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\newlength{\sk}
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\setlength{\sk}{-1.9pt}
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In this section we formalize the ideas we outlined in the introduction. We start by the definition of types followed by the language and its reduction semantics. The static semantics is the core of our work: we first present a declarative type system that deduces (possibly many) types for well-typed expressions and then the algorithms to decide whether an expression is well typed or not. 
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\subsection{Types}

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\begin{definition}[Types]\label{def:types} The set of types \types{} is formed by the terms $t$ coinductively produced by the grammar:\vspace{-1.45mm}
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\[
\begin{array}{lrcl}
\textbf{Types} & t & ::= & b\alt t\to t\alt t\times t\alt t\vee t \alt \neg t \alt \Empty 
\end{array}
\]
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and that satisfy the following conditions\vspace{-.85mm}
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\begin{itemize}
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\item (regularity) every term has a finite number of different sub-terms;
\item (contractivity) every infinite branch of a term contains an infinite number of occurrences of the
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arrow or product type constructors.\vspace{-1mm}
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\end{itemize}
\end{definition}
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We use the following abbreviations: $
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    t_1 \land t_2 \eqdef \neg (\neg t_1 \vee \neg t_2)$, 
    $t_ 1 \setminus t_2 \eqdef t_1 \wedge \neg t_2$, $\Any \eqdef \neg \Empty$.
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$b$ ranges over basic types
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(e.g., \Int, \Bool),
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$\Empty$ and $\Any$ respectively denote the empty (that types no value)
and top (that types all values) types. Coinduction accounts for
recursive types and the condition on infinite branches bars out
ill-formed types such as 
$t = t \lor t$ (which does not carry any information about the set
denoted by the type) or $t = \neg t$ (which cannot represent any
set). 
It also ensures that the binary relation $\vartriangleright
\,\subseteq\!\types^{2}$ defined by $t_1 \lor t_2 \vartriangleright
t_i$, $t_1 \land t_2 \vartriangleright
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t_i$, $\neg t \vartriangleright t$ is Noetherian.
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This gives an induction principle on $\types$ that we
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will use without any further explicit reference to the relation.\footnote{In a nutshell, we can do proofs by induction on the structure of unions and negations---and, thus, intersections---but arrows, products, and basic types are the base cases for the induction.} 
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We refer to $ b $, $\times$, and $ \to $ as \emph{type constructors}
and to $ \lor $, $ \land $, $ \lnot $, and $ \setminus $
as \emph{type connectives}.

The subtyping relation for these types, noted $\leq$, is the one defined
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by~\citet{Frisch2008} to which the reader may refer. A detailed description of the algorithm to decide it can be found in~\cite{Cas15}.
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For this presentation it suffices to consider that
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types are interpreted as sets of \emph{values} ({i.e., either
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constants, $\lambda$-abstractions, or pairs of values: see
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Section~\ref{sec:syntax} right below) that have that type, and that subtyping is set
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containment (i.e., a type $s$ is a subtype of a type $t$ if and only if $t$
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contains all the values of type $s$). In particular, $s\to t$
contains all $\lambda$-abstractions that when applied to a value of
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type $s$, if their computation terminates, then they return a result of
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type $t$ (e.g., $\Empty\to\Any$ is the set of all
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functions\footnote{\label{allfunctions}Actually, for every type $t$,
all types of the form $\Empty{\to}t$ are equivalent and each of them
denotes the set of all functions.} and $\Any\to\Empty$ is the set
of functions that diverge on every argument). Type connectives
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(i.e., union, intersection, negation) are interpreted as the
corresponding set-theoretic operators (e.g.,~$s\vee t$ is the
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union of the values of the two types). We use $\simeq$ to denote the
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symmetric closure of $\leq$: thus $s\simeq t$ (read, $s$ is equivalent to $t$) means that $s$ and $t$ denote the same set of values and, as such, they are semantically the same type.
