem~\ref{suppabs}.
To sum up this section, we have established a reasoning infrastructure for the
types @{text "ty\<AL>"}$_{1..n}$ by first lifting definitions from the
`raw' level to the quotient level and then by proving facts about
these lifted definitions. All necessary proofs are generated automatically
by custom ML-code.
*}
section {* Strong Induction Principles *}
text {*
In the previous section we derived induction principles for alpha-equated
terms (see \eqref{induct} for the general form and \eqref{inductex} for an
example). This was done by lifting the corresponding inductions principles
for `raw' terms. We already employed these induction principles for
deriving several facts about alpha-equated terms, including the property that
the free-atom functions and the notion of support coincide. Still, we
call these induction principles \emph{weak}, because for a term-constructor,
say \mbox{@{text "C\<^sup>\<alpha> x\<^isub>1\<dots>x\<^isub>r"}}, the induction
hypothesis requires us to establish (under some assumptions) a property
@{text "P (C\<^sup>\<alpha> x\<^isub>1\<dots>x\<^isub>r)"} for \emph{all} @{text
"x"}$_{1..r}$. The problem with this is that in the presence of binders we cannot make
any assumptions about the atoms that are bound---for example assuming the variable convention.
One obvious way around this
problem is to rename bound atoms. Unfortunately, this leads to very clunky proofs
and makes formalisations grievous experiences (especially in the context of
multiple bound atoms).
For the older versions of Nominal Isabelle we described in \cite{Urban08} a
method for automatically strengthening weak induction principles. These
stronger induction principles allow the user to make additional assumptions
about bound atoms. The advantage of these assumptions is that they make in
most cases any renaming of bound atoms unnecessary. To explain how the
strengthening works, we use as running example the lambda-calculus with
@{text "Let"}-patterns shown in \eqref{letpat}. Its weak induction principle
is given in \eqref{inductex}. The stronger induction principle is as
follows:
\begin{equation}\label{stronginduct}
\mbox{
\begin{tabular}{@ {}c@ {}}
\infer{@{text "P\<^bsub>trm\<^esub> c y\<^isub>1 \<and> P\<^bsub>pat\<^esub> c y\<^isub>2"}}
{\begin{array}{l}
@{text "\<forall>x c. P\<^bsub>trm\<^esub> c (Var\<^sup>\<alpha> x)"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2 c. (\<forall>d. P\<^bsub>trm\<^esub> d x\<^isub>1) \<and> (\<forall>d. P\<^bsub>trm\<^esub> d x\<^isub>2) \<Rightarrow> P\<^bsub>trm\<^esub> c (App\<^sup>\<alpha> x\<^isub>1 x\<^isub>2)"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2 c. atom x\<^isub>1 # c \<and> (\<forall>d. P\<^bsub>trm\<^esub> d x\<^isub>2) \<Rightarrow> P\<^bsub>trm\<^esub> c (Lam\<^sup>\<alpha> x\<^isub>1 x\<^isub>2)"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2 x\<^isub>3 c. (set (bn\<^sup>\<alpha> x\<^isub>1)) #\<^sup>* c \<and>"}\\
\hspace{10mm}@{text "(\<forall>d. P\<^bsub>pat\<^esub> d x\<^isub>1) \<and> (\<forall>d. P\<^bsub>trm\<^esub> d x\<^isub>2) \<and> (\<forall>d. P\<^bsub>trm\<^esub> d x\<^isub>3) \<Rightarrow> P\<^bsub>trm\<^esub> c (Let_pat\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 x\<^isub>3)"}\\
@{text "\<forall>x c. P\<^bsub>pat\<^esub> c (PVar\<^sup>\<alpha> x)"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2 c. (\<forall>d. P\<^bsub>pat\<^esub> d x\<^isub>1) \<and> (\<forall>d. P\<^bsub>pat\<^esub> d x\<^isub>2) \<Rightarrow> P\<^bsub>pat\<^esub> c (PTup\<^sup>\<alpha> x\<^isub>1 x\<^isub>2)"}
\end{array}}
\end{tabular}}
\end{equation}\smallskip
\noindent
Notice that instead of establishing two properties of the form @{text "
P\<^bsub>trm\<^esub> y\<^isub>1 \<and> P\<^bsub>pat\<^esub> y\<^isub>2"}, as the
weak one does, the stronger induction principle establishes the properties
of the form @{text " P\<^bsub>trm\<^esub> c y\<^isub>1 \<and>
P\<^bsub>pat\<^esub> c y\<^isub>2"} in which the additional parameter @{text
c} is assumed to be of finite support. The purpose of @{text "c"} is to
`control' which freshness assumptions the binders should satisfy in the
@{text "Lam\<^sup>\<alpha>"} and @{text "Let_pat\<^sup>\<alpha>"} cases: for @{text
"Lam\<^sup>\<alpha>"} we can assume the bound atom @{text "x\<^isub>1"} is fresh
for @{text "c"} (third line); for @{text "Let_pat\<^sup>\<alpha>"} we can assume
all bound atoms from an assignment are fresh for @{text "c"} (fourth
line). In order to see how an instantiation for @{text "c"} in the
conclusion `controls' the premises, one has to take into account that
Isabelle/HOL is a typed logic. That means if @{text "c"} is instantiated
with, for example, a pair, then this type-constraint will be propagated to
the premises. The main point is that if @{text "c"} is instantiated
appropriately, then the user can mimic the usual `pencil-and-paper'
reasoning employing the variable convention about bound and free variables
being distinct \cite{Urban08}.