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\subsection{Syntax}\label{sec:syntax}
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The expressions $e$ and values $v$ of our language are inductively generated by the following grammars:\vspace{-1mm}
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\begin{equation}\label{expressions}
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\begin{array}{lrclr}  
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  \textbf{Expr} &e &::=& c\alt x\alt ee\alt\lambda^{\wedge_{i\in I}s_i\to t_i} x.e\alt \pi_j e\alt(e,e)\alt\tcase{e}{t}{e}{e}\\[.3mm]
  \textbf{Values} &v &::=& c\alt\lambda^{\wedge_{i\in I}s_i\to t_i} x.e\alt (v,v)\\[-1mm]
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\end{array}
\end{equation}
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for $j=1,2$. In~\eqref{expressions}, $c$ ranges over constants
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(e.g., \texttt{true}, \texttt{false}, \texttt{1}, \texttt{2},
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...) which are values of basic types (we use $\basic{c}$ to denote the
basic type of the constant $c$); $x$ ranges over variables; $(e,e)$
denote pairs and $\pi_i e$ their projections; $\tcase{e}{t}{e_1}{e_2}$
denotes the type-case expression that evaluates either $e_1$ or $e_2$
according to whether the value returned by $e$ (if any) is of type $t$
or not; $\lambda^{\wedge_{i\in I}s_i\to t_i} x.e$ is a value of type
$\wedge_{i\in I}s_i\to t_i$ and denotes the function of parameter $x$
and body $e$. An expression has an intersection type if and only if it
has all the types that compose the intersection. Therefore,
intuitively, $\lambda^{\wedge_{i\in I}s_i\to t_i} x.e$ is a well-typed
value if for all $i{\in} I$ the hypothesis that $x$ is of type $s_i$
implies that the body $e$ has type $t_i$, that is to say, it is well
typed if $\lambda^{\wedge_{i\in I}s_i\to t_i} x.e$ has type $s_i\to
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t_i$ for all $i\in I$. Every value is associated to a type: the type of $c$ is $\basic c$; the type of
 $\lambda^{\wedge_{i\in I}s_i\to t_i} x.e$ is $\wedge_{i\in I}s_i\to t_i$; and, inductively,
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the type of a pair of values is the product of the types of the
values.
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\subsection{Dynamic semantics}\label{sec:opsem}
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The dynamic semantics is defined as a classic left-to-right call-by-value reduction for a $\lambda$-calculus with pairs, enriched with specific rules for type-cases. We have the following  notions of reduction:\vspace{-1.2mm}
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\[
\begin{array}{rcll}
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  (\lambda^{\wedge_{i\in I}s_i\to t_i} x.e)v &\reduces& e\subst x v\\[-.4mm]
  \pi_i(v_1,v_2) &\reduces& v_i & i=1,2\\[-.4mm]
  \tcase{v}{t}{e_1}{e_2} &\reduces& e_1 &v\in t\\[-.4mm] 
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  \tcase{v}{t}{e_1}{e_2} &\reduces& e_2 &v\not\in t\\[-1.3mm]
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\end{array}
\]
The semantics of type-cases uses the relation $v\in t$ that we
informally defined in the previous section. We delay its formal
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definition to Section~\ref{sec:type-schemes} (where it deals with some corner cases for negated arrow types). Contextual reductions are
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defined by the following evaluation contexts:\\[1mm]
\centerline{\(
%\[
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\Cx[] ::= [\,]\alt \Cx e\alt v\Cx \alt (\Cx,e)\alt (v,\Cx)\alt \pi_i\Cx\alt \tcase{\Cx}tee
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%\]
\)}\\[1mm]
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As usual we denote by $\Cx[e]$ the term obtained by replacing $e$ for
the hole in the context $\Cx$ and we have that $e\reduces e'$ implies
$\Cx[e]\reduces\Cx[e']$.
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\subsection{Static semantics}\label{sec:static}
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While the syntax and reduction semantics are, on the whole, pretty
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standard, for what concerns the type system we will have to introduce several
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unconventional features that we anticipated in
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Section~\ref{sec:challenges} and are at the core of our work. Let
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us start with the standard part, that is the typing of the functional
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core and the use of subtyping, given by the following typing rules:\vspace{-1mm}
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\begin{mathpar}
  \Infer[Const]
      { }
      {\Gamma\vdash c:\basic{c}}
      { }
  \quad
  \Infer[App]
      {
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        \Gamma \vdash e_1: \arrow {t_1}{t_2}\quad
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        \Gamma \vdash e_2: t_1
      }
      { \Gamma \vdash {e_1}{e_2}: t_2 }
      { }
  \quad
  \Infer[Abs+]
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      {{\scriptstyle\forall i\in I}\quad\Gamma,x:s_i\vdash e:t_i}
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      {
      \Gamma\vdash\lambda^{\wedge_{i\in I}\arrow {s_i} {t_i}}x.e:\textstyle \bigwedge_{i\in I}\arrow {s_i} {t_i}
      }
      { }
%      \Infer[If]
%            {\Gamma\vdash e:t_0\\
%            %t_0\not\leq \neg t \Rightarrow
%            \Gamma \cvdash + e t e_1:t'\\
%            %t_0\not\leq t \Rightarrow
%            \Gamma \cvdash - e t e_2:t'}
%            {\Gamma\vdash \ite {e} t {e_1}{e_2}: t'}
%            { }
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\vspace{-2mm} \\
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      \Infer[Sel]
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  {\Gamma \vdash e:\pair{t_1}{t_2}}
  {\Gamma \vdash \pi_i e:t_i}
  { }
  \qquad
  \Infer[Pair]
  {\Gamma \vdash e_1:t_1 \and \Gamma \vdash e_2:t_2}
  {\Gamma \vdash (e_1,e_2):\pair {t_1} {t_2}}
  { }
  \qquad
    \Infer[Subs]
      { \Gamma \vdash e:t\\t\leq t' }
      { \Gamma \vdash e: t' }
      { }
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  \qquad\vspace{-3mm}
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\end{mathpar}
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These rules are quite standard and do not need any particular explanation besides those already given in Section~\ref{sec:syntax}. Just notice that we used a classic subsumption rule (i.e., \Rule{Subs}) to embed subtyping in the type system. Let us next focus on the unconventional aspects of our system, from the simplest to the hardest.
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The first unconventional aspect is that, as explained in
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Section~\ref{sec:challenges}, our type assumptions are about
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expressions. Therefore, in our rules the type environments, ranged over
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by $\Gamma$, map \emph{expressions}---rather than just variables---into
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types. This explains why the classic typing rule for variables is replaced by a more general \Rule{Env} rule defined below:\vspace{-1mm}
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\begin{mathpar}
  \Infer[Env]
      { }
      { \Gamma \vdash e: \Gamma(e) }
      { e\in\dom\Gamma }
  \qquad
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  \Infer[Inter]
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      { \Gamma \vdash e:t_1\\\Gamma \vdash e:t_2 }
      { \Gamma \vdash e: t_1 \wedge t_2 }
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      { }\vspace{-3mm}
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\end{mathpar}
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The \Rule{Env} rule is coupled with the standard intersection introduction rule \Rule{Inter}
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which allows us to deduce for a complex expression the intersection of
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the types recorded by the occurrence typing analysis in the
environment $\Gamma$ with the static type deduced for the same
expression by using the other typing rules. This same intersection
rule is also used to infer the second unconventional aspect of our
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system, that is, the fact that $\lambda$-abstractions can have negated
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arrow types, as long as these negated types do not make the type deduced for the function empty:\vspace{-.5mm}
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\begin{mathpar}
  \Infer[Abs-]
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    {\Gamma \vdash \lambda^{\wedge_{i\in I}\arrow {s_i} {t_i}}x.e:t}
    { \Gamma \vdash\lambda^{\wedge_{i\in I}\arrow {s_i} {t_i}}x.e:\neg(t_1\to t_2)  }
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    { ((\wedge_{i\in I}\arrow {s_i} {t_i})\wedge\neg(t_1\to t_2))\not\simeq\Empty }\vspace{-1mm}
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\end{mathpar}
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%\beppe{I have doubt: is this safe or should we play it safer and
%  deduce $t\wedge\neg(t_1\to t_2)$? In other terms is is possible to
%  deduce two separate negation of arrow types that when intersected
%  with the interface are non empty, but by intersecting everything
%  makes the type empty? It should be safe since otherwise intersection
%  would not be admissible in semantic subtyping (see Theorem 6.15 in
%  JACM), but I think we should doube ckeck it.}
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As explained in Section~\ref{sec:challenges}, we need to be able to
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deduce for, say, the function $\lambda^{\Int\to\Int} x.x$ a type such
as $(\Int\to\Int)\wedge\neg(\Bool\to\Bool)$ (in particular, if this is
the term $e$ in equation \eqref{bistwo} we need to deduce for it the
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type $(\Int\to t)\wedge\neg(\Int\to\neg\Bool)$, that is,
$(\Int\to t)\setminus(\Int\to\neg\Bool)$ ). But the sole rule \Rule{Abs+}
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above does not allow us to deduce  negations of
arrows for abstractions: the rule \Rule{Abs-} makes this possible. As an aside, note that this kind
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of deduction was already present in the system by~\citet{Frisch2008}
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though in that system this presence was motivated by the semantics of types rather than, as in our case,
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by the soundness of the type system.