In what follows we will show that the weak induction principle in
\eqref{inductex} implies the strong one \eqref{stronginduct}. This fact was established for
single binders in \cite{Urban08} by some quite involved, nevertheless
automated, induction proof. In this paper we simplify the proof by
leveraging the automated proving tools from the function package of
Isabelle/HOL \cite{Krauss09}. The reasoning principle behind these tools
is well-founded induction. To use them in our setting, we have to discharge
two proof obligations: one is that we have well-founded measures (one for
each type @{text "ty"}$^\alpha_{1..n}$) that decrease in every induction
step and the other is that we have covered all cases in the induction
principle. Once these two proof obligations are discharged, the reasoning
infrastructure of the function package will automatically derive the
stronger induction principle. This way of establishing the stronger induction
principle is considerably simpler than the earlier work presented in \cite{Urban08}.
As measures we can use the size functions @{text "size_ty"}$^\alpha_{1..n}$,
which we lifted in the previous section and which are all well-founded. It
is straightforward to establish that the sizes decrease in every
induction step. What is left to show is that we covered all cases.
To do so, we have to derive stronger cases lemmas, which look in our
running example are as follows:
\[\mbox{
\begin{tabular}{@ {}c@ {\hspace{4mm}}c@ {}}
\infer{@{text "P\<^bsub>trm\<^esub>"}}
{\begin{array}{@ {}l@ {}}
@{text "\<forall>x. y = Var\<^sup>\<alpha> x \<Rightarrow> P\<^bsub>trm\<^esub>"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. y = App\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 \<Rightarrow> P\<^bsub>trm\<^esub>"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. atom x\<^isub>1 # c \<and> y = Lam\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 \<Rightarrow> P\<^bsub>trm\<^esub>"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2 x\<^isub>3. set (bn\<^sup>\<alpha> x\<^isub>1) #\<^sup>* c \<and> y = Let_pat\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 x\<^isub>3 \<Rightarrow> P\<^bsub>trm\<^esub>"}
\end{array}} &
\infer{@{text "P\<^bsub>pat\<^esub>"}}
{\begin{array}{@ {}l@ {}}
@{text "\<forall>x. y = PVar\<^sup>\<alpha> x \<Rightarrow> P\<^bsub>pat\<^esub>"}\\
@{text "\<forall>x\<^isub>1 x\<^isub>2. y = PTup\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 \<Rightarrow> P\<^bsub>pat\<^esub>"}
\end{array}}
\end{tabular}}
\]\smallskip
\noindent
They are stronger in the sense that they allow us to assume in the @{text
"Lam\<^sup>\<alpha>"} and @{text "Let_pat\<^sup>\<alpha>"} cases that the bound atoms
avoid, or are fresh for, a context @{text "c"} (which is assumed to be finitely supported).
These stronger cases lemmas can be derived from the `weak' cases lemmas
given in \eqref{inductex}. This is trivial in case of patterns (the one on
the right-hand side) since the weak and strong cases lemma coincide (there
is no binding in patterns). Interesting are only the cases for @{text
"Lam\<^sup>\<alpha>"} and @{text "Let_pat\<^sup>\<alpha>"}, where we have some binders and
therefore have an additional assumption about avoiding @{text "c"}. Let us
first establish the case for @{text "Lam\<^sup>\<alpha>"}. By the weak cases lemma
\eqref{inductex} we can assume that
\begin{equation}\label{assm}
@{text "y = Lam\<^sup>\<alpha> x\<^isub>1 x\<^isub>2"}
\end{equation}\smallskip
\noindent
holds, and need to establish @{text "P\<^bsub>trm\<^esub>"}. The stronger cases lemma has the
corresponding implication
\begin{equation}\label{imp}
@{text "\<forall>x\<^isub>1 x\<^isub>2. atom x\<^isub>1 # c \<and> y = Lam\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 \<Rightarrow> P\<^bsub>trm\<^esub>"}
\end{equation}\smallskip
\noindent
which we must use in order to infer @{text "P\<^bsub>trm\<^esub>"}. Clearly, we cannot
use this implication directly, because we have no information whether or not @{text
"x\<^isub>1"} is fresh for @{text "c"}. However, we can use Properties
\ref{supppermeq} and \ref{avoiding} to rename @{text "x\<^isub>1"}. We know
by Theorem~\ref{suppfa} that @{text "{atom x\<^isub>1} #\<^sup>* Lam\<^sup>\<alpha>
x\<^isub>1 x\<^isub>2"} (since its support is @{text "supp x\<^isub>2 -
{atom x\<^isub>1}"}). Property \ref{avoiding} provides us then with a
permutation @{text "\<pi>"}, such that @{text "{atom (\<pi> \<bullet> x\<^isub>1)} #\<^sup>*
c"} and \mbox{@{text "supp (Lam\<^sup>\<alpha> x\<^isub>1 x\<^isub>2) #\<^sup>* \<pi>"}} hold.
By using Property \ref{supppermeq}, we can infer from the latter that
\[
@{text "Lam\<^sup>\<alpha> (\<pi> \<bullet> x\<^isub>1) (\<pi> \<bullet> x\<^isub>2) = Lam\<^sup>\<alpha> x\<^isub>1 x\<^isub>2"}
\]\smallskip
\noindent
holds. We can use this equation in the assumption \eqref{assm}, and hence
use the implication \eqref{imp} with the renamed @{text "\<pi> \<bullet> x\<^isub>1"}
and @{text "\<pi> \<bullet> x\<^isub>2"} for concluding this case.
The @{text "Let_pat\<^sup>\<alpha>"}-case involving a deep binder is slightly more complicated.
We have the assumption
\begin{equation}\label{assmtwo}
@{text "y = Let_pat\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 x\<^isub>3"}
\end{equation}\smallskip
\noindent
and the implication from the stronger cases lemma
\begin{equation}\label{impletpat}
@{text "\<forall>x\<^isub>1 x\<^isub>2 x\<^isub>3. set (bn\<^sup>\<alpha> x\<^isub>1) #\<^sup>* c \<and> y = Let_pat\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 x\<^isub>3 \<Rightarrow> P\<^bsub>trm\<^esub>"}
\end{equation}\smallskip
\noindent
The reason that this case is more complicated is that we cannot directly apply Property
\ref{avoiding} for obtaining a renaming permutation. Property \ref{avoiding} requires
that the binders are fresh for the term in which we want to perform the renaming. But
this is not true in terms such as (using an informal notation)
\[
@{text "Let (x, y) := (x, y) in (x, y)"}
\]\smallskip
\noindent
where @{text x} and @{text y} are bound in the term, but are also free
in the right-hand side of the assignment. We can, however, obtain such a renaming permutation, say
@{text "\<pi>"}, for the abstraction @{term "Abs_lst (bn_al x\<^isub>1)
x\<^isub>3"}. As a result we have \mbox{@{term "set (bn_al (\<pi> \<bullet> x\<^isub>1))
\<sharp>* c"}} and @{term "Abs_lst (bn_al (\<pi> \<bullet> x\<^isub>1)) (\<pi> \<bullet> x\<^isub>3) =
Abs_lst (bn_al x\<^isub>1) x\<^isub>3"} (remember @{text "set"} and @{text
"bn\<^sup>\<alpha>"} are equivariant). Now the quasi-injective property for @{text
"Let_pat\<^sup>\<alpha>"} states that
\[
\infer{@{text "Let_pat\<^sup>\<alpha> p t\<^isub>1 t\<^isub>2 = Let_pat\<^sup>\<alpha> p\<PRIME> t\<PRIME>\<^isub>1 t\<PRIME>\<^isub>2"}}
{@{text "[bn\<^sup>\<alpha> p]\<^bsub>list\<^esub>. t\<^isub>2 = [bn\<^sup>\<alpha> p']\<^bsub>list\<^esub>. t\<PRIME>\<^isub>2"}\;\; &
@{text "p \<approx>\<AL>\<^bsub>bn\<^esub> p\<PRIME>"}\;\; & @{text "t\<^isub>1 = t\<PRIME>\<^isub>1"}}
\]\smallskip
\noindent
Since all atoms in a pattern are bound by @{text "Let_pat\<^sup>\<alpha>"}, we can infer
that @{text "(\<pi> \<bullet> x\<^isub>1) \<approx>\<AL>\<^bsub>bn\<^esub> x\<^isub>1"} holds for every @{text "\<pi>"}. Therefore we have that
\[
@{text "Let_pat\<^sup>\<alpha> (\<pi> \<bullet> x\<^isub>1) x\<^isub>2 (\<pi> \<bullet> x\<^isub>3) = Let_pat\<^sup>\<alpha> x\<^isub>1 x\<^isub>2 x\<^isub>3"}
\]\smallskip
\noindent
Taking the left-hand side in the assumption shown in \eqref{assmtwo}, we can use
the implication \eqref{impletpat} from the stronger cases lemma to infer @{text "P\<^bsub>trm\<^esub>"}, as needed.