Rules \Rule{Abs+} and \Rule{Abs-} are not enough to deduce for
$\lambda$-abstractions all the types we wish. In particular, these
rules alone are not enough to type general overloaded functions. For
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instance, consider this simple example of a function that applied to an
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integer returns its successor and applied to anything else returns
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\textsf{true}:\\[1mm]
\centerline{\(
%\[
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\lambda^{(\Int\to\Int)\wedge(\neg\Int\to\Bool)} x\,.\,\tcase{x}{\Int}{x+1}{\textsf{true}}
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%\]
\)}\\[1mm]
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Clearly, the expression above is well typed, but the rule \Rule{Abs+} alone
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is not enough to type it. In particular, according to \Rule{Abs+} we
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have to prove that under the hypothesis that $x$ is of type $\Int$ the expression
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$(\tcase{x}{\Int}{x+1}{\textsf{true}})$ is of type $\Int$, too.  That is, that under the
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hypothesis that $x$ has type $\Int\wedge\Int$ (we apply occurrence
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typing) the expression $x+1$ is of type \Int{} (which holds) and that under the
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hypothesis that $x$ has type $\Int\setminus\Int$, that is $\Empty$
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(we apply once more occurrence typing), \textsf{true} is of type \Int{}
(which \emph{does not} hold). The problem is that we are trying to type the
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second case of a type-case even if we know that there is no chance that, when $x$ is bound to an integer,
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that case will be ever selected. The fact that it is never selected is witnessed
by the presence of a type hypothesis with  $\Empty$ type. To
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avoid this problem (and type the term above) we add the rule
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\Rule{Efq} (\emph{ex falso quodlibet}) that allows the system to deduce any type
for an expression that will never be selected, that is, for an
expression whose type environment contains an empty assumption:
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\begin{mathpar}
  \Infer[Efq]
  { }
  { \Gamma, (e:\Empty) \vdash e': t }
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  { }\vspace{-3mm}
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\end{mathpar}
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Once more, this kind of deduction was already present in the system
by~\citet{Frisch2008} to type full fledged overloaded functions,
though it was embedded in the typing rule for the type-case. Here we
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need the rule \Rule{Efq}, which is more general, to ensure the
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property of subject reduction.
%\beppe{Example?}
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Finally, there remains one last rule in our type system, the one that
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implements occurrence typing, that is, the rule for the
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type-case:\vspace{-1mm}
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\begin{mathpar}
    \Infer[Case]
        {\Gamma\vdash e:t_0\\
        %t_0\not\leq \neg t \Rightarrow
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        \Gamma \evdash e t \Gamma_1 \\ \Gamma_1 \vdash e_1:t'\\
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        %t_0\not\leq t \Rightarrow
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        \Gamma \evdash e {\neg t} \Gamma_2 \\ \Gamma_2 \vdash e_2:t'}
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        {\Gamma\vdash \tcase {e} t {e_1}{e_2}: t'}
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        { }\vspace{-3mm}
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\end{mathpar}
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The rule \Rule{Case} checks whether the expression $e$, whose type is
being tested, is well-typed and then performs the occurrence typing
analysis that produces the environments $\Gamma_i$'s under whose
hypothesis the expressions $e_i$'s are typed. The production of these
environments is represented by the judgments $\Gamma \evdash e
{(\neg)t} \Gamma_i$. The intuition is that when $\Gamma \evdash e t
\Gamma_1$ is provable then $\Gamma_1$ is a version of $\Gamma$
extended with type hypotheses for all expressions occurring in $e$,
type hypotheses that can be deduced assuming that the test $e\in t$
succeeds. Likewise, $\Gamma \evdash e {\neg t} \Gamma_2$ (notice the negation on $t$) extends
$\Gamma$ with the hypothesis deduced assuming that $e\in\neg t$, that
is, for when the test $e\in t$ fails.
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All it remains to do is to show how to deduce judgments of the form
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$\Gamma \evdash e t \Gamma'$. For that we first define how
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to denote occurrences of an expression. These are identified by paths in the
syntax tree of the expressions, that is, by possibly empty strings of
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characters denoting directions starting from the root of the tree (we
use $\epsilon$ for the empty string/path, which corresponds to the
root of the tree).