The remaining difficulty is when a deep binder contains some atoms that are
bound and some that are free. An example is @{text "Let\<^sup>\<alpha>"} in
\eqref{letrecs}. In such cases @{text "(\<pi> \<bullet> x\<^isub>1)
\<approx>\<AL>\<^bsub>bn\<^esub> x\<^isub>1"} does not hold in general. The idea however is
that @{text "\<pi>"} only renames atoms that become bound. In this way @{text "\<pi>"}
does not affect @{text "\<approx>\<AL>\<^bsub>bn\<^esub>"} (which only tracks alpha-equivalence of terms that are not
under the binder). However, the problem is that the
permutation operation @{text "\<pi> \<bullet> x\<^isub>1"} applies to all atoms in @{text "x\<^isub>1"}. To avoid this
we introduce an auxiliary permutation operations, written @{text "_
\<bullet>\<^bsub>bn\<^esub> _"}, for deep binders that only permutes bound atoms (or
more precisely the atoms specified by the @{text "bn"}-functions) and leaves
the other atoms unchanged. Like the functions @{text "fa_bn"}$_{1..m}$, we
can define these permutation operations over raw terms analysing how the functions @{text
"bn"}$_{1..m}$ are defined. Assuming the user specified a clause
\[
@{text "bn (C x\<^isub>1 \<dots> x\<^isub>r) = rhs"}
\]\smallskip
\noindent
we define @{text "\<pi> \<bullet>\<^bsub>bn\<^esub> (C x\<^isub>1 \<dots> x\<^isub>r) \<equiv> C y\<^isub>1 \<dots> y\<^isub>r"} with @{text "y\<^isub>i"} determined as follows:
\[\mbox{
\begin{tabular}{c@ {\hspace{2mm}}p{0.9\textwidth}}
$\bullet$ & @{text "y\<^isub>i \<equiv> x\<^isub>i"} provided @{text "x\<^isub>i"} does not occur in @{text "rhs"}\\
$\bullet$ & @{text "y\<^isub>i \<equiv> \<pi> \<bullet>\<^bsub>bn\<^esub> x\<^isub>i"} provided @{text "bn x\<^isub>i"} is in @{text "rhs"}\\
$\bullet$ & @{text "y\<^isub>i \<equiv> \<pi> \<bullet> x\<^isub>i"} otherwise
\end{tabular}}
\]\smallskip
\noindent
Using again the quotient package we can lift the auxiliary permutation operations
@{text "_ \<bullet>\<^bsub>bn\<^esub> _"}
to alpha-equated terms. Moreover we can prove the following two properties:
\begin{lem}\label{permutebn}
Given a binding function @{text "bn\<^sup>\<alpha>"} and auxiliary equivalence @{text "\<approx>\<AL>\<^bsub>bn\<^esub>"}
then for all @{text "\<pi>"}\smallskip\\
{\it (i)} @{text "\<pi> \<bullet> (bn\<^sup>\<alpha> x) = bn\<^sup>\<alpha> (\<pi> \<bullet>\<AL>\<^bsub>bn\<^esub> x)"} and\\
{\it (ii)} @{text "(\<pi> \<bullet>\<AL>\<^bsub>bn\<^esub> x) \<approx>\<AL>\<^bsub>bn\<^esub> x"}.
\end{lem}
\begin{proof}
By induction on @{text x}. The properties follow by unfolding of the
definitions.
\end{proof}
\noindent
The first property states that a permutation applied to a binding function
is equivalent to first permuting the binders and then calculating the bound
atoms. The second states that @{text "_ \<bullet>\<AL>\<^bsub>bn\<^esub> _"} preserves
@{text "\<approx>\<AL>\<^bsub>bn\<^esub>"}. The main point of the auxiliary
permutation functions is that they allow us to rename just the bound atoms in a
term, without changing anything else.
Having the auxiliary permutation function in place, we can now solve all remaining cases.
For the @{text "Let\<^sup>\<alpha>"} term-constructor, for example, we can by Property \ref{avoiding}
obtain a @{text "\<pi>"} such that
\[
@{text "(\<pi> \<bullet> (set (bn\<^sup>\<alpha> x\<^isub>1)) #\<^sup>* c"} \hspace{10mm}
@{text "\<pi> \<bullet> [bn\<^sup>\<alpha> x\<^isub>1]\<^bsub>list\<^esub>. x\<^isub>2 = [bn\<^sup>\<alpha> x\<^isub>1]\<^bsub>list\<^esub>. x\<^isub>2"}
\]\smallskip
\noindent
hold. Using the first part of Lemma \ref{permutebn}, we can simplify this
to @{text "set (bn\<^sup>\<alpha> (\<pi> \<bullet>\<AL>\<^bsub>bn\<^esub> x\<^isub>1)) #\<^sup>* c"} and
\mbox{@{text "[bn\<^sup>\<alpha> (\<pi> \<bullet>\<AL>\<^bsub>bn\<^esub> x\<^isub>1)]\<^bsub>list\<^esub>. (\<pi> \<bullet> x\<^isub>2) = [bn\<^sup>\<alpha> x\<^isub>1]\<^bsub>list\<^esub>. x\<^isub>2"}}. Since
@{text "(\<pi> \<bullet>\<AL>\<^bsub>bn\<^esub> x\<^isub>1) \<approx>\<AL>\<^bsub>bn\<^esub> x\<^isub>1"} holds by the second part,
we can infer that
\[
@{text "Let\<^sup>\<alpha> (\<pi> \<bullet>\<AL>\<^bsub>bn\<^esub> x\<^isub>1) (\<pi> \<bullet> x\<^isub>2) = Let\<^sup>\<alpha> x\<^isub>1 x\<^isub>2"}
\]\smallskip
\noindent
holds. This allows us to use the implication from the strong cases
lemma, and we are done.
Consequently, we can discharge all proof-obligations about having covered all
cases. This completes the proof establishing that the weak induction principles imply
the strong induction principles. These strong induction principles have already proved
being very useful in practice, particularly for proving properties about
capture-avoiding substitution \cite{Urban08}.
*}
section {* Related Work\label{related} *}
text {*
To our knowledge the earliest usage of general binders in a theorem prover
is described by Nara\-schew\-ski and Nipkow \cite{NaraschewskiNipkow99} with a
formalisation of the algorithm W. This formalisation implements binding in
type-schemes using a de-Bruijn indices representation. Since type-schemes in
W contain only a single place where variables are bound, different indices
do not refer to different binders (as in the usual de-Bruijn
representation), but to different bound variables. A similar idea has been
recently explored for general binders by Chargu\'eraud \cite{chargueraud09}
in the locally nameless approach to
binding. There, de-Bruijn indices consist of two
numbers, one referring to the place where a variable is bound, and the other
to which variable is bound. The reasoning infrastructure for both
representations of bindings comes for free in theorem provers like
Isabelle/HOL and Coq, since the corresponding term-calculi can be implemented
as `normal' datatypes. However, in both approaches it seems difficult to
achieve our fine-grained control over the `semantics' of bindings
(i.e.~whether the order of binders should matter, or vacuous binders should
be taken into account). To do so, one would require additional predicates
that filter out unwanted terms. Our guess is that such predicates result in
rather intricate formal reasoning. We are not aware of any formalisation of
a non-trivial language that uses Chargu\'eraud's idea.
Another technique for representing binding is higher-order abstract syntax
(HOAS), which for example is implemented in the Twelf system \cite{pfenningsystem}.
This representation technique supports very elegantly many aspects of
\emph{single} binding, and impressive work by Lee et al~\cite{LeeCraryHarper07}
has been done that uses HOAS for mechanising the metatheory of SML. We
are, however, not aware how multiple binders of SML are represented in this
work. Judging from the submitted Twelf-solution for the POPLmark challenge,
HOAS cannot easily deal with binding constructs where the number of bound
variables is not fixed. For example, in the second part of this challenge,
@{text "Let"}s involve patterns that bind multiple variables at once. In
such situations, HOAS seems to have to resort to the
iterated-single-binders-approach with all the unwanted consequences when
reasoning about the resulting terms.
Two formalisations involving general binders have been
performed in older
versions of Nominal Isabelle (one about Psi-calculi and one about algorithm W
\cite{BengtsonParow09,UrbanNipkow09}). Both
use the approach based on iterated single binders. Our experience with
the latter formalisation has been disappointing. The major pain arose from
the need to `unbind' variables. This can be done in one step with our
general binders described in this paper, but needs a cumbersome
iteration with single binders. The resulting formal reasoning turned out to
be rather unpleasant.