Let $e$ be an expression and $\varpi\in\{0,1,l,r,f,s\}^*$ a
\emph{path}; we denote $\occ e\varpi$ the occurrence of $e$ reached by
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the path $\varpi$, that is (for $i=1,2$, and undefined otherwise)\vspace{-.4mm}
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%% \[
%% \begin{array}{l}
%% \begin{array}{r@{\downarrow}l@{\quad=\quad}l}
%% e&\epsilon & e\\
%% e_0e_1& i.\varpi & \occ{e_i}\varpi\qquad i=0,1\\
%% (e_0,e_1)& l.\varpi & \occ{e_0}\varpi\\
%% (e_0,e_1)& r.\varpi & \occ{e_1}\varpi\\
%% \pi_1 e& f.\varpi & \occ{e}\varpi\\
%% \pi_2 e& s.\varpi & \occ{e}\varpi\\
%% \end{array}\\
%% \text{undefined otherwise}
%% \end{array}
%% \]
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\[
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\begin{array}{r@{\downarrow}l@{\quad=\quad}lr@{\downarrow}l@{\quad=\quad}lr@{\downarrow}l@{\quad=\quad}l}
e&\epsilon & e & (e_0,e_1)& l.\varpi & \occ{e_0}\varpi &\pi_1 e& f.\varpi & \occ{e}\varpi\\
e_0e_1& i.\varpi & \occ{e_i}\varpi \quad\qquad& (e_0,e_1)& r.\varpi & \occ{e_1}\varpi \quad\qquad&
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\pi_2 e& s.\varpi & \occ{e}\varpi\\[-.4mm]
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\end{array}
\]
To ease our analysis we used different directions for each kind of
term. So we have $0$ and $1$ for the function and argument of an
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application, $l$ and $r$ for the $l$eft and $r$ight expressions forming a pair,
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and $f$ and $s$ for the argument of a $f$irst or of a $s$econd projection. Note also that we do not consider occurrences
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under $\lambda$'s (since their type is frozen in their annotations) and type-cases (since they reset the analysis).
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%
The judgments  $\Gamma \evdash e t \Gamma'$ are then deduced by the following two rules:\vspace{-1mm} \begin{mathpar}
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%        \Infer[Base]
%            { \Gamma \vdash e':t' }
%            { \Gamma \cvdash p e t e':t' }
%            { }
%            \qquad
%        \Infer[Path]
%            { \pvdash \Gamma p e t \varpi:t_1 \\ \Gamma,(\occ e \varpi:t_1) \cvdash p e t e':t_2 }
%            { \Gamma \cvdash p e t e':t_2 }
%            { }
    \Infer[Base]
      { }
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      { \Gamma \evdash e t \Gamma }
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      { }
    \qquad
    \Infer[Path]
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      { \pvdash {\Gamma'} e t \varpi:t' \\ \Gamma \evdash e t \Gamma' }
      { \Gamma \evdash e t \Gamma',(\occ e \varpi:t') }
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      { }\vspace{-1.5mm}
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\end{mathpar}
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These rules describe how to produce by occurrence typing the type
environments while checking that an expression $e$ has type $t$. They state that $(i)$ we can
deduce from $\Gamma$ all the hypothesis already in $\Gamma$ (rule
\Rule{Base}) and that $(ii)$ if we can deduce a given type $t'$ for a particular
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occurrence $\varpi$ of the expression $e$ being checked, then we can add this
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hypothesis to the produced type environment (rule \Rule{Path}). The rule
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\Rule{Path} uses a (last) auxiliary judgement $\pvdash {\Gamma}  e t
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\varpi:t'$ to deduce the type $t'$ of the occurrence $\occ e \varpi$ when
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checking $e$ against $t$ under the hypotheses $\Gamma$. This rule \Rule{Path} is subtler than it may appear at
first sight, insofar as the deduction of the type for $\varpi$ may already use
some hypothesis on $\occ e \varpi$ (in $\Gamma'$) and, from an
algorithmic viewpoint, this will imply the computation of a fix-point
(see Section~\ref{sec:typenv}). The last ingredient for our type system is the deduction of the
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judgements of the form $\pvdash {\Gamma}  e t \varpi:t'$ where
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$\varpi$ is a path to an expression occurring in $e$. This is given by the following set
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of rules.