The most closely related work to the one presented here is the Ott-tool by
Sewell et al \cite{ott-jfp} and the C$\alpha$ml language by Pottier
\cite{Pottier06}. Ott is a nifty front-end for creating \LaTeX{} documents
from specifications of term-calculi involving general binders. For a subset
of the specifications Ott can also generate theorem prover code using a `raw'
representation of terms, and in Coq also a locally nameless
representation. The developers of this tool have also put forward (on paper)
a definition for alpha-equivalence and free variables for terms that can be
specified in Ott. This definition is rather different from ours, not using
any nominal techniques. To our knowledge there is no concrete mathematical
result concerning this notion of alpha-equivalence and free variables. We
have proved that our definitions lead to alpha-equated terms, whose support
is as expected (that means bound atoms are removed from the support). We
also showed that our specifications lift from `raw' types to types of
alpha-equivalence classes. For this we have established (automatically) that every
term-constructor and function defined for `raw' types
is respectful w.r.t.~alpha-equivalence.
Although we were heavily inspired by the syntax of Ott, its definition of
alpha-equi\-valence is unsuitable for our extension of Nominal
Isabelle. First, it is far too complicated to be a basis for automated
proofs implemented on the ML-level of Isabelle/HOL. Second, it covers cases
of binders depending on other binders, which just do not make sense for our
alpha-equated terms. Third, it allows empty types that have no meaning in a
HOL-based theorem prover. We also had to generalise slightly Ott's binding
clauses. In Ott one specifies binding clauses with a single body; we allow
more than one. We have to do this, because this makes a difference for our
notion of alpha-equivalence in case of \isacommand{binds (set)} and
\isacommand{binds (set+)}. Consider the examples
\[\mbox{
\begin{tabular}{@ {}l@ {\hspace{2mm}}l@ {}}
@{text "Foo\<^isub>1 xs::name fset t::trm s::trm"} &
\isacommand{binds (set)} @{text "xs"} \isacommand{in} @{text "t s"}\\
@{text "Foo\<^isub>2 xs::name fset t::trm s::trm"} &
\isacommand{binds (set)} @{text "xs"} \isacommand{in} @{text "t"},
\isacommand{binds (set)} @{text "xs"} \isacommand{in} @{text "s"}\\
\end{tabular}}
\]\smallskip
\noindent
In the first term-constructor we have a single body that happens to be
`spread' over two arguments; in the second term-constructor we have two
independent bodies in which the same variables are bound. As a result we
have\footnote{Assuming @{term "a \<noteq> b"}, there is no permutation that can
make @{text "(a, b)"} equal with both @{text "(a, b)"} and @{text "(b, a)"}, but
there are two permutations so that we can make @{text "(a, b)"} and @{text
"(a, b)"} equal with one permutation, and @{text "(a, b)"} and @{text "(b,
a)"} with the other.}
\[\mbox{
\begin{tabular}{r@ {\hspace{1.5mm}}c@ {\hspace{1.5mm}}l}
@{text "Foo\<^isub>1 {a, b} (a, b) (a, b)"} & $\not=$ &
@{text "Foo\<^isub>1 {a, b} (a, b) (b, a)"}
\end{tabular}}
\]\smallskip
\noindent
but
\[\mbox{
\begin{tabular}{r@ {\hspace{1.5mm}}c@ {\hspace{1.5mm}}l}
@{text "Foo\<^isub>2 {a, b} (a, b) (a, b)"} & $=$ &
@{text "Foo\<^isub>2 {a, b} (a, b) (b, a)"}\\
\end{tabular}}
\]\smallskip
\noindent
and therefore need the extra generality to be able to distinguish between
both specifications. Because of how we set up our definitions, we also had
to impose some restrictions (like a single binding function for a deep
binder) that are not present in Ott. Our expectation is that we can still
cover many interesting term-calculi from programming language research, for
example the Core-Haskell language from the Introduction. With the work
presented in this paper we can define it formally as shown in
Figure~\ref{nominalcorehas} and then Nominal Isabelle derives automatically
a corresponding reasoning infrastructure.
\begin{figure}[p]
\begin{boxedminipage}{\linewidth}
\small
\begin{tabular}{l}
\isacommand{atom\_decl}~@{text "var cvar tvar"}\\[1mm]
\isacommand{nominal\_datatype}~@{text "tkind ="}~@{text "KStar"}~$|$~@{text "KFun tkind tkind"}\\
\isacommand{and}~@{text "ckind ="}~@{text "CKSim ty ty"}\\
\isacommand{and}~@{text "ty ="}~@{text "TVar tvar"}~$|$~@{text "T string"}~$|$~@{text "TApp ty ty"}\\
$|$~@{text "TFun string ty_list"}~%
$|$~@{text "TAll tv::tvar tkind ty::ty"}\hspace{3mm}\isacommand{binds}~@{text "tv"}~\isacommand{in}~@{text ty}\\
$|$~@{text "TArr ckind ty"}\\
\isacommand{and}~@{text "ty_lst ="}~@{text "TNil"}~$|$~@{text "TCons ty ty_lst"}\\
\isacommand{and}~@{text "cty ="}~@{text "CVar cvar"}~%
$|$~@{text "C string"}~$|$~@{text "CApp cty cty"}~$|$~@{text "CFun string co_lst"}\\
$|$~@{text "CAll cv::cvar ckind cty::cty"}\hspace{3mm}\isacommand{binds}~@{text "cv"}~\isacommand{in}~@{text cty}\\
$|$~@{text "CArr ckind cty"}~$|$~@{text "CRefl ty"}~$|$~@{text "CSym cty"}~$|$~@{text "CCirc cty cty"}\\
$|$~@{text "CAt cty ty"}~$|$~@{text "CLeft cty"}~$|$~@{text "CRight cty"}~$|$~@{text "CSim cty cty"}\\
$|$~@{text "CRightc cty"}~$|$~@{text "CLeftc cty"}~$|$~@{text "Coerce cty cty"}\\
\isacommand{and}~@{text "co_lst ="}~@{text "CNil"}~$|$~@{text "CCons cty co_lst"}\\
\isacommand{and}~@{text "trm ="}~@{text "Var var"}~$|$~@{text "K string"}\\
$|$~@{text "LAM_ty tv::tvar tkind t::trm"}\hspace{3mm}\isacommand{binds}~@{text "tv"}~\isacommand{in}~@{text t}\\
$|$~@{text "LAM_cty cv::cvar ckind t::trm"}\hspace{3mm}\isacommand{binds}~@{text "cv"}~\isacommand{in}~@{text t}\\
$|$~@{text "App_ty trm ty"}~$|$~@{text "App_cty trm cty"}~$|$~@{text "App trm trm"}\\
$|$~@{text "Lam v::var ty t::trm"}\hspace{3mm}\isacommand{binds}~@{text "v"}~\isacommand{in}~@{text t}\\
$|$~@{text "Let x::var ty trm t::trm"}\hspace{3mm}\isacommand{binds}~@{text x}~\isacommand{in}~@{text t}\\
$|$~@{text "Case trm assoc_lst"}~$|$~@{text "Cast trm co"}\\
\isacommand{and}~@{text "assoc_lst ="}~@{text ANil}~%
$|$~@{text "ACons p::pat t::trm assoc_lst"}\hspace{3mm}\isacommand{binds}~@{text "bv p"}~\isacommand{in}~@{text t}\\
\isacommand{and}~@{text "pat ="}~@{text "Kpat string tvtk_lst tvck_lst vt_lst"}\\
\isacommand{and}~@{text "vt_lst ="}~@{text VTNil}~$|$~@{text "VTCons var ty vt_lst"}\\
\isacommand{and}~@{text "tvtk_lst ="}~@{text TVTKNil}~$|$~@{text "TVTKCons tvar tkind tvtk_lst"}\\
\isacommand{and}~@{text "tvck_lst ="}~@{text TVCKNil}~$|$ @{text "TVCKCons cvar ckind tvck_lst"}\\
\isacommand{binder}\\
@{text "bv :: pat \<Rightarrow> atom list"}~\isacommand{and}\\
@{text "bv\<^isub>1 :: vt_lst \<Rightarrow> atom list"}~\isacommand{and}\\
@{text "bv\<^isub>2 :: tvtk_lst \<Rightarrow> atom list"}~\isacommand{and}\\
@{text "bv\<^isub>3 :: tvck_lst \<Rightarrow> atom list"}\\
\isacommand{where}\\
\phantom{$|$}~@{text "bv (K s tvts tvcs vs) = (bv\<^isub>3 tvts) @ (bv\<^isub>2 tvcs) @ (bv\<^isub>1 vs)"}\\
$|$~@{text "bv\<^isub>1 VTNil = []"}\\
$|$~@{text "bv\<^isub>1 (VTCons x ty tl) = (atom x)::(bv\<^isub>1 tl)"}\\
$|$~@{text "bv\<^isub>2 TVTKNil = []"}\\
$|$~@{text "bv\<^isub>2 (TVTKCons a ty tl) = (atom a)::(bv\<^isub>2 tl)"}\\
$|$~@{text "bv\<^isub>3 TVCKNil = []"}\\
$|$~@{text "bv\<^isub>3 (TVCKCons c cty tl) = (atom c)::(bv\<^isub>3 tl)"}\\
\end{tabular}
\end{boxedminipage}
\caption{A definition for Core-Haskell in Nominal Isabelle. For the moment we
do not support nested types; therefore we explicitly have to unfold the
lists @{text "co_lst"}, @{text "assoc_lst"} and so on. Apart from that limitation, the
definition follows closely the original shown in Figure~\ref{corehas}. The
point of our work is that having made such a definition in Nominal Isabelle,
one obtains automatically a reasoning infrastructure for Core-Haskell.