\begin{mathpar}
    \Infer[PSubs]
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        { \pvdash \Gamma e t \varpi:t_1 \\ t_1\leq t_2 }
        { \pvdash \Gamma e t \varpi:t_2 }
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        { }
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        \quad
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    \Infer[PInter]
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        { \pvdash \Gamma e t \varpi:t_1 \\ \pvdash \Gamma e t \varpi:t_2 }
        { \pvdash \Gamma e t \varpi:t_1\land t_2 }
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        { }
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        \quad
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    \Infer[PTypeof]
        { \Gamma \vdash \occ e \varpi:t' }
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        { \pvdash \Gamma e t \varpi:t' }
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        { }
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\vspace{-1.2mm}\\
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    \Infer[PEps]
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        { }
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        { \pvdash \Gamma e t \epsilon:t }
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        { }
        \qquad
    \Infer[PAppR]
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        { \pvdash \Gamma e t \varpi.0:\arrow{t_1}{t_2} \\ \pvdash \Gamma e t \varpi:t_2'}
        { \pvdash \Gamma e t \varpi.1:\neg t_1 }
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        { t_2\land t_2' \simeq \Empty  }
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\ifsubmission\vspace{-1.2mm}\\\else\end{mathpar}\begin{mathpar}\fi
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    \Infer[PAppL]
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        { \pvdash \Gamma e t \varpi.1:t_1 \\ \pvdash \Gamma e t \varpi:t_2 }
        { \pvdash \Gamma e t \varpi.0:\neg (\arrow {t_1} {\neg t_2}) }
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        { }
        \qquad
    \Infer[PPairL]
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        { \pvdash \Gamma e t \varpi:\pair{t_1}{t_2} }
        { \pvdash \Gamma e t \varpi.l:t_1 }
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        { }
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\vspace{-1.2mm}\\
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    \Infer[PPairR]
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        { \pvdash \Gamma e t \varpi:\pair{t_1}{t_2} }
        { \pvdash \Gamma e t \varpi.r:t_2 }
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        { }
        \qquad
    \Infer[PFst]
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        { \pvdash \Gamma e t \varpi:t' }
        { \pvdash \Gamma e t \varpi.f:\pair {t'} \Any }
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        { }
        \qquad
    \Infer[PSnd]
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        { \pvdash \Gamma e t \varpi:t' }
        { \pvdash \Gamma e t \varpi.s:\pair \Any {t'} }
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        { }\vspace{-1.9mm}
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\end{mathpar}
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These rules implement the analysis described in
Section~\ref{sec:ideas} for functions and extend it to products.  Let
us comment each rule in detail. \Rule{PSubs} is just subsumption for
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the deduction $\vdashp$. The rule \Rule{PInter} combined with
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\Rule{PTypeof} allows the system to deduce for an occurrence $\varpi$
the intersection of the static type of $\occ e \varpi$ (deduced by
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\Rule{PTypeof}) with the type deduced for $\varpi$ by the other $\vdashp$ rules. The
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rule \Rule{PEps} is the starting point of the analysis: if we are assuming that the test $e\in t$ succeeds, then we can assume that $e$ (i.e.,
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$\occ e\epsilon$) has type $t$ (recall that assuming that the test $e\in t$ fails corresponds to having $\neg t$ at the index of the turnstyle).
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The rule \Rule{PApprR} implements occurrence typing for
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the arguments of applications, since it states that if a function maps
arguments of type $t_1$ in results of type $t_2$ and an application of
this function yields results (in $t'_2$) that cannot be in $t_2$
(since $t_2\land t_2' \simeq \Empty$), then the argument of this application cannot be of type $t_1$. \Rule{PAppR} performs the
occurrence typing analysis for the function part of an application,
since it states that if an application has type $t_2$ and the argument
of this application has type $t_1$, then the function in this
application cannot have type $t_1\to\neg t_2$. Rules \Rule{PPair\_}
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are straightforward since they state that the $i$-th projection of a pair
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that is of type $\pair{t_1}{t_2}$ must be of type $t_i$. So are the last two
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rules that essentially state that if $\pi_1 e$ (respectively, $\pi_2
e$) is of type $t'$, then the type of $e$ must be of the form
$\pair{t'}\Any$ (respectively, $\pair\Any{t'}$).

This concludes the presentation of our type system, which satisfies
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the property of safety, deduced, as customary, from the properties
of progress and subject reduction (\emph{cf.} Appendix~\ref{app:soundness}).
\begin{theorem}[type safety]
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For every expression $e$ such that $\varnothing\vdash e:t$ either  $e$
diverges or there
exists a value $v$ of type $t$ such that $e\reduces^* v$.
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\end{theorem}
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