\label{nominalcorehas}}
\end{figure}
\afterpage{\clearpage}
Pottier presents a programming language, called C$\alpha$ml, for
representing terms with general binders inside OCaml \cite{Pottier06}. This
language is implemented as a front-end that can be translated to OCaml with
the help of a library. He presents a type-system in which the scope of
general binders can be specified using special markers, written @{text
"inner"} and @{text "outer"}. It seems our and his specifications can be
inter-translated as long as ours use the binding mode \isacommand{binds}
only. However, we have not proved this. Pottier gives a definition for
alpha-equivalence, which also uses a permutation operation (like ours).
Still, this definition is rather different from ours and he only proves that
it defines an equivalence relation. A complete reasoning infrastructure is
well beyond the purposes of his language. Similar work for Haskell with
similar results was reported by Cheney \cite{Cheney05a} and more recently
by Weirich et al \cite{WeirichYorgeySheard11}.
In a slightly different domain (programming with dependent types),
Altenkirch et al \cite{Altenkirch10} present a calculus with a notion of
alpha-equivalence related to our binding mode \isacommand{binds (set+)}.
Their definition is similar to the one by Pottier, except that it has a more
operational flavour and calculates a partial (renaming) map. In this way,
the definition can deal with vacuous binders. However, to our best
knowledge, no concrete mathematical result concerning this definition of
alpha-equivalence has been proved.
*}
section {* Conclusion *}
text {*
We have presented an extension of Nominal Isabelle for dealing with general
binders, that is where term-constructors have multiple bound atoms. For this
extension we introduced new definitions of alpha-equivalence and automated
all necessary proofs in Isabelle/HOL. To specify general binders we used
the syntax from Ott, but extended it in some places and restricted
it in others so that the definitions make sense in the context of alpha-equated
terms. We also introduced two binding modes (set and set+) that do not exist
in Ott. We have tried out the extension with calculi such as Core-Haskell,
type-schemes and approximately a dozen of other typical examples from
programming language research~\cite{SewellBestiary}. The code will
eventually become part of the Isabelle distribution.\footnote{It
can be downloaded already from \href{http://isabelle.in.tum.de/nominal/download}
{http://isabelle.in.tum.de/nominal/download}.}
We have left out a discussion about how functions can be defined over
alpha-equated terms involving general binders. In earlier versions of
Nominal Isabelle this turned out to be a thorny issue. We hope to do better
this time by using the function package \cite{Krauss09} that has recently
been implemented in Isabelle/HOL and also by restricting function
definitions to equivariant functions (for them we can provide more
automation).
There are some restrictions we imposed in this paper that we would like to lift in
future work. One is the exclusion of nested datatype definitions. Nested
datatype definitions would allow one to specify, for instance, the function kinds
in Core-Haskell as @{text "TFun string (ty list)"} instead of the unfolded
version @{text "TFun string ty_list"} (see Figure~\ref{nominalcorehas}). To
achieve this, we need more clever implementation than we have
at the moment. However, really lifting this restriction will involve major
work on the underlying datatype package of Isabelle/HOL.
A more interesting line of investigation is whether we can go beyond the
simple-minded form of binding functions that we adopted from Ott. At the moment, binding
functions can only return the empty set, a singleton atom set or unions
of atom sets (similarly for lists). It remains to be seen whether
properties like
\[
\mbox{@{text "fa_ty x = bn x \<union> fa_bn x"}}
\]\smallskip
\noindent
allow us to support more interesting binding functions.
We have also not yet played with other binding modes. For example we can
imagine that there is need for a binding mode where instead of usual lists,
we abstract lists of distinct elements (the corresponding type @{text
"dlist"} already exists in the library of Isabelle/HOL). We expect the
presented work can be extended to accommodate such binding modes.\medskip
\noindent
{\bf Acknowledgements:} We are very grateful to Andrew Pitts for many
discussions about Nominal Isabelle. We thank Peter Sewell for making the
informal notes \cite{SewellBestiary} available to us and also for patiently
explaining some of the finer points of the Ott-tool. Stephanie Weirich
suggested to separate the subgrammars of kinds and types in our Core-Haskell
example. Ramana Kumar and Andrei Popescu helped us with comments about
an earlier version of this paper.
*}
(*<*)
end
(*>*)